2005-04-17 06:20:36 +08:00
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#ifndef _LINUX_SCHED_H
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#define _LINUX_SCHED_H
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2012-10-13 17:46:48 +08:00
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#include <uapi/linux/sched.h>
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2006-04-27 07:12:56 +08:00
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struct sched_param {
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int sched_priority;
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};
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2005-04-17 06:20:36 +08:00
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#include <asm/param.h> /* for HZ */
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#include <linux/capability.h>
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#include <linux/threads.h>
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#include <linux/kernel.h>
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#include <linux/types.h>
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#include <linux/timex.h>
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#include <linux/jiffies.h>
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#include <linux/rbtree.h>
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#include <linux/thread_info.h>
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#include <linux/cpumask.h>
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#include <linux/errno.h>
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#include <linux/nodemask.h>
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2007-10-16 16:24:43 +08:00
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#include <linux/mm_types.h>
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2013-08-14 20:55:31 +08:00
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#include <linux/preempt.h>
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2005-04-17 06:20:36 +08:00
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#include <asm/page.h>
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#include <asm/ptrace.h>
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#include <asm/cputime.h>
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#include <linux/smp.h>
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#include <linux/sem.h>
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#include <linux/signal.h>
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#include <linux/compiler.h>
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#include <linux/completion.h>
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#include <linux/pid.h>
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#include <linux/percpu.h>
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#include <linux/topology.h>
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2007-10-17 14:25:50 +08:00
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#include <linux/proportions.h>
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2005-04-17 06:20:36 +08:00
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#include <linux/seccomp.h>
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2006-01-08 17:01:37 +08:00
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#include <linux/rcupdate.h>
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2009-04-15 02:17:16 +08:00
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#include <linux/rculist.h>
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2006-06-27 17:54:53 +08:00
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#include <linux/rtmutex.h>
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2005-04-17 06:20:36 +08:00
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2006-04-25 21:54:40 +08:00
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#include <linux/time.h>
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#include <linux/param.h>
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#include <linux/resource.h>
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#include <linux/timer.h>
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#include <linux/hrtimer.h>
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2006-12-10 18:19:19 +08:00
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#include <linux/task_io_accounting.h>
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2008-01-26 04:08:34 +08:00
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#include <linux/latencytop.h>
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2008-08-13 23:20:04 +08:00
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#include <linux/cred.h>
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2011-09-12 19:06:17 +08:00
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#include <linux/llist.h>
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2011-11-17 15:20:58 +08:00
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#include <linux/uidgid.h>
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mm: teach mm by current context info to not do I/O during memory allocation
This patch introduces PF_MEMALLOC_NOIO on process flag('flags' field of
'struct task_struct'), so that the flag can be set by one task to avoid
doing I/O inside memory allocation in the task's context.
The patch trys to solve one deadlock problem caused by block device, and
the problem may happen at least in the below situations:
- during block device runtime resume, if memory allocation with
GFP_KERNEL is called inside runtime resume callback of any one of its
ancestors(or the block device itself), the deadlock may be triggered
inside the memory allocation since it might not complete until the block
device becomes active and the involed page I/O finishes. The situation
is pointed out first by Alan Stern. It is not a good approach to
convert all GFP_KERNEL[1] in the path into GFP_NOIO because several
subsystems may be involved(for example, PCI, USB and SCSI may be
involved for usb mass stoarage device, network devices involved too in
the iSCSI case)
- during block device runtime suspend, because runtime resume need to
wait for completion of concurrent runtime suspend.
- during error handling of usb mass storage deivce, USB bus reset will
be put on the device, so there shouldn't have any memory allocation with
GFP_KERNEL during USB bus reset, otherwise the deadlock similar with
above may be triggered. Unfortunately, any usb device may include one
mass storage interface in theory, so it requires all usb interface
drivers to handle the situation. In fact, most usb drivers don't know
how to handle bus reset on the device and don't provide .pre_set() and
.post_reset() callback at all, so USB core has to unbind and bind driver
for these devices. So it is still not practical to resort to GFP_NOIO
for solving the problem.
Also the introduced solution can be used by block subsystem or block
drivers too, for example, set the PF_MEMALLOC_NOIO flag before doing
actual I/O transfer.
It is not a good idea to convert all these GFP_KERNEL in the affected
path into GFP_NOIO because these functions doing that may be implemented
as library and will be called in many other contexts.
In fact, memalloc_noio_flags() can convert some of current static
GFP_NOIO allocation into GFP_KERNEL back in other non-affected contexts,
at least almost all GFP_NOIO in USB subsystem can be converted into
GFP_KERNEL after applying the approach and make allocation with GFP_NOIO
only happen in runtime resume/bus reset/block I/O transfer contexts
generally.
[1], several GFP_KERNEL allocation examples in runtime resume path
- pci subsystem
acpi_os_allocate
<-acpi_ut_allocate
<-ACPI_ALLOCATE_ZEROED
<-acpi_evaluate_object
<-__acpi_bus_set_power
<-acpi_bus_set_power
<-acpi_pci_set_power_state
<-platform_pci_set_power_state
<-pci_platform_power_transition
<-__pci_complete_power_transition
<-pci_set_power_state
<-pci_restore_standard_config
<-pci_pm_runtime_resume
- usb subsystem
usb_get_status
<-finish_port_resume
<-usb_port_resume
<-generic_resume
<-usb_resume_device
<-usb_resume_both
<-usb_runtime_resume
- some individual usb drivers
usblp, uvc, gspca, most of dvb-usb-v2 media drivers, cpia2, az6007, ....
That is just what I have found. Unfortunately, this allocation can only
be found by human being now, and there should be many not found since
any function in the resume path(call tree) may allocate memory with
GFP_KERNEL.
Signed-off-by: Ming Lei <ming.lei@canonical.com>
Signed-off-by: Minchan Kim <minchan@kernel.org>
Cc: Alan Stern <stern@rowland.harvard.edu>
Cc: Oliver Neukum <oneukum@suse.de>
Cc: Jiri Kosina <jiri.kosina@suse.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Michal Hocko <mhocko@suse.cz>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: "Rafael J. Wysocki" <rjw@sisk.pl>
Cc: Greg KH <greg@kroah.com>
Cc: Jens Axboe <axboe@kernel.dk>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Eric Dumazet <eric.dumazet@gmail.com>
Cc: David Decotigny <david.decotigny@google.com>
Cc: Tom Herbert <therbert@google.com>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-02-23 08:34:08 +08:00
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#include <linux/gfp.h>
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2006-04-25 21:54:40 +08:00
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#include <asm/processor.h>
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2005-09-07 06:16:49 +08:00
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2005-04-17 06:20:36 +08:00
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struct exec_domain;
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2006-06-27 17:54:58 +08:00
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struct futex_pi_state;
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2008-01-26 04:08:34 +08:00
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struct robust_list_head;
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2010-02-23 15:55:42 +08:00
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struct bio_list;
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2009-03-30 07:50:06 +08:00
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struct fs_struct;
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perf: Do the big rename: Performance Counters -> Performance Events
Bye-bye Performance Counters, welcome Performance Events!
In the past few months the perfcounters subsystem has grown out its
initial role of counting hardware events, and has become (and is
becoming) a much broader generic event enumeration, reporting, logging,
monitoring, analysis facility.
Naming its core object 'perf_counter' and naming the subsystem
'perfcounters' has become more and more of a misnomer. With pending
code like hw-breakpoints support the 'counter' name is less and
less appropriate.
All in one, we've decided to rename the subsystem to 'performance
events' and to propagate this rename through all fields, variables
and API names. (in an ABI compatible fashion)
The word 'event' is also a bit shorter than 'counter' - which makes
it slightly more convenient to write/handle as well.
Thanks goes to Stephane Eranian who first observed this misnomer and
suggested a rename.
User-space tooling and ABI compatibility is not affected - this patch
should be function-invariant. (Also, defconfigs were not touched to
keep the size down.)
This patch has been generated via the following script:
FILES=$(find * -type f | grep -vE 'oprofile|[^K]config')
sed -i \
-e 's/PERF_EVENT_/PERF_RECORD_/g' \
-e 's/PERF_COUNTER/PERF_EVENT/g' \
-e 's/perf_counter/perf_event/g' \
-e 's/nb_counters/nb_events/g' \
-e 's/swcounter/swevent/g' \
-e 's/tpcounter_event/tp_event/g' \
$FILES
for N in $(find . -name perf_counter.[ch]); do
M=$(echo $N | sed 's/perf_counter/perf_event/g')
mv $N $M
done
FILES=$(find . -name perf_event.*)
sed -i \
-e 's/COUNTER_MASK/REG_MASK/g' \
-e 's/COUNTER/EVENT/g' \
-e 's/\<event\>/event_id/g' \
-e 's/counter/event/g' \
-e 's/Counter/Event/g' \
$FILES
... to keep it as correct as possible. This script can also be
used by anyone who has pending perfcounters patches - it converts
a Linux kernel tree over to the new naming. We tried to time this
change to the point in time where the amount of pending patches
is the smallest: the end of the merge window.
Namespace clashes were fixed up in a preparatory patch - and some
stylistic fallout will be fixed up in a subsequent patch.
( NOTE: 'counters' are still the proper terminology when we deal
with hardware registers - and these sed scripts are a bit
over-eager in renaming them. I've undone some of that, but
in case there's something left where 'counter' would be
better than 'event' we can undo that on an individual basis
instead of touching an otherwise nicely automated patch. )
Suggested-by: Stephane Eranian <eranian@google.com>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Acked-by: Paul Mackerras <paulus@samba.org>
Reviewed-by: Arjan van de Ven <arjan@linux.intel.com>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Arnaldo Carvalho de Melo <acme@redhat.com>
Cc: Frederic Weisbecker <fweisbec@gmail.com>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Howells <dhowells@redhat.com>
Cc: Kyle McMartin <kyle@mcmartin.ca>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Cc: <linux-arch@vger.kernel.org>
LKML-Reference: <new-submission>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 18:02:48 +08:00
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struct perf_event_context;
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2011-03-08 20:19:51 +08:00
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struct blk_plug;
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2005-04-17 06:20:36 +08:00
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/*
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* List of flags we want to share for kernel threads,
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* if only because they are not used by them anyway.
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*/
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#define CLONE_KERNEL (CLONE_FS | CLONE_FILES | CLONE_SIGHAND)
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/*
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* These are the constant used to fake the fixed-point load-average
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* counting. Some notes:
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* - 11 bit fractions expand to 22 bits by the multiplies: this gives
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* a load-average precision of 10 bits integer + 11 bits fractional
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* - if you want to count load-averages more often, you need more
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* precision, or rounding will get you. With 2-second counting freq,
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* the EXP_n values would be 1981, 2034 and 2043 if still using only
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* 11 bit fractions.
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*/
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extern unsigned long avenrun[]; /* Load averages */
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2009-05-03 02:08:52 +08:00
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extern void get_avenrun(unsigned long *loads, unsigned long offset, int shift);
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2005-04-17 06:20:36 +08:00
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#define FSHIFT 11 /* nr of bits of precision */
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#define FIXED_1 (1<<FSHIFT) /* 1.0 as fixed-point */
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2007-10-08 07:17:38 +08:00
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#define LOAD_FREQ (5*HZ+1) /* 5 sec intervals */
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2005-04-17 06:20:36 +08:00
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#define EXP_1 1884 /* 1/exp(5sec/1min) as fixed-point */
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#define EXP_5 2014 /* 1/exp(5sec/5min) */
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#define EXP_15 2037 /* 1/exp(5sec/15min) */
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#define CALC_LOAD(load,exp,n) \
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load *= exp; \
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load += n*(FIXED_1-exp); \
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load >>= FSHIFT;
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extern unsigned long total_forks;
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extern int nr_threads;
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DECLARE_PER_CPU(unsigned long, process_counts);
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extern int nr_processes(void);
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extern unsigned long nr_running(void);
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extern unsigned long nr_iowait(void);
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2010-07-01 15:07:17 +08:00
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extern unsigned long nr_iowait_cpu(int cpu);
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2009-09-22 08:04:08 +08:00
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extern unsigned long this_cpu_load(void);
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2010-12-01 02:48:45 +08:00
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extern void calc_global_load(unsigned long ticks);
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2012-05-17 23:15:29 +08:00
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extern void update_cpu_load_nohz(void);
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2005-04-17 06:20:36 +08:00
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2009-01-23 08:01:40 +08:00
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extern unsigned long get_parent_ip(unsigned long addr);
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2012-09-20 07:58:38 +08:00
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extern void dump_cpu_task(int cpu);
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2007-07-10 00:52:00 +08:00
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struct seq_file;
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struct cfs_rq;
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2007-10-15 23:00:14 +08:00
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struct task_group;
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2007-07-10 00:52:00 +08:00
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#ifdef CONFIG_SCHED_DEBUG
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extern void proc_sched_show_task(struct task_struct *p, struct seq_file *m);
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extern void proc_sched_set_task(struct task_struct *p);
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extern void
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2007-08-09 17:16:47 +08:00
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print_cfs_rq(struct seq_file *m, int cpu, struct cfs_rq *cfs_rq);
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2007-07-10 00:52:00 +08:00
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#endif
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2005-04-17 06:20:36 +08:00
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2005-09-30 06:18:21 +08:00
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/*
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* Task state bitmask. NOTE! These bits are also
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* encoded in fs/proc/array.c: get_task_state().
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*
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* We have two separate sets of flags: task->state
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* is about runnability, while task->exit_state are
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* about the task exiting. Confusing, but this way
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* modifying one set can't modify the other one by
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* mistake.
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*/
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2005-04-17 06:20:36 +08:00
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#define TASK_RUNNING 0
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#define TASK_INTERRUPTIBLE 1
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#define TASK_UNINTERRUPTIBLE 2
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2007-12-07 00:13:16 +08:00
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#define __TASK_STOPPED 4
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#define __TASK_TRACED 8
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2005-09-30 06:18:21 +08:00
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/* in tsk->exit_state */
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#define EXIT_ZOMBIE 16
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#define EXIT_DEAD 32
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/* in tsk->state again */
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2007-10-15 23:00:13 +08:00
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#define TASK_DEAD 64
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2007-12-07 00:13:16 +08:00
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#define TASK_WAKEKILL 128
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2009-09-15 20:43:03 +08:00
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#define TASK_WAKING 256
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2013-04-09 15:33:34 +08:00
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#define TASK_PARKED 512
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#define TASK_STATE_MAX 1024
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2007-12-07 00:13:16 +08:00
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2013-04-09 15:33:34 +08:00
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#define TASK_STATE_TO_CHAR_STR "RSDTtZXxKWP"
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2009-12-17 20:16:27 +08:00
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2009-12-17 20:16:30 +08:00
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extern char ___assert_task_state[1 - 2*!!(
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sizeof(TASK_STATE_TO_CHAR_STR)-1 != ilog2(TASK_STATE_MAX)+1)];
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2007-12-07 00:13:16 +08:00
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/* Convenience macros for the sake of set_task_state */
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#define TASK_KILLABLE (TASK_WAKEKILL | TASK_UNINTERRUPTIBLE)
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#define TASK_STOPPED (TASK_WAKEKILL | __TASK_STOPPED)
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#define TASK_TRACED (TASK_WAKEKILL | __TASK_TRACED)
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2005-04-17 06:20:36 +08:00
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2007-12-06 23:55:25 +08:00
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/* Convenience macros for the sake of wake_up */
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#define TASK_NORMAL (TASK_INTERRUPTIBLE | TASK_UNINTERRUPTIBLE)
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2007-12-07 00:13:16 +08:00
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#define TASK_ALL (TASK_NORMAL | __TASK_STOPPED | __TASK_TRACED)
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2007-12-06 23:55:25 +08:00
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/* get_task_state() */
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#define TASK_REPORT (TASK_RUNNING | TASK_INTERRUPTIBLE | \
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2007-12-07 00:13:16 +08:00
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TASK_UNINTERRUPTIBLE | __TASK_STOPPED | \
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__TASK_TRACED)
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2007-12-06 23:55:25 +08:00
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2007-12-07 00:13:16 +08:00
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#define task_is_traced(task) ((task->state & __TASK_TRACED) != 0)
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#define task_is_stopped(task) ((task->state & __TASK_STOPPED) != 0)
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2010-07-29 19:45:55 +08:00
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#define task_is_dead(task) ((task)->exit_state != 0)
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2007-12-06 23:55:25 +08:00
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#define task_is_stopped_or_traced(task) \
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2007-12-07 00:13:16 +08:00
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((task->state & (__TASK_STOPPED | __TASK_TRACED)) != 0)
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2007-12-06 23:55:25 +08:00
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#define task_contributes_to_load(task) \
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2009-04-09 08:45:12 +08:00
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((task->state & TASK_UNINTERRUPTIBLE) != 0 && \
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2011-11-22 04:32:24 +08:00
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(task->flags & PF_FROZEN) == 0)
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2005-04-17 06:20:36 +08:00
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|
|
|
|
|
#define __set_task_state(tsk, state_value) \
|
|
|
|
|
do { (tsk)->state = (state_value); } while (0)
|
|
|
|
|
#define set_task_state(tsk, state_value) \
|
|
|
|
|
set_mb((tsk)->state, (state_value))
|
|
|
|
|
|
2005-09-13 16:25:14 +08:00
|
|
|
/*
|
|
|
|
|
* set_current_state() includes a barrier so that the write of current->state
|
|
|
|
|
* is correctly serialised wrt the caller's subsequent test of whether to
|
|
|
|
|
* actually sleep:
|
|
|
|
|
*
|
|
|
|
|
* set_current_state(TASK_UNINTERRUPTIBLE);
|
|
|
|
|
* if (do_i_need_to_sleep())
|
|
|
|
|
* schedule();
|
|
|
|
|
*
|
|
|
|
|
* If the caller does not need such serialisation then use __set_current_state()
|
|
|
|
|
*/
|
2005-04-17 06:20:36 +08:00
|
|
|
#define __set_current_state(state_value) \
|
|
|
|
|
do { current->state = (state_value); } while (0)
|
|
|
|
|
#define set_current_state(state_value) \
|
|
|
|
|
set_mb(current->state, (state_value))
|
|
|
|
|
|
|
|
|
|
/* Task command name length */
|
|
|
|
|
#define TASK_COMM_LEN 16
|
|
|
|
|
|
|
|
|
|
#include <linux/spinlock.h>
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* This serializes "schedule()" and also protects
|
|
|
|
|
* the run-queue from deletions/modifications (but
|
|
|
|
|
* _adding_ to the beginning of the run-queue has
|
|
|
|
|
* a separate lock).
|
|
|
|
|
*/
|
|
|
|
|
extern rwlock_t tasklist_lock;
|
|
|
|
|
extern spinlock_t mmlist_lock;
|
|
|
|
|
|
2006-07-03 15:25:41 +08:00
|
|
|
struct task_struct;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2010-03-03 23:46:56 +08:00
|
|
|
#ifdef CONFIG_PROVE_RCU
|
|
|
|
|
extern int lockdep_tasklist_lock_is_held(void);
|
|
|
|
|
#endif /* #ifdef CONFIG_PROVE_RCU */
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void sched_init(void);
|
|
|
|
|
extern void sched_init_smp(void);
|
2008-02-16 01:56:34 +08:00
|
|
|
extern asmlinkage void schedule_tail(struct task_struct *prev);
|
2006-07-03 15:25:41 +08:00
|
|
|
extern void init_idle(struct task_struct *idle, int cpu);
|
2007-07-10 00:51:58 +08:00
|
|
|
extern void init_idle_bootup_task(struct task_struct *idle);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2009-09-20 02:55:44 +08:00
|
|
|
extern int runqueue_is_locked(int cpu);
|
2008-05-13 03:20:52 +08:00
|
|
|
|
2011-08-11 05:21:01 +08:00
|
|
|
#if defined(CONFIG_SMP) && defined(CONFIG_NO_HZ_COMMON)
|
2012-09-10 15:10:58 +08:00
|
|
|
extern void nohz_balance_enter_idle(int cpu);
|
2011-12-02 09:07:33 +08:00
|
|
|
extern void set_cpu_sd_state_idle(void);
|
2010-05-22 08:09:41 +08:00
|
|
|
extern int get_nohz_timer_target(void);
|
2007-05-08 15:32:51 +08:00
|
|
|
#else
|
2012-09-10 15:10:58 +08:00
|
|
|
static inline void nohz_balance_enter_idle(int cpu) { }
|
2011-12-06 19:47:55 +08:00
|
|
|
static inline void set_cpu_sd_state_idle(void) { }
|
2007-05-08 15:32:51 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2006-12-07 12:35:59 +08:00
|
|
|
/*
|
2007-04-26 11:50:03 +08:00
|
|
|
* Only dump TASK_* tasks. (0 for all tasks)
|
2006-12-07 12:35:59 +08:00
|
|
|
*/
|
|
|
|
|
extern void show_state_filter(unsigned long state_filter);
|
|
|
|
|
|
|
|
|
|
static inline void show_state(void)
|
|
|
|
|
{
|
2007-04-26 11:50:03 +08:00
|
|
|
show_state_filter(0);
|
2006-12-07 12:35:59 +08:00
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void show_regs(struct pt_regs *);
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* TASK is a pointer to the task whose backtrace we want to see (or NULL for current
|
|
|
|
|
* task), SP is the stack pointer of the first frame that should be shown in the back
|
|
|
|
|
* trace (or NULL if the entire call-chain of the task should be shown).
|
|
|
|
|
*/
|
|
|
|
|
extern void show_stack(struct task_struct *task, unsigned long *sp);
|
|
|
|
|
|
|
|
|
|
void io_schedule(void);
|
|
|
|
|
long io_schedule_timeout(long timeout);
|
|
|
|
|
|
|
|
|
|
extern void cpu_init (void);
|
|
|
|
|
extern void trap_init(void);
|
|
|
|
|
extern void update_process_times(int user);
|
|
|
|
|
extern void scheduler_tick(void);
|
|
|
|
|
|
2008-01-26 04:08:02 +08:00
|
|
|
extern void sched_show_task(struct task_struct *p);
|
|
|
|
|
|
2010-05-13 08:30:49 +08:00
|
|
|
#ifdef CONFIG_LOCKUP_DETECTOR
|
2005-09-07 06:16:27 +08:00
|
|
|
extern void touch_softlockup_watchdog(void);
|
2010-01-28 06:25:22 +08:00
|
|
|
extern void touch_softlockup_watchdog_sync(void);
|
2007-05-08 15:28:05 +08:00
|
|
|
extern void touch_all_softlockup_watchdogs(void);
|
2010-05-08 05:11:45 +08:00
|
|
|
extern int proc_dowatchdog_thresh(struct ctl_table *table, int write,
|
|
|
|
|
void __user *buffer,
|
|
|
|
|
size_t *lenp, loff_t *ppos);
|
2008-05-13 03:21:04 +08:00
|
|
|
extern unsigned int softlockup_panic;
|
2010-11-26 01:38:29 +08:00
|
|
|
void lockup_detector_init(void);
|
2005-09-07 06:16:27 +08:00
|
|
|
#else
|
|
|
|
|
static inline void touch_softlockup_watchdog(void)
|
|
|
|
|
{
|
|
|
|
|
}
|
2010-01-28 06:25:22 +08:00
|
|
|
static inline void touch_softlockup_watchdog_sync(void)
|
|
|
|
|
{
|
|
|
|
|
}
|
2007-05-08 15:28:05 +08:00
|
|
|
static inline void touch_all_softlockup_watchdogs(void)
|
|
|
|
|
{
|
|
|
|
|
}
|
2010-11-26 01:38:29 +08:00
|
|
|
static inline void lockup_detector_init(void)
|
|
|
|
|
{
|
|
|
|
|
}
|
2005-09-07 06:16:27 +08:00
|
|
|
#endif
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* Attach to any functions which should be ignored in wchan output. */
|
|
|
|
|
#define __sched __attribute__((__section__(".sched.text")))
|
2007-11-28 22:52:56 +08:00
|
|
|
|
|
|
|
|
/* Linker adds these: start and end of __sched functions */
|
|
|
|
|
extern char __sched_text_start[], __sched_text_end[];
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* Is this address in the __sched functions? */
|
|
|
|
|
extern int in_sched_functions(unsigned long addr);
|
|
|
|
|
|
|
|
|
|
#define MAX_SCHEDULE_TIMEOUT LONG_MAX
|
2008-02-14 07:03:15 +08:00
|
|
|
extern signed long schedule_timeout(signed long timeout);
|
2005-09-10 15:27:21 +08:00
|
|
|
extern signed long schedule_timeout_interruptible(signed long timeout);
|
2007-12-07 00:59:46 +08:00
|
|
|
extern signed long schedule_timeout_killable(signed long timeout);
|
2005-09-10 15:27:21 +08:00
|
|
|
extern signed long schedule_timeout_uninterruptible(signed long timeout);
|
2005-04-17 06:20:36 +08:00
|
|
|
asmlinkage void schedule(void);
|
2011-03-21 19:09:35 +08:00
|
|
|
extern void schedule_preempt_disabled(void);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2006-10-02 17:18:06 +08:00
|
|
|
struct nsproxy;
|
2007-07-16 14:40:59 +08:00
|
|
|
struct user_namespace;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2010-01-16 09:01:35 +08:00
|
|
|
#ifdef CONFIG_MMU
|
|
|
|
|
extern void arch_pick_mmap_layout(struct mm_struct *mm);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern unsigned long
|
|
|
|
|
arch_get_unmapped_area(struct file *, unsigned long, unsigned long,
|
|
|
|
|
unsigned long, unsigned long);
|
|
|
|
|
extern unsigned long
|
|
|
|
|
arch_get_unmapped_area_topdown(struct file *filp, unsigned long addr,
|
|
|
|
|
unsigned long len, unsigned long pgoff,
|
|
|
|
|
unsigned long flags);
|
2010-01-16 09:01:35 +08:00
|
|
|
#else
|
|
|
|
|
static inline void arch_pick_mmap_layout(struct mm_struct *mm) {}
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2009-01-07 06:40:29 +08:00
|
|
|
|
2007-07-19 16:48:27 +08:00
|
|
|
extern void set_dumpable(struct mm_struct *mm, int value);
|
|
|
|
|
extern int get_dumpable(struct mm_struct *mm);
|
|
|
|
|
|
|
|
|
|
/* mm flags */
|
2007-07-19 16:48:28 +08:00
|
|
|
/* dumpable bits */
|
2007-07-19 16:48:27 +08:00
|
|
|
#define MMF_DUMPABLE 0 /* core dump is permitted */
|
|
|
|
|
#define MMF_DUMP_SECURELY 1 /* core file is readable only by root */
|
2009-09-22 08:01:57 +08:00
|
|
|
|
2007-07-19 16:48:28 +08:00
|
|
|
#define MMF_DUMPABLE_BITS 2
|
2009-09-22 08:01:57 +08:00
|
|
|
#define MMF_DUMPABLE_MASK ((1 << MMF_DUMPABLE_BITS) - 1)
|
2007-07-19 16:48:28 +08:00
|
|
|
|
|
|
|
|
/* coredump filter bits */
|
|
|
|
|
#define MMF_DUMP_ANON_PRIVATE 2
|
|
|
|
|
#define MMF_DUMP_ANON_SHARED 3
|
|
|
|
|
#define MMF_DUMP_MAPPED_PRIVATE 4
|
|
|
|
|
#define MMF_DUMP_MAPPED_SHARED 5
|
2007-10-17 14:27:02 +08:00
|
|
|
#define MMF_DUMP_ELF_HEADERS 6
|
coredump_filter: add hugepage dumping
Presently hugepage's vma has a VM_RESERVED flag in order not to be
swapped. But a VM_RESERVED vma isn't core dumped because this flag is
often used for some kernel vmas (e.g. vmalloc, sound related).
Thus hugepages are never dumped and it can't be debugged easily. Many
developers want hugepages to be included into core-dump.
However, We can't read generic VM_RESERVED area because this area is often
IO mapping area. then these area reading may change device state. it is
definitly undesiable side-effect.
So adding a hugepage specific bit to the coredump filter is better. It
will be able to hugepage core dumping and doesn't cause any side-effect to
any i/o devices.
In additional, libhugetlb use hugetlb private mapping pages as anonymous
page. Then, hugepage private mapping pages should be core dumped by
default.
Then, /proc/[pid]/core_dump_filter has two new bits.
- bit 5 mean hugetlb private mapping pages are dumped or not. (default: yes)
- bit 6 mean hugetlb shared mapping pages are dumped or not. (default: no)
I tested by following method.
% ulimit -c unlimited
% ./crash_hugepage 50
% ./crash_hugepage 50 -p
% ls -lh
% gdb ./crash_hugepage core
%
% echo 0x43 > /proc/self/coredump_filter
% ./crash_hugepage 50
% ./crash_hugepage 50 -p
% ls -lh
% gdb ./crash_hugepage core
#include <stdlib.h>
#include <stdio.h>
#include <unistd.h>
#include <sys/mman.h>
#include <string.h>
#include "hugetlbfs.h"
int main(int argc, char** argv){
char* p;
int ch;
int mmap_flags = MAP_SHARED;
int fd;
int nr_pages;
while((ch = getopt(argc, argv, "p")) != -1) {
switch (ch) {
case 'p':
mmap_flags &= ~MAP_SHARED;
mmap_flags |= MAP_PRIVATE;
break;
default:
/* nothing*/
break;
}
}
argc -= optind;
argv += optind;
if (argc == 0){
printf("need # of pages\n");
exit(1);
}
nr_pages = atoi(argv[0]);
if (nr_pages < 2) {
printf("nr_pages must >2\n");
exit(1);
}
fd = hugetlbfs_unlinked_fd();
p = mmap(NULL, nr_pages * gethugepagesize(),
PROT_READ|PROT_WRITE, mmap_flags, fd, 0);
sleep(2);
*(p + gethugepagesize()) = 1; /* COW */
sleep(2);
/* crash! */
*(int*)0 = 1;
return 0;
}
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Reviewed-by: Kawai Hidehiro <hidehiro.kawai.ez@hitachi.com>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: William Irwin <wli@holomorphy.com>
Cc: Adam Litke <agl@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 11:27:08 +08:00
|
|
|
#define MMF_DUMP_HUGETLB_PRIVATE 7
|
|
|
|
|
#define MMF_DUMP_HUGETLB_SHARED 8
|
2009-09-22 08:01:57 +08:00
|
|
|
|
2007-07-19 16:48:28 +08:00
|
|
|
#define MMF_DUMP_FILTER_SHIFT MMF_DUMPABLE_BITS
|
coredump_filter: add hugepage dumping
Presently hugepage's vma has a VM_RESERVED flag in order not to be
swapped. But a VM_RESERVED vma isn't core dumped because this flag is
often used for some kernel vmas (e.g. vmalloc, sound related).
Thus hugepages are never dumped and it can't be debugged easily. Many
developers want hugepages to be included into core-dump.
However, We can't read generic VM_RESERVED area because this area is often
IO mapping area. then these area reading may change device state. it is
definitly undesiable side-effect.
So adding a hugepage specific bit to the coredump filter is better. It
will be able to hugepage core dumping and doesn't cause any side-effect to
any i/o devices.
In additional, libhugetlb use hugetlb private mapping pages as anonymous
page. Then, hugepage private mapping pages should be core dumped by
default.
Then, /proc/[pid]/core_dump_filter has two new bits.
- bit 5 mean hugetlb private mapping pages are dumped or not. (default: yes)
- bit 6 mean hugetlb shared mapping pages are dumped or not. (default: no)
I tested by following method.
% ulimit -c unlimited
% ./crash_hugepage 50
% ./crash_hugepage 50 -p
% ls -lh
% gdb ./crash_hugepage core
%
% echo 0x43 > /proc/self/coredump_filter
% ./crash_hugepage 50
% ./crash_hugepage 50 -p
% ls -lh
% gdb ./crash_hugepage core
#include <stdlib.h>
#include <stdio.h>
#include <unistd.h>
#include <sys/mman.h>
#include <string.h>
#include "hugetlbfs.h"
int main(int argc, char** argv){
char* p;
int ch;
int mmap_flags = MAP_SHARED;
int fd;
int nr_pages;
while((ch = getopt(argc, argv, "p")) != -1) {
switch (ch) {
case 'p':
mmap_flags &= ~MAP_SHARED;
mmap_flags |= MAP_PRIVATE;
break;
default:
/* nothing*/
break;
}
}
argc -= optind;
argv += optind;
if (argc == 0){
printf("need # of pages\n");
exit(1);
}
nr_pages = atoi(argv[0]);
if (nr_pages < 2) {
printf("nr_pages must >2\n");
exit(1);
}
fd = hugetlbfs_unlinked_fd();
p = mmap(NULL, nr_pages * gethugepagesize(),
PROT_READ|PROT_WRITE, mmap_flags, fd, 0);
sleep(2);
*(p + gethugepagesize()) = 1; /* COW */
sleep(2);
/* crash! */
*(int*)0 = 1;
return 0;
}
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Reviewed-by: Kawai Hidehiro <hidehiro.kawai.ez@hitachi.com>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: William Irwin <wli@holomorphy.com>
Cc: Adam Litke <agl@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 11:27:08 +08:00
|
|
|
#define MMF_DUMP_FILTER_BITS 7
|
2007-07-19 16:48:28 +08:00
|
|
|
#define MMF_DUMP_FILTER_MASK \
|
|
|
|
|
(((1 << MMF_DUMP_FILTER_BITS) - 1) << MMF_DUMP_FILTER_SHIFT)
|
|
|
|
|
#define MMF_DUMP_FILTER_DEFAULT \
|
coredump_filter: add hugepage dumping
Presently hugepage's vma has a VM_RESERVED flag in order not to be
swapped. But a VM_RESERVED vma isn't core dumped because this flag is
often used for some kernel vmas (e.g. vmalloc, sound related).
Thus hugepages are never dumped and it can't be debugged easily. Many
developers want hugepages to be included into core-dump.
However, We can't read generic VM_RESERVED area because this area is often
IO mapping area. then these area reading may change device state. it is
definitly undesiable side-effect.
So adding a hugepage specific bit to the coredump filter is better. It
will be able to hugepage core dumping and doesn't cause any side-effect to
any i/o devices.
In additional, libhugetlb use hugetlb private mapping pages as anonymous
page. Then, hugepage private mapping pages should be core dumped by
default.
Then, /proc/[pid]/core_dump_filter has two new bits.
- bit 5 mean hugetlb private mapping pages are dumped or not. (default: yes)
- bit 6 mean hugetlb shared mapping pages are dumped or not. (default: no)
I tested by following method.
% ulimit -c unlimited
% ./crash_hugepage 50
% ./crash_hugepage 50 -p
% ls -lh
% gdb ./crash_hugepage core
%
% echo 0x43 > /proc/self/coredump_filter
% ./crash_hugepage 50
% ./crash_hugepage 50 -p
% ls -lh
% gdb ./crash_hugepage core
#include <stdlib.h>
#include <stdio.h>
#include <unistd.h>
#include <sys/mman.h>
#include <string.h>
#include "hugetlbfs.h"
int main(int argc, char** argv){
char* p;
int ch;
int mmap_flags = MAP_SHARED;
int fd;
int nr_pages;
while((ch = getopt(argc, argv, "p")) != -1) {
switch (ch) {
case 'p':
mmap_flags &= ~MAP_SHARED;
mmap_flags |= MAP_PRIVATE;
break;
default:
/* nothing*/
break;
}
}
argc -= optind;
argv += optind;
if (argc == 0){
printf("need # of pages\n");
exit(1);
}
nr_pages = atoi(argv[0]);
if (nr_pages < 2) {
printf("nr_pages must >2\n");
exit(1);
}
fd = hugetlbfs_unlinked_fd();
p = mmap(NULL, nr_pages * gethugepagesize(),
PROT_READ|PROT_WRITE, mmap_flags, fd, 0);
sleep(2);
*(p + gethugepagesize()) = 1; /* COW */
sleep(2);
/* crash! */
*(int*)0 = 1;
return 0;
}
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Reviewed-by: Kawai Hidehiro <hidehiro.kawai.ez@hitachi.com>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: William Irwin <wli@holomorphy.com>
Cc: Adam Litke <agl@us.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-10-19 11:27:08 +08:00
|
|
|
((1 << MMF_DUMP_ANON_PRIVATE) | (1 << MMF_DUMP_ANON_SHARED) |\
|
2008-10-19 11:28:23 +08:00
|
|
|
(1 << MMF_DUMP_HUGETLB_PRIVATE) | MMF_DUMP_MASK_DEFAULT_ELF)
|
|
|
|
|
|
|
|
|
|
#ifdef CONFIG_CORE_DUMP_DEFAULT_ELF_HEADERS
|
|
|
|
|
# define MMF_DUMP_MASK_DEFAULT_ELF (1 << MMF_DUMP_ELF_HEADERS)
|
|
|
|
|
#else
|
|
|
|
|
# define MMF_DUMP_MASK_DEFAULT_ELF 0
|
|
|
|
|
#endif
|
2009-09-22 08:01:57 +08:00
|
|
|
/* leave room for more dump flags */
|
|
|
|
|
#define MMF_VM_MERGEABLE 16 /* KSM may merge identical pages */
|
2011-01-14 07:46:58 +08:00
|
|
|
#define MMF_VM_HUGEPAGE 17 /* set when VM_HUGEPAGE is set on vma */
|
2012-06-08 05:21:11 +08:00
|
|
|
#define MMF_EXE_FILE_CHANGED 18 /* see prctl_set_mm_exe_file() */
|
2009-09-22 08:01:57 +08:00
|
|
|
|
2012-08-19 22:15:09 +08:00
|
|
|
#define MMF_HAS_UPROBES 19 /* has uprobes */
|
|
|
|
|
#define MMF_RECALC_UPROBES 20 /* MMF_HAS_UPROBES can be wrong */
|
2012-08-08 23:11:42 +08:00
|
|
|
|
2009-09-22 08:01:57 +08:00
|
|
|
#define MMF_INIT_MASK (MMF_DUMPABLE_MASK | MMF_DUMP_FILTER_MASK)
|
2007-07-19 16:48:27 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
struct sighand_struct {
|
|
|
|
|
atomic_t count;
|
|
|
|
|
struct k_sigaction action[_NSIG];
|
|
|
|
|
spinlock_t siglock;
|
2007-09-21 03:40:16 +08:00
|
|
|
wait_queue_head_t signalfd_wqh;
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
|
|
|
|
|
2006-06-25 20:49:24 +08:00
|
|
|
struct pacct_struct {
|
2006-06-25 20:49:25 +08:00
|
|
|
int ac_flag;
|
|
|
|
|
long ac_exitcode;
|
2006-06-25 20:49:24 +08:00
|
|
|
unsigned long ac_mem;
|
2006-06-25 20:49:26 +08:00
|
|
|
cputime_t ac_utime, ac_stime;
|
|
|
|
|
unsigned long ac_minflt, ac_majflt;
|
2006-06-25 20:49:24 +08:00
|
|
|
};
|
|
|
|
|
|
2009-07-29 18:15:26 +08:00
|
|
|
struct cpu_itimer {
|
|
|
|
|
cputime_t expires;
|
|
|
|
|
cputime_t incr;
|
2009-07-29 18:15:27 +08:00
|
|
|
u32 error;
|
|
|
|
|
u32 incr_error;
|
2009-07-29 18:15:26 +08:00
|
|
|
};
|
|
|
|
|
|
2012-11-22 07:58:35 +08:00
|
|
|
/**
|
|
|
|
|
* struct cputime - snaphsot of system and user cputime
|
|
|
|
|
* @utime: time spent in user mode
|
|
|
|
|
* @stime: time spent in system mode
|
|
|
|
|
*
|
|
|
|
|
* Gathers a generic snapshot of user and system time.
|
|
|
|
|
*/
|
|
|
|
|
struct cputime {
|
|
|
|
|
cputime_t utime;
|
|
|
|
|
cputime_t stime;
|
|
|
|
|
};
|
|
|
|
|
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
/**
|
|
|
|
|
* struct task_cputime - collected CPU time counts
|
|
|
|
|
* @utime: time spent in user mode, in &cputime_t units
|
|
|
|
|
* @stime: time spent in kernel mode, in &cputime_t units
|
|
|
|
|
* @sum_exec_runtime: total time spent on the CPU, in nanoseconds
|
2008-09-14 23:11:46 +08:00
|
|
|
*
|
2012-11-22 07:58:35 +08:00
|
|
|
* This is an extension of struct cputime that includes the total runtime
|
|
|
|
|
* spent by the task from the scheduler point of view.
|
|
|
|
|
*
|
|
|
|
|
* As a result, this structure groups together three kinds of CPU time
|
|
|
|
|
* that are tracked for threads and thread groups. Most things considering
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
* CPU time want to group these counts together and treat all three
|
|
|
|
|
* of them in parallel.
|
|
|
|
|
*/
|
|
|
|
|
struct task_cputime {
|
|
|
|
|
cputime_t utime;
|
|
|
|
|
cputime_t stime;
|
|
|
|
|
unsigned long long sum_exec_runtime;
|
|
|
|
|
};
|
|
|
|
|
/* Alternate field names when used to cache expirations. */
|
|
|
|
|
#define prof_exp stime
|
|
|
|
|
#define virt_exp utime
|
|
|
|
|
#define sched_exp sum_exec_runtime
|
|
|
|
|
|
2009-02-05 19:24:16 +08:00
|
|
|
#define INIT_CPUTIME \
|
|
|
|
|
(struct task_cputime) { \
|
2011-12-15 21:56:09 +08:00
|
|
|
.utime = 0, \
|
|
|
|
|
.stime = 0, \
|
2009-02-05 19:24:16 +08:00
|
|
|
.sum_exec_runtime = 0, \
|
|
|
|
|
}
|
|
|
|
|
|
2013-09-24 01:04:26 +08:00
|
|
|
#define PREEMPT_ENABLED (PREEMPT_NEED_RESCHED)
|
|
|
|
|
|
|
|
|
|
#ifdef CONFIG_PREEMPT_COUNT
|
|
|
|
|
#define PREEMPT_DISABLED (1 + PREEMPT_ENABLED)
|
|
|
|
|
#else
|
|
|
|
|
#define PREEMPT_DISABLED PREEMPT_ENABLED
|
|
|
|
|
#endif
|
|
|
|
|
|
2009-07-10 20:57:56 +08:00
|
|
|
/*
|
|
|
|
|
* Disable preemption until the scheduler is running.
|
|
|
|
|
* Reset by start_kernel()->sched_init()->init_idle().
|
2009-07-10 20:57:57 +08:00
|
|
|
*
|
|
|
|
|
* We include PREEMPT_ACTIVE to avoid cond_resched() from working
|
|
|
|
|
* before the scheduler is active -- see should_resched().
|
2009-07-10 20:57:56 +08:00
|
|
|
*/
|
2013-09-24 01:04:26 +08:00
|
|
|
#define INIT_PREEMPT_COUNT (PREEMPT_DISABLED + PREEMPT_ACTIVE)
|
2009-07-10 20:57:56 +08:00
|
|
|
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
/**
|
2009-02-05 19:24:16 +08:00
|
|
|
* struct thread_group_cputimer - thread group interval timer counts
|
|
|
|
|
* @cputime: thread group interval timers.
|
|
|
|
|
* @running: non-zero when there are timers running and
|
|
|
|
|
* @cputime receives updates.
|
|
|
|
|
* @lock: lock for fields in this struct.
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
*
|
|
|
|
|
* This structure contains the version of task_cputime, above, that is
|
2009-02-05 19:24:16 +08:00
|
|
|
* used for thread group CPU timer calculations.
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
*/
|
2009-02-05 19:24:16 +08:00
|
|
|
struct thread_group_cputimer {
|
|
|
|
|
struct task_cputime cputime;
|
|
|
|
|
int running;
|
2009-07-26 00:56:56 +08:00
|
|
|
raw_spinlock_t lock;
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
};
|
|
|
|
|
|
2011-05-27 07:25:18 +08:00
|
|
|
#include <linux/rwsem.h>
|
sched: Add 'autogroup' scheduling feature: automated per session task groups
A recurring complaint from CFS users is that parallel kbuild has
a negative impact on desktop interactivity. This patch
implements an idea from Linus, to automatically create task
groups. Currently, only per session autogroups are implemented,
but the patch leaves the way open for enhancement.
Implementation: each task's signal struct contains an inherited
pointer to a refcounted autogroup struct containing a task group
pointer, the default for all tasks pointing to the
init_task_group. When a task calls setsid(), a new task group
is created, the process is moved into the new task group, and a
reference to the preveious task group is dropped. Child
processes inherit this task group thereafter, and increase it's
refcount. When the last thread of a process exits, the
process's reference is dropped, such that when the last process
referencing an autogroup exits, the autogroup is destroyed.
At runqueue selection time, IFF a task has no cgroup assignment,
its current autogroup is used.
Autogroup bandwidth is controllable via setting it's nice level
through the proc filesystem:
cat /proc/<pid>/autogroup
Displays the task's group and the group's nice level.
echo <nice level> > /proc/<pid>/autogroup
Sets the task group's shares to the weight of nice <level> task.
Setting nice level is rate limited for !admin users due to the
abuse risk of task group locking.
The feature is enabled from boot by default if
CONFIG_SCHED_AUTOGROUP=y is selected, but can be disabled via
the boot option noautogroup, and can also be turned on/off on
the fly via:
echo [01] > /proc/sys/kernel/sched_autogroup_enabled
... which will automatically move tasks to/from the root task group.
Signed-off-by: Mike Galbraith <efault@gmx.de>
Acked-by: Linus Torvalds <torvalds@linux-foundation.org>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Markus Trippelsdorf <markus@trippelsdorf.de>
Cc: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Cc: Paul Turner <pjt@google.com>
Cc: Oleg Nesterov <oleg@redhat.com>
[ Removed the task_group_path() debug code, and fixed !EVENTFD build failure. ]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
LKML-Reference: <1290281700.28711.9.camel@maggy.simson.net>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2010-11-30 21:18:03 +08:00
|
|
|
struct autogroup;
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
2011-03-22 03:24:47 +08:00
|
|
|
* NOTE! "signal_struct" does not have its own
|
2005-04-17 06:20:36 +08:00
|
|
|
* locking, because a shared signal_struct always
|
|
|
|
|
* implies a shared sighand_struct, so locking
|
|
|
|
|
* sighand_struct is always a proper superset of
|
|
|
|
|
* the locking of signal_struct.
|
|
|
|
|
*/
|
|
|
|
|
struct signal_struct {
|
signals: make task_struct->signal immutable/refcountable
We have a lot of problems with accessing task_struct->signal, it can
"disappear" at any moment. Even current can't use its ->signal safely
after exit_notify(). ->siglock helps, but it is not convenient, not
always possible, and sometimes it makes sense to use task->signal even
after this task has already dead.
This patch adds the reference counter, sigcnt, into signal_struct. This
reference is owned by task_struct and it is dropped in
__put_task_struct(). Perhaps it makes sense to export
get/put_signal_struct() later, but currently I don't see the immediate
reason.
Rename __cleanup_signal() to free_signal_struct() and unexport it. With
the previous changes it does nothing except kmem_cache_free().
Change __exit_signal() to not clear/free ->signal, it will be freed when
the last reference to any thread in the thread group goes away.
Note:
- when the last thead exits signal->tty can point to nowhere, see
the next patch.
- with or without this patch signal_struct->count should go away,
or at least it should be "int nr_threads" for fs/proc. This will
be addressed later.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Alan Cox <alan@linux.intel.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Peter Zijlstra <peterz@infradead.org>
Acked-by: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-05-27 05:43:16 +08:00
|
|
|
atomic_t sigcnt;
|
2005-04-17 06:20:36 +08:00
|
|
|
atomic_t live;
|
2010-05-27 05:43:24 +08:00
|
|
|
int nr_threads;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
wait_queue_head_t wait_chldexit; /* for wait4() */
|
|
|
|
|
|
|
|
|
|
/* current thread group signal load-balancing target: */
|
2006-07-03 15:25:41 +08:00
|
|
|
struct task_struct *curr_target;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* shared signal handling: */
|
|
|
|
|
struct sigpending shared_pending;
|
|
|
|
|
|
|
|
|
|
/* thread group exit support */
|
|
|
|
|
int group_exit_code;
|
|
|
|
|
/* overloaded:
|
|
|
|
|
* - notify group_exit_task when ->count is equal to notify_count
|
|
|
|
|
* - everyone except group_exit_task is stopped during signal delivery
|
|
|
|
|
* of fatal signals, group_exit_task processes the signal.
|
|
|
|
|
*/
|
|
|
|
|
int notify_count;
|
2008-08-01 20:18:04 +08:00
|
|
|
struct task_struct *group_exit_task;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* thread group stop support, overloads group_exit_code too */
|
|
|
|
|
int group_stop_count;
|
|
|
|
|
unsigned int flags; /* see SIGNAL_* flags below */
|
|
|
|
|
|
prctl: add PR_{SET,GET}_CHILD_SUBREAPER to allow simple process supervision
Userspace service managers/supervisors need to track their started
services. Many services daemonize by double-forking and get implicitly
re-parented to PID 1. The service manager will no longer be able to
receive the SIGCHLD signals for them, and is no longer in charge of
reaping the children with wait(). All information about the children is
lost at the moment PID 1 cleans up the re-parented processes.
With this prctl, a service manager process can mark itself as a sort of
'sub-init', able to stay as the parent for all orphaned processes
created by the started services. All SIGCHLD signals will be delivered
to the service manager.
Receiving SIGCHLD and doing wait() is in cases of a service-manager much
preferred over any possible asynchronous notification about specific
PIDs, because the service manager has full access to the child process
data in /proc and the PID can not be re-used until the wait(), the
service-manager itself is in charge of, has happened.
As a side effect, the relevant parent PID information does not get lost
by a double-fork, which results in a more elaborate process tree and
'ps' output:
before:
# ps afx
253 ? Ss 0:00 /bin/dbus-daemon --system --nofork
294 ? Sl 0:00 /usr/libexec/polkit-1/polkitd
328 ? S 0:00 /usr/sbin/modem-manager
608 ? Sl 0:00 /usr/libexec/colord
658 ? Sl 0:00 /usr/libexec/upowerd
819 ? Sl 0:00 /usr/libexec/imsettings-daemon
916 ? Sl 0:00 /usr/libexec/udisks-daemon
917 ? S 0:00 \_ udisks-daemon: not polling any devices
after:
# ps afx
294 ? Ss 0:00 /bin/dbus-daemon --system --nofork
426 ? Sl 0:00 \_ /usr/libexec/polkit-1/polkitd
449 ? S 0:00 \_ /usr/sbin/modem-manager
635 ? Sl 0:00 \_ /usr/libexec/colord
705 ? Sl 0:00 \_ /usr/libexec/upowerd
959 ? Sl 0:00 \_ /usr/libexec/udisks-daemon
960 ? S 0:00 | \_ udisks-daemon: not polling any devices
977 ? Sl 0:00 \_ /usr/libexec/packagekitd
This prctl is orthogonal to PID namespaces. PID namespaces are isolated
from each other, while a service management process usually requires the
services to live in the same namespace, to be able to talk to each
other.
Users of this will be the systemd per-user instance, which provides
init-like functionality for the user's login session and D-Bus, which
activates bus services on-demand. Both need init-like capabilities to
be able to properly keep track of the services they start.
Many thanks to Oleg for several rounds of review and insights.
[akpm@linux-foundation.org: fix comment layout and spelling]
[akpm@linux-foundation.org: add lengthy code comment from Oleg]
Reviewed-by: Oleg Nesterov <oleg@redhat.com>
Signed-off-by: Lennart Poettering <lennart@poettering.net>
Signed-off-by: Kay Sievers <kay.sievers@vrfy.org>
Acked-by: Valdis Kletnieks <Valdis.Kletnieks@vt.edu>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-03-24 06:01:54 +08:00
|
|
|
/*
|
|
|
|
|
* PR_SET_CHILD_SUBREAPER marks a process, like a service
|
|
|
|
|
* manager, to re-parent orphan (double-forking) child processes
|
|
|
|
|
* to this process instead of 'init'. The service manager is
|
|
|
|
|
* able to receive SIGCHLD signals and is able to investigate
|
|
|
|
|
* the process until it calls wait(). All children of this
|
|
|
|
|
* process will inherit a flag if they should look for a
|
|
|
|
|
* child_subreaper process at exit.
|
|
|
|
|
*/
|
|
|
|
|
unsigned int is_child_subreaper:1;
|
|
|
|
|
unsigned int has_child_subreaper:1;
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* POSIX.1b Interval Timers */
|
2013-03-11 17:12:21 +08:00
|
|
|
int posix_timer_id;
|
|
|
|
|
struct list_head posix_timers;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* ITIMER_REAL timer for the process */
|
2006-01-10 12:52:34 +08:00
|
|
|
struct hrtimer real_timer;
|
2008-02-08 20:19:19 +08:00
|
|
|
struct pid *leader_pid;
|
2006-01-10 12:52:34 +08:00
|
|
|
ktime_t it_real_incr;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2009-07-29 18:15:26 +08:00
|
|
|
/*
|
|
|
|
|
* ITIMER_PROF and ITIMER_VIRTUAL timers for the process, we use
|
|
|
|
|
* CPUCLOCK_PROF and CPUCLOCK_VIRT for indexing array as these
|
|
|
|
|
* values are defined to 0 and 1 respectively
|
|
|
|
|
*/
|
|
|
|
|
struct cpu_itimer it[2];
|
2005-04-17 06:20:36 +08:00
|
|
|
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
/*
|
2009-02-05 19:24:16 +08:00
|
|
|
* Thread group totals for process CPU timers.
|
|
|
|
|
* See thread_group_cputimer(), et al, for details.
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
*/
|
2009-02-05 19:24:16 +08:00
|
|
|
struct thread_group_cputimer cputimer;
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
|
|
|
|
|
/* Earliest-expiration cache. */
|
|
|
|
|
struct task_cputime cputime_expires;
|
|
|
|
|
|
|
|
|
|
struct list_head cpu_timers[3];
|
|
|
|
|
|
2007-02-12 16:53:00 +08:00
|
|
|
struct pid *tty_old_pgrp;
|
2006-12-08 18:37:55 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* boolean value for session group leader */
|
|
|
|
|
int leader;
|
|
|
|
|
|
|
|
|
|
struct tty_struct *tty; /* NULL if no tty */
|
|
|
|
|
|
sched: Add 'autogroup' scheduling feature: automated per session task groups
A recurring complaint from CFS users is that parallel kbuild has
a negative impact on desktop interactivity. This patch
implements an idea from Linus, to automatically create task
groups. Currently, only per session autogroups are implemented,
but the patch leaves the way open for enhancement.
Implementation: each task's signal struct contains an inherited
pointer to a refcounted autogroup struct containing a task group
pointer, the default for all tasks pointing to the
init_task_group. When a task calls setsid(), a new task group
is created, the process is moved into the new task group, and a
reference to the preveious task group is dropped. Child
processes inherit this task group thereafter, and increase it's
refcount. When the last thread of a process exits, the
process's reference is dropped, such that when the last process
referencing an autogroup exits, the autogroup is destroyed.
At runqueue selection time, IFF a task has no cgroup assignment,
its current autogroup is used.
Autogroup bandwidth is controllable via setting it's nice level
through the proc filesystem:
cat /proc/<pid>/autogroup
Displays the task's group and the group's nice level.
echo <nice level> > /proc/<pid>/autogroup
Sets the task group's shares to the weight of nice <level> task.
Setting nice level is rate limited for !admin users due to the
abuse risk of task group locking.
The feature is enabled from boot by default if
CONFIG_SCHED_AUTOGROUP=y is selected, but can be disabled via
the boot option noautogroup, and can also be turned on/off on
the fly via:
echo [01] > /proc/sys/kernel/sched_autogroup_enabled
... which will automatically move tasks to/from the root task group.
Signed-off-by: Mike Galbraith <efault@gmx.de>
Acked-by: Linus Torvalds <torvalds@linux-foundation.org>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Markus Trippelsdorf <markus@trippelsdorf.de>
Cc: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Cc: Paul Turner <pjt@google.com>
Cc: Oleg Nesterov <oleg@redhat.com>
[ Removed the task_group_path() debug code, and fixed !EVENTFD build failure. ]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
LKML-Reference: <1290281700.28711.9.camel@maggy.simson.net>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2010-11-30 21:18:03 +08:00
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#ifdef CONFIG_SCHED_AUTOGROUP
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struct autogroup *autogroup;
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#endif
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2005-04-17 06:20:36 +08:00
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/*
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* Cumulative resource counters for dead threads in the group,
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* and for reaped dead child processes forked by this group.
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* Live threads maintain their own counters and add to these
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* in __exit_signal, except for the group leader.
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*/
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2009-02-05 19:24:15 +08:00
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cputime_t utime, stime, cutime, cstime;
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2007-10-15 23:00:19 +08:00
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cputime_t gtime;
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cputime_t cgtime;
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2013-02-26 00:25:39 +08:00
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#ifndef CONFIG_VIRT_CPU_ACCOUNTING_NATIVE
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2012-11-22 07:58:35 +08:00
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struct cputime prev_cputime;
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sched, cputime: Introduce thread_group_times()
This is a real fix for problem of utime/stime values decreasing
described in the thread:
http://lkml.org/lkml/2009/11/3/522
Now cputime is accounted in the following way:
- {u,s}time in task_struct are increased every time when the thread
is interrupted by a tick (timer interrupt).
- When a thread exits, its {u,s}time are added to signal->{u,s}time,
after adjusted by task_times().
- When all threads in a thread_group exits, accumulated {u,s}time
(and also c{u,s}time) in signal struct are added to c{u,s}time
in signal struct of the group's parent.
So {u,s}time in task struct are "raw" tick count, while
{u,s}time and c{u,s}time in signal struct are "adjusted" values.
And accounted values are used by:
- task_times(), to get cputime of a thread:
This function returns adjusted values that originates from raw
{u,s}time and scaled by sum_exec_runtime that accounted by CFS.
- thread_group_cputime(), to get cputime of a thread group:
This function returns sum of all {u,s}time of living threads in
the group, plus {u,s}time in the signal struct that is sum of
adjusted cputimes of all exited threads belonged to the group.
The problem is the return value of thread_group_cputime(),
because it is mixed sum of "raw" value and "adjusted" value:
group's {u,s}time = foreach(thread){{u,s}time} + exited({u,s}time)
This misbehavior can break {u,s}time monotonicity.
Assume that if there is a thread that have raw values greater
than adjusted values (e.g. interrupted by 1000Hz ticks 50 times
but only runs 45ms) and if it exits, cputime will decrease (e.g.
-5ms).
To fix this, we could do:
group's {u,s}time = foreach(t){task_times(t)} + exited({u,s}time)
But task_times() contains hard divisions, so applying it for
every thread should be avoided.
This patch fixes the above problem in the following way:
- Modify thread's exit (= __exit_signal()) not to use task_times().
It means {u,s}time in signal struct accumulates raw values instead
of adjusted values. As the result it makes thread_group_cputime()
to return pure sum of "raw" values.
- Introduce a new function thread_group_times(*task, *utime, *stime)
that converts "raw" values of thread_group_cputime() to "adjusted"
values, in same calculation procedure as task_times().
- Modify group's exit (= wait_task_zombie()) to use this introduced
thread_group_times(). It make c{u,s}time in signal struct to
have adjusted values like before this patch.
- Replace some thread_group_cputime() by thread_group_times().
This replacements are only applied where conveys the "adjusted"
cputime to users, and where already uses task_times() near by it.
(i.e. sys_times(), getrusage(), and /proc/<PID>/stat.)
This patch have a positive side effect:
- Before this patch, if a group contains many short-life threads
(e.g. runs 0.9ms and not interrupted by ticks), the group's
cputime could be invisible since thread's cputime was accumulated
after adjusted: imagine adjustment function as adj(ticks, runtime),
{adj(0, 0.9) + adj(0, 0.9) + ....} = {0 + 0 + ....} = 0.
After this patch it will not happen because the adjustment is
applied after accumulated.
v2:
- remove if()s, put new variables into signal_struct.
Signed-off-by: Hidetoshi Seto <seto.hidetoshi@jp.fujitsu.com>
Acked-by: Peter Zijlstra <peterz@infradead.org>
Cc: Spencer Candland <spencer@bluehost.com>
Cc: Americo Wang <xiyou.wangcong@gmail.com>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Stanislaw Gruszka <sgruszka@redhat.com>
LKML-Reference: <4B162517.8040909@jp.fujitsu.com>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-12-02 16:28:07 +08:00
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#endif
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2005-04-17 06:20:36 +08:00
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unsigned long nvcsw, nivcsw, cnvcsw, cnivcsw;
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unsigned long min_flt, maj_flt, cmin_flt, cmaj_flt;
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2007-05-11 13:22:37 +08:00
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unsigned long inblock, oublock, cinblock, coublock;
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getrusage: fill ru_maxrss value
Make ->ru_maxrss value in struct rusage filled accordingly to rss hiwater
mark. This struct is filled as a parameter to getrusage syscall.
->ru_maxrss value is set to KBs which is the way it is done in BSD
systems. /usr/bin/time (gnu time) application converts ->ru_maxrss to KBs
which seems to be incorrect behavior. Maintainer of this util was
notified by me with the patch which corrects it and cc'ed.
To make this happen we extend struct signal_struct by two fields. The
first one is ->maxrss which we use to store rss hiwater of the task. The
second one is ->cmaxrss which we use to store highest rss hiwater of all
task childs. These values are used in k_getrusage() to actually fill
->ru_maxrss. k_getrusage() uses current rss hiwater value directly if mm
struct exists.
Note:
exec() clear mm->hiwater_rss, but doesn't clear sig->maxrss.
it is intetionally behavior. *BSD getrusage have exec() inheriting.
test programs
========================================================
getrusage.c
===========
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <sys/types.h>
#include <sys/time.h>
#include <sys/resource.h>
#include <sys/types.h>
#include <sys/wait.h>
#include <unistd.h>
#include <signal.h>
#include <sys/mman.h>
#include "common.h"
#define err(str) perror(str), exit(1)
int main(int argc, char** argv)
{
int status;
printf("allocate 100MB\n");
consume(100);
printf("testcase1: fork inherit? \n");
printf(" expect: initial.self ~= child.self\n");
show_rusage("initial");
if (__fork()) {
wait(&status);
} else {
show_rusage("fork child");
_exit(0);
}
printf("\n");
printf("testcase2: fork inherit? (cont.) \n");
printf(" expect: initial.children ~= 100MB, but child.children = 0\n");
show_rusage("initial");
if (__fork()) {
wait(&status);
} else {
show_rusage("child");
_exit(0);
}
printf("\n");
printf("testcase3: fork + malloc \n");
printf(" expect: child.self ~= initial.self + 50MB\n");
show_rusage("initial");
if (__fork()) {
wait(&status);
} else {
printf("allocate +50MB\n");
consume(50);
show_rusage("fork child");
_exit(0);
}
printf("\n");
printf("testcase4: grandchild maxrss\n");
printf(" expect: post_wait.children ~= 300MB\n");
show_rusage("initial");
if (__fork()) {
wait(&status);
show_rusage("post_wait");
} else {
system("./child -n 0 -g 300");
_exit(0);
}
printf("\n");
printf("testcase5: zombie\n");
printf(" expect: pre_wait ~= initial, IOW the zombie process is not accounted.\n");
printf(" post_wait ~= 400MB, IOW wait() collect child's max_rss. \n");
show_rusage("initial");
if (__fork()) {
sleep(1); /* children become zombie */
show_rusage("pre_wait");
wait(&status);
show_rusage("post_wait");
} else {
system("./child -n 400");
_exit(0);
}
printf("\n");
printf("testcase6: SIG_IGN\n");
printf(" expect: initial ~= after_zombie (child's 500MB alloc should be ignored).\n");
show_rusage("initial");
signal(SIGCHLD, SIG_IGN);
if (__fork()) {
sleep(1); /* children become zombie */
show_rusage("after_zombie");
} else {
system("./child -n 500");
_exit(0);
}
printf("\n");
signal(SIGCHLD, SIG_DFL);
printf("testcase7: exec (without fork) \n");
printf(" expect: initial ~= exec \n");
show_rusage("initial");
execl("./child", "child", "-v", NULL);
return 0;
}
child.c
=======
#include <sys/types.h>
#include <unistd.h>
#include <sys/types.h>
#include <sys/wait.h>
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <sys/types.h>
#include <sys/time.h>
#include <sys/resource.h>
#include "common.h"
int main(int argc, char** argv)
{
int status;
int c;
long consume_size = 0;
long grandchild_consume_size = 0;
int show = 0;
while ((c = getopt(argc, argv, "n:g:v")) != -1) {
switch (c) {
case 'n':
consume_size = atol(optarg);
break;
case 'v':
show = 1;
break;
case 'g':
grandchild_consume_size = atol(optarg);
break;
default:
break;
}
}
if (show)
show_rusage("exec");
if (consume_size) {
printf("child alloc %ldMB\n", consume_size);
consume(consume_size);
}
if (grandchild_consume_size) {
if (fork()) {
wait(&status);
} else {
printf("grandchild alloc %ldMB\n", grandchild_consume_size);
consume(grandchild_consume_size);
exit(0);
}
}
return 0;
}
common.c
========
#include <stdio.h>
#include <stdlib.h>
#include <string.h>
#include <sys/types.h>
#include <sys/time.h>
#include <sys/resource.h>
#include <sys/types.h>
#include <sys/wait.h>
#include <unistd.h>
#include <signal.h>
#include <sys/mman.h>
#include "common.h"
#define err(str) perror(str), exit(1)
void show_rusage(char *prefix)
{
int err, err2;
struct rusage rusage_self;
struct rusage rusage_children;
printf("%s: ", prefix);
err = getrusage(RUSAGE_SELF, &rusage_self);
if (!err)
printf("self %ld ", rusage_self.ru_maxrss);
err2 = getrusage(RUSAGE_CHILDREN, &rusage_children);
if (!err2)
printf("children %ld ", rusage_children.ru_maxrss);
printf("\n");
}
/* Some buggy OS need this worthless CPU waste. */
void make_pagefault(void)
{
void *addr;
int size = getpagesize();
int i;
for (i=0; i<1000; i++) {
addr = mmap(NULL, size, PROT_READ | PROT_WRITE, MAP_PRIVATE | MAP_ANON, -1, 0);
if (addr == MAP_FAILED)
err("make_pagefault");
memset(addr, 0, size);
munmap(addr, size);
}
}
void consume(int mega)
{
size_t sz = mega * 1024 * 1024;
void *ptr;
ptr = malloc(sz);
memset(ptr, 0, sz);
make_pagefault();
}
pid_t __fork(void)
{
pid_t pid;
pid = fork();
make_pagefault();
return pid;
}
common.h
========
void show_rusage(char *prefix);
void make_pagefault(void);
void consume(int mega);
pid_t __fork(void);
FreeBSD result (expected result)
========================================================
allocate 100MB
testcase1: fork inherit?
expect: initial.self ~= child.self
initial: self 103492 children 0
fork child: self 103540 children 0
testcase2: fork inherit? (cont.)
expect: initial.children ~= 100MB, but child.children = 0
initial: self 103540 children 103540
child: self 103564 children 0
testcase3: fork + malloc
expect: child.self ~= initial.self + 50MB
initial: self 103564 children 103564
allocate +50MB
fork child: self 154860 children 0
testcase4: grandchild maxrss
expect: post_wait.children ~= 300MB
initial: self 103564 children 154860
grandchild alloc 300MB
post_wait: self 103564 children 308720
testcase5: zombie
expect: pre_wait ~= initial, IOW the zombie process is not accounted.
post_wait ~= 400MB, IOW wait() collect child's max_rss.
initial: self 103564 children 308720
child alloc 400MB
pre_wait: self 103564 children 308720
post_wait: self 103564 children 411312
testcase6: SIG_IGN
expect: initial ~= after_zombie (child's 500MB alloc should be ignored).
initial: self 103564 children 411312
child alloc 500MB
after_zombie: self 103624 children 411312
testcase7: exec (without fork)
expect: initial ~= exec
initial: self 103624 children 411312
exec: self 103624 children 411312
Linux result (actual test result)
========================================================
allocate 100MB
testcase1: fork inherit?
expect: initial.self ~= child.self
initial: self 102848 children 0
fork child: self 102572 children 0
testcase2: fork inherit? (cont.)
expect: initial.children ~= 100MB, but child.children = 0
initial: self 102876 children 102644
child: self 102572 children 0
testcase3: fork + malloc
expect: child.self ~= initial.self + 50MB
initial: self 102876 children 102644
allocate +50MB
fork child: self 153804 children 0
testcase4: grandchild maxrss
expect: post_wait.children ~= 300MB
initial: self 102876 children 153864
grandchild alloc 300MB
post_wait: self 102876 children 307536
testcase5: zombie
expect: pre_wait ~= initial, IOW the zombie process is not accounted.
post_wait ~= 400MB, IOW wait() collect child's max_rss.
initial: self 102876 children 307536
child alloc 400MB
pre_wait: self 102876 children 307536
post_wait: self 102876 children 410076
testcase6: SIG_IGN
expect: initial ~= after_zombie (child's 500MB alloc should be ignored).
initial: self 102876 children 410076
child alloc 500MB
after_zombie: self 102880 children 410076
testcase7: exec (without fork)
expect: initial ~= exec
initial: self 102880 children 410076
exec: self 102880 children 410076
Signed-off-by: Jiri Pirko <jpirko@redhat.com>
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Hugh Dickins <hugh.dickins@tiscali.co.uk>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-23 07:44:10 +08:00
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unsigned long maxrss, cmaxrss;
|
2008-07-28 06:48:12 +08:00
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struct task_io_accounting ioac;
|
2005-04-17 06:20:36 +08:00
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2009-02-05 19:24:15 +08:00
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/*
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* Cumulative ns of schedule CPU time fo dead threads in the
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* group, not including a zombie group leader, (This only differs
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* from jiffies_to_ns(utime + stime) if sched_clock uses something
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* other than jiffies.)
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*/
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unsigned long long sum_sched_runtime;
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2005-04-17 06:20:36 +08:00
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/*
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* We don't bother to synchronize most readers of this at all,
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* because there is no reader checking a limit that actually needs
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* to get both rlim_cur and rlim_max atomically, and either one
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* alone is a single word that can safely be read normally.
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* getrlimit/setrlimit use task_lock(current->group_leader) to
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* protect this instead of the siglock, because they really
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* have no need to disable irqs.
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*/
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struct rlimit rlim[RLIM_NLIMITS];
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2006-06-25 20:49:24 +08:00
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#ifdef CONFIG_BSD_PROCESS_ACCT
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struct pacct_struct pacct; /* per-process accounting information */
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#endif
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2006-07-14 15:24:44 +08:00
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#ifdef CONFIG_TASKSTATS
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struct taskstats *stats;
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#endif
|
Audit: add TTY input auditing
Add TTY input auditing, used to audit system administrator's actions. This is
required by various security standards such as DCID 6/3 and PCI to provide
non-repudiation of administrator's actions and to allow a review of past
actions if the administrator seems to overstep their duties or if the system
becomes misconfigured for unknown reasons. These requirements do not make it
necessary to audit TTY output as well.
Compared to an user-space keylogger, this approach records TTY input using the
audit subsystem, correlated with other audit events, and it is completely
transparent to the user-space application (e.g. the console ioctls still
work).
TTY input auditing works on a higher level than auditing all system calls
within the session, which would produce an overwhelming amount of mostly
useless audit events.
Add an "audit_tty" attribute, inherited across fork (). Data read from TTYs
by process with the attribute is sent to the audit subsystem by the kernel.
The audit netlink interface is extended to allow modifying the audit_tty
attribute, and to allow sending explanatory audit events from user-space (for
example, a shell might send an event containing the final command, after the
interactive command-line editing and history expansion is performed, which
might be difficult to decipher from the TTY input alone).
Because the "audit_tty" attribute is inherited across fork (), it would be set
e.g. for sshd restarted within an audited session. To prevent this, the
audit_tty attribute is cleared when a process with no open TTY file
descriptors (e.g. after daemon startup) opens a TTY.
See https://www.redhat.com/archives/linux-audit/2007-June/msg00000.html for a
more detailed rationale document for an older version of this patch.
[akpm@linux-foundation.org: build fix]
Signed-off-by: Miloslav Trmac <mitr@redhat.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Alan Cox <alan@lxorguk.ukuu.org.uk>
Cc: Paul Fulghum <paulkf@microgate.com>
Cc: Casey Schaufler <casey@schaufler-ca.com>
Cc: Steve Grubb <sgrubb@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-07-16 14:40:56 +08:00
|
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#ifdef CONFIG_AUDIT
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unsigned audit_tty;
|
2013-05-04 02:03:50 +08:00
|
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unsigned audit_tty_log_passwd;
|
Audit: add TTY input auditing
Add TTY input auditing, used to audit system administrator's actions. This is
required by various security standards such as DCID 6/3 and PCI to provide
non-repudiation of administrator's actions and to allow a review of past
actions if the administrator seems to overstep their duties or if the system
becomes misconfigured for unknown reasons. These requirements do not make it
necessary to audit TTY output as well.
Compared to an user-space keylogger, this approach records TTY input using the
audit subsystem, correlated with other audit events, and it is completely
transparent to the user-space application (e.g. the console ioctls still
work).
TTY input auditing works on a higher level than auditing all system calls
within the session, which would produce an overwhelming amount of mostly
useless audit events.
Add an "audit_tty" attribute, inherited across fork (). Data read from TTYs
by process with the attribute is sent to the audit subsystem by the kernel.
The audit netlink interface is extended to allow modifying the audit_tty
attribute, and to allow sending explanatory audit events from user-space (for
example, a shell might send an event containing the final command, after the
interactive command-line editing and history expansion is performed, which
might be difficult to decipher from the TTY input alone).
Because the "audit_tty" attribute is inherited across fork (), it would be set
e.g. for sshd restarted within an audited session. To prevent this, the
audit_tty attribute is cleared when a process with no open TTY file
descriptors (e.g. after daemon startup) opens a TTY.
See https://www.redhat.com/archives/linux-audit/2007-June/msg00000.html for a
more detailed rationale document for an older version of this patch.
[akpm@linux-foundation.org: build fix]
Signed-off-by: Miloslav Trmac <mitr@redhat.com>
Cc: Al Viro <viro@zeniv.linux.org.uk>
Cc: Alan Cox <alan@lxorguk.ukuu.org.uk>
Cc: Paul Fulghum <paulkf@microgate.com>
Cc: Casey Schaufler <casey@schaufler-ca.com>
Cc: Steve Grubb <sgrubb@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-07-16 14:40:56 +08:00
|
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struct tty_audit_buf *tty_audit_buf;
|
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#endif
|
2011-05-27 07:25:18 +08:00
|
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|
#ifdef CONFIG_CGROUPS
|
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/*
|
2011-12-13 10:12:21 +08:00
|
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* group_rwsem prevents new tasks from entering the threadgroup and
|
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* member tasks from exiting,a more specifically, setting of
|
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* PF_EXITING. fork and exit paths are protected with this rwsem
|
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* using threadgroup_change_begin/end(). Users which require
|
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* threadgroup to remain stable should use threadgroup_[un]lock()
|
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* which also takes care of exec path. Currently, cgroup is the
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* only user.
|
2011-05-27 07:25:18 +08:00
|
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*/
|
2011-12-13 10:12:21 +08:00
|
|
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struct rw_semaphore group_rwsem;
|
2011-05-27 07:25:18 +08:00
|
|
|
#endif
|
oom: move oom_adj value from task_struct to signal_struct
Currently, OOM logic callflow is here.
__out_of_memory()
select_bad_process() for each task
badness() calculate badness of one task
oom_kill_process() search child
oom_kill_task() kill target task and mm shared tasks with it
example, process-A have two thread, thread-A and thread-B and it have very
fat memory and each thread have following oom_adj and oom_score.
thread-A: oom_adj = OOM_DISABLE, oom_score = 0
thread-B: oom_adj = 0, oom_score = very-high
Then, select_bad_process() select thread-B, but oom_kill_task() refuse
kill the task because thread-A have OOM_DISABLE. Thus __out_of_memory()
call select_bad_process() again. but select_bad_process() select the same
task. It mean kernel fall in livelock.
The fact is, select_bad_process() must select killable task. otherwise
OOM logic go into livelock.
And root cause is, oom_adj shouldn't be per-thread value. it should be
per-process value because OOM-killer kill a process, not thread. Thus
This patch moves oomkilladj (now more appropriately named oom_adj) from
struct task_struct to struct signal_struct. it naturally prevent
select_bad_process() choose wrong task.
Signed-off-by: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Cc: Paul Menage <menage@google.com>
Cc: David Rientjes <rientjes@google.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Rik van Riel <riel@redhat.com>
Cc: Oleg Nesterov <oleg@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-22 08:03:13 +08:00
|
|
|
|
2012-12-12 08:02:56 +08:00
|
|
|
oom_flags_t oom_flags;
|
2012-12-12 08:02:54 +08:00
|
|
|
short oom_score_adj; /* OOM kill score adjustment */
|
|
|
|
|
short oom_score_adj_min; /* OOM kill score adjustment min value.
|
|
|
|
|
* Only settable by CAP_SYS_RESOURCE. */
|
2010-10-28 06:34:08 +08:00
|
|
|
|
|
|
|
|
struct mutex cred_guard_mutex; /* guard against foreign influences on
|
|
|
|
|
* credential calculations
|
|
|
|
|
* (notably. ptrace) */
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Bits in flags field of signal_struct.
|
|
|
|
|
*/
|
|
|
|
|
#define SIGNAL_STOP_STOPPED 0x00000001 /* job control stop in effect */
|
signal: Turn SIGNAL_STOP_DEQUEUED into GROUP_STOP_DEQUEUED
This patch moves SIGNAL_STOP_DEQUEUED from signal_struct->flags to
task_struct->group_stop, and thus makes it per-thread.
Like SIGNAL_STOP_DEQUEUED, GROUP_STOP_DEQUEUED can be false-positive
after return from get_signal_to_deliver(), this is fine. The only
purpose of this bit is: we can drop ->siglock after __dequeue_signal()
returns the sig_kernel_stop() signal and before we call
do_signal_stop(), in this case we must not miss SIGCONT if it comes in
between.
But, unlike SIGNAL_STOP_DEQUEUED, GROUP_STOP_DEQUEUED can not be
false-positive in do_signal_stop() if multiple threads dequeue the
sig_kernel_stop() signal at the same time.
Consider two threads T1 and T2, SIGTTIN has a hanlder.
- T1 dequeues SIGTSTP and sets SIGNAL_STOP_DEQUEUED, then
it drops ->siglock
- SIGCONT comes and clears SIGNAL_STOP_DEQUEUED, SIGTSTP
should be cancelled.
- T2 dequeues SIGTTIN and sets SIGNAL_STOP_DEQUEUED again.
Since we have a handler we should not stop, T2 returns
to usermode to run the handler.
- T1 continues, calls do_signal_stop() and wrongly starts
the group stop because SIGNAL_STOP_DEQUEUED was restored
in between.
With or without this change:
- we need to do something with ptrace_signal() which can
return SIGSTOP, but this needs another discussion
- SIGSTOP can be lost if it races with the mt exec, will
be fixed later.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Signed-off-by: Tejun Heo <tj@kernel.org>
2011-04-02 02:12:38 +08:00
|
|
|
#define SIGNAL_STOP_CONTINUED 0x00000002 /* SIGCONT since WCONTINUED reap */
|
|
|
|
|
#define SIGNAL_GROUP_EXIT 0x00000004 /* group exit in progress */
|
2013-05-01 06:28:10 +08:00
|
|
|
#define SIGNAL_GROUP_COREDUMP 0x00000008 /* coredump in progress */
|
2008-04-30 15:52:44 +08:00
|
|
|
/*
|
|
|
|
|
* Pending notifications to parent.
|
|
|
|
|
*/
|
|
|
|
|
#define SIGNAL_CLD_STOPPED 0x00000010
|
|
|
|
|
#define SIGNAL_CLD_CONTINUED 0x00000020
|
|
|
|
|
#define SIGNAL_CLD_MASK (SIGNAL_CLD_STOPPED|SIGNAL_CLD_CONTINUED)
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2008-04-30 15:53:03 +08:00
|
|
|
#define SIGNAL_UNKILLABLE 0x00000040 /* for init: ignore fatal signals */
|
|
|
|
|
|
2008-02-05 14:27:24 +08:00
|
|
|
/* If true, all threads except ->group_exit_task have pending SIGKILL */
|
|
|
|
|
static inline int signal_group_exit(const struct signal_struct *sig)
|
|
|
|
|
{
|
|
|
|
|
return (sig->flags & SIGNAL_GROUP_EXIT) ||
|
|
|
|
|
(sig->group_exit_task != NULL);
|
|
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
|
|
|
|
* Some day this will be a full-fledged user tracking system..
|
|
|
|
|
*/
|
|
|
|
|
struct user_struct {
|
|
|
|
|
atomic_t __count; /* reference count */
|
|
|
|
|
atomic_t processes; /* How many processes does this user have? */
|
|
|
|
|
atomic_t files; /* How many open files does this user have? */
|
|
|
|
|
atomic_t sigpending; /* How many pending signals does this user have? */
|
2006-06-02 04:10:59 +08:00
|
|
|
#ifdef CONFIG_INOTIFY_USER
|
[PATCH] inotify
inotify is intended to correct the deficiencies of dnotify, particularly
its inability to scale and its terrible user interface:
* dnotify requires the opening of one fd per each directory
that you intend to watch. This quickly results in too many
open files and pins removable media, preventing unmount.
* dnotify is directory-based. You only learn about changes to
directories. Sure, a change to a file in a directory affects
the directory, but you are then forced to keep a cache of
stat structures.
* dnotify's interface to user-space is awful. Signals?
inotify provides a more usable, simple, powerful solution to file change
notification:
* inotify's interface is a system call that returns a fd, not SIGIO.
You get a single fd, which is select()-able.
* inotify has an event that says "the filesystem that the item
you were watching is on was unmounted."
* inotify can watch directories or files.
Inotify is currently used by Beagle (a desktop search infrastructure),
Gamin (a FAM replacement), and other projects.
See Documentation/filesystems/inotify.txt.
Signed-off-by: Robert Love <rml@novell.com>
Cc: John McCutchan <ttb@tentacle.dhs.org>
Cc: Christoph Hellwig <hch@lst.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2005-07-13 05:06:03 +08:00
|
|
|
atomic_t inotify_watches; /* How many inotify watches does this user have? */
|
|
|
|
|
atomic_t inotify_devs; /* How many inotify devs does this user have opened? */
|
|
|
|
|
#endif
|
2010-10-29 05:21:58 +08:00
|
|
|
#ifdef CONFIG_FANOTIFY
|
|
|
|
|
atomic_t fanotify_listeners;
|
|
|
|
|
#endif
|
epoll: introduce resource usage limits
It has been thought that the per-user file descriptors limit would also
limit the resources that a normal user can request via the epoll
interface. Vegard Nossum reported a very simple program (a modified
version attached) that can make a normal user to request a pretty large
amount of kernel memory, well within the its maximum number of fds. To
solve such problem, default limits are now imposed, and /proc based
configuration has been introduced. A new directory has been created,
named /proc/sys/fs/epoll/ and inside there, there are two configuration
points:
max_user_instances = Maximum number of devices - per user
max_user_watches = Maximum number of "watched" fds - per user
The current default for "max_user_watches" limits the memory used by epoll
to store "watches", to 1/32 of the amount of the low RAM. As example, a
256MB 32bit machine, will have "max_user_watches" set to roughly 90000.
That should be enough to not break existing heavy epoll users. The
default value for "max_user_instances" is set to 128, that should be
enough too.
This also changes the userspace, because a new error code can now come out
from EPOLL_CTL_ADD (-ENOSPC). The EMFILE from epoll_create() was already
listed, so that should be ok.
[akpm@linux-foundation.org: use get_current_user()]
Signed-off-by: Davide Libenzi <davidel@xmailserver.org>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: <stable@kernel.org>
Cc: Cyrill Gorcunov <gorcunov@gmail.com>
Reported-by: Vegard Nossum <vegardno@ifi.uio.no>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-12-02 05:13:55 +08:00
|
|
|
#ifdef CONFIG_EPOLL
|
2011-01-13 09:00:01 +08:00
|
|
|
atomic_long_t epoll_watches; /* The number of file descriptors currently watched */
|
epoll: introduce resource usage limits
It has been thought that the per-user file descriptors limit would also
limit the resources that a normal user can request via the epoll
interface. Vegard Nossum reported a very simple program (a modified
version attached) that can make a normal user to request a pretty large
amount of kernel memory, well within the its maximum number of fds. To
solve such problem, default limits are now imposed, and /proc based
configuration has been introduced. A new directory has been created,
named /proc/sys/fs/epoll/ and inside there, there are two configuration
points:
max_user_instances = Maximum number of devices - per user
max_user_watches = Maximum number of "watched" fds - per user
The current default for "max_user_watches" limits the memory used by epoll
to store "watches", to 1/32 of the amount of the low RAM. As example, a
256MB 32bit machine, will have "max_user_watches" set to roughly 90000.
That should be enough to not break existing heavy epoll users. The
default value for "max_user_instances" is set to 128, that should be
enough too.
This also changes the userspace, because a new error code can now come out
from EPOLL_CTL_ADD (-ENOSPC). The EMFILE from epoll_create() was already
listed, so that should be ok.
[akpm@linux-foundation.org: use get_current_user()]
Signed-off-by: Davide Libenzi <davidel@xmailserver.org>
Cc: Michael Kerrisk <mtk.manpages@gmail.com>
Cc: <stable@kernel.org>
Cc: Cyrill Gorcunov <gorcunov@gmail.com>
Reported-by: Vegard Nossum <vegardno@ifi.uio.no>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-12-02 05:13:55 +08:00
|
|
|
#endif
|
2007-10-17 14:30:09 +08:00
|
|
|
#ifdef CONFIG_POSIX_MQUEUE
|
2005-04-17 06:20:36 +08:00
|
|
|
/* protected by mq_lock */
|
|
|
|
|
unsigned long mq_bytes; /* How many bytes can be allocated to mqueue? */
|
2007-10-17 14:30:09 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
unsigned long locked_shm; /* How many pages of mlocked shm ? */
|
|
|
|
|
|
|
|
|
|
#ifdef CONFIG_KEYS
|
|
|
|
|
struct key *uid_keyring; /* UID specific keyring */
|
|
|
|
|
struct key *session_keyring; /* UID's default session keyring */
|
|
|
|
|
#endif
|
|
|
|
|
|
|
|
|
|
/* Hash table maintenance information */
|
2007-09-19 13:46:44 +08:00
|
|
|
struct hlist_node uidhash_node;
|
2011-11-17 15:20:58 +08:00
|
|
|
kuid_t uid;
|
2007-10-15 23:00:09 +08:00
|
|
|
|
perf: Do the big rename: Performance Counters -> Performance Events
Bye-bye Performance Counters, welcome Performance Events!
In the past few months the perfcounters subsystem has grown out its
initial role of counting hardware events, and has become (and is
becoming) a much broader generic event enumeration, reporting, logging,
monitoring, analysis facility.
Naming its core object 'perf_counter' and naming the subsystem
'perfcounters' has become more and more of a misnomer. With pending
code like hw-breakpoints support the 'counter' name is less and
less appropriate.
All in one, we've decided to rename the subsystem to 'performance
events' and to propagate this rename through all fields, variables
and API names. (in an ABI compatible fashion)
The word 'event' is also a bit shorter than 'counter' - which makes
it slightly more convenient to write/handle as well.
Thanks goes to Stephane Eranian who first observed this misnomer and
suggested a rename.
User-space tooling and ABI compatibility is not affected - this patch
should be function-invariant. (Also, defconfigs were not touched to
keep the size down.)
This patch has been generated via the following script:
FILES=$(find * -type f | grep -vE 'oprofile|[^K]config')
sed -i \
-e 's/PERF_EVENT_/PERF_RECORD_/g' \
-e 's/PERF_COUNTER/PERF_EVENT/g' \
-e 's/perf_counter/perf_event/g' \
-e 's/nb_counters/nb_events/g' \
-e 's/swcounter/swevent/g' \
-e 's/tpcounter_event/tp_event/g' \
$FILES
for N in $(find . -name perf_counter.[ch]); do
M=$(echo $N | sed 's/perf_counter/perf_event/g')
mv $N $M
done
FILES=$(find . -name perf_event.*)
sed -i \
-e 's/COUNTER_MASK/REG_MASK/g' \
-e 's/COUNTER/EVENT/g' \
-e 's/\<event\>/event_id/g' \
-e 's/counter/event/g' \
-e 's/Counter/Event/g' \
$FILES
... to keep it as correct as possible. This script can also be
used by anyone who has pending perfcounters patches - it converts
a Linux kernel tree over to the new naming. We tried to time this
change to the point in time where the amount of pending patches
is the smallest: the end of the merge window.
Namespace clashes were fixed up in a preparatory patch - and some
stylistic fallout will be fixed up in a subsequent patch.
( NOTE: 'counters' are still the proper terminology when we deal
with hardware registers - and these sed scripts are a bit
over-eager in renaming them. I've undone some of that, but
in case there's something left where 'counter' would be
better than 'event' we can undo that on an individual basis
instead of touching an otherwise nicely automated patch. )
Suggested-by: Stephane Eranian <eranian@google.com>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Acked-by: Paul Mackerras <paulus@samba.org>
Reviewed-by: Arjan van de Ven <arjan@linux.intel.com>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Arnaldo Carvalho de Melo <acme@redhat.com>
Cc: Frederic Weisbecker <fweisbec@gmail.com>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Howells <dhowells@redhat.com>
Cc: Kyle McMartin <kyle@mcmartin.ca>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Cc: <linux-arch@vger.kernel.org>
LKML-Reference: <new-submission>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 18:02:48 +08:00
|
|
|
#ifdef CONFIG_PERF_EVENTS
|
2009-05-15 21:19:27 +08:00
|
|
|
atomic_long_t locked_vm;
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
|
|
|
|
|
2007-11-02 20:47:53 +08:00
|
|
|
extern int uids_sysfs_init(void);
|
2007-10-15 23:00:14 +08:00
|
|
|
|
2011-11-17 15:20:58 +08:00
|
|
|
extern struct user_struct *find_user(kuid_t);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
extern struct user_struct root_user;
|
|
|
|
|
#define INIT_USER (&root_user)
|
|
|
|
|
|
2008-11-14 07:39:16 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
struct backing_dev_info;
|
|
|
|
|
struct reclaim_state;
|
|
|
|
|
|
2006-07-14 15:24:38 +08:00
|
|
|
#if defined(CONFIG_SCHEDSTATS) || defined(CONFIG_TASK_DELAY_ACCT)
|
2005-04-17 06:20:36 +08:00
|
|
|
struct sched_info {
|
|
|
|
|
/* cumulative counters */
|
2007-10-15 23:00:12 +08:00
|
|
|
unsigned long pcount; /* # of times run on this cpu */
|
2008-12-17 15:41:22 +08:00
|
|
|
unsigned long long run_delay; /* time spent waiting on a runqueue */
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* timestamps */
|
2007-07-10 00:52:00 +08:00
|
|
|
unsigned long long last_arrival,/* when we last ran on a cpu */
|
|
|
|
|
last_queued; /* when we were last queued to run */
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
2006-07-14 15:24:38 +08:00
|
|
|
#endif /* defined(CONFIG_SCHEDSTATS) || defined(CONFIG_TASK_DELAY_ACCT) */
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2006-07-14 15:24:36 +08:00
|
|
|
#ifdef CONFIG_TASK_DELAY_ACCT
|
|
|
|
|
struct task_delay_info {
|
|
|
|
|
spinlock_t lock;
|
|
|
|
|
unsigned int flags; /* Private per-task flags */
|
|
|
|
|
|
|
|
|
|
/* For each stat XXX, add following, aligned appropriately
|
|
|
|
|
*
|
|
|
|
|
* struct timespec XXX_start, XXX_end;
|
|
|
|
|
* u64 XXX_delay;
|
|
|
|
|
* u32 XXX_count;
|
|
|
|
|
*
|
|
|
|
|
* Atomicity of updates to XXX_delay, XXX_count protected by
|
|
|
|
|
* single lock above (split into XXX_lock if contention is an issue).
|
|
|
|
|
*/
|
2006-07-14 15:24:37 +08:00
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* XXX_count is incremented on every XXX operation, the delay
|
|
|
|
|
* associated with the operation is added to XXX_delay.
|
|
|
|
|
* XXX_delay contains the accumulated delay time in nanoseconds.
|
|
|
|
|
*/
|
|
|
|
|
struct timespec blkio_start, blkio_end; /* Shared by blkio, swapin */
|
|
|
|
|
u64 blkio_delay; /* wait for sync block io completion */
|
|
|
|
|
u64 swapin_delay; /* wait for swapin block io completion */
|
|
|
|
|
u32 blkio_count; /* total count of the number of sync block */
|
|
|
|
|
/* io operations performed */
|
|
|
|
|
u32 swapin_count; /* total count of the number of swapin block */
|
|
|
|
|
/* io operations performed */
|
2008-07-25 16:48:52 +08:00
|
|
|
|
|
|
|
|
struct timespec freepages_start, freepages_end;
|
|
|
|
|
u64 freepages_delay; /* wait for memory reclaim */
|
|
|
|
|
u32 freepages_count; /* total count of memory reclaim */
|
2006-07-14 15:24:36 +08:00
|
|
|
};
|
2006-07-14 15:24:38 +08:00
|
|
|
#endif /* CONFIG_TASK_DELAY_ACCT */
|
|
|
|
|
|
|
|
|
|
static inline int sched_info_on(void)
|
|
|
|
|
{
|
|
|
|
|
#ifdef CONFIG_SCHEDSTATS
|
|
|
|
|
return 1;
|
|
|
|
|
#elif defined(CONFIG_TASK_DELAY_ACCT)
|
|
|
|
|
extern int delayacct_on;
|
|
|
|
|
return delayacct_on;
|
|
|
|
|
#else
|
|
|
|
|
return 0;
|
2006-07-14 15:24:36 +08:00
|
|
|
#endif
|
2006-07-14 15:24:38 +08:00
|
|
|
}
|
2006-07-14 15:24:36 +08:00
|
|
|
|
2007-07-10 00:51:57 +08:00
|
|
|
enum cpu_idle_type {
|
|
|
|
|
CPU_IDLE,
|
|
|
|
|
CPU_NOT_IDLE,
|
|
|
|
|
CPU_NEWLY_IDLE,
|
|
|
|
|
CPU_MAX_IDLE_TYPES
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
|
|
|
|
|
2011-05-19 01:09:39 +08:00
|
|
|
/*
|
|
|
|
|
* Increase resolution of cpu_power calculations
|
|
|
|
|
*/
|
|
|
|
|
#define SCHED_POWER_SHIFT 10
|
|
|
|
|
#define SCHED_POWER_SCALE (1L << SCHED_POWER_SHIFT)
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2011-05-19 01:09:39 +08:00
|
|
|
/*
|
|
|
|
|
* sched-domains (multiprocessor balancing) declarations:
|
|
|
|
|
*/
|
[PATCH] sched: implement smpnice
Problem:
The introduction of separate run queues per CPU has brought with it "nice"
enforcement problems that are best described by a simple example.
For the sake of argument suppose that on a single CPU machine with a
nice==19 hard spinner and a nice==0 hard spinner running that the nice==0
task gets 95% of the CPU and the nice==19 task gets 5% of the CPU. Now
suppose that there is a system with 2 CPUs and 2 nice==19 hard spinners and
2 nice==0 hard spinners running. The user of this system would be entitled
to expect that the nice==0 tasks each get 95% of a CPU and the nice==19
tasks only get 5% each. However, whether this expectation is met is pretty
much down to luck as there are four equally likely distributions of the
tasks to the CPUs that the load balancing code will consider to be balanced
with loads of 2.0 for each CPU. Two of these distributions involve one
nice==0 and one nice==19 task per CPU and in these circumstances the users
expectations will be met. The other two distributions both involve both
nice==0 tasks being on one CPU and both nice==19 being on the other CPU and
each task will get 50% of a CPU and the user's expectations will not be
met.
Solution:
The solution to this problem that is implemented in the attached patch is
to use weighted loads when determining if the system is balanced and, when
an imbalance is detected, to move an amount of weighted load between run
queues (as opposed to a number of tasks) to restore the balance. Once
again, the easiest way to explain why both of these measures are necessary
is to use a simple example. Suppose that (in a slight variation of the
above example) that we have a two CPU system with 4 nice==0 and 4 nice=19
hard spinning tasks running and that the 4 nice==0 tasks are on one CPU and
the 4 nice==19 tasks are on the other CPU. The weighted loads for the two
CPUs would be 4.0 and 0.2 respectively and the load balancing code would
move 2 tasks resulting in one CPU with a load of 2.0 and the other with
load of 2.2. If this was considered to be a big enough imbalance to
justify moving a task and that task was moved using the current
move_tasks() then it would move the highest priority task that it found and
this would result in one CPU with a load of 3.0 and the other with a load
of 1.2 which would result in the movement of a task in the opposite
direction and so on -- infinite loop. If, on the other hand, an amount of
load to be moved is calculated from the imbalance (in this case 0.1) and
move_tasks() skips tasks until it find ones whose contributions to the
weighted load are less than this amount it would move two of the nice==19
tasks resulting in a system with 2 nice==0 and 2 nice=19 on each CPU with
loads of 2.1 for each CPU.
One of the advantages of this mechanism is that on a system where all tasks
have nice==0 the load balancing calculations would be mathematically
identical to the current load balancing code.
Notes:
struct task_struct:
has a new field load_weight which (in a trade off of space for speed)
stores the contribution that this task makes to a CPU's weighted load when
it is runnable.
struct runqueue:
has a new field raw_weighted_load which is the sum of the load_weight
values for the currently runnable tasks on this run queue. This field
always needs to be updated when nr_running is updated so two new inline
functions inc_nr_running() and dec_nr_running() have been created to make
sure that this happens. This also offers a convenient way to optimize away
this part of the smpnice mechanism when CONFIG_SMP is not defined.
int try_to_wake_up():
in this function the value SCHED_LOAD_BALANCE is used to represent the load
contribution of a single task in various calculations in the code that
decides which CPU to put the waking task on. While this would be a valid
on a system where the nice values for the runnable tasks were distributed
evenly around zero it will lead to anomalous load balancing if the
distribution is skewed in either direction. To overcome this problem
SCHED_LOAD_SCALE has been replaced by the load_weight for the relevant task
or by the average load_weight per task for the queue in question (as
appropriate).
int move_tasks():
The modifications to this function were complicated by the fact that
active_load_balance() uses it to move exactly one task without checking
whether an imbalance actually exists. This precluded the simple
overloading of max_nr_move with max_load_move and necessitated the addition
of the latter as an extra argument to the function. The internal
implementation is then modified to move up to max_nr_move tasks and
max_load_move of weighted load. This slightly complicates the code where
move_tasks() is called and if ever active_load_balance() is changed to not
use move_tasks() the implementation of move_tasks() should be simplified
accordingly.
struct sched_group *find_busiest_group():
Similar to try_to_wake_up(), there are places in this function where
SCHED_LOAD_SCALE is used to represent the load contribution of a single
task and the same issues are created. A similar solution is adopted except
that it is now the average per task contribution to a group's load (as
opposed to a run queue) that is required. As this value is not directly
available from the group it is calculated on the fly as the queues in the
groups are visited when determining the busiest group.
A key change to this function is that it is no longer to scale down
*imbalance on exit as move_tasks() uses the load in its scaled form.
void set_user_nice():
has been modified to update the task's load_weight field when it's nice
value and also to ensure that its run queue's raw_weighted_load field is
updated if it was runnable.
From: "Siddha, Suresh B" <suresh.b.siddha@intel.com>
With smpnice, sched groups with highest priority tasks can mask the imbalance
between the other sched groups with in the same domain. This patch fixes some
of the listed down scenarios by not considering the sched groups which are
lightly loaded.
a) on a simple 4-way MP system, if we have one high priority and 4 normal
priority tasks, with smpnice we would like to see the high priority task
scheduled on one cpu, two other cpus getting one normal task each and the
fourth cpu getting the remaining two normal tasks. but with current
smpnice extra normal priority task keeps jumping from one cpu to another
cpu having the normal priority task. This is because of the
busiest_has_loaded_cpus, nr_loaded_cpus logic.. We are not including the
cpu with high priority task in max_load calculations but including that in
total and avg_load calcuations.. leading to max_load < avg_load and load
balance between cpus running normal priority tasks(2 Vs 1) will always show
imbalanace as one normal priority and the extra normal priority task will
keep moving from one cpu to another cpu having normal priority task..
b) 4-way system with HT (8 logical processors). Package-P0 T0 has a
highest priority task, T1 is idle. Package-P1 Both T0 and T1 have 1 normal
priority task each.. P2 and P3 are idle. With this patch, one of the
normal priority tasks on P1 will be moved to P2 or P3..
c) With the current weighted smp nice calculations, it doesn't always make
sense to look at the highest weighted runqueue in the busy group..
Consider a load balance scenario on a DP with HT system, with Package-0
containing one high priority and one low priority, Package-1 containing one
low priority(with other thread being idle).. Package-1 thinks that it need
to take the low priority thread from Package-0. And find_busiest_queue()
returns the cpu thread with highest priority task.. And ultimately(with
help of active load balance) we move high priority task to Package-1. And
same continues with Package-0 now, moving high priority task from package-1
to package-0.. Even without the presence of active load balance, load
balance will fail to balance the above scenario.. Fix find_busiest_queue
to use "imbalance" when it is lightly loaded.
[kernel@kolivas.org: sched: store weighted load on up]
[kernel@kolivas.org: sched: add discrete weighted cpu load function]
[suresh.b.siddha@intel.com: sched: remove dead code]
Signed-off-by: Peter Williams <pwil3058@bigpond.com.au>
Cc: "Siddha, Suresh B" <suresh.b.siddha@intel.com>
Cc: "Chen, Kenneth W" <kenneth.w.chen@intel.com>
Acked-by: Ingo Molnar <mingo@elte.hu>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Signed-off-by: Con Kolivas <kernel@kolivas.org>
Cc: John Hawkes <hawkes@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-06-27 17:54:34 +08:00
|
|
|
#ifdef CONFIG_SMP
|
2009-09-01 16:34:33 +08:00
|
|
|
#define SD_LOAD_BALANCE 0x0001 /* Do load balancing on this domain. */
|
|
|
|
|
#define SD_BALANCE_NEWIDLE 0x0002 /* Balance when about to become idle */
|
|
|
|
|
#define SD_BALANCE_EXEC 0x0004 /* Balance on exec */
|
|
|
|
|
#define SD_BALANCE_FORK 0x0008 /* Balance on fork, clone */
|
sched: Merge select_task_rq_fair() and sched_balance_self()
The problem with wake_idle() is that is doesn't respect things like
cpu_power, which means it doesn't deal well with SMT nor the recent
RT interaction.
To cure this, it needs to do what sched_balance_self() does, which
leads to the possibility of merging select_task_rq_fair() and
sched_balance_self().
Modify sched_balance_self() to:
- update_shares() when walking up the domain tree,
(it only called it for the top domain, but it should
have done this anyway), which allows us to remove
this ugly bit from try_to_wake_up().
- do wake_affine() on the smallest domain that contains
both this (the waking) and the prev (the wakee) cpu for
WAKE invocations.
Then use the top-down balance steps it had to replace wake_idle().
This leads to the dissapearance of SD_WAKE_BALANCE and
SD_WAKE_IDLE_FAR, with SD_WAKE_IDLE replaced with SD_BALANCE_WAKE.
SD_WAKE_AFFINE needs SD_BALANCE_WAKE to be effective.
Touch all topology bits to replace the old with new SD flags --
platforms might need re-tuning, enabling SD_BALANCE_WAKE
conditionally on a NUMA distance seems like a good additional
feature, magny-core and small nehalem systems would want this
enabled, systems with slow interconnects would not.
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
LKML-Reference: <new-submission>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-10 19:50:02 +08:00
|
|
|
#define SD_BALANCE_WAKE 0x0010 /* Balance on wakeup */
|
2009-09-01 16:34:33 +08:00
|
|
|
#define SD_WAKE_AFFINE 0x0020 /* Wake task to waking CPU */
|
|
|
|
|
#define SD_SHARE_CPUPOWER 0x0080 /* Domain members share cpu power */
|
|
|
|
|
#define SD_SHARE_PKG_RESOURCES 0x0200 /* Domain members share cpu pkg resources */
|
|
|
|
|
#define SD_SERIALIZE 0x0400 /* Only a single load balancing instance */
|
2010-06-08 12:57:02 +08:00
|
|
|
#define SD_ASYM_PACKING 0x0800 /* Place busy groups earlier in the domain */
|
2009-09-01 16:34:33 +08:00
|
|
|
#define SD_PREFER_SIBLING 0x1000 /* Prefer to place tasks in a sibling domain */
|
2011-07-15 16:35:52 +08:00
|
|
|
#define SD_OVERLAP 0x2000 /* sched_domains of this level overlap */
|
2013-10-07 18:29:00 +08:00
|
|
|
#define SD_NUMA 0x4000 /* cross-node balancing */
|
2006-06-27 17:54:42 +08:00
|
|
|
|
2010-06-08 12:57:02 +08:00
|
|
|
extern int __weak arch_sd_sibiling_asym_packing(void);
|
|
|
|
|
|
2008-04-15 13:04:23 +08:00
|
|
|
struct sched_domain_attr {
|
|
|
|
|
int relax_domain_level;
|
|
|
|
|
};
|
|
|
|
|
|
|
|
|
|
#define SD_ATTR_INIT (struct sched_domain_attr) { \
|
|
|
|
|
.relax_domain_level = -1, \
|
|
|
|
|
}
|
|
|
|
|
|
2011-04-07 20:10:04 +08:00
|
|
|
extern int sched_domain_level_max;
|
|
|
|
|
|
2013-03-05 16:06:23 +08:00
|
|
|
struct sched_group;
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
struct sched_domain {
|
|
|
|
|
/* These fields must be setup */
|
|
|
|
|
struct sched_domain *parent; /* top domain must be null terminated */
|
2006-10-03 16:14:08 +08:00
|
|
|
struct sched_domain *child; /* bottom domain must be null terminated */
|
2005-04-17 06:20:36 +08:00
|
|
|
struct sched_group *groups; /* the balancing groups of the domain */
|
|
|
|
|
unsigned long min_interval; /* Minimum balance interval ms */
|
|
|
|
|
unsigned long max_interval; /* Maximum balance interval ms */
|
|
|
|
|
unsigned int busy_factor; /* less balancing by factor if busy */
|
|
|
|
|
unsigned int imbalance_pct; /* No balance until over watermark */
|
|
|
|
|
unsigned int cache_nice_tries; /* Leave cache hot tasks for # tries */
|
2005-06-26 05:57:13 +08:00
|
|
|
unsigned int busy_idx;
|
|
|
|
|
unsigned int idle_idx;
|
|
|
|
|
unsigned int newidle_idx;
|
|
|
|
|
unsigned int wake_idx;
|
2005-06-26 05:57:19 +08:00
|
|
|
unsigned int forkexec_idx;
|
2009-09-01 16:34:35 +08:00
|
|
|
unsigned int smt_gain;
|
2013-04-23 22:59:02 +08:00
|
|
|
|
|
|
|
|
int nohz_idle; /* NOHZ IDLE status */
|
2005-04-17 06:20:36 +08:00
|
|
|
int flags; /* See SD_* */
|
2011-04-07 20:10:04 +08:00
|
|
|
int level;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* Runtime fields. */
|
|
|
|
|
unsigned long last_balance; /* init to jiffies. units in jiffies */
|
|
|
|
|
unsigned int balance_interval; /* initialise to 1. units in ms. */
|
|
|
|
|
unsigned int nr_balance_failed; /* initialise to 0 */
|
|
|
|
|
|
2008-06-27 19:41:35 +08:00
|
|
|
u64 last_update;
|
2013-09-14 02:26:53 +08:00
|
|
|
|
|
|
|
|
/* idle_balance() stats */
|
2013-09-14 02:26:52 +08:00
|
|
|
u64 max_newidle_lb_cost;
|
2013-09-14 02:26:53 +08:00
|
|
|
unsigned long next_decay_max_lb_cost;
|
2008-06-27 19:41:35 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
#ifdef CONFIG_SCHEDSTATS
|
|
|
|
|
/* load_balance() stats */
|
2007-10-19 03:32:56 +08:00
|
|
|
unsigned int lb_count[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_failed[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_balanced[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_imbalance[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_gained[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_hot_gained[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_nobusyg[CPU_MAX_IDLE_TYPES];
|
|
|
|
|
unsigned int lb_nobusyq[CPU_MAX_IDLE_TYPES];
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* Active load balancing */
|
2007-10-19 03:32:56 +08:00
|
|
|
unsigned int alb_count;
|
|
|
|
|
unsigned int alb_failed;
|
|
|
|
|
unsigned int alb_pushed;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2005-06-26 05:57:20 +08:00
|
|
|
/* SD_BALANCE_EXEC stats */
|
2007-10-19 03:32:56 +08:00
|
|
|
unsigned int sbe_count;
|
|
|
|
|
unsigned int sbe_balanced;
|
|
|
|
|
unsigned int sbe_pushed;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2005-06-26 05:57:20 +08:00
|
|
|
/* SD_BALANCE_FORK stats */
|
2007-10-19 03:32:56 +08:00
|
|
|
unsigned int sbf_count;
|
|
|
|
|
unsigned int sbf_balanced;
|
|
|
|
|
unsigned int sbf_pushed;
|
2005-06-26 05:57:20 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* try_to_wake_up() stats */
|
2007-10-19 03:32:56 +08:00
|
|
|
unsigned int ttwu_wake_remote;
|
|
|
|
|
unsigned int ttwu_move_affine;
|
|
|
|
|
unsigned int ttwu_move_balance;
|
2005-04-17 06:20:36 +08:00
|
|
|
#endif
|
2008-10-09 17:35:51 +08:00
|
|
|
#ifdef CONFIG_SCHED_DEBUG
|
|
|
|
|
char *name;
|
|
|
|
|
#endif
|
2011-04-07 20:09:50 +08:00
|
|
|
union {
|
|
|
|
|
void *private; /* used during construction */
|
|
|
|
|
struct rcu_head rcu; /* used during destruction */
|
|
|
|
|
};
|
2008-11-25 00:05:04 +08:00
|
|
|
|
2010-04-16 20:59:29 +08:00
|
|
|
unsigned int span_weight;
|
2009-05-19 15:22:19 +08:00
|
|
|
/*
|
|
|
|
|
* Span of all CPUs in this domain.
|
|
|
|
|
*
|
|
|
|
|
* NOTE: this field is variable length. (Allocated dynamically
|
|
|
|
|
* by attaching extra space to the end of the structure,
|
|
|
|
|
* depending on how many CPUs the kernel has booted up with)
|
|
|
|
|
*/
|
|
|
|
|
unsigned long span[0];
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
|
|
|
|
|
2008-11-25 00:05:04 +08:00
|
|
|
static inline struct cpumask *sched_domain_span(struct sched_domain *sd)
|
|
|
|
|
{
|
2008-11-25 00:05:04 +08:00
|
|
|
return to_cpumask(sd->span);
|
2008-11-25 00:05:04 +08:00
|
|
|
}
|
|
|
|
|
|
2009-11-03 12:23:40 +08:00
|
|
|
extern void partition_sched_domains(int ndoms_new, cpumask_var_t doms_new[],
|
2008-04-15 13:04:23 +08:00
|
|
|
struct sched_domain_attr *dattr_new);
|
2007-10-19 14:40:20 +08:00
|
|
|
|
2009-11-03 12:23:40 +08:00
|
|
|
/* Allocate an array of sched domains, for partition_sched_domains(). */
|
|
|
|
|
cpumask_var_t *alloc_sched_domains(unsigned int ndoms);
|
|
|
|
|
void free_sched_domains(cpumask_var_t doms[], unsigned int ndoms);
|
|
|
|
|
|
2012-01-26 19:44:34 +08:00
|
|
|
bool cpus_share_cache(int this_cpu, int that_cpu);
|
|
|
|
|
|
2008-07-18 20:01:39 +08:00
|
|
|
#else /* CONFIG_SMP */
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2008-07-18 20:01:39 +08:00
|
|
|
struct sched_domain_attr;
|
2007-07-26 19:40:43 +08:00
|
|
|
|
2008-07-18 20:01:39 +08:00
|
|
|
static inline void
|
2009-11-03 12:23:40 +08:00
|
|
|
partition_sched_domains(int ndoms_new, cpumask_var_t doms_new[],
|
2008-07-18 20:01:39 +08:00
|
|
|
struct sched_domain_attr *dattr_new)
|
|
|
|
|
{
|
2007-07-26 19:40:43 +08:00
|
|
|
}
|
2012-01-26 19:44:34 +08:00
|
|
|
|
|
|
|
|
static inline bool cpus_share_cache(int this_cpu, int that_cpu)
|
|
|
|
|
{
|
|
|
|
|
return true;
|
|
|
|
|
}
|
|
|
|
|
|
2008-07-18 20:01:39 +08:00
|
|
|
#endif /* !CONFIG_SMP */
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2009-09-02 19:49:18 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
struct io_context; /* See blkdev.h */
|
|
|
|
|
|
|
|
|
|
|
2005-09-10 04:02:02 +08:00
|
|
|
#ifdef ARCH_HAS_PREFETCH_SWITCH_STACK
|
2006-07-03 15:25:41 +08:00
|
|
|
extern void prefetch_stack(struct task_struct *t);
|
2005-09-10 04:02:02 +08:00
|
|
|
#else
|
|
|
|
|
static inline void prefetch_stack(struct task_struct *t) { }
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
struct audit_context; /* See audit.c */
|
|
|
|
|
struct mempolicy;
|
2006-04-11 19:52:07 +08:00
|
|
|
struct pipe_inode_info;
|
2006-10-02 17:18:14 +08:00
|
|
|
struct uts_namespace;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2007-07-10 00:51:58 +08:00
|
|
|
struct load_weight {
|
|
|
|
|
unsigned long weight, inv_weight;
|
|
|
|
|
};
|
|
|
|
|
|
2012-10-04 19:18:29 +08:00
|
|
|
struct sched_avg {
|
|
|
|
|
/*
|
|
|
|
|
* These sums represent an infinite geometric series and so are bound
|
2013-06-27 14:04:09 +08:00
|
|
|
* above by 1024/(1-y). Thus we only need a u32 to store them for all
|
2012-10-04 19:18:29 +08:00
|
|
|
* choices of y < 1-2^(-32)*1024.
|
|
|
|
|
*/
|
|
|
|
|
u32 runnable_avg_sum, runnable_avg_period;
|
|
|
|
|
u64 last_runnable_update;
|
2012-10-04 19:18:30 +08:00
|
|
|
s64 decay_count;
|
2012-10-04 19:18:30 +08:00
|
|
|
unsigned long load_avg_contrib;
|
2012-10-04 19:18:29 +08:00
|
|
|
};
|
|
|
|
|
|
2007-08-02 23:41:40 +08:00
|
|
|
#ifdef CONFIG_SCHEDSTATS
|
2010-03-11 10:37:45 +08:00
|
|
|
struct sched_statistics {
|
2007-07-10 00:51:58 +08:00
|
|
|
u64 wait_start;
|
2007-08-02 23:41:40 +08:00
|
|
|
u64 wait_max;
|
2008-01-26 04:08:35 +08:00
|
|
|
u64 wait_count;
|
|
|
|
|
u64 wait_sum;
|
2009-07-21 02:26:58 +08:00
|
|
|
u64 iowait_count;
|
|
|
|
|
u64 iowait_sum;
|
2007-08-02 23:41:40 +08:00
|
|
|
|
2007-07-10 00:51:58 +08:00
|
|
|
u64 sleep_start;
|
|
|
|
|
u64 sleep_max;
|
2007-08-02 23:41:40 +08:00
|
|
|
s64 sum_sleep_runtime;
|
|
|
|
|
|
|
|
|
|
u64 block_start;
|
2007-07-10 00:51:58 +08:00
|
|
|
u64 block_max;
|
|
|
|
|
u64 exec_max;
|
2007-10-15 23:00:02 +08:00
|
|
|
u64 slice_max;
|
2007-10-15 23:00:18 +08:00
|
|
|
|
|
|
|
|
u64 nr_migrations_cold;
|
|
|
|
|
u64 nr_failed_migrations_affine;
|
|
|
|
|
u64 nr_failed_migrations_running;
|
|
|
|
|
u64 nr_failed_migrations_hot;
|
|
|
|
|
u64 nr_forced_migrations;
|
|
|
|
|
|
|
|
|
|
u64 nr_wakeups;
|
|
|
|
|
u64 nr_wakeups_sync;
|
|
|
|
|
u64 nr_wakeups_migrate;
|
|
|
|
|
u64 nr_wakeups_local;
|
|
|
|
|
u64 nr_wakeups_remote;
|
|
|
|
|
u64 nr_wakeups_affine;
|
|
|
|
|
u64 nr_wakeups_affine_attempts;
|
|
|
|
|
u64 nr_wakeups_passive;
|
|
|
|
|
u64 nr_wakeups_idle;
|
2010-03-11 10:37:45 +08:00
|
|
|
};
|
|
|
|
|
#endif
|
|
|
|
|
|
|
|
|
|
struct sched_entity {
|
|
|
|
|
struct load_weight load; /* for load-balancing */
|
|
|
|
|
struct rb_node run_node;
|
|
|
|
|
struct list_head group_node;
|
|
|
|
|
unsigned int on_rq;
|
|
|
|
|
|
|
|
|
|
u64 exec_start;
|
|
|
|
|
u64 sum_exec_runtime;
|
|
|
|
|
u64 vruntime;
|
|
|
|
|
u64 prev_sum_exec_runtime;
|
|
|
|
|
|
|
|
|
|
u64 nr_migrations;
|
|
|
|
|
|
|
|
|
|
#ifdef CONFIG_SCHEDSTATS
|
|
|
|
|
struct sched_statistics statistics;
|
2007-08-02 23:41:40 +08:00
|
|
|
#endif
|
|
|
|
|
|
2007-07-10 00:51:58 +08:00
|
|
|
#ifdef CONFIG_FAIR_GROUP_SCHED
|
|
|
|
|
struct sched_entity *parent;
|
|
|
|
|
/* rq on which this entity is (to be) queued: */
|
|
|
|
|
struct cfs_rq *cfs_rq;
|
|
|
|
|
/* rq "owned" by this entity/group: */
|
|
|
|
|
struct cfs_rq *my_q;
|
|
|
|
|
#endif
|
2013-02-07 23:47:07 +08:00
|
|
|
|
2013-06-26 13:05:39 +08:00
|
|
|
#ifdef CONFIG_SMP
|
2012-10-04 19:18:32 +08:00
|
|
|
/* Per-entity load-tracking */
|
2012-10-04 19:18:29 +08:00
|
|
|
struct sched_avg avg;
|
|
|
|
|
#endif
|
2007-07-10 00:51:58 +08:00
|
|
|
};
|
2006-07-03 15:25:42 +08:00
|
|
|
|
2008-01-26 04:08:27 +08:00
|
|
|
struct sched_rt_entity {
|
|
|
|
|
struct list_head run_list;
|
2008-01-26 04:08:27 +08:00
|
|
|
unsigned long timeout;
|
sched/rt: Avoid updating RT entry timeout twice within one tick period
The issue below was found in 2.6.34-rt rather than mainline rt
kernel, but the issue still exists upstream as well.
So please let me describe how it was noticed on 2.6.34-rt:
On this version, each softirq has its own thread, it means there
is at least one RT FIFO task per cpu. The priority of these
tasks is set to 49 by default. If user launches an RT FIFO task
with priority lower than 49 of softirq RT tasks, it's possible
there are two RT FIFO tasks enqueued one cpu runqueue at one
moment. By current strategy of balancing RT tasks, when it comes
to RT tasks, we really need to put them off to a CPU that they
can run on as soon as possible. Even if it means a bit of cache
line flushing, we want RT tasks to be run with the least latency.
When the user RT FIFO task which just launched before is
running, the sched timer tick of the current cpu happens. In this
tick period, the timeout value of the user RT task will be
updated once. Subsequently, we try to wake up one softirq RT
task on its local cpu. As the priority of current user RT task
is lower than the softirq RT task, the current task will be
preempted by the higher priority softirq RT task. Before
preemption, we check to see if current can readily move to a
different cpu. If so, we will reschedule to allow the RT push logic
to try to move current somewhere else. Whenever the woken
softirq RT task runs, it first tries to migrate the user FIFO RT
task over to a cpu that is running a task of lesser priority. If
migration is done, it will send a reschedule request to the found
cpu by IPI interrupt. Once the target cpu responds the IPI
interrupt, it will pick the migrated user RT task to preempt its
current task. When the user RT task is running on the new cpu,
the sched timer tick of the cpu fires. So it will tick the user
RT task again. This also means the RT task timeout value will be
updated again. As the migration may be done in one tick period,
it means the user RT task timeout value will be updated twice
within one tick.
If we set a limit on the amount of cpu time for the user RT task
by setrlimit(RLIMIT_RTTIME), the SIGXCPU signal should be posted
upon reaching the soft limit.
But exactly when the SIGXCPU signal should be sent depends on the
RT task timeout value. In fact the timeout mechanism of sending
the SIGXCPU signal assumes the RT task timeout is increased once
every tick.
However, currently the timeout value may be added twice per
tick. So it results in the SIGXCPU signal being sent earlier
than expected.
To solve this issue, we prevent the timeout value from increasing
twice within one tick time by remembering the jiffies value of
last updating the timeout. As long as the RT task's jiffies is
different with the global jiffies value, we allow its timeout to
be updated.
Signed-off-by: Ying Xue <ying.xue@windriver.com>
Signed-off-by: Fan Du <fan.du@windriver.com>
Reviewed-by: Yong Zhang <yong.zhang0@gmail.com>
Acked-by: Steven Rostedt <rostedt@goodmis.org>
Cc: <peterz@infradead.org>
Link: http://lkml.kernel.org/r/1342508623-2887-1-git-send-email-ying.xue@windriver.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
2012-07-17 15:03:43 +08:00
|
|
|
unsigned long watchdog_stamp;
|
2008-08-01 20:24:08 +08:00
|
|
|
unsigned int time_slice;
|
2008-01-26 04:08:30 +08:00
|
|
|
|
2008-04-20 01:45:00 +08:00
|
|
|
struct sched_rt_entity *back;
|
2008-02-13 22:45:40 +08:00
|
|
|
#ifdef CONFIG_RT_GROUP_SCHED
|
2008-01-26 04:08:30 +08:00
|
|
|
struct sched_rt_entity *parent;
|
|
|
|
|
/* rq on which this entity is (to be) queued: */
|
|
|
|
|
struct rt_rq *rt_rq;
|
|
|
|
|
/* rq "owned" by this entity/group: */
|
|
|
|
|
struct rt_rq *my_q;
|
|
|
|
|
#endif
|
2008-01-26 04:08:27 +08:00
|
|
|
};
|
|
|
|
|
|
2013-02-07 23:47:07 +08:00
|
|
|
|
2009-08-28 06:00:12 +08:00
|
|
|
struct rcu_node;
|
|
|
|
|
|
2010-09-02 22:50:03 +08:00
|
|
|
enum perf_event_task_context {
|
|
|
|
|
perf_invalid_context = -1,
|
|
|
|
|
perf_hw_context = 0,
|
2010-09-07 23:34:50 +08:00
|
|
|
perf_sw_context,
|
2010-09-02 22:50:03 +08:00
|
|
|
perf_nr_task_contexts,
|
|
|
|
|
};
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
struct task_struct {
|
|
|
|
|
volatile long state; /* -1 unrunnable, 0 runnable, >0 stopped */
|
2007-05-09 17:35:17 +08:00
|
|
|
void *stack;
|
2005-04-17 06:20:36 +08:00
|
|
|
atomic_t usage;
|
2007-05-08 15:23:41 +08:00
|
|
|
unsigned int flags; /* per process flags, defined below */
|
|
|
|
|
unsigned int ptrace;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
[PATCH] sched: implement smpnice
Problem:
The introduction of separate run queues per CPU has brought with it "nice"
enforcement problems that are best described by a simple example.
For the sake of argument suppose that on a single CPU machine with a
nice==19 hard spinner and a nice==0 hard spinner running that the nice==0
task gets 95% of the CPU and the nice==19 task gets 5% of the CPU. Now
suppose that there is a system with 2 CPUs and 2 nice==19 hard spinners and
2 nice==0 hard spinners running. The user of this system would be entitled
to expect that the nice==0 tasks each get 95% of a CPU and the nice==19
tasks only get 5% each. However, whether this expectation is met is pretty
much down to luck as there are four equally likely distributions of the
tasks to the CPUs that the load balancing code will consider to be balanced
with loads of 2.0 for each CPU. Two of these distributions involve one
nice==0 and one nice==19 task per CPU and in these circumstances the users
expectations will be met. The other two distributions both involve both
nice==0 tasks being on one CPU and both nice==19 being on the other CPU and
each task will get 50% of a CPU and the user's expectations will not be
met.
Solution:
The solution to this problem that is implemented in the attached patch is
to use weighted loads when determining if the system is balanced and, when
an imbalance is detected, to move an amount of weighted load between run
queues (as opposed to a number of tasks) to restore the balance. Once
again, the easiest way to explain why both of these measures are necessary
is to use a simple example. Suppose that (in a slight variation of the
above example) that we have a two CPU system with 4 nice==0 and 4 nice=19
hard spinning tasks running and that the 4 nice==0 tasks are on one CPU and
the 4 nice==19 tasks are on the other CPU. The weighted loads for the two
CPUs would be 4.0 and 0.2 respectively and the load balancing code would
move 2 tasks resulting in one CPU with a load of 2.0 and the other with
load of 2.2. If this was considered to be a big enough imbalance to
justify moving a task and that task was moved using the current
move_tasks() then it would move the highest priority task that it found and
this would result in one CPU with a load of 3.0 and the other with a load
of 1.2 which would result in the movement of a task in the opposite
direction and so on -- infinite loop. If, on the other hand, an amount of
load to be moved is calculated from the imbalance (in this case 0.1) and
move_tasks() skips tasks until it find ones whose contributions to the
weighted load are less than this amount it would move two of the nice==19
tasks resulting in a system with 2 nice==0 and 2 nice=19 on each CPU with
loads of 2.1 for each CPU.
One of the advantages of this mechanism is that on a system where all tasks
have nice==0 the load balancing calculations would be mathematically
identical to the current load balancing code.
Notes:
struct task_struct:
has a new field load_weight which (in a trade off of space for speed)
stores the contribution that this task makes to a CPU's weighted load when
it is runnable.
struct runqueue:
has a new field raw_weighted_load which is the sum of the load_weight
values for the currently runnable tasks on this run queue. This field
always needs to be updated when nr_running is updated so two new inline
functions inc_nr_running() and dec_nr_running() have been created to make
sure that this happens. This also offers a convenient way to optimize away
this part of the smpnice mechanism when CONFIG_SMP is not defined.
int try_to_wake_up():
in this function the value SCHED_LOAD_BALANCE is used to represent the load
contribution of a single task in various calculations in the code that
decides which CPU to put the waking task on. While this would be a valid
on a system where the nice values for the runnable tasks were distributed
evenly around zero it will lead to anomalous load balancing if the
distribution is skewed in either direction. To overcome this problem
SCHED_LOAD_SCALE has been replaced by the load_weight for the relevant task
or by the average load_weight per task for the queue in question (as
appropriate).
int move_tasks():
The modifications to this function were complicated by the fact that
active_load_balance() uses it to move exactly one task without checking
whether an imbalance actually exists. This precluded the simple
overloading of max_nr_move with max_load_move and necessitated the addition
of the latter as an extra argument to the function. The internal
implementation is then modified to move up to max_nr_move tasks and
max_load_move of weighted load. This slightly complicates the code where
move_tasks() is called and if ever active_load_balance() is changed to not
use move_tasks() the implementation of move_tasks() should be simplified
accordingly.
struct sched_group *find_busiest_group():
Similar to try_to_wake_up(), there are places in this function where
SCHED_LOAD_SCALE is used to represent the load contribution of a single
task and the same issues are created. A similar solution is adopted except
that it is now the average per task contribution to a group's load (as
opposed to a run queue) that is required. As this value is not directly
available from the group it is calculated on the fly as the queues in the
groups are visited when determining the busiest group.
A key change to this function is that it is no longer to scale down
*imbalance on exit as move_tasks() uses the load in its scaled form.
void set_user_nice():
has been modified to update the task's load_weight field when it's nice
value and also to ensure that its run queue's raw_weighted_load field is
updated if it was runnable.
From: "Siddha, Suresh B" <suresh.b.siddha@intel.com>
With smpnice, sched groups with highest priority tasks can mask the imbalance
between the other sched groups with in the same domain. This patch fixes some
of the listed down scenarios by not considering the sched groups which are
lightly loaded.
a) on a simple 4-way MP system, if we have one high priority and 4 normal
priority tasks, with smpnice we would like to see the high priority task
scheduled on one cpu, two other cpus getting one normal task each and the
fourth cpu getting the remaining two normal tasks. but with current
smpnice extra normal priority task keeps jumping from one cpu to another
cpu having the normal priority task. This is because of the
busiest_has_loaded_cpus, nr_loaded_cpus logic.. We are not including the
cpu with high priority task in max_load calculations but including that in
total and avg_load calcuations.. leading to max_load < avg_load and load
balance between cpus running normal priority tasks(2 Vs 1) will always show
imbalanace as one normal priority and the extra normal priority task will
keep moving from one cpu to another cpu having normal priority task..
b) 4-way system with HT (8 logical processors). Package-P0 T0 has a
highest priority task, T1 is idle. Package-P1 Both T0 and T1 have 1 normal
priority task each.. P2 and P3 are idle. With this patch, one of the
normal priority tasks on P1 will be moved to P2 or P3..
c) With the current weighted smp nice calculations, it doesn't always make
sense to look at the highest weighted runqueue in the busy group..
Consider a load balance scenario on a DP with HT system, with Package-0
containing one high priority and one low priority, Package-1 containing one
low priority(with other thread being idle).. Package-1 thinks that it need
to take the low priority thread from Package-0. And find_busiest_queue()
returns the cpu thread with highest priority task.. And ultimately(with
help of active load balance) we move high priority task to Package-1. And
same continues with Package-0 now, moving high priority task from package-1
to package-0.. Even without the presence of active load balance, load
balance will fail to balance the above scenario.. Fix find_busiest_queue
to use "imbalance" when it is lightly loaded.
[kernel@kolivas.org: sched: store weighted load on up]
[kernel@kolivas.org: sched: add discrete weighted cpu load function]
[suresh.b.siddha@intel.com: sched: remove dead code]
Signed-off-by: Peter Williams <pwil3058@bigpond.com.au>
Cc: "Siddha, Suresh B" <suresh.b.siddha@intel.com>
Cc: "Chen, Kenneth W" <kenneth.w.chen@intel.com>
Acked-by: Ingo Molnar <mingo@elte.hu>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Signed-off-by: Con Kolivas <kernel@kolivas.org>
Cc: John Hawkes <hawkes@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-06-27 17:54:34 +08:00
|
|
|
#ifdef CONFIG_SMP
|
2011-09-12 19:06:17 +08:00
|
|
|
struct llist_node wake_entry;
|
2011-04-05 23:23:40 +08:00
|
|
|
int on_cpu;
|
sched: Implement smarter wake-affine logic
The wake-affine scheduler feature is currently always trying to pull
the wakee close to the waker. In theory this should be beneficial if
the waker's CPU caches hot data for the wakee, and it's also beneficial
in the extreme ping-pong high context switch rate case.
Testing shows it can benefit hackbench up to 15%.
However, the feature is somewhat blind, from which some workloads
such as pgbench suffer. It's also time-consuming algorithmically.
Testing shows it can damage pgbench up to 50% - far more than the
benefit it brings in the best case.
So wake-affine should be smarter and it should realize when to
stop its thankless effort at trying to find a suitable CPU to wake on.
This patch introduces 'wakee_flips', which will be increased each
time the task flips (switches) its wakee target.
So a high 'wakee_flips' value means the task has more than one
wakee, and the bigger the number, the higher the wakeup frequency.
Now when making the decision on whether to pull or not, pay attention to
the wakee with a high 'wakee_flips', pulling such a task may benefit
the wakee. Also imply that the waker will face cruel competition later,
it could be very cruel or very fast depends on the story behind
'wakee_flips', waker therefore suffers.
Furthermore, if waker also has a high 'wakee_flips', that implies that
multiple tasks rely on it, then waker's higher latency will damage all
of them, so pulling wakee seems to be a bad deal.
Thus, when 'waker->wakee_flips / wakee->wakee_flips' becomes
higher and higher, the cost of pulling seems to be worse and worse.
The patch therefore helps the wake-affine feature to stop its pulling
work when:
wakee->wakee_flips > factor &&
waker->wakee_flips > (factor * wakee->wakee_flips)
The 'factor' here is the number of CPUs in the current CPU's NUMA node,
so a bigger node will lead to more pulling since the trial becomes more
severe.
After applying the patch, pgbench shows up to 40% improvements and no regressions.
Tested with 12 cpu x86 server and tip 3.10.0-rc7.
The percentages in the final column highlight the areas with the biggest wins,
all other areas improved as well:
pgbench base smart
| db_size | clients | tps | | tps |
+---------+---------+-------+ +-------+
| 22 MB | 1 | 10598 | | 10796 |
| 22 MB | 2 | 21257 | | 21336 |
| 22 MB | 4 | 41386 | | 41622 |
| 22 MB | 8 | 51253 | | 57932 |
| 22 MB | 12 | 48570 | | 54000 |
| 22 MB | 16 | 46748 | | 55982 | +19.75%
| 22 MB | 24 | 44346 | | 55847 | +25.93%
| 22 MB | 32 | 43460 | | 54614 | +25.66%
| 7484 MB | 1 | 8951 | | 9193 |
| 7484 MB | 2 | 19233 | | 19240 |
| 7484 MB | 4 | 37239 | | 37302 |
| 7484 MB | 8 | 46087 | | 50018 |
| 7484 MB | 12 | 42054 | | 48763 |
| 7484 MB | 16 | 40765 | | 51633 | +26.66%
| 7484 MB | 24 | 37651 | | 52377 | +39.11%
| 7484 MB | 32 | 37056 | | 51108 | +37.92%
| 15 GB | 1 | 8845 | | 9104 |
| 15 GB | 2 | 19094 | | 19162 |
| 15 GB | 4 | 36979 | | 36983 |
| 15 GB | 8 | 46087 | | 49977 |
| 15 GB | 12 | 41901 | | 48591 |
| 15 GB | 16 | 40147 | | 50651 | +26.16%
| 15 GB | 24 | 37250 | | 52365 | +40.58%
| 15 GB | 32 | 36470 | | 50015 | +37.14%
Signed-off-by: Michael Wang <wangyun@linux.vnet.ibm.com>
Cc: Mike Galbraith <efault@gmx.de>
Signed-off-by: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/51D50057.9000809@linux.vnet.ibm.com
[ Improved the changelog. ]
Signed-off-by: Ingo Molnar <mingo@kernel.org>
2013-07-04 12:55:51 +08:00
|
|
|
struct task_struct *last_wakee;
|
|
|
|
|
unsigned long wakee_flips;
|
|
|
|
|
unsigned long wakee_flip_decay_ts;
|
[PATCH] sched: implement smpnice
Problem:
The introduction of separate run queues per CPU has brought with it "nice"
enforcement problems that are best described by a simple example.
For the sake of argument suppose that on a single CPU machine with a
nice==19 hard spinner and a nice==0 hard spinner running that the nice==0
task gets 95% of the CPU and the nice==19 task gets 5% of the CPU. Now
suppose that there is a system with 2 CPUs and 2 nice==19 hard spinners and
2 nice==0 hard spinners running. The user of this system would be entitled
to expect that the nice==0 tasks each get 95% of a CPU and the nice==19
tasks only get 5% each. However, whether this expectation is met is pretty
much down to luck as there are four equally likely distributions of the
tasks to the CPUs that the load balancing code will consider to be balanced
with loads of 2.0 for each CPU. Two of these distributions involve one
nice==0 and one nice==19 task per CPU and in these circumstances the users
expectations will be met. The other two distributions both involve both
nice==0 tasks being on one CPU and both nice==19 being on the other CPU and
each task will get 50% of a CPU and the user's expectations will not be
met.
Solution:
The solution to this problem that is implemented in the attached patch is
to use weighted loads when determining if the system is balanced and, when
an imbalance is detected, to move an amount of weighted load between run
queues (as opposed to a number of tasks) to restore the balance. Once
again, the easiest way to explain why both of these measures are necessary
is to use a simple example. Suppose that (in a slight variation of the
above example) that we have a two CPU system with 4 nice==0 and 4 nice=19
hard spinning tasks running and that the 4 nice==0 tasks are on one CPU and
the 4 nice==19 tasks are on the other CPU. The weighted loads for the two
CPUs would be 4.0 and 0.2 respectively and the load balancing code would
move 2 tasks resulting in one CPU with a load of 2.0 and the other with
load of 2.2. If this was considered to be a big enough imbalance to
justify moving a task and that task was moved using the current
move_tasks() then it would move the highest priority task that it found and
this would result in one CPU with a load of 3.0 and the other with a load
of 1.2 which would result in the movement of a task in the opposite
direction and so on -- infinite loop. If, on the other hand, an amount of
load to be moved is calculated from the imbalance (in this case 0.1) and
move_tasks() skips tasks until it find ones whose contributions to the
weighted load are less than this amount it would move two of the nice==19
tasks resulting in a system with 2 nice==0 and 2 nice=19 on each CPU with
loads of 2.1 for each CPU.
One of the advantages of this mechanism is that on a system where all tasks
have nice==0 the load balancing calculations would be mathematically
identical to the current load balancing code.
Notes:
struct task_struct:
has a new field load_weight which (in a trade off of space for speed)
stores the contribution that this task makes to a CPU's weighted load when
it is runnable.
struct runqueue:
has a new field raw_weighted_load which is the sum of the load_weight
values for the currently runnable tasks on this run queue. This field
always needs to be updated when nr_running is updated so two new inline
functions inc_nr_running() and dec_nr_running() have been created to make
sure that this happens. This also offers a convenient way to optimize away
this part of the smpnice mechanism when CONFIG_SMP is not defined.
int try_to_wake_up():
in this function the value SCHED_LOAD_BALANCE is used to represent the load
contribution of a single task in various calculations in the code that
decides which CPU to put the waking task on. While this would be a valid
on a system where the nice values for the runnable tasks were distributed
evenly around zero it will lead to anomalous load balancing if the
distribution is skewed in either direction. To overcome this problem
SCHED_LOAD_SCALE has been replaced by the load_weight for the relevant task
or by the average load_weight per task for the queue in question (as
appropriate).
int move_tasks():
The modifications to this function were complicated by the fact that
active_load_balance() uses it to move exactly one task without checking
whether an imbalance actually exists. This precluded the simple
overloading of max_nr_move with max_load_move and necessitated the addition
of the latter as an extra argument to the function. The internal
implementation is then modified to move up to max_nr_move tasks and
max_load_move of weighted load. This slightly complicates the code where
move_tasks() is called and if ever active_load_balance() is changed to not
use move_tasks() the implementation of move_tasks() should be simplified
accordingly.
struct sched_group *find_busiest_group():
Similar to try_to_wake_up(), there are places in this function where
SCHED_LOAD_SCALE is used to represent the load contribution of a single
task and the same issues are created. A similar solution is adopted except
that it is now the average per task contribution to a group's load (as
opposed to a run queue) that is required. As this value is not directly
available from the group it is calculated on the fly as the queues in the
groups are visited when determining the busiest group.
A key change to this function is that it is no longer to scale down
*imbalance on exit as move_tasks() uses the load in its scaled form.
void set_user_nice():
has been modified to update the task's load_weight field when it's nice
value and also to ensure that its run queue's raw_weighted_load field is
updated if it was runnable.
From: "Siddha, Suresh B" <suresh.b.siddha@intel.com>
With smpnice, sched groups with highest priority tasks can mask the imbalance
between the other sched groups with in the same domain. This patch fixes some
of the listed down scenarios by not considering the sched groups which are
lightly loaded.
a) on a simple 4-way MP system, if we have one high priority and 4 normal
priority tasks, with smpnice we would like to see the high priority task
scheduled on one cpu, two other cpus getting one normal task each and the
fourth cpu getting the remaining two normal tasks. but with current
smpnice extra normal priority task keeps jumping from one cpu to another
cpu having the normal priority task. This is because of the
busiest_has_loaded_cpus, nr_loaded_cpus logic.. We are not including the
cpu with high priority task in max_load calculations but including that in
total and avg_load calcuations.. leading to max_load < avg_load and load
balance between cpus running normal priority tasks(2 Vs 1) will always show
imbalanace as one normal priority and the extra normal priority task will
keep moving from one cpu to another cpu having normal priority task..
b) 4-way system with HT (8 logical processors). Package-P0 T0 has a
highest priority task, T1 is idle. Package-P1 Both T0 and T1 have 1 normal
priority task each.. P2 and P3 are idle. With this patch, one of the
normal priority tasks on P1 will be moved to P2 or P3..
c) With the current weighted smp nice calculations, it doesn't always make
sense to look at the highest weighted runqueue in the busy group..
Consider a load balance scenario on a DP with HT system, with Package-0
containing one high priority and one low priority, Package-1 containing one
low priority(with other thread being idle).. Package-1 thinks that it need
to take the low priority thread from Package-0. And find_busiest_queue()
returns the cpu thread with highest priority task.. And ultimately(with
help of active load balance) we move high priority task to Package-1. And
same continues with Package-0 now, moving high priority task from package-1
to package-0.. Even without the presence of active load balance, load
balance will fail to balance the above scenario.. Fix find_busiest_queue
to use "imbalance" when it is lightly loaded.
[kernel@kolivas.org: sched: store weighted load on up]
[kernel@kolivas.org: sched: add discrete weighted cpu load function]
[suresh.b.siddha@intel.com: sched: remove dead code]
Signed-off-by: Peter Williams <pwil3058@bigpond.com.au>
Cc: "Siddha, Suresh B" <suresh.b.siddha@intel.com>
Cc: "Chen, Kenneth W" <kenneth.w.chen@intel.com>
Acked-by: Ingo Molnar <mingo@elte.hu>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Signed-off-by: Con Kolivas <kernel@kolivas.org>
Cc: John Hawkes <hawkes@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-06-27 17:54:34 +08:00
|
|
|
#endif
|
2011-04-05 23:23:44 +08:00
|
|
|
int on_rq;
|
2007-07-10 00:52:00 +08:00
|
|
|
|
2006-06-27 17:54:51 +08:00
|
|
|
int prio, static_prio, normal_prio;
|
2008-05-15 19:09:15 +08:00
|
|
|
unsigned int rt_priority;
|
2007-10-15 23:00:12 +08:00
|
|
|
const struct sched_class *sched_class;
|
2007-07-10 00:51:58 +08:00
|
|
|
struct sched_entity se;
|
2008-01-26 04:08:27 +08:00
|
|
|
struct sched_rt_entity rt;
|
2012-06-22 19:36:05 +08:00
|
|
|
#ifdef CONFIG_CGROUP_SCHED
|
|
|
|
|
struct task_group *sched_task_group;
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2007-07-26 19:40:43 +08:00
|
|
|
#ifdef CONFIG_PREEMPT_NOTIFIERS
|
|
|
|
|
/* list of struct preempt_notifier: */
|
|
|
|
|
struct hlist_head preempt_notifiers;
|
|
|
|
|
#endif
|
|
|
|
|
|
2007-10-17 14:30:26 +08:00
|
|
|
/*
|
|
|
|
|
* fpu_counter contains the number of consecutive context switches
|
|
|
|
|
* that the FPU is used. If this is over a threshold, the lazy fpu
|
|
|
|
|
* saving becomes unlazy to save the trap. This is an unsigned char
|
|
|
|
|
* so that after 256 times the counter wraps and the behavior turns
|
|
|
|
|
* lazy again; this to deal with bursty apps that only use FPU for
|
|
|
|
|
* a short time
|
|
|
|
|
*/
|
|
|
|
|
unsigned char fpu_counter;
|
2006-09-29 16:59:40 +08:00
|
|
|
#ifdef CONFIG_BLK_DEV_IO_TRACE
|
2006-03-24 03:00:26 +08:00
|
|
|
unsigned int btrace_seq;
|
2006-09-29 16:59:40 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2007-05-08 15:23:41 +08:00
|
|
|
unsigned int policy;
|
2012-04-23 18:11:21 +08:00
|
|
|
int nr_cpus_allowed;
|
2005-04-17 06:20:36 +08:00
|
|
|
cpumask_t cpus_allowed;
|
|
|
|
|
|
2010-06-30 07:49:16 +08:00
|
|
|
#ifdef CONFIG_PREEMPT_RCU
|
2008-01-26 04:08:24 +08:00
|
|
|
int rcu_read_lock_nesting;
|
rcu: Merge preemptable-RCU functionality into hierarchical RCU
Create a kernel/rcutree_plugin.h file that contains definitions
for preemptable RCU (or, under the #else branch of the #ifdef,
empty definitions for the classic non-preemptable semantics).
These definitions fit into plugins defined in kernel/rcutree.c
for this purpose.
This variant of preemptable RCU uses a new algorithm whose
read-side expense is roughly that of classic hierarchical RCU
under CONFIG_PREEMPT. This new algorithm's update-side expense
is similar to that of classic hierarchical RCU, and, in absence
of read-side preemption or blocking, is exactly that of classic
hierarchical RCU. Perhaps more important, this new algorithm
has a much simpler implementation, saving well over 1,000 lines
of code compared to mainline's implementation of preemptable
RCU, which will hopefully be retired in favor of this new
algorithm.
The simplifications are obtained by maintaining per-task
nesting state for running tasks, and using a simple
lock-protected algorithm to handle accounting when tasks block
within RCU read-side critical sections, making use of lessons
learned while creating numerous user-level RCU implementations
over the past 18 months.
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: laijs@cn.fujitsu.com
Cc: dipankar@in.ibm.com
Cc: akpm@linux-foundation.org
Cc: mathieu.desnoyers@polymtl.ca
Cc: josht@linux.vnet.ibm.com
Cc: dvhltc@us.ibm.com
Cc: niv@us.ibm.com
Cc: peterz@infradead.org
Cc: rostedt@goodmis.org
LKML-Reference: <12509746134003-git-send-email->
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-08-23 04:56:52 +08:00
|
|
|
char rcu_read_unlock_special;
|
|
|
|
|
struct list_head rcu_node_entry;
|
2010-06-30 07:49:16 +08:00
|
|
|
#endif /* #ifdef CONFIG_PREEMPT_RCU */
|
|
|
|
|
#ifdef CONFIG_TREE_PREEMPT_RCU
|
|
|
|
|
struct rcu_node *rcu_blocked_node;
|
rcu: Merge preemptable-RCU functionality into hierarchical RCU
Create a kernel/rcutree_plugin.h file that contains definitions
for preemptable RCU (or, under the #else branch of the #ifdef,
empty definitions for the classic non-preemptable semantics).
These definitions fit into plugins defined in kernel/rcutree.c
for this purpose.
This variant of preemptable RCU uses a new algorithm whose
read-side expense is roughly that of classic hierarchical RCU
under CONFIG_PREEMPT. This new algorithm's update-side expense
is similar to that of classic hierarchical RCU, and, in absence
of read-side preemption or blocking, is exactly that of classic
hierarchical RCU. Perhaps more important, this new algorithm
has a much simpler implementation, saving well over 1,000 lines
of code compared to mainline's implementation of preemptable
RCU, which will hopefully be retired in favor of this new
algorithm.
The simplifications are obtained by maintaining per-task
nesting state for running tasks, and using a simple
lock-protected algorithm to handle accounting when tasks block
within RCU read-side critical sections, making use of lessons
learned while creating numerous user-level RCU implementations
over the past 18 months.
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: laijs@cn.fujitsu.com
Cc: dipankar@in.ibm.com
Cc: akpm@linux-foundation.org
Cc: mathieu.desnoyers@polymtl.ca
Cc: josht@linux.vnet.ibm.com
Cc: dvhltc@us.ibm.com
Cc: niv@us.ibm.com
Cc: peterz@infradead.org
Cc: rostedt@goodmis.org
LKML-Reference: <12509746134003-git-send-email->
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-08-23 04:56:52 +08:00
|
|
|
#endif /* #ifdef CONFIG_TREE_PREEMPT_RCU */
|
2010-09-28 08:25:23 +08:00
|
|
|
#ifdef CONFIG_RCU_BOOST
|
|
|
|
|
struct rt_mutex *rcu_boost_mutex;
|
|
|
|
|
#endif /* #ifdef CONFIG_RCU_BOOST */
|
2008-01-26 04:08:24 +08:00
|
|
|
|
2006-07-14 15:24:38 +08:00
|
|
|
#if defined(CONFIG_SCHEDSTATS) || defined(CONFIG_TASK_DELAY_ACCT)
|
2005-04-17 06:20:36 +08:00
|
|
|
struct sched_info sched_info;
|
|
|
|
|
#endif
|
|
|
|
|
|
|
|
|
|
struct list_head tasks;
|
2010-12-01 02:51:33 +08:00
|
|
|
#ifdef CONFIG_SMP
|
sched: create "pushable_tasks" list to limit pushing to one attempt
The RT scheduler employs a "push/pull" design to actively balance tasks
within the system (on a per disjoint cpuset basis). When a task is
awoken, it is immediately determined if there are any lower priority
cpus which should be preempted. This is opposed to the way normal
SCHED_OTHER tasks behave, which will wait for a periodic rebalancing
operation to occur before spreading out load.
When a particular RQ has more than 1 active RT task, it is said to
be in an "overloaded" state. Once this occurs, the system enters
the active balancing mode, where it will try to push the task away,
or persuade a different cpu to pull it over. The system will stay
in this state until the system falls back below the <= 1 queued RT
task per RQ.
However, the current implementation suffers from a limitation in the
push logic. Once overloaded, all tasks (other than current) on the
RQ are analyzed on every push operation, even if it was previously
unpushable (due to affinity, etc). Whats more, the operation stops
at the first task that is unpushable and will not look at items
lower in the queue. This causes two problems:
1) We can have the same tasks analyzed over and over again during each
push, which extends out the fast path in the scheduler for no
gain. Consider a RQ that has dozens of tasks that are bound to a
core. Each one of those tasks will be encountered and skipped
for each push operation while they are queued.
2) There may be lower-priority tasks under the unpushable task that
could have been successfully pushed, but will never be considered
until either the unpushable task is cleared, or a pull operation
succeeds. The net result is a potential latency source for mid
priority tasks.
This patch aims to rectify these two conditions by introducing a new
priority sorted list: "pushable_tasks". A task is added to the list
each time a task is activated or preempted. It is removed from the
list any time it is deactivated, made current, or fails to push.
This works because a task only needs to be attempted to push once.
After an initial failure to push, the other cpus will eventually try to
pull the task when the conditions are proper. This also solves the
problem that we don't completely analyze all tasks due to encountering
an unpushable tasks. Now every task will have a push attempted (when
appropriate).
This reduces latency both by shorting the critical section of the
rq->lock for certain workloads, and by making sure the algorithm
considers all eligible tasks in the system.
[ rostedt: added a couple more BUG_ONs ]
Signed-off-by: Gregory Haskins <ghaskins@novell.com>
Acked-by: Steven Rostedt <srostedt@redhat.com>
2008-12-29 22:39:53 +08:00
|
|
|
struct plist_node pushable_tasks;
|
2010-12-01 02:51:33 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
struct mm_struct *mm, *active_mm;
|
2011-04-15 06:22:09 +08:00
|
|
|
#ifdef CONFIG_COMPAT_BRK
|
|
|
|
|
unsigned brk_randomized:1;
|
|
|
|
|
#endif
|
2010-03-06 05:41:40 +08:00
|
|
|
#if defined(SPLIT_RSS_COUNTING)
|
|
|
|
|
struct task_rss_stat rss_stat;
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
/* task state */
|
2007-05-08 15:23:41 +08:00
|
|
|
int exit_state;
|
2005-04-17 06:20:36 +08:00
|
|
|
int exit_code, exit_signal;
|
|
|
|
|
int pdeath_signal; /* The signal sent when the parent dies */
|
2011-06-02 17:13:59 +08:00
|
|
|
unsigned int jobctl; /* JOBCTL_*, siglock protected */
|
2013-04-12 01:30:29 +08:00
|
|
|
|
|
|
|
|
/* Used for emulating ABI behavior of previous Linux versions */
|
2007-05-08 15:23:41 +08:00
|
|
|
unsigned int personality;
|
2013-04-12 01:30:29 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
unsigned did_exec:1;
|
2009-02-05 16:18:11 +08:00
|
|
|
unsigned in_execve:1; /* Tell the LSMs that the process is doing an
|
|
|
|
|
* execve */
|
2009-07-21 02:26:58 +08:00
|
|
|
unsigned in_iowait:1;
|
|
|
|
|
|
Add PR_{GET,SET}_NO_NEW_PRIVS to prevent execve from granting privs
With this change, calling
prctl(PR_SET_NO_NEW_PRIVS, 1, 0, 0, 0)
disables privilege granting operations at execve-time. For example, a
process will not be able to execute a setuid binary to change their uid
or gid if this bit is set. The same is true for file capabilities.
Additionally, LSM_UNSAFE_NO_NEW_PRIVS is defined to ensure that
LSMs respect the requested behavior.
To determine if the NO_NEW_PRIVS bit is set, a task may call
prctl(PR_GET_NO_NEW_PRIVS, 0, 0, 0, 0);
It returns 1 if set and 0 if it is not set. If any of the arguments are
non-zero, it will return -1 and set errno to -EINVAL.
(PR_SET_NO_NEW_PRIVS behaves similarly.)
This functionality is desired for the proposed seccomp filter patch
series. By using PR_SET_NO_NEW_PRIVS, it allows a task to modify the
system call behavior for itself and its child tasks without being
able to impact the behavior of a more privileged task.
Another potential use is making certain privileged operations
unprivileged. For example, chroot may be considered "safe" if it cannot
affect privileged tasks.
Note, this patch causes execve to fail when PR_SET_NO_NEW_PRIVS is
set and AppArmor is in use. It is fixed in a subsequent patch.
Signed-off-by: Andy Lutomirski <luto@amacapital.net>
Signed-off-by: Will Drewry <wad@chromium.org>
Acked-by: Eric Paris <eparis@redhat.com>
Acked-by: Kees Cook <keescook@chromium.org>
v18: updated change desc
v17: using new define values as per 3.4
Signed-off-by: James Morris <james.l.morris@oracle.com>
2012-04-13 05:47:50 +08:00
|
|
|
/* task may not gain privileges */
|
|
|
|
|
unsigned no_new_privs:1;
|
2009-06-15 23:17:47 +08:00
|
|
|
|
|
|
|
|
/* Revert to default priority/policy when forking */
|
|
|
|
|
unsigned sched_reset_on_fork:1;
|
2011-04-05 23:23:49 +08:00
|
|
|
unsigned sched_contributes_to_load:1;
|
2009-06-15 23:17:47 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
pid_t pid;
|
|
|
|
|
pid_t tgid;
|
2006-09-26 16:52:38 +08:00
|
|
|
|
2009-08-18 14:06:02 +08:00
|
|
|
#ifdef CONFIG_CC_STACKPROTECTOR
|
2006-09-26 16:52:38 +08:00
|
|
|
/* Canary value for the -fstack-protector gcc feature */
|
|
|
|
|
unsigned long stack_canary;
|
2009-08-18 14:06:02 +08:00
|
|
|
#endif
|
2012-05-11 08:59:08 +08:00
|
|
|
/*
|
2005-04-17 06:20:36 +08:00
|
|
|
* pointers to (original) parent process, youngest child, younger sibling,
|
2012-05-11 08:59:08 +08:00
|
|
|
* older sibling, respectively. (p->father can be replaced with
|
2008-03-25 09:36:23 +08:00
|
|
|
* p->real_parent->pid)
|
2005-04-17 06:20:36 +08:00
|
|
|
*/
|
2011-12-15 06:39:26 +08:00
|
|
|
struct task_struct __rcu *real_parent; /* real parent process */
|
|
|
|
|
struct task_struct __rcu *parent; /* recipient of SIGCHLD, wait4() reports */
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
2008-03-25 09:36:23 +08:00
|
|
|
* children/sibling forms the list of my natural children
|
2005-04-17 06:20:36 +08:00
|
|
|
*/
|
|
|
|
|
struct list_head children; /* list of my children */
|
|
|
|
|
struct list_head sibling; /* linkage in my parent's children list */
|
|
|
|
|
struct task_struct *group_leader; /* threadgroup leader */
|
|
|
|
|
|
2008-03-25 09:36:23 +08:00
|
|
|
/*
|
|
|
|
|
* ptraced is the list of tasks this task is using ptrace on.
|
|
|
|
|
* This includes both natural children and PTRACE_ATTACH targets.
|
|
|
|
|
* p->ptrace_entry is p's link on the p->parent->ptraced list.
|
|
|
|
|
*/
|
|
|
|
|
struct list_head ptraced;
|
|
|
|
|
struct list_head ptrace_entry;
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* PID/PID hash table linkage. */
|
[PATCH] pidhash: Refactor the pid hash table
Simplifies the code, reduces the need for 4 pid hash tables, and makes the
code more capable.
In the discussions I had with Oleg it was felt that to a large extent the
cleanup itself justified the work. With struct pid being dynamically
allocated meant we could create the hash table entry when the pid was
allocated and free the hash table entry when the pid was freed. Instead of
playing with the hash lists when ever a process would attach or detach to a
process.
For myself the fact that it gave what my previous task_ref patch gave for free
with simpler code was a big win. The problem is that if you hold a reference
to struct task_struct you lock in 10K of low memory. If you do that in a user
controllable way like /proc does, with an unprivileged but hostile user space
application with typical resource limits of 1000 fds and 100 processes I can
trigger the OOM killer by consuming all of low memory with task structs, on a
machine wight 1GB of low memory.
If I instead hold a reference to struct pid which holds a pointer to my
task_struct, I don't suffer from that problem because struct pid is 2 orders
of magnitude smaller. In fact struct pid is small enough that most other
kernel data structures dwarf it, so simply limiting the number of referring
data structures is enough to prevent exhaustion of low memory.
This splits the current struct pid into two structures, struct pid and struct
pid_link, and reduces our number of hash tables from PIDTYPE_MAX to just one.
struct pid_link is the per process linkage into the hash tables and lives in
struct task_struct. struct pid is given an indepedent lifetime, and holds
pointers to each of the pid types.
The independent life of struct pid simplifies attach_pid, and detach_pid,
because we are always manipulating the list of pids and not the hash table.
In addition in giving struct pid an indpendent life it makes the concept much
more powerful.
Kernel data structures can now embed a struct pid * instead of a pid_t and
not suffer from pid wrap around problems or from keeping unnecessarily
large amounts of memory allocated.
Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-31 18:31:42 +08:00
|
|
|
struct pid_link pids[PIDTYPE_MAX];
|
2006-03-29 08:11:25 +08:00
|
|
|
struct list_head thread_group;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
struct completion *vfork_done; /* for vfork() */
|
|
|
|
|
int __user *set_child_tid; /* CLONE_CHILD_SETTID */
|
|
|
|
|
int __user *clear_child_tid; /* CLONE_CHILD_CLEARTID */
|
|
|
|
|
|
2007-10-18 18:06:34 +08:00
|
|
|
cputime_t utime, stime, utimescaled, stimescaled;
|
2007-10-15 23:00:19 +08:00
|
|
|
cputime_t gtime;
|
2013-02-26 00:25:39 +08:00
|
|
|
#ifndef CONFIG_VIRT_CPU_ACCOUNTING_NATIVE
|
2012-11-22 07:58:35 +08:00
|
|
|
struct cputime prev_cputime;
|
2012-12-17 03:00:34 +08:00
|
|
|
#endif
|
|
|
|
|
#ifdef CONFIG_VIRT_CPU_ACCOUNTING_GEN
|
|
|
|
|
seqlock_t vtime_seqlock;
|
|
|
|
|
unsigned long long vtime_snap;
|
|
|
|
|
enum {
|
|
|
|
|
VTIME_SLEEPING = 0,
|
|
|
|
|
VTIME_USER,
|
|
|
|
|
VTIME_SYS,
|
|
|
|
|
} vtime_snap_whence;
|
2009-12-02 16:26:47 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
unsigned long nvcsw, nivcsw; /* context switch counts */
|
2007-07-16 14:39:42 +08:00
|
|
|
struct timespec start_time; /* monotonic time */
|
|
|
|
|
struct timespec real_start_time; /* boot based time */
|
2005-04-17 06:20:36 +08:00
|
|
|
/* mm fault and swap info: this can arguably be seen as either mm-specific or thread-specific */
|
|
|
|
|
unsigned long min_flt, maj_flt;
|
|
|
|
|
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
struct task_cputime cputime_expires;
|
2005-04-17 06:20:36 +08:00
|
|
|
struct list_head cpu_timers[3];
|
|
|
|
|
|
|
|
|
|
/* process credentials */
|
2010-02-25 02:45:09 +08:00
|
|
|
const struct cred __rcu *real_cred; /* objective and real subjective task
|
2008-11-14 07:39:26 +08:00
|
|
|
* credentials (COW) */
|
2010-02-25 02:45:09 +08:00
|
|
|
const struct cred __rcu *cred; /* effective (overridable) subjective task
|
2008-11-14 07:39:26 +08:00
|
|
|
* credentials (COW) */
|
2005-05-06 07:16:12 +08:00
|
|
|
char comm[TASK_COMM_LEN]; /* executable name excluding path
|
|
|
|
|
- access with [gs]et_task_comm (which lock
|
|
|
|
|
it with task_lock())
|
Split 'flush_old_exec' into two functions
'flush_old_exec()' is the point of no return when doing an execve(), and
it is pretty badly misnamed. It doesn't just flush the old executable
environment, it also starts up the new one.
Which is very inconvenient for things like setting up the new
personality, because we want the new personality to affect the starting
of the new environment, but at the same time we do _not_ want the new
personality to take effect if flushing the old one fails.
As a result, the x86-64 '32-bit' personality is actually done using this
insane "I'm going to change the ABI, but I haven't done it yet" bit
(TIF_ABI_PENDING), with SET_PERSONALITY() not actually setting the
personality, but just the "pending" bit, so that "flush_thread()" can do
the actual personality magic.
This patch in no way changes any of that insanity, but it does split the
'flush_old_exec()' function up into a preparatory part that can fail
(still called flush_old_exec()), and a new part that will actually set
up the new exec environment (setup_new_exec()). All callers are changed
to trivially comply with the new world order.
Signed-off-by: H. Peter Anvin <hpa@zytor.com>
Cc: stable@kernel.org
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2010-01-29 14:14:42 +08:00
|
|
|
- initialized normally by setup_new_exec */
|
2005-04-17 06:20:36 +08:00
|
|
|
/* file system info */
|
|
|
|
|
int link_count, total_link_count;
|
2006-09-29 16:59:40 +08:00
|
|
|
#ifdef CONFIG_SYSVIPC
|
2005-04-17 06:20:36 +08:00
|
|
|
/* ipc stuff */
|
|
|
|
|
struct sysv_sem sysvsem;
|
2006-09-29 16:59:40 +08:00
|
|
|
#endif
|
2009-01-16 03:08:40 +08:00
|
|
|
#ifdef CONFIG_DETECT_HUNG_TASK
|
2008-01-26 04:08:02 +08:00
|
|
|
/* hung task detection */
|
|
|
|
|
unsigned long last_switch_count;
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
/* CPU-specific state of this task */
|
|
|
|
|
struct thread_struct thread;
|
|
|
|
|
/* filesystem information */
|
|
|
|
|
struct fs_struct *fs;
|
|
|
|
|
/* open file information */
|
|
|
|
|
struct files_struct *files;
|
2006-10-02 17:18:08 +08:00
|
|
|
/* namespaces */
|
2006-10-02 17:18:06 +08:00
|
|
|
struct nsproxy *nsproxy;
|
2005-04-17 06:20:36 +08:00
|
|
|
/* signal handlers */
|
|
|
|
|
struct signal_struct *signal;
|
|
|
|
|
struct sighand_struct *sighand;
|
|
|
|
|
|
|
|
|
|
sigset_t blocked, real_blocked;
|
2008-04-30 15:53:09 +08:00
|
|
|
sigset_t saved_sigmask; /* restored if set_restore_sigmask() was used */
|
2005-04-17 06:20:36 +08:00
|
|
|
struct sigpending pending;
|
|
|
|
|
|
|
|
|
|
unsigned long sas_ss_sp;
|
|
|
|
|
size_t sas_ss_size;
|
|
|
|
|
int (*notifier)(void *priv);
|
|
|
|
|
void *notifier_data;
|
|
|
|
|
sigset_t *notifier_mask;
|
2012-06-27 15:07:19 +08:00
|
|
|
struct callback_head *task_works;
|
2012-05-11 08:59:07 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
struct audit_context *audit_context;
|
2008-01-10 17:53:18 +08:00
|
|
|
#ifdef CONFIG_AUDITSYSCALL
|
2012-09-11 13:39:43 +08:00
|
|
|
kuid_t loginuid;
|
2008-01-08 23:06:53 +08:00
|
|
|
unsigned int sessionid;
|
2008-01-10 17:53:18 +08:00
|
|
|
#endif
|
2012-04-13 05:47:54 +08:00
|
|
|
struct seccomp seccomp;
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* Thread group tracking */
|
|
|
|
|
u32 parent_exec_id;
|
|
|
|
|
u32 self_exec_id;
|
cpuset,mm: update tasks' mems_allowed in time
Fix allocating page cache/slab object on the unallowed node when memory
spread is set by updating tasks' mems_allowed after its cpuset's mems is
changed.
In order to update tasks' mems_allowed in time, we must modify the code of
memory policy. Because the memory policy is applied in the process's
context originally. After applying this patch, one task directly
manipulates anothers mems_allowed, and we use alloc_lock in the
task_struct to protect mems_allowed and memory policy of the task.
But in the fast path, we didn't use lock to protect them, because adding a
lock may lead to performance regression. But if we don't add a lock,the
task might see no nodes when changing cpuset's mems_allowed to some
non-overlapping set. In order to avoid it, we set all new allowed nodes,
then clear newly disallowed ones.
[lee.schermerhorn@hp.com:
The rework of mpol_new() to extract the adjusting of the node mask to
apply cpuset and mpol flags "context" breaks set_mempolicy() and mbind()
with MPOL_PREFERRED and a NULL nodemask--i.e., explicit local
allocation. Fix this by adding the check for MPOL_PREFERRED and empty
node mask to mpol_new_mpolicy().
Remove the now unneeded 'nodes = NULL' from mpol_new().
Note that mpol_new_mempolicy() is always called with a non-NULL
'nodes' parameter now that it has been removed from mpol_new().
Therefore, we don't need to test nodes for NULL before testing it for
'empty'. However, just to be extra paranoid, add a VM_BUG_ON() to
verify this assumption.]
[lee.schermerhorn@hp.com:
I don't think the function name 'mpol_new_mempolicy' is descriptive
enough to differentiate it from mpol_new().
This function applies cpuset set context, usually constraining nodes
to those allowed by the cpuset. However, when the 'RELATIVE_NODES flag
is set, it also translates the nodes. So I settled on
'mpol_set_nodemask()', because the comment block for mpol_new() mentions
that we need to call this function to "set nodes".
Some additional minor line length, whitespace and typo cleanup.]
Signed-off-by: Miao Xie <miaox@cn.fujitsu.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Paul Menage <menage@google.com>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Cc: Yasunori Goto <y-goto@jp.fujitsu.com>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-06-17 06:31:49 +08:00
|
|
|
/* Protection of (de-)allocation: mm, files, fs, tty, keyrings, mems_allowed,
|
|
|
|
|
* mempolicy */
|
2005-04-17 06:20:36 +08:00
|
|
|
spinlock_t alloc_lock;
|
|
|
|
|
|
2006-06-27 17:54:51 +08:00
|
|
|
/* Protection of the PI data structures: */
|
2009-11-17 21:54:03 +08:00
|
|
|
raw_spinlock_t pi_lock;
|
2006-06-27 17:54:51 +08:00
|
|
|
|
2006-06-27 17:54:53 +08:00
|
|
|
#ifdef CONFIG_RT_MUTEXES
|
|
|
|
|
/* PI waiters blocked on a rt_mutex held by this task */
|
|
|
|
|
struct plist_head pi_waiters;
|
|
|
|
|
/* Deadlock detection and priority inheritance handling */
|
|
|
|
|
struct rt_mutex_waiter *pi_blocked_on;
|
|
|
|
|
#endif
|
|
|
|
|
|
2006-01-10 07:59:20 +08:00
|
|
|
#ifdef CONFIG_DEBUG_MUTEXES
|
|
|
|
|
/* mutex deadlock detection */
|
|
|
|
|
struct mutex_waiter *blocked_on;
|
|
|
|
|
#endif
|
2006-07-03 15:24:42 +08:00
|
|
|
#ifdef CONFIG_TRACE_IRQFLAGS
|
|
|
|
|
unsigned int irq_events;
|
|
|
|
|
unsigned long hardirq_enable_ip;
|
|
|
|
|
unsigned long hardirq_disable_ip;
|
2009-11-30 13:59:44 +08:00
|
|
|
unsigned int hardirq_enable_event;
|
2006-07-03 15:24:42 +08:00
|
|
|
unsigned int hardirq_disable_event;
|
2009-11-30 13:59:44 +08:00
|
|
|
int hardirqs_enabled;
|
|
|
|
|
int hardirq_context;
|
2006-07-03 15:24:42 +08:00
|
|
|
unsigned long softirq_disable_ip;
|
|
|
|
|
unsigned long softirq_enable_ip;
|
2009-11-30 13:59:44 +08:00
|
|
|
unsigned int softirq_disable_event;
|
2006-07-03 15:24:42 +08:00
|
|
|
unsigned int softirq_enable_event;
|
2009-11-30 13:59:44 +08:00
|
|
|
int softirqs_enabled;
|
2006-07-03 15:24:42 +08:00
|
|
|
int softirq_context;
|
|
|
|
|
#endif
|
[PATCH] lockdep: core
Do 'make oldconfig' and accept all the defaults for new config options -
reboot into the kernel and if everything goes well it should boot up fine and
you should have /proc/lockdep and /proc/lockdep_stats files.
Typically if the lock validator finds some problem it will print out
voluminous debug output that begins with "BUG: ..." and which syslog output
can be used by kernel developers to figure out the precise locking scenario.
What does the lock validator do? It "observes" and maps all locking rules as
they occur dynamically (as triggered by the kernel's natural use of spinlocks,
rwlocks, mutexes and rwsems). Whenever the lock validator subsystem detects a
new locking scenario, it validates this new rule against the existing set of
rules. If this new rule is consistent with the existing set of rules then the
new rule is added transparently and the kernel continues as normal. If the
new rule could create a deadlock scenario then this condition is printed out.
When determining validity of locking, all possible "deadlock scenarios" are
considered: assuming arbitrary number of CPUs, arbitrary irq context and task
context constellations, running arbitrary combinations of all the existing
locking scenarios. In a typical system this means millions of separate
scenarios. This is why we call it a "locking correctness" validator - for all
rules that are observed the lock validator proves it with mathematical
certainty that a deadlock could not occur (assuming that the lock validator
implementation itself is correct and its internal data structures are not
corrupted by some other kernel subsystem). [see more details and conditionals
of this statement in include/linux/lockdep.h and
Documentation/lockdep-design.txt]
Furthermore, this "all possible scenarios" property of the validator also
enables the finding of complex, highly unlikely multi-CPU multi-context races
via single single-context rules, increasing the likelyhood of finding bugs
drastically. In practical terms: the lock validator already found a bug in
the upstream kernel that could only occur on systems with 3 or more CPUs, and
which needed 3 very unlikely code sequences to occur at once on the 3 CPUs.
That bug was found and reported on a single-CPU system (!). So in essence a
race will be found "piecemail-wise", triggering all the necessary components
for the race, without having to reproduce the race scenario itself! In its
short existence the lock validator found and reported many bugs before they
actually caused a real deadlock.
To further increase the efficiency of the validator, the mapping is not per
"lock instance", but per "lock-class". For example, all struct inode objects
in the kernel have inode->inotify_mutex. If there are 10,000 inodes cached,
then there are 10,000 lock objects. But ->inotify_mutex is a single "lock
type", and all locking activities that occur against ->inotify_mutex are
"unified" into this single lock-class. The advantage of the lock-class
approach is that all historical ->inotify_mutex uses are mapped into a single
(and as narrow as possible) set of locking rules - regardless of how many
different tasks or inode structures it took to build this set of rules. The
set of rules persist during the lifetime of the kernel.
To see the rough magnitude of checking that the lock validator does, here's a
portion of /proc/lockdep_stats, fresh after bootup:
lock-classes: 694 [max: 2048]
direct dependencies: 1598 [max: 8192]
indirect dependencies: 17896
all direct dependencies: 16206
dependency chains: 1910 [max: 8192]
in-hardirq chains: 17
in-softirq chains: 105
in-process chains: 1065
stack-trace entries: 38761 [max: 131072]
combined max dependencies: 2033928
hardirq-safe locks: 24
hardirq-unsafe locks: 176
softirq-safe locks: 53
softirq-unsafe locks: 137
irq-safe locks: 59
irq-unsafe locks: 176
The lock validator has observed 1598 actual single-thread locking patterns,
and has validated all possible 2033928 distinct locking scenarios.
More details about the design of the lock validator can be found in
Documentation/lockdep-design.txt, which can also found at:
http://redhat.com/~mingo/lockdep-patches/lockdep-design.txt
[bunk@stusta.de: cleanups]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Arjan van de Ven <arjan@linux.intel.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-07-03 15:24:50 +08:00
|
|
|
#ifdef CONFIG_LOCKDEP
|
2008-02-26 06:02:48 +08:00
|
|
|
# define MAX_LOCK_DEPTH 48UL
|
[PATCH] lockdep: core
Do 'make oldconfig' and accept all the defaults for new config options -
reboot into the kernel and if everything goes well it should boot up fine and
you should have /proc/lockdep and /proc/lockdep_stats files.
Typically if the lock validator finds some problem it will print out
voluminous debug output that begins with "BUG: ..." and which syslog output
can be used by kernel developers to figure out the precise locking scenario.
What does the lock validator do? It "observes" and maps all locking rules as
they occur dynamically (as triggered by the kernel's natural use of spinlocks,
rwlocks, mutexes and rwsems). Whenever the lock validator subsystem detects a
new locking scenario, it validates this new rule against the existing set of
rules. If this new rule is consistent with the existing set of rules then the
new rule is added transparently and the kernel continues as normal. If the
new rule could create a deadlock scenario then this condition is printed out.
When determining validity of locking, all possible "deadlock scenarios" are
considered: assuming arbitrary number of CPUs, arbitrary irq context and task
context constellations, running arbitrary combinations of all the existing
locking scenarios. In a typical system this means millions of separate
scenarios. This is why we call it a "locking correctness" validator - for all
rules that are observed the lock validator proves it with mathematical
certainty that a deadlock could not occur (assuming that the lock validator
implementation itself is correct and its internal data structures are not
corrupted by some other kernel subsystem). [see more details and conditionals
of this statement in include/linux/lockdep.h and
Documentation/lockdep-design.txt]
Furthermore, this "all possible scenarios" property of the validator also
enables the finding of complex, highly unlikely multi-CPU multi-context races
via single single-context rules, increasing the likelyhood of finding bugs
drastically. In practical terms: the lock validator already found a bug in
the upstream kernel that could only occur on systems with 3 or more CPUs, and
which needed 3 very unlikely code sequences to occur at once on the 3 CPUs.
That bug was found and reported on a single-CPU system (!). So in essence a
race will be found "piecemail-wise", triggering all the necessary components
for the race, without having to reproduce the race scenario itself! In its
short existence the lock validator found and reported many bugs before they
actually caused a real deadlock.
To further increase the efficiency of the validator, the mapping is not per
"lock instance", but per "lock-class". For example, all struct inode objects
in the kernel have inode->inotify_mutex. If there are 10,000 inodes cached,
then there are 10,000 lock objects. But ->inotify_mutex is a single "lock
type", and all locking activities that occur against ->inotify_mutex are
"unified" into this single lock-class. The advantage of the lock-class
approach is that all historical ->inotify_mutex uses are mapped into a single
(and as narrow as possible) set of locking rules - regardless of how many
different tasks or inode structures it took to build this set of rules. The
set of rules persist during the lifetime of the kernel.
To see the rough magnitude of checking that the lock validator does, here's a
portion of /proc/lockdep_stats, fresh after bootup:
lock-classes: 694 [max: 2048]
direct dependencies: 1598 [max: 8192]
indirect dependencies: 17896
all direct dependencies: 16206
dependency chains: 1910 [max: 8192]
in-hardirq chains: 17
in-softirq chains: 105
in-process chains: 1065
stack-trace entries: 38761 [max: 131072]
combined max dependencies: 2033928
hardirq-safe locks: 24
hardirq-unsafe locks: 176
softirq-safe locks: 53
softirq-unsafe locks: 137
irq-safe locks: 59
irq-unsafe locks: 176
The lock validator has observed 1598 actual single-thread locking patterns,
and has validated all possible 2033928 distinct locking scenarios.
More details about the design of the lock validator can be found in
Documentation/lockdep-design.txt, which can also found at:
http://redhat.com/~mingo/lockdep-patches/lockdep-design.txt
[bunk@stusta.de: cleanups]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Arjan van de Ven <arjan@linux.intel.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-07-03 15:24:50 +08:00
|
|
|
u64 curr_chain_key;
|
|
|
|
|
int lockdep_depth;
|
|
|
|
|
unsigned int lockdep_recursion;
|
2008-05-15 19:09:15 +08:00
|
|
|
struct held_lock held_locks[MAX_LOCK_DEPTH];
|
lockdep: annotate reclaim context (__GFP_NOFS)
Here is another version, with the incremental patch rolled up, and
added reclaim context annotation to kswapd, and allocation tracing
to slab allocators (which may only ever reach the page allocator
in rare cases, so it is good to put annotations here too).
Haven't tested this version as such, but it should be getting closer
to merge worthy ;)
--
After noticing some code in mm/filemap.c accidentally perform a __GFP_FS
allocation when it should not have been, I thought it might be a good idea to
try to catch this kind of thing with lockdep.
I coded up a little idea that seems to work. Unfortunately the system has to
actually be in __GFP_FS page reclaim, then take the lock, before it will mark
it. But at least that might still be some orders of magnitude more common
(and more debuggable) than an actual deadlock condition, so we have some
improvement I hope (the concept is no less complete than discovery of a lock's
interrupt contexts).
I guess we could even do the same thing with __GFP_IO (normal reclaim), and
even GFP_NOIO locks too... but filesystems will have the most locks and fiddly
code paths, so let's start there and see how it goes.
It *seems* to work. I did a quick test.
=================================
[ INFO: inconsistent lock state ]
2.6.28-rc6-00007-ged31348-dirty #26
---------------------------------
inconsistent {in-reclaim-W} -> {ov-reclaim-W} usage.
modprobe/8526 [HC0[0]:SC0[0]:HE1:SE1] takes:
(testlock){--..}, at: [<ffffffffa0020055>] brd_init+0x55/0x216 [brd]
{in-reclaim-W} state was registered at:
[<ffffffff80267bdb>] __lock_acquire+0x75b/0x1a60
[<ffffffff80268f71>] lock_acquire+0x91/0xc0
[<ffffffff8070f0e1>] mutex_lock_nested+0xb1/0x310
[<ffffffffa002002b>] brd_init+0x2b/0x216 [brd]
[<ffffffff8020903b>] _stext+0x3b/0x170
[<ffffffff80272ebf>] sys_init_module+0xaf/0x1e0
[<ffffffff8020c3fb>] system_call_fastpath+0x16/0x1b
[<ffffffffffffffff>] 0xffffffffffffffff
irq event stamp: 3929
hardirqs last enabled at (3929): [<ffffffff8070f2b5>] mutex_lock_nested+0x285/0x310
hardirqs last disabled at (3928): [<ffffffff8070f089>] mutex_lock_nested+0x59/0x310
softirqs last enabled at (3732): [<ffffffff8061f623>] sk_filter+0x83/0xe0
softirqs last disabled at (3730): [<ffffffff8061f5b6>] sk_filter+0x16/0xe0
other info that might help us debug this:
1 lock held by modprobe/8526:
#0: (testlock){--..}, at: [<ffffffffa0020055>] brd_init+0x55/0x216 [brd]
stack backtrace:
Pid: 8526, comm: modprobe Not tainted 2.6.28-rc6-00007-ged31348-dirty #26
Call Trace:
[<ffffffff80265483>] print_usage_bug+0x193/0x1d0
[<ffffffff80266530>] mark_lock+0xaf0/0xca0
[<ffffffff80266735>] mark_held_locks+0x55/0xc0
[<ffffffffa0020000>] ? brd_init+0x0/0x216 [brd]
[<ffffffff802667ca>] trace_reclaim_fs+0x2a/0x60
[<ffffffff80285005>] __alloc_pages_internal+0x475/0x580
[<ffffffff8070f29e>] ? mutex_lock_nested+0x26e/0x310
[<ffffffffa0020000>] ? brd_init+0x0/0x216 [brd]
[<ffffffffa002006a>] brd_init+0x6a/0x216 [brd]
[<ffffffffa0020000>] ? brd_init+0x0/0x216 [brd]
[<ffffffff8020903b>] _stext+0x3b/0x170
[<ffffffff8070f8b9>] ? mutex_unlock+0x9/0x10
[<ffffffff8070f83d>] ? __mutex_unlock_slowpath+0x10d/0x180
[<ffffffff802669ec>] ? trace_hardirqs_on_caller+0x12c/0x190
[<ffffffff80272ebf>] sys_init_module+0xaf/0x1e0
[<ffffffff8020c3fb>] system_call_fastpath+0x16/0x1b
Signed-off-by: Nick Piggin <npiggin@suse.de>
Signed-off-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-01-21 15:12:39 +08:00
|
|
|
gfp_t lockdep_reclaim_gfp;
|
[PATCH] lockdep: core
Do 'make oldconfig' and accept all the defaults for new config options -
reboot into the kernel and if everything goes well it should boot up fine and
you should have /proc/lockdep and /proc/lockdep_stats files.
Typically if the lock validator finds some problem it will print out
voluminous debug output that begins with "BUG: ..." and which syslog output
can be used by kernel developers to figure out the precise locking scenario.
What does the lock validator do? It "observes" and maps all locking rules as
they occur dynamically (as triggered by the kernel's natural use of spinlocks,
rwlocks, mutexes and rwsems). Whenever the lock validator subsystem detects a
new locking scenario, it validates this new rule against the existing set of
rules. If this new rule is consistent with the existing set of rules then the
new rule is added transparently and the kernel continues as normal. If the
new rule could create a deadlock scenario then this condition is printed out.
When determining validity of locking, all possible "deadlock scenarios" are
considered: assuming arbitrary number of CPUs, arbitrary irq context and task
context constellations, running arbitrary combinations of all the existing
locking scenarios. In a typical system this means millions of separate
scenarios. This is why we call it a "locking correctness" validator - for all
rules that are observed the lock validator proves it with mathematical
certainty that a deadlock could not occur (assuming that the lock validator
implementation itself is correct and its internal data structures are not
corrupted by some other kernel subsystem). [see more details and conditionals
of this statement in include/linux/lockdep.h and
Documentation/lockdep-design.txt]
Furthermore, this "all possible scenarios" property of the validator also
enables the finding of complex, highly unlikely multi-CPU multi-context races
via single single-context rules, increasing the likelyhood of finding bugs
drastically. In practical terms: the lock validator already found a bug in
the upstream kernel that could only occur on systems with 3 or more CPUs, and
which needed 3 very unlikely code sequences to occur at once on the 3 CPUs.
That bug was found and reported on a single-CPU system (!). So in essence a
race will be found "piecemail-wise", triggering all the necessary components
for the race, without having to reproduce the race scenario itself! In its
short existence the lock validator found and reported many bugs before they
actually caused a real deadlock.
To further increase the efficiency of the validator, the mapping is not per
"lock instance", but per "lock-class". For example, all struct inode objects
in the kernel have inode->inotify_mutex. If there are 10,000 inodes cached,
then there are 10,000 lock objects. But ->inotify_mutex is a single "lock
type", and all locking activities that occur against ->inotify_mutex are
"unified" into this single lock-class. The advantage of the lock-class
approach is that all historical ->inotify_mutex uses are mapped into a single
(and as narrow as possible) set of locking rules - regardless of how many
different tasks or inode structures it took to build this set of rules. The
set of rules persist during the lifetime of the kernel.
To see the rough magnitude of checking that the lock validator does, here's a
portion of /proc/lockdep_stats, fresh after bootup:
lock-classes: 694 [max: 2048]
direct dependencies: 1598 [max: 8192]
indirect dependencies: 17896
all direct dependencies: 16206
dependency chains: 1910 [max: 8192]
in-hardirq chains: 17
in-softirq chains: 105
in-process chains: 1065
stack-trace entries: 38761 [max: 131072]
combined max dependencies: 2033928
hardirq-safe locks: 24
hardirq-unsafe locks: 176
softirq-safe locks: 53
softirq-unsafe locks: 137
irq-safe locks: 59
irq-unsafe locks: 176
The lock validator has observed 1598 actual single-thread locking patterns,
and has validated all possible 2033928 distinct locking scenarios.
More details about the design of the lock validator can be found in
Documentation/lockdep-design.txt, which can also found at:
http://redhat.com/~mingo/lockdep-patches/lockdep-design.txt
[bunk@stusta.de: cleanups]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Arjan van de Ven <arjan@linux.intel.com>
Signed-off-by: Adrian Bunk <bunk@stusta.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-07-03 15:24:50 +08:00
|
|
|
#endif
|
2006-01-10 07:59:20 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* journalling filesystem info */
|
|
|
|
|
void *journal_info;
|
|
|
|
|
|
When stacked block devices are in-use (e.g. md or dm), the recursive calls
to generic_make_request can use up a lot of space, and we would rather they
didn't.
As generic_make_request is a void function, and as it is generally not
expected that it will have any effect immediately, it is safe to delay any
call to generic_make_request until there is sufficient stack space
available.
As ->bi_next is reserved for the driver to use, it can have no valid value
when generic_make_request is called, and as __make_request implicitly
assumes it will be NULL (ELEVATOR_BACK_MERGE fork of switch) we can be
certain that all callers set it to NULL. We can therefore safely use
bi_next to link pending requests together, providing we clear it before
making the real call.
So, we choose to allow each thread to only be active in one
generic_make_request at a time. If a subsequent (recursive) call is made,
the bio is linked into a per-thread list, and is handled when the active
call completes.
As the list of pending bios is per-thread, there are no locking issues to
worry about.
I say above that it is "safe to delay any call...". There are, however,
some behaviours of a make_request_fn which would make it unsafe. These
include any behaviour that assumes anything will have changed after a
recursive call to generic_make_request.
These could include:
- waiting for that call to finish and call it's bi_end_io function.
md use to sometimes do this (marking the superblock dirty before
completing a write) but doesn't any more
- inspecting the bio for fields that generic_make_request might
change, such as bi_sector or bi_bdev. It is hard to see a good
reason for this, and I don't think anyone actually does it.
- inspecing the queue to see if, e.g. it is 'full' yet. Again, I
think this is very unlikely to be useful, or to be done.
Signed-off-by: Neil Brown <neilb@suse.de>
Cc: Jens Axboe <axboe@kernel.dk>
Cc: <dm-devel@redhat.com>
Alasdair G Kergon <agk@redhat.com> said:
I can see nothing wrong with this in principle.
For device-mapper at the moment though it's essential that, while the bio
mappings may now get delayed, they still get processed in exactly
the same order as they were passed to generic_make_request().
My main concern is whether the timing changes implicit in this patch
will make the rare data-corrupting races in the existing snapshot code
more likely. (I'm working on a fix for these races, but the unfinished
patch is already several hundred lines long.)
It would be helpful if some people on this mailing list would test
this patch in various scenarios and report back.
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Jens Axboe <jens.axboe@oracle.com>
2007-05-01 15:53:42 +08:00
|
|
|
/* stacked block device info */
|
2010-02-23 15:55:42 +08:00
|
|
|
struct bio_list *bio_list;
|
When stacked block devices are in-use (e.g. md or dm), the recursive calls
to generic_make_request can use up a lot of space, and we would rather they
didn't.
As generic_make_request is a void function, and as it is generally not
expected that it will have any effect immediately, it is safe to delay any
call to generic_make_request until there is sufficient stack space
available.
As ->bi_next is reserved for the driver to use, it can have no valid value
when generic_make_request is called, and as __make_request implicitly
assumes it will be NULL (ELEVATOR_BACK_MERGE fork of switch) we can be
certain that all callers set it to NULL. We can therefore safely use
bi_next to link pending requests together, providing we clear it before
making the real call.
So, we choose to allow each thread to only be active in one
generic_make_request at a time. If a subsequent (recursive) call is made,
the bio is linked into a per-thread list, and is handled when the active
call completes.
As the list of pending bios is per-thread, there are no locking issues to
worry about.
I say above that it is "safe to delay any call...". There are, however,
some behaviours of a make_request_fn which would make it unsafe. These
include any behaviour that assumes anything will have changed after a
recursive call to generic_make_request.
These could include:
- waiting for that call to finish and call it's bi_end_io function.
md use to sometimes do this (marking the superblock dirty before
completing a write) but doesn't any more
- inspecting the bio for fields that generic_make_request might
change, such as bi_sector or bi_bdev. It is hard to see a good
reason for this, and I don't think anyone actually does it.
- inspecing the queue to see if, e.g. it is 'full' yet. Again, I
think this is very unlikely to be useful, or to be done.
Signed-off-by: Neil Brown <neilb@suse.de>
Cc: Jens Axboe <axboe@kernel.dk>
Cc: <dm-devel@redhat.com>
Alasdair G Kergon <agk@redhat.com> said:
I can see nothing wrong with this in principle.
For device-mapper at the moment though it's essential that, while the bio
mappings may now get delayed, they still get processed in exactly
the same order as they were passed to generic_make_request().
My main concern is whether the timing changes implicit in this patch
will make the rare data-corrupting races in the existing snapshot code
more likely. (I'm working on a fix for these races, but the unfinished
patch is already several hundred lines long.)
It would be helpful if some people on this mailing list would test
this patch in various scenarios and report back.
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Jens Axboe <jens.axboe@oracle.com>
2007-05-01 15:53:42 +08:00
|
|
|
|
2011-03-08 20:19:51 +08:00
|
|
|
#ifdef CONFIG_BLOCK
|
|
|
|
|
/* stack plugging */
|
|
|
|
|
struct blk_plug *plug;
|
|
|
|
|
#endif
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* VM state */
|
|
|
|
|
struct reclaim_state *reclaim_state;
|
|
|
|
|
|
|
|
|
|
struct backing_dev_info *backing_dev_info;
|
|
|
|
|
|
|
|
|
|
struct io_context *io_context;
|
|
|
|
|
|
|
|
|
|
unsigned long ptrace_message;
|
|
|
|
|
siginfo_t *last_siginfo; /* For ptrace use. */
|
2006-12-10 18:19:19 +08:00
|
|
|
struct task_io_accounting ioac;
|
2006-10-01 14:28:59 +08:00
|
|
|
#if defined(CONFIG_TASK_XACCT)
|
2005-04-17 06:20:36 +08:00
|
|
|
u64 acct_rss_mem1; /* accumulated rss usage */
|
|
|
|
|
u64 acct_vm_mem1; /* accumulated virtual memory usage */
|
2008-07-25 16:48:40 +08:00
|
|
|
cputime_t acct_timexpd; /* stime + utime since last update */
|
2005-04-17 06:20:36 +08:00
|
|
|
#endif
|
|
|
|
|
#ifdef CONFIG_CPUSETS
|
cpuset,mm: update tasks' mems_allowed in time
Fix allocating page cache/slab object on the unallowed node when memory
spread is set by updating tasks' mems_allowed after its cpuset's mems is
changed.
In order to update tasks' mems_allowed in time, we must modify the code of
memory policy. Because the memory policy is applied in the process's
context originally. After applying this patch, one task directly
manipulates anothers mems_allowed, and we use alloc_lock in the
task_struct to protect mems_allowed and memory policy of the task.
But in the fast path, we didn't use lock to protect them, because adding a
lock may lead to performance regression. But if we don't add a lock,the
task might see no nodes when changing cpuset's mems_allowed to some
non-overlapping set. In order to avoid it, we set all new allowed nodes,
then clear newly disallowed ones.
[lee.schermerhorn@hp.com:
The rework of mpol_new() to extract the adjusting of the node mask to
apply cpuset and mpol flags "context" breaks set_mempolicy() and mbind()
with MPOL_PREFERRED and a NULL nodemask--i.e., explicit local
allocation. Fix this by adding the check for MPOL_PREFERRED and empty
node mask to mpol_new_mpolicy().
Remove the now unneeded 'nodes = NULL' from mpol_new().
Note that mpol_new_mempolicy() is always called with a non-NULL
'nodes' parameter now that it has been removed from mpol_new().
Therefore, we don't need to test nodes for NULL before testing it for
'empty'. However, just to be extra paranoid, add a VM_BUG_ON() to
verify this assumption.]
[lee.schermerhorn@hp.com:
I don't think the function name 'mpol_new_mempolicy' is descriptive
enough to differentiate it from mpol_new().
This function applies cpuset set context, usually constraining nodes
to those allowed by the cpuset. However, when the 'RELATIVE_NODES flag
is set, it also translates the nodes. So I settled on
'mpol_set_nodemask()', because the comment block for mpol_new() mentions
that we need to call this function to "set nodes".
Some additional minor line length, whitespace and typo cleanup.]
Signed-off-by: Miao Xie <miaox@cn.fujitsu.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Paul Menage <menage@google.com>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Cc: Yasunori Goto <y-goto@jp.fujitsu.com>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-06-17 06:31:49 +08:00
|
|
|
nodemask_t mems_allowed; /* Protected by alloc_lock */
|
cpuset: mm: reduce large amounts of memory barrier related damage v3
Commit c0ff7453bb5c ("cpuset,mm: fix no node to alloc memory when
changing cpuset's mems") wins a super prize for the largest number of
memory barriers entered into fast paths for one commit.
[get|put]_mems_allowed is incredibly heavy with pairs of full memory
barriers inserted into a number of hot paths. This was detected while
investigating at large page allocator slowdown introduced some time
after 2.6.32. The largest portion of this overhead was shown by
oprofile to be at an mfence introduced by this commit into the page
allocator hot path.
For extra style points, the commit introduced the use of yield() in an
implementation of what looks like a spinning mutex.
This patch replaces the full memory barriers on both read and write
sides with a sequence counter with just read barriers on the fast path
side. This is much cheaper on some architectures, including x86. The
main bulk of the patch is the retry logic if the nodemask changes in a
manner that can cause a false failure.
While updating the nodemask, a check is made to see if a false failure
is a risk. If it is, the sequence number gets bumped and parallel
allocators will briefly stall while the nodemask update takes place.
In a page fault test microbenchmark, oprofile samples from
__alloc_pages_nodemask went from 4.53% of all samples to 1.15%. The
actual results were
3.3.0-rc3 3.3.0-rc3
rc3-vanilla nobarrier-v2r1
Clients 1 UserTime 0.07 ( 0.00%) 0.08 (-14.19%)
Clients 2 UserTime 0.07 ( 0.00%) 0.07 ( 2.72%)
Clients 4 UserTime 0.08 ( 0.00%) 0.07 ( 3.29%)
Clients 1 SysTime 0.70 ( 0.00%) 0.65 ( 6.65%)
Clients 2 SysTime 0.85 ( 0.00%) 0.82 ( 3.65%)
Clients 4 SysTime 1.41 ( 0.00%) 1.41 ( 0.32%)
Clients 1 WallTime 0.77 ( 0.00%) 0.74 ( 4.19%)
Clients 2 WallTime 0.47 ( 0.00%) 0.45 ( 3.73%)
Clients 4 WallTime 0.38 ( 0.00%) 0.37 ( 1.58%)
Clients 1 Flt/sec/cpu 497620.28 ( 0.00%) 520294.53 ( 4.56%)
Clients 2 Flt/sec/cpu 414639.05 ( 0.00%) 429882.01 ( 3.68%)
Clients 4 Flt/sec/cpu 257959.16 ( 0.00%) 258761.48 ( 0.31%)
Clients 1 Flt/sec 495161.39 ( 0.00%) 517292.87 ( 4.47%)
Clients 2 Flt/sec 820325.95 ( 0.00%) 850289.77 ( 3.65%)
Clients 4 Flt/sec 1020068.93 ( 0.00%) 1022674.06 ( 0.26%)
MMTests Statistics: duration
Sys Time Running Test (seconds) 135.68 132.17
User+Sys Time Running Test (seconds) 164.2 160.13
Total Elapsed Time (seconds) 123.46 120.87
The overall improvement is small but the System CPU time is much
improved and roughly in correlation to what oprofile reported (these
performance figures are without profiling so skew is expected). The
actual number of page faults is noticeably improved.
For benchmarks like kernel builds, the overall benefit is marginal but
the system CPU time is slightly reduced.
To test the actual bug the commit fixed I opened two terminals. The
first ran within a cpuset and continually ran a small program that
faulted 100M of anonymous data. In a second window, the nodemask of the
cpuset was continually randomised in a loop.
Without the commit, the program would fail every so often (usually
within 10 seconds) and obviously with the commit everything worked fine.
With this patch applied, it also worked fine so the fix should be
functionally equivalent.
Signed-off-by: Mel Gorman <mgorman@suse.de>
Cc: Miao Xie <miaox@cn.fujitsu.com>
Cc: David Rientjes <rientjes@google.com>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Christoph Lameter <cl@linux.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-03-22 07:34:11 +08:00
|
|
|
seqcount_t mems_allowed_seq; /* Seqence no to catch updates */
|
[PATCH] cpuset memory spread basic implementation
This patch provides the implementation and cpuset interface for an alternative
memory allocation policy that can be applied to certain kinds of memory
allocations, such as the page cache (file system buffers) and some slab caches
(such as inode caches).
The policy is called "memory spreading." If enabled, it spreads out these
kinds of memory allocations over all the nodes allowed to a task, instead of
preferring to place them on the node where the task is executing.
All other kinds of allocations, including anonymous pages for a tasks stack
and data regions, are not affected by this policy choice, and continue to be
allocated preferring the node local to execution, as modified by the NUMA
mempolicy.
There are two boolean flag files per cpuset that control where the kernel
allocates pages for the file system buffers and related in kernel data
structures. They are called 'memory_spread_page' and 'memory_spread_slab'.
If the per-cpuset boolean flag file 'memory_spread_page' is set, then the
kernel will spread the file system buffers (page cache) evenly over all the
nodes that the faulting task is allowed to use, instead of preferring to put
those pages on the node where the task is running.
If the per-cpuset boolean flag file 'memory_spread_slab' is set, then the
kernel will spread some file system related slab caches, such as for inodes
and dentries evenly over all the nodes that the faulting task is allowed to
use, instead of preferring to put those pages on the node where the task is
running.
The implementation is simple. Setting the cpuset flags 'memory_spread_page'
or 'memory_spread_cache' turns on the per-process flags PF_SPREAD_PAGE or
PF_SPREAD_SLAB, respectively, for each task that is in the cpuset or
subsequently joins that cpuset. In subsequent patches, the page allocation
calls for the affected page cache and slab caches are modified to perform an
inline check for these flags, and if set, a call to a new routine
cpuset_mem_spread_node() returns the node to prefer for the allocation.
The cpuset_mem_spread_node() routine is also simple. It uses the value of a
per-task rotor cpuset_mem_spread_rotor to select the next node in the current
tasks mems_allowed to prefer for the allocation.
This policy can provide substantial improvements for jobs that need to place
thread local data on the corresponding node, but that need to access large
file system data sets that need to be spread across the several nodes in the
jobs cpuset in order to fit. Without this patch, especially for jobs that
might have one thread reading in the data set, the memory allocation across
the nodes in the jobs cpuset can become very uneven.
A couple of Copyright year ranges are updated as well. And a couple of email
addresses that can be found in the MAINTAINERS file are removed.
Signed-off-by: Paul Jackson <pj@sgi.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-24 19:16:03 +08:00
|
|
|
int cpuset_mem_spread_rotor;
|
2010-05-27 05:42:49 +08:00
|
|
|
int cpuset_slab_spread_rotor;
|
2005-04-17 06:20:36 +08:00
|
|
|
#endif
|
Task Control Groups: basic task cgroup framework
Generic Process Control Groups
--------------------------
There have recently been various proposals floating around for
resource management/accounting and other task grouping subsystems in
the kernel, including ResGroups, User BeanCounters, NSProxy
cgroups, and others. These all need the basic abstraction of being
able to group together multiple processes in an aggregate, in order to
track/limit the resources permitted to those processes, or control
other behaviour of the processes, and all implement this grouping in
different ways.
This patchset provides a framework for tracking and grouping processes
into arbitrary "cgroups" and assigning arbitrary state to those
groupings, in order to control the behaviour of the cgroup as an
aggregate.
The intention is that the various resource management and
virtualization/cgroup efforts can also become task cgroup
clients, with the result that:
- the userspace APIs are (somewhat) normalised
- it's easier to test e.g. the ResGroups CPU controller in
conjunction with the BeanCounters memory controller, or use either of
them as the resource-control portion of a virtual server system.
- the additional kernel footprint of any of the competing resource
management systems is substantially reduced, since it doesn't need
to provide process grouping/containment, hence improving their
chances of getting into the kernel
This patch:
Add the main task cgroups framework - the cgroup filesystem, and the
basic structures for tracking membership and associating subsystem state
objects to tasks.
Signed-off-by: Paul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-19 14:39:30 +08:00
|
|
|
#ifdef CONFIG_CGROUPS
|
2007-10-19 14:39:36 +08:00
|
|
|
/* Control Group info protected by css_set_lock */
|
2010-02-25 02:41:39 +08:00
|
|
|
struct css_set __rcu *cgroups;
|
2007-10-19 14:39:36 +08:00
|
|
|
/* cg_list protected by css_set_lock and tsk->alloc_lock */
|
|
|
|
|
struct list_head cg_list;
|
Task Control Groups: basic task cgroup framework
Generic Process Control Groups
--------------------------
There have recently been various proposals floating around for
resource management/accounting and other task grouping subsystems in
the kernel, including ResGroups, User BeanCounters, NSProxy
cgroups, and others. These all need the basic abstraction of being
able to group together multiple processes in an aggregate, in order to
track/limit the resources permitted to those processes, or control
other behaviour of the processes, and all implement this grouping in
different ways.
This patchset provides a framework for tracking and grouping processes
into arbitrary "cgroups" and assigning arbitrary state to those
groupings, in order to control the behaviour of the cgroup as an
aggregate.
The intention is that the various resource management and
virtualization/cgroup efforts can also become task cgroup
clients, with the result that:
- the userspace APIs are (somewhat) normalised
- it's easier to test e.g. the ResGroups CPU controller in
conjunction with the BeanCounters memory controller, or use either of
them as the resource-control portion of a virtual server system.
- the additional kernel footprint of any of the competing resource
management systems is substantially reduced, since it doesn't need
to provide process grouping/containment, hence improving their
chances of getting into the kernel
This patch:
Add the main task cgroups framework - the cgroup filesystem, and the
basic structures for tracking membership and associating subsystem state
objects to tasks.
Signed-off-by: Paul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-19 14:39:30 +08:00
|
|
|
#endif
|
2007-10-17 14:27:30 +08:00
|
|
|
#ifdef CONFIG_FUTEX
|
2006-03-27 17:16:22 +08:00
|
|
|
struct robust_list_head __user *robust_list;
|
2006-03-27 17:16:24 +08:00
|
|
|
#ifdef CONFIG_COMPAT
|
|
|
|
|
struct compat_robust_list_head __user *compat_robust_list;
|
|
|
|
|
#endif
|
2006-06-27 17:54:58 +08:00
|
|
|
struct list_head pi_state_list;
|
|
|
|
|
struct futex_pi_state *pi_state_cache;
|
2008-05-15 19:09:15 +08:00
|
|
|
#endif
|
perf: Do the big rename: Performance Counters -> Performance Events
Bye-bye Performance Counters, welcome Performance Events!
In the past few months the perfcounters subsystem has grown out its
initial role of counting hardware events, and has become (and is
becoming) a much broader generic event enumeration, reporting, logging,
monitoring, analysis facility.
Naming its core object 'perf_counter' and naming the subsystem
'perfcounters' has become more and more of a misnomer. With pending
code like hw-breakpoints support the 'counter' name is less and
less appropriate.
All in one, we've decided to rename the subsystem to 'performance
events' and to propagate this rename through all fields, variables
and API names. (in an ABI compatible fashion)
The word 'event' is also a bit shorter than 'counter' - which makes
it slightly more convenient to write/handle as well.
Thanks goes to Stephane Eranian who first observed this misnomer and
suggested a rename.
User-space tooling and ABI compatibility is not affected - this patch
should be function-invariant. (Also, defconfigs were not touched to
keep the size down.)
This patch has been generated via the following script:
FILES=$(find * -type f | grep -vE 'oprofile|[^K]config')
sed -i \
-e 's/PERF_EVENT_/PERF_RECORD_/g' \
-e 's/PERF_COUNTER/PERF_EVENT/g' \
-e 's/perf_counter/perf_event/g' \
-e 's/nb_counters/nb_events/g' \
-e 's/swcounter/swevent/g' \
-e 's/tpcounter_event/tp_event/g' \
$FILES
for N in $(find . -name perf_counter.[ch]); do
M=$(echo $N | sed 's/perf_counter/perf_event/g')
mv $N $M
done
FILES=$(find . -name perf_event.*)
sed -i \
-e 's/COUNTER_MASK/REG_MASK/g' \
-e 's/COUNTER/EVENT/g' \
-e 's/\<event\>/event_id/g' \
-e 's/counter/event/g' \
-e 's/Counter/Event/g' \
$FILES
... to keep it as correct as possible. This script can also be
used by anyone who has pending perfcounters patches - it converts
a Linux kernel tree over to the new naming. We tried to time this
change to the point in time where the amount of pending patches
is the smallest: the end of the merge window.
Namespace clashes were fixed up in a preparatory patch - and some
stylistic fallout will be fixed up in a subsequent patch.
( NOTE: 'counters' are still the proper terminology when we deal
with hardware registers - and these sed scripts are a bit
over-eager in renaming them. I've undone some of that, but
in case there's something left where 'counter' would be
better than 'event' we can undo that on an individual basis
instead of touching an otherwise nicely automated patch. )
Suggested-by: Stephane Eranian <eranian@google.com>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Acked-by: Paul Mackerras <paulus@samba.org>
Reviewed-by: Arjan van de Ven <arjan@linux.intel.com>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Arnaldo Carvalho de Melo <acme@redhat.com>
Cc: Frederic Weisbecker <fweisbec@gmail.com>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Howells <dhowells@redhat.com>
Cc: Kyle McMartin <kyle@mcmartin.ca>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Cc: <linux-arch@vger.kernel.org>
LKML-Reference: <new-submission>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 18:02:48 +08:00
|
|
|
#ifdef CONFIG_PERF_EVENTS
|
2010-09-02 22:50:03 +08:00
|
|
|
struct perf_event_context *perf_event_ctxp[perf_nr_task_contexts];
|
perf: Do the big rename: Performance Counters -> Performance Events
Bye-bye Performance Counters, welcome Performance Events!
In the past few months the perfcounters subsystem has grown out its
initial role of counting hardware events, and has become (and is
becoming) a much broader generic event enumeration, reporting, logging,
monitoring, analysis facility.
Naming its core object 'perf_counter' and naming the subsystem
'perfcounters' has become more and more of a misnomer. With pending
code like hw-breakpoints support the 'counter' name is less and
less appropriate.
All in one, we've decided to rename the subsystem to 'performance
events' and to propagate this rename through all fields, variables
and API names. (in an ABI compatible fashion)
The word 'event' is also a bit shorter than 'counter' - which makes
it slightly more convenient to write/handle as well.
Thanks goes to Stephane Eranian who first observed this misnomer and
suggested a rename.
User-space tooling and ABI compatibility is not affected - this patch
should be function-invariant. (Also, defconfigs were not touched to
keep the size down.)
This patch has been generated via the following script:
FILES=$(find * -type f | grep -vE 'oprofile|[^K]config')
sed -i \
-e 's/PERF_EVENT_/PERF_RECORD_/g' \
-e 's/PERF_COUNTER/PERF_EVENT/g' \
-e 's/perf_counter/perf_event/g' \
-e 's/nb_counters/nb_events/g' \
-e 's/swcounter/swevent/g' \
-e 's/tpcounter_event/tp_event/g' \
$FILES
for N in $(find . -name perf_counter.[ch]); do
M=$(echo $N | sed 's/perf_counter/perf_event/g')
mv $N $M
done
FILES=$(find . -name perf_event.*)
sed -i \
-e 's/COUNTER_MASK/REG_MASK/g' \
-e 's/COUNTER/EVENT/g' \
-e 's/\<event\>/event_id/g' \
-e 's/counter/event/g' \
-e 's/Counter/Event/g' \
$FILES
... to keep it as correct as possible. This script can also be
used by anyone who has pending perfcounters patches - it converts
a Linux kernel tree over to the new naming. We tried to time this
change to the point in time where the amount of pending patches
is the smallest: the end of the merge window.
Namespace clashes were fixed up in a preparatory patch - and some
stylistic fallout will be fixed up in a subsequent patch.
( NOTE: 'counters' are still the proper terminology when we deal
with hardware registers - and these sed scripts are a bit
over-eager in renaming them. I've undone some of that, but
in case there's something left where 'counter' would be
better than 'event' we can undo that on an individual basis
instead of touching an otherwise nicely automated patch. )
Suggested-by: Stephane Eranian <eranian@google.com>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Acked-by: Paul Mackerras <paulus@samba.org>
Reviewed-by: Arjan van de Ven <arjan@linux.intel.com>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Arnaldo Carvalho de Melo <acme@redhat.com>
Cc: Frederic Weisbecker <fweisbec@gmail.com>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Benjamin Herrenschmidt <benh@kernel.crashing.org>
Cc: David Howells <dhowells@redhat.com>
Cc: Kyle McMartin <kyle@mcmartin.ca>
Cc: Martin Schwidefsky <schwidefsky@de.ibm.com>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: "H. Peter Anvin" <hpa@zytor.com>
Cc: <linux-arch@vger.kernel.org>
LKML-Reference: <new-submission>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-09-21 18:02:48 +08:00
|
|
|
struct mutex perf_event_mutex;
|
|
|
|
|
struct list_head perf_event_list;
|
perf_counter: Dynamically allocate tasks' perf_counter_context struct
This replaces the struct perf_counter_context in the task_struct with
a pointer to a dynamically allocated perf_counter_context struct. The
main reason for doing is this is to allow us to transfer a
perf_counter_context from one task to another when we do lazy PMU
switching in a later patch.
This has a few side-benefits: the task_struct becomes a little smaller,
we save some memory because only tasks that have perf_counters attached
get a perf_counter_context allocated for them, and we can remove the
inclusion of <linux/perf_counter.h> in sched.h, meaning that we don't
end up recompiling nearly everything whenever perf_counter.h changes.
The perf_counter_context structures are reference-counted and freed
when the last reference is dropped. A context can have references
from its task and the counters on its task. Counters can outlive the
task so it is possible that a context will be freed well after its
task has exited.
Contexts are allocated on fork if the parent had a context, or
otherwise the first time that a per-task counter is created on a task.
In the latter case, we set the context pointer in the task struct
locklessly using an atomic compare-and-exchange operation in case we
raced with some other task in creating a context for the subject task.
This also removes the task pointer from the perf_counter struct. The
task pointer was not used anywhere and would make it harder to move a
context from one task to another. Anything that needed to know which
task a counter was attached to was already using counter->ctx->task.
The __perf_counter_init_context function moves up in perf_counter.c
so that it can be called from find_get_context, and now initializes
the refcount, but is otherwise unchanged.
We were potentially calling list_del_counter twice: once from
__perf_counter_exit_task when the task exits and once from
__perf_counter_remove_from_context when the counter's fd gets closed.
This adds a check in list_del_counter so it doesn't do anything if
the counter has already been removed from the lists.
Since perf_counter_task_sched_in doesn't do anything if the task doesn't
have a context, and leaves cpuctx->task_ctx = NULL, this adds code to
__perf_install_in_context to set cpuctx->task_ctx if necessary, i.e. in
the case where the current task adds the first counter to itself and
thus creates a context for itself.
This also adds similar code to __perf_counter_enable to handle a
similar situation which can arise when the counters have been disabled
using prctl; that also leaves cpuctx->task_ctx = NULL.
[ Impact: refactor counter context management to prepare for new feature ]
Signed-off-by: Paul Mackerras <paulus@samba.org>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Corey Ashford <cjashfor@linux.vnet.ibm.com>
Cc: Marcelo Tosatti <mtosatti@redhat.com>
Cc: Arnaldo Carvalho de Melo <acme@redhat.com>
LKML-Reference: <18966.10075.781053.231153@cargo.ozlabs.ibm.com>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-05-22 12:17:31 +08:00
|
|
|
#endif
|
2008-05-15 19:09:15 +08:00
|
|
|
#ifdef CONFIG_NUMA
|
cpuset,mm: update tasks' mems_allowed in time
Fix allocating page cache/slab object on the unallowed node when memory
spread is set by updating tasks' mems_allowed after its cpuset's mems is
changed.
In order to update tasks' mems_allowed in time, we must modify the code of
memory policy. Because the memory policy is applied in the process's
context originally. After applying this patch, one task directly
manipulates anothers mems_allowed, and we use alloc_lock in the
task_struct to protect mems_allowed and memory policy of the task.
But in the fast path, we didn't use lock to protect them, because adding a
lock may lead to performance regression. But if we don't add a lock,the
task might see no nodes when changing cpuset's mems_allowed to some
non-overlapping set. In order to avoid it, we set all new allowed nodes,
then clear newly disallowed ones.
[lee.schermerhorn@hp.com:
The rework of mpol_new() to extract the adjusting of the node mask to
apply cpuset and mpol flags "context" breaks set_mempolicy() and mbind()
with MPOL_PREFERRED and a NULL nodemask--i.e., explicit local
allocation. Fix this by adding the check for MPOL_PREFERRED and empty
node mask to mpol_new_mpolicy().
Remove the now unneeded 'nodes = NULL' from mpol_new().
Note that mpol_new_mempolicy() is always called with a non-NULL
'nodes' parameter now that it has been removed from mpol_new().
Therefore, we don't need to test nodes for NULL before testing it for
'empty'. However, just to be extra paranoid, add a VM_BUG_ON() to
verify this assumption.]
[lee.schermerhorn@hp.com:
I don't think the function name 'mpol_new_mempolicy' is descriptive
enough to differentiate it from mpol_new().
This function applies cpuset set context, usually constraining nodes
to those allowed by the cpuset. However, when the 'RELATIVE_NODES flag
is set, it also translates the nodes. So I settled on
'mpol_set_nodemask()', because the comment block for mpol_new() mentions
that we need to call this function to "set nodes".
Some additional minor line length, whitespace and typo cleanup.]
Signed-off-by: Miao Xie <miaox@cn.fujitsu.com>
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Christoph Lameter <cl@linux-foundation.org>
Cc: Paul Menage <menage@google.com>
Cc: Nick Piggin <nickpiggin@yahoo.com.au>
Cc: Yasunori Goto <y-goto@jp.fujitsu.com>
Cc: Pekka Enberg <penberg@cs.helsinki.fi>
Cc: David Rientjes <rientjes@google.com>
Signed-off-by: Lee Schermerhorn <lee.schermerhorn@hp.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-06-17 06:31:49 +08:00
|
|
|
struct mempolicy *mempolicy; /* Protected by alloc_lock */
|
2008-05-15 19:09:15 +08:00
|
|
|
short il_next;
|
2011-03-23 07:30:44 +08:00
|
|
|
short pref_node_fork;
|
2007-10-17 14:27:30 +08:00
|
|
|
#endif
|
2012-10-25 20:16:43 +08:00
|
|
|
#ifdef CONFIG_NUMA_BALANCING
|
|
|
|
|
int numa_scan_seq;
|
|
|
|
|
int numa_migrate_seq;
|
|
|
|
|
unsigned int numa_scan_period;
|
2013-10-07 18:28:55 +08:00
|
|
|
unsigned int numa_scan_period_max;
|
2012-10-25 20:16:43 +08:00
|
|
|
u64 node_stamp; /* migration stamp */
|
|
|
|
|
struct callback_head numa_work;
|
2013-10-07 18:28:57 +08:00
|
|
|
|
2013-10-07 18:28:59 +08:00
|
|
|
/*
|
|
|
|
|
* Exponential decaying average of faults on a per-node basis.
|
|
|
|
|
* Scheduling placement decisions are made based on the these counts.
|
|
|
|
|
* The values remain static for the duration of a PTE scan
|
|
|
|
|
*/
|
2013-10-07 18:28:57 +08:00
|
|
|
unsigned long *numa_faults;
|
2013-10-07 18:28:59 +08:00
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* numa_faults_buffer records faults per node during the current
|
|
|
|
|
* scan window. When the scan completes, the counts in numa_faults
|
|
|
|
|
* decay and these values are copied.
|
|
|
|
|
*/
|
|
|
|
|
unsigned long *numa_faults_buffer;
|
|
|
|
|
|
2013-10-07 18:28:58 +08:00
|
|
|
int numa_preferred_nid;
|
2012-10-25 20:16:43 +08:00
|
|
|
#endif /* CONFIG_NUMA_BALANCING */
|
|
|
|
|
|
2006-01-08 17:01:37 +08:00
|
|
|
struct rcu_head rcu;
|
2006-04-11 19:52:07 +08:00
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* cache last used pipe for splice
|
|
|
|
|
*/
|
|
|
|
|
struct pipe_inode_info *splice_pipe;
|
net: use a per task frag allocator
We currently use a per socket order-0 page cache for tcp_sendmsg()
operations.
This page is used to build fragments for skbs.
Its done to increase probability of coalescing small write() into
single segments in skbs still in write queue (not yet sent)
But it wastes a lot of memory for applications handling many mostly
idle sockets, since each socket holds one page in sk->sk_sndmsg_page
Its also quite inefficient to build TSO 64KB packets, because we need
about 16 pages per skb on arches where PAGE_SIZE = 4096, so we hit
page allocator more than wanted.
This patch adds a per task frag allocator and uses bigger pages,
if available. An automatic fallback is done in case of memory pressure.
(up to 32768 bytes per frag, thats order-3 pages on x86)
This increases TCP stream performance by 20% on loopback device,
but also benefits on other network devices, since 8x less frags are
mapped on transmit and unmapped on tx completion. Alexander Duyck
mentioned a probable performance win on systems with IOMMU enabled.
Its possible some SG enabled hardware cant cope with bigger fragments,
but their ndo_start_xmit() should already handle this, splitting a
fragment in sub fragments, since some arches have PAGE_SIZE=65536
Successfully tested on various ethernet devices.
(ixgbe, igb, bnx2x, tg3, mellanox mlx4)
Signed-off-by: Eric Dumazet <edumazet@google.com>
Cc: Ben Hutchings <bhutchings@solarflare.com>
Cc: Vijay Subramanian <subramanian.vijay@gmail.com>
Cc: Alexander Duyck <alexander.h.duyck@intel.com>
Tested-by: Vijay Subramanian <subramanian.vijay@gmail.com>
Signed-off-by: David S. Miller <davem@davemloft.net>
2012-09-24 07:04:42 +08:00
|
|
|
|
|
|
|
|
struct page_frag task_frag;
|
|
|
|
|
|
2006-07-14 15:24:36 +08:00
|
|
|
#ifdef CONFIG_TASK_DELAY_ACCT
|
|
|
|
|
struct task_delay_info *delays;
|
2006-12-08 18:39:47 +08:00
|
|
|
#endif
|
|
|
|
|
#ifdef CONFIG_FAULT_INJECTION
|
|
|
|
|
int make_it_fail;
|
2006-07-14 15:24:36 +08:00
|
|
|
#endif
|
writeback: per task dirty rate limit
Add two fields to task_struct.
1) account dirtied pages in the individual tasks, for accuracy
2) per-task balance_dirty_pages() call intervals, for flexibility
The balance_dirty_pages() call interval (ie. nr_dirtied_pause) will
scale near-sqrt to the safety gap between dirty pages and threshold.
The main problem of per-task nr_dirtied is, if 1k+ tasks start dirtying
pages at exactly the same time, each task will be assigned a large
initial nr_dirtied_pause, so that the dirty threshold will be exceeded
long before each task reached its nr_dirtied_pause and hence call
balance_dirty_pages().
The solution is to watch for the number of pages dirtied on each CPU in
between the calls into balance_dirty_pages(). If it exceeds ratelimit_pages
(3% dirty threshold), force call balance_dirty_pages() for a chance to
set bdi->dirty_exceeded. In normal situations, this safeguarding
condition is not expected to trigger at all.
On the sqrt in dirty_poll_interval():
It will serve as an initial guess when dirty pages are still in the
freerun area.
When dirty pages are floating inside the dirty control scope [freerun,
limit], a followup patch will use some refined dirty poll interval to
get the desired pause time.
thresh-dirty (MB) sqrt
1 16
2 22
4 32
8 45
16 64
32 90
64 128
128 181
256 256
512 362
1024 512
The above table means, given 1MB (or 1GB) gap and the dd tasks polling
balance_dirty_pages() on every 16 (or 512) pages, the dirty limit won't
be exceeded as long as there are less than 16 (or 512) concurrent dd's.
So sqrt naturally leads to less overheads and more safe concurrent tasks
for large memory servers, which have large (thresh-freerun) gaps.
peter: keep the per-CPU ratelimit for safeguarding the 1k+ tasks case
CC: Peter Zijlstra <a.p.zijlstra@chello.nl>
Reviewed-by: Andrea Righi <andrea@betterlinux.com>
Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-06-12 08:10:12 +08:00
|
|
|
/*
|
|
|
|
|
* when (nr_dirtied >= nr_dirtied_pause), it's time to call
|
|
|
|
|
* balance_dirty_pages() for some dirty throttling pause
|
|
|
|
|
*/
|
|
|
|
|
int nr_dirtied;
|
|
|
|
|
int nr_dirtied_pause;
|
2011-06-12 09:25:42 +08:00
|
|
|
unsigned long dirty_paused_when; /* start of a write-and-pause period */
|
writeback: per task dirty rate limit
Add two fields to task_struct.
1) account dirtied pages in the individual tasks, for accuracy
2) per-task balance_dirty_pages() call intervals, for flexibility
The balance_dirty_pages() call interval (ie. nr_dirtied_pause) will
scale near-sqrt to the safety gap between dirty pages and threshold.
The main problem of per-task nr_dirtied is, if 1k+ tasks start dirtying
pages at exactly the same time, each task will be assigned a large
initial nr_dirtied_pause, so that the dirty threshold will be exceeded
long before each task reached its nr_dirtied_pause and hence call
balance_dirty_pages().
The solution is to watch for the number of pages dirtied on each CPU in
between the calls into balance_dirty_pages(). If it exceeds ratelimit_pages
(3% dirty threshold), force call balance_dirty_pages() for a chance to
set bdi->dirty_exceeded. In normal situations, this safeguarding
condition is not expected to trigger at all.
On the sqrt in dirty_poll_interval():
It will serve as an initial guess when dirty pages are still in the
freerun area.
When dirty pages are floating inside the dirty control scope [freerun,
limit], a followup patch will use some refined dirty poll interval to
get the desired pause time.
thresh-dirty (MB) sqrt
1 16
2 22
4 32
8 45
16 64
32 90
64 128
128 181
256 256
512 362
1024 512
The above table means, given 1MB (or 1GB) gap and the dd tasks polling
balance_dirty_pages() on every 16 (or 512) pages, the dirty limit won't
be exceeded as long as there are less than 16 (or 512) concurrent dd's.
So sqrt naturally leads to less overheads and more safe concurrent tasks
for large memory servers, which have large (thresh-freerun) gaps.
peter: keep the per-CPU ratelimit for safeguarding the 1k+ tasks case
CC: Peter Zijlstra <a.p.zijlstra@chello.nl>
Reviewed-by: Andrea Righi <andrea@betterlinux.com>
Signed-off-by: Wu Fengguang <fengguang.wu@intel.com>
2011-06-12 08:10:12 +08:00
|
|
|
|
2008-01-26 04:08:34 +08:00
|
|
|
#ifdef CONFIG_LATENCYTOP
|
|
|
|
|
int latency_record_count;
|
|
|
|
|
struct latency_record latency_record[LT_SAVECOUNT];
|
|
|
|
|
#endif
|
2008-09-02 06:52:40 +08:00
|
|
|
/*
|
|
|
|
|
* time slack values; these are used to round up poll() and
|
|
|
|
|
* select() etc timeout values. These are in nanoseconds.
|
|
|
|
|
*/
|
|
|
|
|
unsigned long timer_slack_ns;
|
|
|
|
|
unsigned long default_timer_slack_ns;
|
2008-11-06 16:37:40 +08:00
|
|
|
|
2008-11-26 04:07:04 +08:00
|
|
|
#ifdef CONFIG_FUNCTION_GRAPH_TRACER
|
tree-wide: Assorted spelling fixes
In particular, several occurances of funny versions of 'success',
'unknown', 'therefore', 'acknowledge', 'argument', 'achieve', 'address',
'beginning', 'desirable', 'separate' and 'necessary' are fixed.
Signed-off-by: Daniel Mack <daniel@caiaq.de>
Cc: Joe Perches <joe@perches.com>
Cc: Junio C Hamano <gitster@pobox.com>
Signed-off-by: Jiri Kosina <jkosina@suse.cz>
2010-02-03 08:01:28 +08:00
|
|
|
/* Index of current stored address in ret_stack */
|
2008-11-23 13:22:56 +08:00
|
|
|
int curr_ret_stack;
|
|
|
|
|
/* Stack of return addresses for return function tracing */
|
|
|
|
|
struct ftrace_ret_stack *ret_stack;
|
2009-03-24 13:10:15 +08:00
|
|
|
/* time stamp for last schedule */
|
|
|
|
|
unsigned long long ftrace_timestamp;
|
2008-11-23 13:22:56 +08:00
|
|
|
/*
|
|
|
|
|
* Number of functions that haven't been traced
|
|
|
|
|
* because of depth overrun.
|
|
|
|
|
*/
|
|
|
|
|
atomic_t trace_overrun;
|
2008-12-06 10:43:41 +08:00
|
|
|
/* Pause for the tracing */
|
|
|
|
|
atomic_t tracing_graph_pause;
|
2008-11-23 13:22:56 +08:00
|
|
|
#endif
|
2008-12-04 04:36:57 +08:00
|
|
|
#ifdef CONFIG_TRACING
|
|
|
|
|
/* state flags for use by tracers */
|
|
|
|
|
unsigned long trace;
|
ftrace: Add internal recursive checks
Witold reported a reboot caused by the selftests of the dynamic function
tracer. He sent me a config and I used ktest to do a config_bisect on it
(as my config did not cause the crash). It pointed out that the problem
config was CONFIG_PROVE_RCU.
What happened was that if multiple callbacks are attached to the
function tracer, we iterate a list of callbacks. Because the list is
managed by synchronize_sched() and preempt_disable, the access to the
pointers uses rcu_dereference_raw().
When PROVE_RCU is enabled, the rcu_dereference_raw() calls some
debugging functions, which happen to be traced. The tracing of the debug
function would then call rcu_dereference_raw() which would then call the
debug function and then... well you get the idea.
I first wrote two different patches to solve this bug.
1) add a __rcu_dereference_raw() that would not do any checks.
2) add notrace to the offending debug functions.
Both of these patches worked.
Talking with Paul McKenney on IRC, he suggested to add recursion
detection instead. This seemed to be a better solution, so I decided to
implement it. As the task_struct already has a trace_recursion to detect
recursion in the ring buffer, and that has a very small number it
allows, I decided to use that same variable to add flags that can detect
the recursion inside the infrastructure of the function tracer.
I plan to change it so that the task struct bit can be checked in
mcount, but as that requires changes to all archs, I will hold that off
to the next merge window.
Cc: Ingo Molnar <mingo@elte.hu>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Frederic Weisbecker <fweisbec@gmail.com>
Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Link: http://lkml.kernel.org/r/1306348063.1465.116.camel@gandalf.stny.rr.com
Reported-by: Witold Baryluk <baryluk@smp.if.uj.edu.pl>
Signed-off-by: Steven Rostedt <rostedt@goodmis.org>
2011-05-26 02:27:43 +08:00
|
|
|
/* bitmask and counter of trace recursion */
|
tracing: add same level recursion detection
The tracing infrastructure allows for recursion. That is, an interrupt
may interrupt the act of tracing an event, and that interrupt may very well
perform its own trace. This is a recursive trace, and is fine to do.
The problem arises when there is a bug, and the utility doing the trace
calls something that recurses back into the tracer. This recursion is not
caused by an external event like an interrupt, but by code that is not
expected to recurse. The result could be a lockup.
This patch adds a bitmask to the task structure that keeps track
of the trace recursion. To find the interrupt depth, the following
algorithm is used:
level = hardirq_count() + softirq_count() + in_nmi;
Here, level will be the depth of interrutps and softirqs, and even handles
the nmi. Then the corresponding bit is set in the recursion bitmask.
If the bit was already set, we know we had a recursion at the same level
and we warn about it and fail the writing to the buffer.
After the data has been committed to the buffer, we clear the bit.
No atomics are needed. The only races are with interrupts and they reset
the bitmask before returning anywy.
[ Impact: detect same irq level trace recursion ]
Signed-off-by: Steven Rostedt <rostedt@goodmis.org>
2009-04-17 09:41:52 +08:00
|
|
|
unsigned long trace_recursion;
|
|
|
|
|
#endif /* CONFIG_TRACING */
|
2012-08-01 07:43:02 +08:00
|
|
|
#ifdef CONFIG_MEMCG /* memcg uses this to do batch job */
|
2009-12-16 08:47:03 +08:00
|
|
|
struct memcg_batch_info {
|
|
|
|
|
int do_batch; /* incremented when batch uncharge started */
|
|
|
|
|
struct mem_cgroup *memcg; /* target memcg of uncharge */
|
2011-03-24 07:42:35 +08:00
|
|
|
unsigned long nr_pages; /* uncharged usage */
|
|
|
|
|
unsigned long memsw_nr_pages; /* uncharged mem+swap usage */
|
2009-12-16 08:47:03 +08:00
|
|
|
} memcg_batch;
|
2012-12-19 06:22:42 +08:00
|
|
|
unsigned int memcg_kmem_skip_account;
|
2013-09-13 06:13:42 +08:00
|
|
|
struct memcg_oom_info {
|
|
|
|
|
unsigned int may_oom:1;
|
mm: memcg: do not trap chargers with full callstack on OOM
The memcg OOM handling is incredibly fragile and can deadlock. When a
task fails to charge memory, it invokes the OOM killer and loops right
there in the charge code until it succeeds. Comparably, any other task
that enters the charge path at this point will go to a waitqueue right
then and there and sleep until the OOM situation is resolved. The problem
is that these tasks may hold filesystem locks and the mmap_sem; locks that
the selected OOM victim may need to exit.
For example, in one reported case, the task invoking the OOM killer was
about to charge a page cache page during a write(), which holds the
i_mutex. The OOM killer selected a task that was just entering truncate()
and trying to acquire the i_mutex:
OOM invoking task:
mem_cgroup_handle_oom+0x241/0x3b0
mem_cgroup_cache_charge+0xbe/0xe0
add_to_page_cache_locked+0x4c/0x140
add_to_page_cache_lru+0x22/0x50
grab_cache_page_write_begin+0x8b/0xe0
ext3_write_begin+0x88/0x270
generic_file_buffered_write+0x116/0x290
__generic_file_aio_write+0x27c/0x480
generic_file_aio_write+0x76/0xf0 # takes ->i_mutex
do_sync_write+0xea/0x130
vfs_write+0xf3/0x1f0
sys_write+0x51/0x90
system_call_fastpath+0x18/0x1d
OOM kill victim:
do_truncate+0x58/0xa0 # takes i_mutex
do_last+0x250/0xa30
path_openat+0xd7/0x440
do_filp_open+0x49/0xa0
do_sys_open+0x106/0x240
sys_open+0x20/0x30
system_call_fastpath+0x18/0x1d
The OOM handling task will retry the charge indefinitely while the OOM
killed task is not releasing any resources.
A similar scenario can happen when the kernel OOM killer for a memcg is
disabled and a userspace task is in charge of resolving OOM situations.
In this case, ALL tasks that enter the OOM path will be made to sleep on
the OOM waitqueue and wait for userspace to free resources or increase
the group's limit. But a userspace OOM handler is prone to deadlock
itself on the locks held by the waiting tasks. For example one of the
sleeping tasks may be stuck in a brk() call with the mmap_sem held for
writing but the userspace handler, in order to pick an optimal victim,
may need to read files from /proc/<pid>, which tries to acquire the same
mmap_sem for reading and deadlocks.
This patch changes the way tasks behave after detecting a memcg OOM and
makes sure nobody loops or sleeps with locks held:
1. When OOMing in a user fault, invoke the OOM killer and restart the
fault instead of looping on the charge attempt. This way, the OOM
victim can not get stuck on locks the looping task may hold.
2. When OOMing in a user fault but somebody else is handling it
(either the kernel OOM killer or a userspace handler), don't go to
sleep in the charge context. Instead, remember the OOMing memcg in
the task struct and then fully unwind the page fault stack with
-ENOMEM. pagefault_out_of_memory() will then call back into the
memcg code to check if the -ENOMEM came from the memcg, and then
either put the task to sleep on the memcg's OOM waitqueue or just
restart the fault. The OOM victim can no longer get stuck on any
lock a sleeping task may hold.
Debugged by Michal Hocko.
Signed-off-by: Johannes Weiner <hannes@cmpxchg.org>
Reported-by: azurIt <azurit@pobox.sk>
Acked-by: Michal Hocko <mhocko@suse.cz>
Cc: David Rientjes <rientjes@google.com>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-09-13 06:13:44 +08:00
|
|
|
unsigned int in_memcg_oom:1;
|
|
|
|
|
unsigned int oom_locked:1;
|
|
|
|
|
int wakeups;
|
|
|
|
|
struct mem_cgroup *wait_on_memcg;
|
2013-09-13 06:13:42 +08:00
|
|
|
} memcg_oom;
|
2009-12-16 08:47:03 +08:00
|
|
|
#endif
|
uprobes/core: Handle breakpoint and singlestep exceptions
Uprobes uses exception notifiers to get to know if a thread hit
a breakpoint or a singlestep exception.
When a thread hits a uprobe or is singlestepping post a uprobe
hit, the uprobe exception notifier sets its TIF_UPROBE bit,
which will then be checked on its return to userspace path
(do_notify_resume() ->uprobe_notify_resume()), where the
consumers handlers are run (in task context) based on the
defined filters.
Uprobe hits are thread specific and hence we need to maintain
information about if a task hit a uprobe, what uprobe was hit,
the slot where the original instruction was copied for xol so
that it can be singlestepped with appropriate fixups.
In some cases, special care is needed for instructions that are
executed out of line (xol). These are architecture specific
artefacts, such as handling RIP relative instructions on x86_64.
Since the instruction at which the uprobe was inserted is
executed out of line, architecture specific fixups are added so
that the thread continues normal execution in the presence of a
uprobe.
Postpone the signals until we execute the probed insn.
post_xol() path does a recalc_sigpending() before return to
user-mode, this ensures the signal can't be lost.
Uprobes relies on DIE_DEBUG notification to notify if a
singlestep is complete.
Adds x86 specific uprobe exception notifiers and appropriate
hooks needed to determine a uprobe hit and subsequent post
processing.
Add requisite x86 fixups for xol for uprobes. Specific cases
needing fixups include relative jumps (x86_64), calls, etc.
Where possible, we check and skip singlestepping the
breakpointed instructions. For now we skip single byte as well
as few multibyte nop instructions. However this can be extended
to other instructions too.
Credits to Oleg Nesterov for suggestions/patches related to
signal, breakpoint, singlestep handling code.
Signed-off-by: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Ananth N Mavinakayanahalli <ananth@in.ibm.com>
Cc: Jim Keniston <jkenisto@linux.vnet.ibm.com>
Cc: Linux-mm <linux-mm@kvack.org>
Cc: Oleg Nesterov <oleg@redhat.com>
Cc: Andi Kleen <andi@firstfloor.org>
Cc: Christoph Hellwig <hch@infradead.org>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Arnaldo Carvalho de Melo <acme@infradead.org>
Cc: Masami Hiramatsu <masami.hiramatsu.pt@hitachi.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/20120313180011.29771.89027.sendpatchset@srdronam.in.ibm.com
[ Performed various cleanliness edits ]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2012-03-14 02:00:11 +08:00
|
|
|
#ifdef CONFIG_UPROBES
|
|
|
|
|
struct uprobe_task *utask;
|
|
|
|
|
#endif
|
2013-03-24 07:11:31 +08:00
|
|
|
#if defined(CONFIG_BCACHE) || defined(CONFIG_BCACHE_MODULE)
|
|
|
|
|
unsigned int sequential_io;
|
|
|
|
|
unsigned int sequential_io_avg;
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
};
|
|
|
|
|
|
2009-03-13 04:35:43 +08:00
|
|
|
/* Future-safe accessor for struct task_struct's cpus_allowed. */
|
2009-12-18 01:43:29 +08:00
|
|
|
#define tsk_cpus_allowed(tsk) (&(tsk)->cpus_allowed)
|
2009-03-13 04:35:43 +08:00
|
|
|
|
2012-10-25 20:16:43 +08:00
|
|
|
#ifdef CONFIG_NUMA_BALANCING
|
2012-11-21 09:18:23 +08:00
|
|
|
extern void task_numa_fault(int node, int pages, bool migrated);
|
2012-11-22 19:16:36 +08:00
|
|
|
extern void set_numabalancing_state(bool enabled);
|
2012-10-25 20:16:43 +08:00
|
|
|
#else
|
2012-11-21 09:18:23 +08:00
|
|
|
static inline void task_numa_fault(int node, int pages, bool migrated)
|
2012-10-25 20:16:43 +08:00
|
|
|
{
|
|
|
|
|
}
|
2012-11-22 19:16:36 +08:00
|
|
|
static inline void set_numabalancing_state(bool enabled)
|
|
|
|
|
{
|
|
|
|
|
}
|
2012-10-25 20:16:43 +08:00
|
|
|
#endif
|
|
|
|
|
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline struct pid *task_pid(struct task_struct *task)
|
2006-10-02 17:17:09 +08:00
|
|
|
{
|
|
|
|
|
return task->pids[PIDTYPE_PID].pid;
|
|
|
|
|
}
|
|
|
|
|
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline struct pid *task_tgid(struct task_struct *task)
|
2006-10-02 17:17:09 +08:00
|
|
|
{
|
|
|
|
|
return task->group_leader->pids[PIDTYPE_PID].pid;
|
|
|
|
|
}
|
|
|
|
|
|
2009-04-03 07:58:35 +08:00
|
|
|
/*
|
|
|
|
|
* Without tasklist or rcu lock it is not safe to dereference
|
|
|
|
|
* the result of task_pgrp/task_session even if task == current,
|
|
|
|
|
* we can race with another thread doing sys_setsid/sys_setpgid.
|
|
|
|
|
*/
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline struct pid *task_pgrp(struct task_struct *task)
|
2006-10-02 17:17:09 +08:00
|
|
|
{
|
|
|
|
|
return task->group_leader->pids[PIDTYPE_PGID].pid;
|
|
|
|
|
}
|
|
|
|
|
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline struct pid *task_session(struct task_struct *task)
|
2006-10-02 17:17:09 +08:00
|
|
|
{
|
|
|
|
|
return task->group_leader->pids[PIDTYPE_SID].pid;
|
|
|
|
|
}
|
|
|
|
|
|
2007-10-19 14:40:06 +08:00
|
|
|
struct pid_namespace;
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* the helpers to get the task's different pids as they are seen
|
|
|
|
|
* from various namespaces
|
|
|
|
|
*
|
|
|
|
|
* task_xid_nr() : global id, i.e. the id seen from the init namespace;
|
2008-02-08 20:19:15 +08:00
|
|
|
* task_xid_vnr() : virtual id, i.e. the id seen from the pid namespace of
|
|
|
|
|
* current.
|
2007-10-19 14:40:06 +08:00
|
|
|
* task_xid_nr_ns() : id seen from the ns specified;
|
|
|
|
|
*
|
|
|
|
|
* set_task_vxid() : assigns a virtual id to a task;
|
|
|
|
|
*
|
|
|
|
|
* see also pid_nr() etc in include/linux/pid.h
|
|
|
|
|
*/
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
pid_t __task_pid_nr_ns(struct task_struct *task, enum pid_type type,
|
|
|
|
|
struct pid_namespace *ns);
|
2007-10-19 14:40:06 +08:00
|
|
|
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline pid_t task_pid_nr(struct task_struct *tsk)
|
2007-10-19 14:40:06 +08:00
|
|
|
{
|
|
|
|
|
return tsk->pid;
|
|
|
|
|
}
|
|
|
|
|
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
static inline pid_t task_pid_nr_ns(struct task_struct *tsk,
|
|
|
|
|
struct pid_namespace *ns)
|
|
|
|
|
{
|
|
|
|
|
return __task_pid_nr_ns(tsk, PIDTYPE_PID, ns);
|
|
|
|
|
}
|
2007-10-19 14:40:06 +08:00
|
|
|
|
|
|
|
|
static inline pid_t task_pid_vnr(struct task_struct *tsk)
|
|
|
|
|
{
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
return __task_pid_nr_ns(tsk, PIDTYPE_PID, NULL);
|
2007-10-19 14:40:06 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline pid_t task_tgid_nr(struct task_struct *tsk)
|
2007-10-19 14:40:06 +08:00
|
|
|
{
|
|
|
|
|
return tsk->tgid;
|
|
|
|
|
}
|
|
|
|
|
|
2007-10-19 14:40:19 +08:00
|
|
|
pid_t task_tgid_nr_ns(struct task_struct *tsk, struct pid_namespace *ns);
|
2007-10-19 14:40:06 +08:00
|
|
|
|
|
|
|
|
static inline pid_t task_tgid_vnr(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
return pid_vnr(task_tgid(tsk));
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
static inline pid_t task_pgrp_nr_ns(struct task_struct *tsk,
|
|
|
|
|
struct pid_namespace *ns)
|
2007-10-19 14:40:06 +08:00
|
|
|
{
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
return __task_pid_nr_ns(tsk, PIDTYPE_PGID, ns);
|
2007-10-19 14:40:06 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline pid_t task_pgrp_vnr(struct task_struct *tsk)
|
|
|
|
|
{
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
return __task_pid_nr_ns(tsk, PIDTYPE_PGID, NULL);
|
2007-10-19 14:40:06 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
static inline pid_t task_session_nr_ns(struct task_struct *tsk,
|
|
|
|
|
struct pid_namespace *ns)
|
2007-10-19 14:40:06 +08:00
|
|
|
{
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
return __task_pid_nr_ns(tsk, PIDTYPE_SID, ns);
|
2007-10-19 14:40:06 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline pid_t task_session_vnr(struct task_struct *tsk)
|
|
|
|
|
{
|
pids: refactor vnr/nr_ns helpers to make them safe
Inho, the safety rules for vnr/nr_ns helpers are horrible and buggy.
task_pid_nr_ns(task) needs rcu/tasklist depending on task == current.
As for "special" pids, vnr/nr_ns helpers always need rcu. However, if
task != current, they are unsafe even under rcu lock, we can't trust
task->group_leader without the special checks.
And almost every helper has a callsite which needs a fix.
Also, it is a bit annoying that the implementations of, say,
task_pgrp_vnr() and task_pgrp_nr_ns() are not "symmetrical".
This patch introduces the new helper, __task_pid_nr_ns(), which is always
safe to use, and turns all other helpers into the trivial wrappers.
After this I'll send another patch which converts task_tgid_xxx() as well,
they're are a bit special.
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Louis Rilling <Louis.Rilling@kerlabs.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Pavel Emelyanov <xemul@openvz.org>
Cc: Sukadev Bhattiprolu <sukadev@linux.vnet.ibm.com>
Cc: Roland McGrath <roland@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-04-03 07:58:38 +08:00
|
|
|
return __task_pid_nr_ns(tsk, PIDTYPE_SID, NULL);
|
2007-10-19 14:40:06 +08:00
|
|
|
}
|
|
|
|
|
|
2009-04-03 07:58:39 +08:00
|
|
|
/* obsolete, do not use */
|
|
|
|
|
static inline pid_t task_pgrp_nr(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
return task_pgrp_nr_ns(tsk, &init_pid_ns);
|
|
|
|
|
}
|
2007-10-19 14:40:06 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/**
|
|
|
|
|
* pid_alive - check that a task structure is not stale
|
|
|
|
|
* @p: Task structure to be checked.
|
|
|
|
|
*
|
|
|
|
|
* Test if a process is not yet dead (at most zombie state)
|
|
|
|
|
* If pid_alive fails, then pointers within the task structure
|
|
|
|
|
* can be stale and must not be dereferenced.
|
2013-07-13 02:45:47 +08:00
|
|
|
*
|
|
|
|
|
* Return: 1 if the process is alive. 0 otherwise.
|
2005-04-17 06:20:36 +08:00
|
|
|
*/
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline int pid_alive(struct task_struct *p)
|
2005-04-17 06:20:36 +08:00
|
|
|
{
|
[PATCH] pidhash: Refactor the pid hash table
Simplifies the code, reduces the need for 4 pid hash tables, and makes the
code more capable.
In the discussions I had with Oleg it was felt that to a large extent the
cleanup itself justified the work. With struct pid being dynamically
allocated meant we could create the hash table entry when the pid was
allocated and free the hash table entry when the pid was freed. Instead of
playing with the hash lists when ever a process would attach or detach to a
process.
For myself the fact that it gave what my previous task_ref patch gave for free
with simpler code was a big win. The problem is that if you hold a reference
to struct task_struct you lock in 10K of low memory. If you do that in a user
controllable way like /proc does, with an unprivileged but hostile user space
application with typical resource limits of 1000 fds and 100 processes I can
trigger the OOM killer by consuming all of low memory with task structs, on a
machine wight 1GB of low memory.
If I instead hold a reference to struct pid which holds a pointer to my
task_struct, I don't suffer from that problem because struct pid is 2 orders
of magnitude smaller. In fact struct pid is small enough that most other
kernel data structures dwarf it, so simply limiting the number of referring
data structures is enough to prevent exhaustion of low memory.
This splits the current struct pid into two structures, struct pid and struct
pid_link, and reduces our number of hash tables from PIDTYPE_MAX to just one.
struct pid_link is the per process linkage into the hash tables and lives in
struct task_struct. struct pid is given an indepedent lifetime, and holds
pointers to each of the pid types.
The independent life of struct pid simplifies attach_pid, and detach_pid,
because we are always manipulating the list of pids and not the hash table.
In addition in giving struct pid an indpendent life it makes the concept much
more powerful.
Kernel data structures can now embed a struct pid * instead of a pid_t and
not suffer from pid wrap around problems or from keeping unnecessarily
large amounts of memory allocated.
Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-03-31 18:31:42 +08:00
|
|
|
return p->pids[PIDTYPE_PID].pid != NULL;
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
2006-09-29 17:00:07 +08:00
|
|
|
/**
|
2007-10-19 14:39:52 +08:00
|
|
|
* is_global_init - check if a task structure is init
|
2006-10-06 15:44:01 +08:00
|
|
|
* @tsk: Task structure to be checked.
|
|
|
|
|
*
|
|
|
|
|
* Check if a task structure is the first user space task the kernel created.
|
2013-07-13 02:45:47 +08:00
|
|
|
*
|
|
|
|
|
* Return: 1 if the task structure is init. 0 otherwise.
|
2007-10-19 14:39:52 +08:00
|
|
|
*/
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline int is_global_init(struct task_struct *tsk)
|
2007-10-19 14:40:09 +08:00
|
|
|
{
|
|
|
|
|
return tsk->pid == 1;
|
|
|
|
|
}
|
2007-10-19 14:39:52 +08:00
|
|
|
|
2006-10-02 17:19:00 +08:00
|
|
|
extern struct pid *cad_pid;
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void free_task(struct task_struct *tsk);
|
|
|
|
|
#define get_task_struct(tsk) do { atomic_inc(&(tsk)->usage); } while(0)
|
2006-01-08 17:01:37 +08:00
|
|
|
|
2006-03-31 18:31:34 +08:00
|
|
|
extern void __put_task_struct(struct task_struct *t);
|
2006-01-08 17:01:37 +08:00
|
|
|
|
|
|
|
|
static inline void put_task_struct(struct task_struct *t)
|
|
|
|
|
{
|
|
|
|
|
if (atomic_dec_and_test(&t->usage))
|
2006-03-31 18:31:37 +08:00
|
|
|
__put_task_struct(t);
|
2006-01-08 17:01:37 +08:00
|
|
|
}
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2012-12-17 03:00:34 +08:00
|
|
|
#ifdef CONFIG_VIRT_CPU_ACCOUNTING_GEN
|
|
|
|
|
extern void task_cputime(struct task_struct *t,
|
|
|
|
|
cputime_t *utime, cputime_t *stime);
|
|
|
|
|
extern void task_cputime_scaled(struct task_struct *t,
|
|
|
|
|
cputime_t *utimescaled, cputime_t *stimescaled);
|
|
|
|
|
extern cputime_t task_gtime(struct task_struct *t);
|
|
|
|
|
#else
|
2012-11-13 21:20:55 +08:00
|
|
|
static inline void task_cputime(struct task_struct *t,
|
|
|
|
|
cputime_t *utime, cputime_t *stime)
|
|
|
|
|
{
|
|
|
|
|
if (utime)
|
|
|
|
|
*utime = t->utime;
|
|
|
|
|
if (stime)
|
|
|
|
|
*stime = t->stime;
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void task_cputime_scaled(struct task_struct *t,
|
|
|
|
|
cputime_t *utimescaled,
|
|
|
|
|
cputime_t *stimescaled)
|
|
|
|
|
{
|
|
|
|
|
if (utimescaled)
|
|
|
|
|
*utimescaled = t->utimescaled;
|
|
|
|
|
if (stimescaled)
|
|
|
|
|
*stimescaled = t->stimescaled;
|
|
|
|
|
}
|
2012-12-17 03:00:34 +08:00
|
|
|
|
|
|
|
|
static inline cputime_t task_gtime(struct task_struct *t)
|
|
|
|
|
{
|
|
|
|
|
return t->gtime;
|
|
|
|
|
}
|
|
|
|
|
#endif
|
2012-11-21 23:26:44 +08:00
|
|
|
extern void task_cputime_adjusted(struct task_struct *p, cputime_t *ut, cputime_t *st);
|
|
|
|
|
extern void thread_group_cputime_adjusted(struct task_struct *p, cputime_t *ut, cputime_t *st);
|
2008-09-06 00:12:23 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
|
|
|
|
* Per process flags
|
|
|
|
|
*/
|
|
|
|
|
#define PF_EXITING 0x00000004 /* getting shut down */
|
2007-06-09 04:47:00 +08:00
|
|
|
#define PF_EXITPIDONE 0x00000008 /* pi exit done on shut down */
|
2007-10-15 23:00:19 +08:00
|
|
|
#define PF_VCPU 0x00000010 /* I'm a virtual CPU */
|
2010-06-09 03:40:37 +08:00
|
|
|
#define PF_WQ_WORKER 0x00000020 /* I'm a workqueue worker */
|
2005-04-17 06:20:36 +08:00
|
|
|
#define PF_FORKNOEXEC 0x00000040 /* forked but didn't exec */
|
2009-09-16 17:50:14 +08:00
|
|
|
#define PF_MCE_PROCESS 0x00000080 /* process policy on mce errors */
|
2005-04-17 06:20:36 +08:00
|
|
|
#define PF_SUPERPRIV 0x00000100 /* used super-user privileges */
|
|
|
|
|
#define PF_DUMPCORE 0x00000200 /* dumped core */
|
|
|
|
|
#define PF_SIGNALED 0x00000400 /* killed by a signal */
|
|
|
|
|
#define PF_MEMALLOC 0x00000800 /* Allocating memory */
|
2011-08-08 23:02:04 +08:00
|
|
|
#define PF_NPROC_EXCEEDED 0x00001000 /* set_user noticed that RLIMIT_NPROC was exceeded */
|
2005-04-17 06:20:36 +08:00
|
|
|
#define PF_USED_MATH 0x00002000 /* if unset the fpu must be initialized before use */
|
2013-01-16 10:52:51 +08:00
|
|
|
#define PF_USED_ASYNC 0x00004000 /* used async_schedule*(), used by module init */
|
2005-04-17 06:20:36 +08:00
|
|
|
#define PF_NOFREEZE 0x00008000 /* this thread should not be frozen */
|
|
|
|
|
#define PF_FROZEN 0x00010000 /* frozen for system suspend */
|
|
|
|
|
#define PF_FSTRANS 0x00020000 /* inside a filesystem transaction */
|
|
|
|
|
#define PF_KSWAPD 0x00040000 /* I am kswapd */
|
mm: teach mm by current context info to not do I/O during memory allocation
This patch introduces PF_MEMALLOC_NOIO on process flag('flags' field of
'struct task_struct'), so that the flag can be set by one task to avoid
doing I/O inside memory allocation in the task's context.
The patch trys to solve one deadlock problem caused by block device, and
the problem may happen at least in the below situations:
- during block device runtime resume, if memory allocation with
GFP_KERNEL is called inside runtime resume callback of any one of its
ancestors(or the block device itself), the deadlock may be triggered
inside the memory allocation since it might not complete until the block
device becomes active and the involed page I/O finishes. The situation
is pointed out first by Alan Stern. It is not a good approach to
convert all GFP_KERNEL[1] in the path into GFP_NOIO because several
subsystems may be involved(for example, PCI, USB and SCSI may be
involved for usb mass stoarage device, network devices involved too in
the iSCSI case)
- during block device runtime suspend, because runtime resume need to
wait for completion of concurrent runtime suspend.
- during error handling of usb mass storage deivce, USB bus reset will
be put on the device, so there shouldn't have any memory allocation with
GFP_KERNEL during USB bus reset, otherwise the deadlock similar with
above may be triggered. Unfortunately, any usb device may include one
mass storage interface in theory, so it requires all usb interface
drivers to handle the situation. In fact, most usb drivers don't know
how to handle bus reset on the device and don't provide .pre_set() and
.post_reset() callback at all, so USB core has to unbind and bind driver
for these devices. So it is still not practical to resort to GFP_NOIO
for solving the problem.
Also the introduced solution can be used by block subsystem or block
drivers too, for example, set the PF_MEMALLOC_NOIO flag before doing
actual I/O transfer.
It is not a good idea to convert all these GFP_KERNEL in the affected
path into GFP_NOIO because these functions doing that may be implemented
as library and will be called in many other contexts.
In fact, memalloc_noio_flags() can convert some of current static
GFP_NOIO allocation into GFP_KERNEL back in other non-affected contexts,
at least almost all GFP_NOIO in USB subsystem can be converted into
GFP_KERNEL after applying the approach and make allocation with GFP_NOIO
only happen in runtime resume/bus reset/block I/O transfer contexts
generally.
[1], several GFP_KERNEL allocation examples in runtime resume path
- pci subsystem
acpi_os_allocate
<-acpi_ut_allocate
<-ACPI_ALLOCATE_ZEROED
<-acpi_evaluate_object
<-__acpi_bus_set_power
<-acpi_bus_set_power
<-acpi_pci_set_power_state
<-platform_pci_set_power_state
<-pci_platform_power_transition
<-__pci_complete_power_transition
<-pci_set_power_state
<-pci_restore_standard_config
<-pci_pm_runtime_resume
- usb subsystem
usb_get_status
<-finish_port_resume
<-usb_port_resume
<-generic_resume
<-usb_resume_device
<-usb_resume_both
<-usb_runtime_resume
- some individual usb drivers
usblp, uvc, gspca, most of dvb-usb-v2 media drivers, cpia2, az6007, ....
That is just what I have found. Unfortunately, this allocation can only
be found by human being now, and there should be many not found since
any function in the resume path(call tree) may allocate memory with
GFP_KERNEL.
Signed-off-by: Ming Lei <ming.lei@canonical.com>
Signed-off-by: Minchan Kim <minchan@kernel.org>
Cc: Alan Stern <stern@rowland.harvard.edu>
Cc: Oliver Neukum <oneukum@suse.de>
Cc: Jiri Kosina <jiri.kosina@suse.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Michal Hocko <mhocko@suse.cz>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: "Rafael J. Wysocki" <rjw@sisk.pl>
Cc: Greg KH <greg@kroah.com>
Cc: Jens Axboe <axboe@kernel.dk>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Eric Dumazet <eric.dumazet@gmail.com>
Cc: David Decotigny <david.decotigny@google.com>
Cc: Tom Herbert <therbert@google.com>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-02-23 08:34:08 +08:00
|
|
|
#define PF_MEMALLOC_NOIO 0x00080000 /* Allocating memory without IO involved */
|
2005-04-17 06:20:36 +08:00
|
|
|
#define PF_LESS_THROTTLE 0x00100000 /* Throttle me less: I clean memory */
|
2008-07-25 16:47:38 +08:00
|
|
|
#define PF_KTHREAD 0x00200000 /* I am a kernel thread */
|
2006-06-13 14:26:10 +08:00
|
|
|
#define PF_RANDOMIZE 0x00400000 /* randomize virtual address space */
|
|
|
|
|
#define PF_SWAPWRITE 0x00800000 /* Allowed to write to swap */
|
|
|
|
|
#define PF_SPREAD_PAGE 0x01000000 /* Spread page cache over cpuset */
|
|
|
|
|
#define PF_SPREAD_SLAB 0x02000000 /* Spread some slab caches over cpuset */
|
2013-03-20 04:45:20 +08:00
|
|
|
#define PF_NO_SETAFFINITY 0x04000000 /* Userland is not allowed to meddle with cpus_allowed */
|
2009-09-16 17:50:14 +08:00
|
|
|
#define PF_MCE_EARLY 0x08000000 /* Early kill for mce process policy */
|
2006-03-24 19:16:08 +08:00
|
|
|
#define PF_MEMPOLICY 0x10000000 /* Non-default NUMA mempolicy */
|
2006-06-27 17:54:56 +08:00
|
|
|
#define PF_MUTEX_TESTER 0x20000000 /* Thread belongs to the rt mutex tester */
|
2011-02-16 16:25:31 +08:00
|
|
|
#define PF_FREEZER_SKIP 0x40000000 /* Freezer should not count it as freezable */
|
2013-07-25 08:41:33 +08:00
|
|
|
#define PF_SUSPEND_TASK 0x80000000 /* this thread called freeze_processes and should not be frozen */
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Only the _current_ task can read/write to tsk->flags, but other
|
|
|
|
|
* tasks can access tsk->flags in readonly mode for example
|
|
|
|
|
* with tsk_used_math (like during threaded core dumping).
|
|
|
|
|
* There is however an exception to this rule during ptrace
|
|
|
|
|
* or during fork: the ptracer task is allowed to write to the
|
|
|
|
|
* child->flags of its traced child (same goes for fork, the parent
|
|
|
|
|
* can write to the child->flags), because we're guaranteed the
|
|
|
|
|
* child is not running and in turn not changing child->flags
|
|
|
|
|
* at the same time the parent does it.
|
|
|
|
|
*/
|
|
|
|
|
#define clear_stopped_child_used_math(child) do { (child)->flags &= ~PF_USED_MATH; } while (0)
|
|
|
|
|
#define set_stopped_child_used_math(child) do { (child)->flags |= PF_USED_MATH; } while (0)
|
|
|
|
|
#define clear_used_math() clear_stopped_child_used_math(current)
|
|
|
|
|
#define set_used_math() set_stopped_child_used_math(current)
|
|
|
|
|
#define conditional_stopped_child_used_math(condition, child) \
|
|
|
|
|
do { (child)->flags &= ~PF_USED_MATH, (child)->flags |= (condition) ? PF_USED_MATH : 0; } while (0)
|
|
|
|
|
#define conditional_used_math(condition) \
|
|
|
|
|
conditional_stopped_child_used_math(condition, current)
|
|
|
|
|
#define copy_to_stopped_child_used_math(child) \
|
|
|
|
|
do { (child)->flags &= ~PF_USED_MATH, (child)->flags |= current->flags & PF_USED_MATH; } while (0)
|
|
|
|
|
/* NOTE: this will return 0 or PF_USED_MATH, it will never return 1 */
|
|
|
|
|
#define tsk_used_math(p) ((p)->flags & PF_USED_MATH)
|
|
|
|
|
#define used_math() tsk_used_math(current)
|
|
|
|
|
|
mm: teach mm by current context info to not do I/O during memory allocation
This patch introduces PF_MEMALLOC_NOIO on process flag('flags' field of
'struct task_struct'), so that the flag can be set by one task to avoid
doing I/O inside memory allocation in the task's context.
The patch trys to solve one deadlock problem caused by block device, and
the problem may happen at least in the below situations:
- during block device runtime resume, if memory allocation with
GFP_KERNEL is called inside runtime resume callback of any one of its
ancestors(or the block device itself), the deadlock may be triggered
inside the memory allocation since it might not complete until the block
device becomes active and the involed page I/O finishes. The situation
is pointed out first by Alan Stern. It is not a good approach to
convert all GFP_KERNEL[1] in the path into GFP_NOIO because several
subsystems may be involved(for example, PCI, USB and SCSI may be
involved for usb mass stoarage device, network devices involved too in
the iSCSI case)
- during block device runtime suspend, because runtime resume need to
wait for completion of concurrent runtime suspend.
- during error handling of usb mass storage deivce, USB bus reset will
be put on the device, so there shouldn't have any memory allocation with
GFP_KERNEL during USB bus reset, otherwise the deadlock similar with
above may be triggered. Unfortunately, any usb device may include one
mass storage interface in theory, so it requires all usb interface
drivers to handle the situation. In fact, most usb drivers don't know
how to handle bus reset on the device and don't provide .pre_set() and
.post_reset() callback at all, so USB core has to unbind and bind driver
for these devices. So it is still not practical to resort to GFP_NOIO
for solving the problem.
Also the introduced solution can be used by block subsystem or block
drivers too, for example, set the PF_MEMALLOC_NOIO flag before doing
actual I/O transfer.
It is not a good idea to convert all these GFP_KERNEL in the affected
path into GFP_NOIO because these functions doing that may be implemented
as library and will be called in many other contexts.
In fact, memalloc_noio_flags() can convert some of current static
GFP_NOIO allocation into GFP_KERNEL back in other non-affected contexts,
at least almost all GFP_NOIO in USB subsystem can be converted into
GFP_KERNEL after applying the approach and make allocation with GFP_NOIO
only happen in runtime resume/bus reset/block I/O transfer contexts
generally.
[1], several GFP_KERNEL allocation examples in runtime resume path
- pci subsystem
acpi_os_allocate
<-acpi_ut_allocate
<-ACPI_ALLOCATE_ZEROED
<-acpi_evaluate_object
<-__acpi_bus_set_power
<-acpi_bus_set_power
<-acpi_pci_set_power_state
<-platform_pci_set_power_state
<-pci_platform_power_transition
<-__pci_complete_power_transition
<-pci_set_power_state
<-pci_restore_standard_config
<-pci_pm_runtime_resume
- usb subsystem
usb_get_status
<-finish_port_resume
<-usb_port_resume
<-generic_resume
<-usb_resume_device
<-usb_resume_both
<-usb_runtime_resume
- some individual usb drivers
usblp, uvc, gspca, most of dvb-usb-v2 media drivers, cpia2, az6007, ....
That is just what I have found. Unfortunately, this allocation can only
be found by human being now, and there should be many not found since
any function in the resume path(call tree) may allocate memory with
GFP_KERNEL.
Signed-off-by: Ming Lei <ming.lei@canonical.com>
Signed-off-by: Minchan Kim <minchan@kernel.org>
Cc: Alan Stern <stern@rowland.harvard.edu>
Cc: Oliver Neukum <oneukum@suse.de>
Cc: Jiri Kosina <jiri.kosina@suse.com>
Cc: Mel Gorman <mel@csn.ul.ie>
Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Cc: Michal Hocko <mhocko@suse.cz>
Cc: Ingo Molnar <mingo@redhat.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: "Rafael J. Wysocki" <rjw@sisk.pl>
Cc: Greg KH <greg@kroah.com>
Cc: Jens Axboe <axboe@kernel.dk>
Cc: "David S. Miller" <davem@davemloft.net>
Cc: Eric Dumazet <eric.dumazet@gmail.com>
Cc: David Decotigny <david.decotigny@google.com>
Cc: Tom Herbert <therbert@google.com>
Cc: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2013-02-23 08:34:08 +08:00
|
|
|
/* __GFP_IO isn't allowed if PF_MEMALLOC_NOIO is set in current->flags */
|
|
|
|
|
static inline gfp_t memalloc_noio_flags(gfp_t flags)
|
|
|
|
|
{
|
|
|
|
|
if (unlikely(current->flags & PF_MEMALLOC_NOIO))
|
|
|
|
|
flags &= ~__GFP_IO;
|
|
|
|
|
return flags;
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline unsigned int memalloc_noio_save(void)
|
|
|
|
|
{
|
|
|
|
|
unsigned int flags = current->flags & PF_MEMALLOC_NOIO;
|
|
|
|
|
current->flags |= PF_MEMALLOC_NOIO;
|
|
|
|
|
return flags;
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void memalloc_noio_restore(unsigned int flags)
|
|
|
|
|
{
|
|
|
|
|
current->flags = (current->flags & ~PF_MEMALLOC_NOIO) | flags;
|
|
|
|
|
}
|
|
|
|
|
|
signal: Fix premature completion of group stop when interfered by ptrace
task->signal->group_stop_count is used to track the progress of group
stop. It's initialized to the number of tasks which need to stop for
group stop to finish and each stopping or trapping task decrements.
However, each task doesn't keep track of whether it decremented the
counter or not and if woken up before the group stop is complete and
stops again, it can decrement the counter multiple times.
Please consider the following example code.
static void *worker(void *arg)
{
while (1) ;
return NULL;
}
int main(void)
{
pthread_t thread;
pid_t pid;
int i;
pid = fork();
if (!pid) {
for (i = 0; i < 5; i++)
pthread_create(&thread, NULL, worker, NULL);
while (1) ;
return 0;
}
ptrace(PTRACE_ATTACH, pid, NULL, NULL);
while (1) {
waitid(P_PID, pid, NULL, WSTOPPED);
ptrace(PTRACE_SINGLESTEP, pid, NULL, (void *)(long)SIGSTOP);
}
return 0;
}
The child creates five threads and the parent continuously traps the
first thread and whenever the child gets a signal, SIGSTOP is
delivered. If an external process sends SIGSTOP to the child, all
other threads in the process should reliably stop. However, due to
the above bug, the first thread will often end up consuming
group_stop_count multiple times and SIGSTOP often ends up stopping
none or part of the other four threads.
This patch adds a new field task->group_stop which is protected by
siglock and uses GROUP_STOP_CONSUME flag to track which task is still
to consume group_stop_count to fix this bug.
task_clear_group_stop_pending() and task_participate_group_stop() are
added to help manipulating group stop states. As ptrace_stop() now
also uses task_participate_group_stop(), it will set
SIGNAL_STOP_STOPPED if it completes a group stop.
There still are many issues regarding the interaction between group
stop and ptrace. Patches to address them will follow.
- Oleg spotted duplicate GROUP_STOP_CONSUME. Dropped.
Signed-off-by: Tejun Heo <tj@kernel.org>
Acked-by: Oleg Nesterov <oleg@redhat.com>
Cc: Roland McGrath <roland@redhat.com>
2011-03-23 17:37:00 +08:00
|
|
|
/*
|
2011-06-02 17:13:59 +08:00
|
|
|
* task->jobctl flags
|
signal: Fix premature completion of group stop when interfered by ptrace
task->signal->group_stop_count is used to track the progress of group
stop. It's initialized to the number of tasks which need to stop for
group stop to finish and each stopping or trapping task decrements.
However, each task doesn't keep track of whether it decremented the
counter or not and if woken up before the group stop is complete and
stops again, it can decrement the counter multiple times.
Please consider the following example code.
static void *worker(void *arg)
{
while (1) ;
return NULL;
}
int main(void)
{
pthread_t thread;
pid_t pid;
int i;
pid = fork();
if (!pid) {
for (i = 0; i < 5; i++)
pthread_create(&thread, NULL, worker, NULL);
while (1) ;
return 0;
}
ptrace(PTRACE_ATTACH, pid, NULL, NULL);
while (1) {
waitid(P_PID, pid, NULL, WSTOPPED);
ptrace(PTRACE_SINGLESTEP, pid, NULL, (void *)(long)SIGSTOP);
}
return 0;
}
The child creates five threads and the parent continuously traps the
first thread and whenever the child gets a signal, SIGSTOP is
delivered. If an external process sends SIGSTOP to the child, all
other threads in the process should reliably stop. However, due to
the above bug, the first thread will often end up consuming
group_stop_count multiple times and SIGSTOP often ends up stopping
none or part of the other four threads.
This patch adds a new field task->group_stop which is protected by
siglock and uses GROUP_STOP_CONSUME flag to track which task is still
to consume group_stop_count to fix this bug.
task_clear_group_stop_pending() and task_participate_group_stop() are
added to help manipulating group stop states. As ptrace_stop() now
also uses task_participate_group_stop(), it will set
SIGNAL_STOP_STOPPED if it completes a group stop.
There still are many issues regarding the interaction between group
stop and ptrace. Patches to address them will follow.
- Oleg spotted duplicate GROUP_STOP_CONSUME. Dropped.
Signed-off-by: Tejun Heo <tj@kernel.org>
Acked-by: Oleg Nesterov <oleg@redhat.com>
Cc: Roland McGrath <roland@redhat.com>
2011-03-23 17:37:00 +08:00
|
|
|
*/
|
2011-06-02 17:13:59 +08:00
|
|
|
#define JOBCTL_STOP_SIGMASK 0xffff /* signr of the last group stop */
|
signal: Fix premature completion of group stop when interfered by ptrace
task->signal->group_stop_count is used to track the progress of group
stop. It's initialized to the number of tasks which need to stop for
group stop to finish and each stopping or trapping task decrements.
However, each task doesn't keep track of whether it decremented the
counter or not and if woken up before the group stop is complete and
stops again, it can decrement the counter multiple times.
Please consider the following example code.
static void *worker(void *arg)
{
while (1) ;
return NULL;
}
int main(void)
{
pthread_t thread;
pid_t pid;
int i;
pid = fork();
if (!pid) {
for (i = 0; i < 5; i++)
pthread_create(&thread, NULL, worker, NULL);
while (1) ;
return 0;
}
ptrace(PTRACE_ATTACH, pid, NULL, NULL);
while (1) {
waitid(P_PID, pid, NULL, WSTOPPED);
ptrace(PTRACE_SINGLESTEP, pid, NULL, (void *)(long)SIGSTOP);
}
return 0;
}
The child creates five threads and the parent continuously traps the
first thread and whenever the child gets a signal, SIGSTOP is
delivered. If an external process sends SIGSTOP to the child, all
other threads in the process should reliably stop. However, due to
the above bug, the first thread will often end up consuming
group_stop_count multiple times and SIGSTOP often ends up stopping
none or part of the other four threads.
This patch adds a new field task->group_stop which is protected by
siglock and uses GROUP_STOP_CONSUME flag to track which task is still
to consume group_stop_count to fix this bug.
task_clear_group_stop_pending() and task_participate_group_stop() are
added to help manipulating group stop states. As ptrace_stop() now
also uses task_participate_group_stop(), it will set
SIGNAL_STOP_STOPPED if it completes a group stop.
There still are many issues regarding the interaction between group
stop and ptrace. Patches to address them will follow.
- Oleg spotted duplicate GROUP_STOP_CONSUME. Dropped.
Signed-off-by: Tejun Heo <tj@kernel.org>
Acked-by: Oleg Nesterov <oleg@redhat.com>
Cc: Roland McGrath <roland@redhat.com>
2011-03-23 17:37:00 +08:00
|
|
|
|
2011-06-02 17:13:59 +08:00
|
|
|
#define JOBCTL_STOP_DEQUEUED_BIT 16 /* stop signal dequeued */
|
|
|
|
|
#define JOBCTL_STOP_PENDING_BIT 17 /* task should stop for group stop */
|
|
|
|
|
#define JOBCTL_STOP_CONSUME_BIT 18 /* consume group stop count */
|
2011-06-14 17:20:14 +08:00
|
|
|
#define JOBCTL_TRAP_STOP_BIT 19 /* trap for STOP */
|
ptrace: implement TRAP_NOTIFY and use it for group stop events
Currently there's no way for ptracer to find out whether group stop
finished other than polling with INTERRUPT - GETSIGINFO - CONT
sequence. This patch implements group stop notification for ptracer
using STOP traps.
When group stop state of a seized tracee changes, JOBCTL_TRAP_NOTIFY
is set, which schedules a STOP trap which is sticky - it isn't cleared
by other traps and at least one STOP trap will happen eventually.
STOP trap is synchronization point for event notification and the
tracer can determine the current group stop state by looking at the
signal number portion of exit code (si_status from waitid(2) or
si_code from PTRACE_GETSIGINFO).
Notifications are generated both on start and end of group stops but,
because group stop participation always happens before STOP trap, this
doesn't cause an extra trap while tracee is participating in group
stop. The symmetry will be useful later.
Note that this notification works iff tracee is not trapped.
Currently there is no way to be notified of group stop state changes
while tracee is trapped. This will be addressed by a later patch.
An example program follows.
#define PTRACE_SEIZE 0x4206
#define PTRACE_INTERRUPT 0x4207
#define PTRACE_SEIZE_DEVEL 0x80000000
static const struct timespec ts1s = { .tv_sec = 1 };
int main(int argc, char **argv)
{
pid_t tracee, tracer;
int i;
tracee = fork();
if (!tracee)
while (1)
pause();
tracer = fork();
if (!tracer) {
siginfo_t si;
ptrace(PTRACE_SEIZE, tracee, NULL,
(void *)(unsigned long)PTRACE_SEIZE_DEVEL);
ptrace(PTRACE_INTERRUPT, tracee, NULL, NULL);
repeat:
waitid(P_PID, tracee, NULL, WSTOPPED);
ptrace(PTRACE_GETSIGINFO, tracee, NULL, &si);
if (!si.si_code) {
printf("tracer: SIG %d\n", si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL,
(void *)(unsigned long)si.si_signo);
goto repeat;
}
printf("tracer: stopped=%d signo=%d\n",
si.si_signo != SIGTRAP, si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL, NULL);
goto repeat;
}
for (i = 0; i < 3; i++) {
nanosleep(&ts1s, NULL);
printf("mother: SIGSTOP\n");
kill(tracee, SIGSTOP);
nanosleep(&ts1s, NULL);
printf("mother: SIGCONT\n");
kill(tracee, SIGCONT);
}
nanosleep(&ts1s, NULL);
kill(tracer, SIGKILL);
kill(tracee, SIGKILL);
return 0;
}
In the above program, tracer keeps tracee running and gets
notification of each group stop state changes.
# ./test-notify
tracer: stopped=0 signo=5
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
Signed-off-by: Tejun Heo <tj@kernel.org>
Cc: Oleg Nesterov <oleg@redhat.com>
2011-06-14 17:20:17 +08:00
|
|
|
#define JOBCTL_TRAP_NOTIFY_BIT 20 /* trap for NOTIFY */
|
2011-06-02 17:13:59 +08:00
|
|
|
#define JOBCTL_TRAPPING_BIT 21 /* switching to TRACED */
|
ptrace: implement PTRACE_LISTEN
The previous patch implemented async notification for ptrace but it
only worked while trace is running. This patch introduces
PTRACE_LISTEN which is suggested by Oleg Nestrov.
It's allowed iff tracee is in STOP trap and puts tracee into
quasi-running state - tracee never really runs but wait(2) and
ptrace(2) consider it to be running. While ptracer is listening,
tracee is allowed to re-enter STOP to notify an async event.
Listening state is cleared on the first notification. Ptracer can
also clear it by issuing INTERRUPT - tracee will re-trap into STOP
with listening state cleared.
This allows ptracer to monitor group stop state without running tracee
- use INTERRUPT to put tracee into STOP trap, issue LISTEN and then
wait(2) to wait for the next group stop event. When it happens,
PTRACE_GETSIGINFO provides information to determine the current state.
Test program follows.
#define PTRACE_SEIZE 0x4206
#define PTRACE_INTERRUPT 0x4207
#define PTRACE_LISTEN 0x4208
#define PTRACE_SEIZE_DEVEL 0x80000000
static const struct timespec ts1s = { .tv_sec = 1 };
int main(int argc, char **argv)
{
pid_t tracee, tracer;
int i;
tracee = fork();
if (!tracee)
while (1)
pause();
tracer = fork();
if (!tracer) {
siginfo_t si;
ptrace(PTRACE_SEIZE, tracee, NULL,
(void *)(unsigned long)PTRACE_SEIZE_DEVEL);
ptrace(PTRACE_INTERRUPT, tracee, NULL, NULL);
repeat:
waitid(P_PID, tracee, NULL, WSTOPPED);
ptrace(PTRACE_GETSIGINFO, tracee, NULL, &si);
if (!si.si_code) {
printf("tracer: SIG %d\n", si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL,
(void *)(unsigned long)si.si_signo);
goto repeat;
}
printf("tracer: stopped=%d signo=%d\n",
si.si_signo != SIGTRAP, si.si_signo);
if (si.si_signo != SIGTRAP)
ptrace(PTRACE_LISTEN, tracee, NULL, NULL);
else
ptrace(PTRACE_CONT, tracee, NULL, NULL);
goto repeat;
}
for (i = 0; i < 3; i++) {
nanosleep(&ts1s, NULL);
printf("mother: SIGSTOP\n");
kill(tracee, SIGSTOP);
nanosleep(&ts1s, NULL);
printf("mother: SIGCONT\n");
kill(tracee, SIGCONT);
}
nanosleep(&ts1s, NULL);
kill(tracer, SIGKILL);
kill(tracee, SIGKILL);
return 0;
}
This is identical to the program to test TRAP_NOTIFY except that
tracee is PTRACE_LISTEN'd instead of PTRACE_CONT'd when group stopped.
This allows ptracer to monitor when group stop ends without running
tracee.
# ./test-listen
tracer: stopped=0 signo=5
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
-v2: Moved JOBCTL_LISTENING check in wait_task_stopped() into
task_stopped_code() as suggested by Oleg.
Signed-off-by: Tejun Heo <tj@kernel.org>
Cc: Oleg Nesterov <oleg@redhat.com>
2011-06-14 17:20:18 +08:00
|
|
|
#define JOBCTL_LISTENING_BIT 22 /* ptracer is listening for events */
|
2011-06-02 17:13:59 +08:00
|
|
|
|
|
|
|
|
#define JOBCTL_STOP_DEQUEUED (1 << JOBCTL_STOP_DEQUEUED_BIT)
|
|
|
|
|
#define JOBCTL_STOP_PENDING (1 << JOBCTL_STOP_PENDING_BIT)
|
|
|
|
|
#define JOBCTL_STOP_CONSUME (1 << JOBCTL_STOP_CONSUME_BIT)
|
2011-06-14 17:20:14 +08:00
|
|
|
#define JOBCTL_TRAP_STOP (1 << JOBCTL_TRAP_STOP_BIT)
|
ptrace: implement TRAP_NOTIFY and use it for group stop events
Currently there's no way for ptracer to find out whether group stop
finished other than polling with INTERRUPT - GETSIGINFO - CONT
sequence. This patch implements group stop notification for ptracer
using STOP traps.
When group stop state of a seized tracee changes, JOBCTL_TRAP_NOTIFY
is set, which schedules a STOP trap which is sticky - it isn't cleared
by other traps and at least one STOP trap will happen eventually.
STOP trap is synchronization point for event notification and the
tracer can determine the current group stop state by looking at the
signal number portion of exit code (si_status from waitid(2) or
si_code from PTRACE_GETSIGINFO).
Notifications are generated both on start and end of group stops but,
because group stop participation always happens before STOP trap, this
doesn't cause an extra trap while tracee is participating in group
stop. The symmetry will be useful later.
Note that this notification works iff tracee is not trapped.
Currently there is no way to be notified of group stop state changes
while tracee is trapped. This will be addressed by a later patch.
An example program follows.
#define PTRACE_SEIZE 0x4206
#define PTRACE_INTERRUPT 0x4207
#define PTRACE_SEIZE_DEVEL 0x80000000
static const struct timespec ts1s = { .tv_sec = 1 };
int main(int argc, char **argv)
{
pid_t tracee, tracer;
int i;
tracee = fork();
if (!tracee)
while (1)
pause();
tracer = fork();
if (!tracer) {
siginfo_t si;
ptrace(PTRACE_SEIZE, tracee, NULL,
(void *)(unsigned long)PTRACE_SEIZE_DEVEL);
ptrace(PTRACE_INTERRUPT, tracee, NULL, NULL);
repeat:
waitid(P_PID, tracee, NULL, WSTOPPED);
ptrace(PTRACE_GETSIGINFO, tracee, NULL, &si);
if (!si.si_code) {
printf("tracer: SIG %d\n", si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL,
(void *)(unsigned long)si.si_signo);
goto repeat;
}
printf("tracer: stopped=%d signo=%d\n",
si.si_signo != SIGTRAP, si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL, NULL);
goto repeat;
}
for (i = 0; i < 3; i++) {
nanosleep(&ts1s, NULL);
printf("mother: SIGSTOP\n");
kill(tracee, SIGSTOP);
nanosleep(&ts1s, NULL);
printf("mother: SIGCONT\n");
kill(tracee, SIGCONT);
}
nanosleep(&ts1s, NULL);
kill(tracer, SIGKILL);
kill(tracee, SIGKILL);
return 0;
}
In the above program, tracer keeps tracee running and gets
notification of each group stop state changes.
# ./test-notify
tracer: stopped=0 signo=5
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
Signed-off-by: Tejun Heo <tj@kernel.org>
Cc: Oleg Nesterov <oleg@redhat.com>
2011-06-14 17:20:17 +08:00
|
|
|
#define JOBCTL_TRAP_NOTIFY (1 << JOBCTL_TRAP_NOTIFY_BIT)
|
2011-06-02 17:13:59 +08:00
|
|
|
#define JOBCTL_TRAPPING (1 << JOBCTL_TRAPPING_BIT)
|
ptrace: implement PTRACE_LISTEN
The previous patch implemented async notification for ptrace but it
only worked while trace is running. This patch introduces
PTRACE_LISTEN which is suggested by Oleg Nestrov.
It's allowed iff tracee is in STOP trap and puts tracee into
quasi-running state - tracee never really runs but wait(2) and
ptrace(2) consider it to be running. While ptracer is listening,
tracee is allowed to re-enter STOP to notify an async event.
Listening state is cleared on the first notification. Ptracer can
also clear it by issuing INTERRUPT - tracee will re-trap into STOP
with listening state cleared.
This allows ptracer to monitor group stop state without running tracee
- use INTERRUPT to put tracee into STOP trap, issue LISTEN and then
wait(2) to wait for the next group stop event. When it happens,
PTRACE_GETSIGINFO provides information to determine the current state.
Test program follows.
#define PTRACE_SEIZE 0x4206
#define PTRACE_INTERRUPT 0x4207
#define PTRACE_LISTEN 0x4208
#define PTRACE_SEIZE_DEVEL 0x80000000
static const struct timespec ts1s = { .tv_sec = 1 };
int main(int argc, char **argv)
{
pid_t tracee, tracer;
int i;
tracee = fork();
if (!tracee)
while (1)
pause();
tracer = fork();
if (!tracer) {
siginfo_t si;
ptrace(PTRACE_SEIZE, tracee, NULL,
(void *)(unsigned long)PTRACE_SEIZE_DEVEL);
ptrace(PTRACE_INTERRUPT, tracee, NULL, NULL);
repeat:
waitid(P_PID, tracee, NULL, WSTOPPED);
ptrace(PTRACE_GETSIGINFO, tracee, NULL, &si);
if (!si.si_code) {
printf("tracer: SIG %d\n", si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL,
(void *)(unsigned long)si.si_signo);
goto repeat;
}
printf("tracer: stopped=%d signo=%d\n",
si.si_signo != SIGTRAP, si.si_signo);
if (si.si_signo != SIGTRAP)
ptrace(PTRACE_LISTEN, tracee, NULL, NULL);
else
ptrace(PTRACE_CONT, tracee, NULL, NULL);
goto repeat;
}
for (i = 0; i < 3; i++) {
nanosleep(&ts1s, NULL);
printf("mother: SIGSTOP\n");
kill(tracee, SIGSTOP);
nanosleep(&ts1s, NULL);
printf("mother: SIGCONT\n");
kill(tracee, SIGCONT);
}
nanosleep(&ts1s, NULL);
kill(tracer, SIGKILL);
kill(tracee, SIGKILL);
return 0;
}
This is identical to the program to test TRAP_NOTIFY except that
tracee is PTRACE_LISTEN'd instead of PTRACE_CONT'd when group stopped.
This allows ptracer to monitor when group stop ends without running
tracee.
# ./test-listen
tracer: stopped=0 signo=5
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
-v2: Moved JOBCTL_LISTENING check in wait_task_stopped() into
task_stopped_code() as suggested by Oleg.
Signed-off-by: Tejun Heo <tj@kernel.org>
Cc: Oleg Nesterov <oleg@redhat.com>
2011-06-14 17:20:18 +08:00
|
|
|
#define JOBCTL_LISTENING (1 << JOBCTL_LISTENING_BIT)
|
2011-06-02 17:13:59 +08:00
|
|
|
|
ptrace: implement TRAP_NOTIFY and use it for group stop events
Currently there's no way for ptracer to find out whether group stop
finished other than polling with INTERRUPT - GETSIGINFO - CONT
sequence. This patch implements group stop notification for ptracer
using STOP traps.
When group stop state of a seized tracee changes, JOBCTL_TRAP_NOTIFY
is set, which schedules a STOP trap which is sticky - it isn't cleared
by other traps and at least one STOP trap will happen eventually.
STOP trap is synchronization point for event notification and the
tracer can determine the current group stop state by looking at the
signal number portion of exit code (si_status from waitid(2) or
si_code from PTRACE_GETSIGINFO).
Notifications are generated both on start and end of group stops but,
because group stop participation always happens before STOP trap, this
doesn't cause an extra trap while tracee is participating in group
stop. The symmetry will be useful later.
Note that this notification works iff tracee is not trapped.
Currently there is no way to be notified of group stop state changes
while tracee is trapped. This will be addressed by a later patch.
An example program follows.
#define PTRACE_SEIZE 0x4206
#define PTRACE_INTERRUPT 0x4207
#define PTRACE_SEIZE_DEVEL 0x80000000
static const struct timespec ts1s = { .tv_sec = 1 };
int main(int argc, char **argv)
{
pid_t tracee, tracer;
int i;
tracee = fork();
if (!tracee)
while (1)
pause();
tracer = fork();
if (!tracer) {
siginfo_t si;
ptrace(PTRACE_SEIZE, tracee, NULL,
(void *)(unsigned long)PTRACE_SEIZE_DEVEL);
ptrace(PTRACE_INTERRUPT, tracee, NULL, NULL);
repeat:
waitid(P_PID, tracee, NULL, WSTOPPED);
ptrace(PTRACE_GETSIGINFO, tracee, NULL, &si);
if (!si.si_code) {
printf("tracer: SIG %d\n", si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL,
(void *)(unsigned long)si.si_signo);
goto repeat;
}
printf("tracer: stopped=%d signo=%d\n",
si.si_signo != SIGTRAP, si.si_signo);
ptrace(PTRACE_CONT, tracee, NULL, NULL);
goto repeat;
}
for (i = 0; i < 3; i++) {
nanosleep(&ts1s, NULL);
printf("mother: SIGSTOP\n");
kill(tracee, SIGSTOP);
nanosleep(&ts1s, NULL);
printf("mother: SIGCONT\n");
kill(tracee, SIGCONT);
}
nanosleep(&ts1s, NULL);
kill(tracer, SIGKILL);
kill(tracee, SIGKILL);
return 0;
}
In the above program, tracer keeps tracee running and gets
notification of each group stop state changes.
# ./test-notify
tracer: stopped=0 signo=5
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
mother: SIGSTOP
tracer: SIG 19
tracer: stopped=1 signo=19
mother: SIGCONT
tracer: stopped=0 signo=5
tracer: SIG 18
Signed-off-by: Tejun Heo <tj@kernel.org>
Cc: Oleg Nesterov <oleg@redhat.com>
2011-06-14 17:20:17 +08:00
|
|
|
#define JOBCTL_TRAP_MASK (JOBCTL_TRAP_STOP | JOBCTL_TRAP_NOTIFY)
|
2011-06-14 17:20:14 +08:00
|
|
|
#define JOBCTL_PENDING_MASK (JOBCTL_STOP_PENDING | JOBCTL_TRAP_MASK)
|
2011-06-02 17:14:00 +08:00
|
|
|
|
2011-06-02 17:14:00 +08:00
|
|
|
extern bool task_set_jobctl_pending(struct task_struct *task,
|
|
|
|
|
unsigned int mask);
|
2011-06-14 17:20:14 +08:00
|
|
|
extern void task_clear_jobctl_trapping(struct task_struct *task);
|
2011-06-02 17:14:00 +08:00
|
|
|
extern void task_clear_jobctl_pending(struct task_struct *task,
|
|
|
|
|
unsigned int mask);
|
2011-03-23 17:37:00 +08:00
|
|
|
|
2010-06-30 07:49:16 +08:00
|
|
|
#ifdef CONFIG_PREEMPT_RCU
|
rcu: Merge preemptable-RCU functionality into hierarchical RCU
Create a kernel/rcutree_plugin.h file that contains definitions
for preemptable RCU (or, under the #else branch of the #ifdef,
empty definitions for the classic non-preemptable semantics).
These definitions fit into plugins defined in kernel/rcutree.c
for this purpose.
This variant of preemptable RCU uses a new algorithm whose
read-side expense is roughly that of classic hierarchical RCU
under CONFIG_PREEMPT. This new algorithm's update-side expense
is similar to that of classic hierarchical RCU, and, in absence
of read-side preemption or blocking, is exactly that of classic
hierarchical RCU. Perhaps more important, this new algorithm
has a much simpler implementation, saving well over 1,000 lines
of code compared to mainline's implementation of preemptable
RCU, which will hopefully be retired in favor of this new
algorithm.
The simplifications are obtained by maintaining per-task
nesting state for running tasks, and using a simple
lock-protected algorithm to handle accounting when tasks block
within RCU read-side critical sections, making use of lessons
learned while creating numerous user-level RCU implementations
over the past 18 months.
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: laijs@cn.fujitsu.com
Cc: dipankar@in.ibm.com
Cc: akpm@linux-foundation.org
Cc: mathieu.desnoyers@polymtl.ca
Cc: josht@linux.vnet.ibm.com
Cc: dvhltc@us.ibm.com
Cc: niv@us.ibm.com
Cc: peterz@infradead.org
Cc: rostedt@goodmis.org
LKML-Reference: <12509746134003-git-send-email->
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-08-23 04:56:52 +08:00
|
|
|
|
|
|
|
|
#define RCU_READ_UNLOCK_BLOCKED (1 << 0) /* blocked while in RCU read-side. */
|
2012-01-12 09:25:17 +08:00
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|
|
#define RCU_READ_UNLOCK_NEED_QS (1 << 1) /* RCU core needs CPU response. */
|
rcu: Merge preemptable-RCU functionality into hierarchical RCU
Create a kernel/rcutree_plugin.h file that contains definitions
for preemptable RCU (or, under the #else branch of the #ifdef,
empty definitions for the classic non-preemptable semantics).
These definitions fit into plugins defined in kernel/rcutree.c
for this purpose.
This variant of preemptable RCU uses a new algorithm whose
read-side expense is roughly that of classic hierarchical RCU
under CONFIG_PREEMPT. This new algorithm's update-side expense
is similar to that of classic hierarchical RCU, and, in absence
of read-side preemption or blocking, is exactly that of classic
hierarchical RCU. Perhaps more important, this new algorithm
has a much simpler implementation, saving well over 1,000 lines
of code compared to mainline's implementation of preemptable
RCU, which will hopefully be retired in favor of this new
algorithm.
The simplifications are obtained by maintaining per-task
nesting state for running tasks, and using a simple
lock-protected algorithm to handle accounting when tasks block
within RCU read-side critical sections, making use of lessons
learned while creating numerous user-level RCU implementations
over the past 18 months.
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: laijs@cn.fujitsu.com
Cc: dipankar@in.ibm.com
Cc: akpm@linux-foundation.org
Cc: mathieu.desnoyers@polymtl.ca
Cc: josht@linux.vnet.ibm.com
Cc: dvhltc@us.ibm.com
Cc: niv@us.ibm.com
Cc: peterz@infradead.org
Cc: rostedt@goodmis.org
LKML-Reference: <12509746134003-git-send-email->
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-08-23 04:56:52 +08:00
|
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static inline void rcu_copy_process(struct task_struct *p)
|
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|
|
{
|
|
|
|
|
p->rcu_read_lock_nesting = 0;
|
|
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p->rcu_read_unlock_special = 0;
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2010-06-30 07:49:16 +08:00
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#ifdef CONFIG_TREE_PREEMPT_RCU
|
2009-08-28 05:58:16 +08:00
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p->rcu_blocked_node = NULL;
|
2010-09-28 08:25:23 +08:00
|
|
|
#endif /* #ifdef CONFIG_TREE_PREEMPT_RCU */
|
|
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#ifdef CONFIG_RCU_BOOST
|
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p->rcu_boost_mutex = NULL;
|
|
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|
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#endif /* #ifdef CONFIG_RCU_BOOST */
|
rcu: Merge preemptable-RCU functionality into hierarchical RCU
Create a kernel/rcutree_plugin.h file that contains definitions
for preemptable RCU (or, under the #else branch of the #ifdef,
empty definitions for the classic non-preemptable semantics).
These definitions fit into plugins defined in kernel/rcutree.c
for this purpose.
This variant of preemptable RCU uses a new algorithm whose
read-side expense is roughly that of classic hierarchical RCU
under CONFIG_PREEMPT. This new algorithm's update-side expense
is similar to that of classic hierarchical RCU, and, in absence
of read-side preemption or blocking, is exactly that of classic
hierarchical RCU. Perhaps more important, this new algorithm
has a much simpler implementation, saving well over 1,000 lines
of code compared to mainline's implementation of preemptable
RCU, which will hopefully be retired in favor of this new
algorithm.
The simplifications are obtained by maintaining per-task
nesting state for running tasks, and using a simple
lock-protected algorithm to handle accounting when tasks block
within RCU read-side critical sections, making use of lessons
learned while creating numerous user-level RCU implementations
over the past 18 months.
Signed-off-by: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: laijs@cn.fujitsu.com
Cc: dipankar@in.ibm.com
Cc: akpm@linux-foundation.org
Cc: mathieu.desnoyers@polymtl.ca
Cc: josht@linux.vnet.ibm.com
Cc: dvhltc@us.ibm.com
Cc: niv@us.ibm.com
Cc: peterz@infradead.org
Cc: rostedt@goodmis.org
LKML-Reference: <12509746134003-git-send-email->
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2009-08-23 04:56:52 +08:00
|
|
|
INIT_LIST_HEAD(&p->rcu_node_entry);
|
|
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}
|
|
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#else
|
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static inline void rcu_copy_process(struct task_struct *p)
|
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|
|
{
|
|
|
|
|
}
|
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#endif
|
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|
2012-08-01 07:44:07 +08:00
|
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static inline void tsk_restore_flags(struct task_struct *task,
|
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unsigned long orig_flags, unsigned long flags)
|
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|
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{
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task->flags &= ~flags;
|
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task->flags |= orig_flags & flags;
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}
|
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|
|
2005-04-17 06:20:36 +08:00
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|
#ifdef CONFIG_SMP
|
2011-05-19 14:08:58 +08:00
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|
extern void do_set_cpus_allowed(struct task_struct *p,
|
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const struct cpumask *new_mask);
|
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2008-03-27 05:23:49 +08:00
|
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|
extern int set_cpus_allowed_ptr(struct task_struct *p,
|
2008-11-25 00:05:14 +08:00
|
|
|
const struct cpumask *new_mask);
|
2005-04-17 06:20:36 +08:00
|
|
|
#else
|
2011-05-19 14:08:58 +08:00
|
|
|
static inline void do_set_cpus_allowed(struct task_struct *p,
|
|
|
|
|
const struct cpumask *new_mask)
|
|
|
|
|
{
|
|
|
|
|
}
|
2008-03-27 05:23:49 +08:00
|
|
|
static inline int set_cpus_allowed_ptr(struct task_struct *p,
|
2008-11-25 00:05:14 +08:00
|
|
|
const struct cpumask *new_mask)
|
2005-04-17 06:20:36 +08:00
|
|
|
{
|
2008-11-25 00:05:14 +08:00
|
|
|
if (!cpumask_test_cpu(0, new_mask))
|
2005-04-17 06:20:36 +08:00
|
|
|
return -EINVAL;
|
|
|
|
|
return 0;
|
|
|
|
|
}
|
|
|
|
|
#endif
|
2009-09-24 23:34:38 +08:00
|
|
|
|
2011-08-11 05:21:01 +08:00
|
|
|
#ifdef CONFIG_NO_HZ_COMMON
|
2012-06-22 21:52:09 +08:00
|
|
|
void calc_load_enter_idle(void);
|
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|
|
void calc_load_exit_idle(void);
|
|
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|
|
#else
|
|
|
|
|
static inline void calc_load_enter_idle(void) { }
|
|
|
|
|
static inline void calc_load_exit_idle(void) { }
|
2011-08-11 05:21:01 +08:00
|
|
|
#endif /* CONFIG_NO_HZ_COMMON */
|
2012-06-22 21:52:09 +08:00
|
|
|
|
2009-09-24 23:34:38 +08:00
|
|
|
#ifndef CONFIG_CPUMASK_OFFSTACK
|
2008-03-27 05:23:49 +08:00
|
|
|
static inline int set_cpus_allowed(struct task_struct *p, cpumask_t new_mask)
|
|
|
|
|
{
|
|
|
|
|
return set_cpus_allowed_ptr(p, &new_mask);
|
|
|
|
|
}
|
2009-09-24 23:34:38 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2009-02-27 03:20:29 +08:00
|
|
|
/*
|
2010-05-25 16:48:51 +08:00
|
|
|
* Do not use outside of architecture code which knows its limitations.
|
|
|
|
|
*
|
|
|
|
|
* sched_clock() has no promise of monotonicity or bounded drift between
|
|
|
|
|
* CPUs, use (which you should not) requires disabling IRQs.
|
|
|
|
|
*
|
|
|
|
|
* Please use one of the three interfaces below.
|
2009-02-27 03:20:29 +08:00
|
|
|
*/
|
2009-12-10 09:07:03 +08:00
|
|
|
extern unsigned long long notrace sched_clock(void);
|
2010-05-25 16:48:51 +08:00
|
|
|
/*
|
2012-04-02 16:00:44 +08:00
|
|
|
* See the comment in kernel/sched/clock.c
|
2010-05-25 16:48:51 +08:00
|
|
|
*/
|
|
|
|
|
extern u64 cpu_clock(int cpu);
|
|
|
|
|
extern u64 local_clock(void);
|
|
|
|
|
extern u64 sched_clock_cpu(int cpu);
|
|
|
|
|
|
2007-07-20 03:28:35 +08:00
|
|
|
|
2008-08-11 14:59:03 +08:00
|
|
|
extern void sched_clock_init(void);
|
2008-05-04 00:29:28 +08:00
|
|
|
|
2008-08-11 14:59:03 +08:00
|
|
|
#ifndef CONFIG_HAVE_UNSTABLE_SCHED_CLOCK
|
2008-05-04 00:29:28 +08:00
|
|
|
static inline void sched_clock_tick(void)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void sched_clock_idle_sleep_event(void)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void sched_clock_idle_wakeup_event(u64 delta_ns)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
#else
|
2010-05-25 16:48:51 +08:00
|
|
|
/*
|
|
|
|
|
* Architectures can set this to 1 if they have specified
|
|
|
|
|
* CONFIG_HAVE_UNSTABLE_SCHED_CLOCK in their arch Kconfig,
|
|
|
|
|
* but then during bootup it turns out that sched_clock()
|
|
|
|
|
* is reliable after all:
|
|
|
|
|
*/
|
|
|
|
|
extern int sched_clock_stable;
|
|
|
|
|
|
2008-05-04 00:29:28 +08:00
|
|
|
extern void sched_clock_tick(void);
|
|
|
|
|
extern void sched_clock_idle_sleep_event(void);
|
|
|
|
|
extern void sched_clock_idle_wakeup_event(u64 delta_ns);
|
|
|
|
|
#endif
|
|
|
|
|
|
2010-10-05 08:03:19 +08:00
|
|
|
#ifdef CONFIG_IRQ_TIME_ACCOUNTING
|
|
|
|
|
/*
|
|
|
|
|
* An i/f to runtime opt-in for irq time accounting based off of sched_clock.
|
|
|
|
|
* The reason for this explicit opt-in is not to have perf penalty with
|
|
|
|
|
* slow sched_clocks.
|
|
|
|
|
*/
|
|
|
|
|
extern void enable_sched_clock_irqtime(void);
|
|
|
|
|
extern void disable_sched_clock_irqtime(void);
|
|
|
|
|
#else
|
|
|
|
|
static inline void enable_sched_clock_irqtime(void) {}
|
|
|
|
|
static inline void disable_sched_clock_irqtime(void) {}
|
|
|
|
|
#endif
|
|
|
|
|
|
2006-07-03 15:25:41 +08:00
|
|
|
extern unsigned long long
|
2007-07-10 00:51:58 +08:00
|
|
|
task_sched_runtime(struct task_struct *task);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/* sched_exec is called by processes performing an exec */
|
|
|
|
|
#ifdef CONFIG_SMP
|
|
|
|
|
extern void sched_exec(void);
|
|
|
|
|
#else
|
|
|
|
|
#define sched_exec() {}
|
|
|
|
|
#endif
|
|
|
|
|
|
2007-08-23 21:18:02 +08:00
|
|
|
extern void sched_clock_idle_sleep_event(void);
|
|
|
|
|
extern void sched_clock_idle_wakeup_event(u64 delta_ns);
|
2007-07-10 00:51:59 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
#ifdef CONFIG_HOTPLUG_CPU
|
|
|
|
|
extern void idle_task_exit(void);
|
|
|
|
|
#else
|
|
|
|
|
static inline void idle_task_exit(void) {}
|
|
|
|
|
#endif
|
|
|
|
|
|
2011-08-11 05:21:01 +08:00
|
|
|
#if defined(CONFIG_NO_HZ_COMMON) && defined(CONFIG_SMP)
|
2011-08-11 05:21:01 +08:00
|
|
|
extern void wake_up_nohz_cpu(int cpu);
|
2008-03-22 16:20:24 +08:00
|
|
|
#else
|
2011-08-11 05:21:01 +08:00
|
|
|
static inline void wake_up_nohz_cpu(int cpu) { }
|
2008-03-22 16:20:24 +08:00
|
|
|
#endif
|
|
|
|
|
|
2013-04-20 21:15:35 +08:00
|
|
|
#ifdef CONFIG_NO_HZ_FULL
|
|
|
|
|
extern bool sched_can_stop_tick(void);
|
2013-05-03 09:39:05 +08:00
|
|
|
extern u64 scheduler_tick_max_deferment(void);
|
2013-04-20 21:15:35 +08:00
|
|
|
#else
|
|
|
|
|
static inline bool sched_can_stop_tick(void) { return false; }
|
2008-03-22 16:20:24 +08:00
|
|
|
#endif
|
|
|
|
|
|
sched: Add 'autogroup' scheduling feature: automated per session task groups
A recurring complaint from CFS users is that parallel kbuild has
a negative impact on desktop interactivity. This patch
implements an idea from Linus, to automatically create task
groups. Currently, only per session autogroups are implemented,
but the patch leaves the way open for enhancement.
Implementation: each task's signal struct contains an inherited
pointer to a refcounted autogroup struct containing a task group
pointer, the default for all tasks pointing to the
init_task_group. When a task calls setsid(), a new task group
is created, the process is moved into the new task group, and a
reference to the preveious task group is dropped. Child
processes inherit this task group thereafter, and increase it's
refcount. When the last thread of a process exits, the
process's reference is dropped, such that when the last process
referencing an autogroup exits, the autogroup is destroyed.
At runqueue selection time, IFF a task has no cgroup assignment,
its current autogroup is used.
Autogroup bandwidth is controllable via setting it's nice level
through the proc filesystem:
cat /proc/<pid>/autogroup
Displays the task's group and the group's nice level.
echo <nice level> > /proc/<pid>/autogroup
Sets the task group's shares to the weight of nice <level> task.
Setting nice level is rate limited for !admin users due to the
abuse risk of task group locking.
The feature is enabled from boot by default if
CONFIG_SCHED_AUTOGROUP=y is selected, but can be disabled via
the boot option noautogroup, and can also be turned on/off on
the fly via:
echo [01] > /proc/sys/kernel/sched_autogroup_enabled
... which will automatically move tasks to/from the root task group.
Signed-off-by: Mike Galbraith <efault@gmx.de>
Acked-by: Linus Torvalds <torvalds@linux-foundation.org>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Markus Trippelsdorf <markus@trippelsdorf.de>
Cc: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Cc: Paul Turner <pjt@google.com>
Cc: Oleg Nesterov <oleg@redhat.com>
[ Removed the task_group_path() debug code, and fixed !EVENTFD build failure. ]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
LKML-Reference: <1290281700.28711.9.camel@maggy.simson.net>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2010-11-30 21:18:03 +08:00
|
|
|
#ifdef CONFIG_SCHED_AUTOGROUP
|
|
|
|
|
extern void sched_autogroup_create_attach(struct task_struct *p);
|
|
|
|
|
extern void sched_autogroup_detach(struct task_struct *p);
|
|
|
|
|
extern void sched_autogroup_fork(struct signal_struct *sig);
|
|
|
|
|
extern void sched_autogroup_exit(struct signal_struct *sig);
|
|
|
|
|
#ifdef CONFIG_PROC_FS
|
|
|
|
|
extern void proc_sched_autogroup_show_task(struct task_struct *p, struct seq_file *m);
|
2012-02-23 16:41:27 +08:00
|
|
|
extern int proc_sched_autogroup_set_nice(struct task_struct *p, int nice);
|
sched: Add 'autogroup' scheduling feature: automated per session task groups
A recurring complaint from CFS users is that parallel kbuild has
a negative impact on desktop interactivity. This patch
implements an idea from Linus, to automatically create task
groups. Currently, only per session autogroups are implemented,
but the patch leaves the way open for enhancement.
Implementation: each task's signal struct contains an inherited
pointer to a refcounted autogroup struct containing a task group
pointer, the default for all tasks pointing to the
init_task_group. When a task calls setsid(), a new task group
is created, the process is moved into the new task group, and a
reference to the preveious task group is dropped. Child
processes inherit this task group thereafter, and increase it's
refcount. When the last thread of a process exits, the
process's reference is dropped, such that when the last process
referencing an autogroup exits, the autogroup is destroyed.
At runqueue selection time, IFF a task has no cgroup assignment,
its current autogroup is used.
Autogroup bandwidth is controllable via setting it's nice level
through the proc filesystem:
cat /proc/<pid>/autogroup
Displays the task's group and the group's nice level.
echo <nice level> > /proc/<pid>/autogroup
Sets the task group's shares to the weight of nice <level> task.
Setting nice level is rate limited for !admin users due to the
abuse risk of task group locking.
The feature is enabled from boot by default if
CONFIG_SCHED_AUTOGROUP=y is selected, but can be disabled via
the boot option noautogroup, and can also be turned on/off on
the fly via:
echo [01] > /proc/sys/kernel/sched_autogroup_enabled
... which will automatically move tasks to/from the root task group.
Signed-off-by: Mike Galbraith <efault@gmx.de>
Acked-by: Linus Torvalds <torvalds@linux-foundation.org>
Acked-by: Peter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Markus Trippelsdorf <markus@trippelsdorf.de>
Cc: Mathieu Desnoyers <mathieu.desnoyers@efficios.com>
Cc: Paul Turner <pjt@google.com>
Cc: Oleg Nesterov <oleg@redhat.com>
[ Removed the task_group_path() debug code, and fixed !EVENTFD build failure. ]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
LKML-Reference: <1290281700.28711.9.camel@maggy.simson.net>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2010-11-30 21:18:03 +08:00
|
|
|
#endif
|
|
|
|
|
#else
|
|
|
|
|
static inline void sched_autogroup_create_attach(struct task_struct *p) { }
|
|
|
|
|
static inline void sched_autogroup_detach(struct task_struct *p) { }
|
|
|
|
|
static inline void sched_autogroup_fork(struct signal_struct *sig) { }
|
|
|
|
|
static inline void sched_autogroup_exit(struct signal_struct *sig) { }
|
|
|
|
|
#endif
|
|
|
|
|
|
2011-02-01 22:50:51 +08:00
|
|
|
extern bool yield_to(struct task_struct *p, bool preempt);
|
2006-07-03 15:25:41 +08:00
|
|
|
extern void set_user_nice(struct task_struct *p, long nice);
|
|
|
|
|
extern int task_prio(const struct task_struct *p);
|
|
|
|
|
extern int task_nice(const struct task_struct *p);
|
|
|
|
|
extern int can_nice(const struct task_struct *p, const int nice);
|
|
|
|
|
extern int task_curr(const struct task_struct *p);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern int idle_cpu(int cpu);
|
2010-10-21 07:01:12 +08:00
|
|
|
extern int sched_setscheduler(struct task_struct *, int,
|
|
|
|
|
const struct sched_param *);
|
2008-06-23 11:55:38 +08:00
|
|
|
extern int sched_setscheduler_nocheck(struct task_struct *, int,
|
2010-10-21 07:01:12 +08:00
|
|
|
const struct sched_param *);
|
2006-07-03 15:25:41 +08:00
|
|
|
extern struct task_struct *idle_task(int cpu);
|
2011-11-11 04:41:56 +08:00
|
|
|
/**
|
|
|
|
|
* is_idle_task - is the specified task an idle task?
|
2012-01-22 03:03:13 +08:00
|
|
|
* @p: the task in question.
|
2013-07-13 02:45:47 +08:00
|
|
|
*
|
|
|
|
|
* Return: 1 if @p is an idle task. 0 otherwise.
|
2011-11-11 04:41:56 +08:00
|
|
|
*/
|
2011-12-21 00:20:46 +08:00
|
|
|
static inline bool is_idle_task(const struct task_struct *p)
|
2011-11-11 04:41:56 +08:00
|
|
|
{
|
|
|
|
|
return p->pid == 0;
|
|
|
|
|
}
|
2006-07-03 15:25:41 +08:00
|
|
|
extern struct task_struct *curr_task(int cpu);
|
|
|
|
|
extern void set_curr_task(int cpu, struct task_struct *p);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
void yield(void);
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* The default (Linux) execution domain.
|
|
|
|
|
*/
|
|
|
|
|
extern struct exec_domain default_exec_domain;
|
|
|
|
|
|
|
|
|
|
union thread_union {
|
|
|
|
|
struct thread_info thread_info;
|
|
|
|
|
unsigned long stack[THREAD_SIZE/sizeof(long)];
|
|
|
|
|
};
|
|
|
|
|
|
|
|
|
|
#ifndef __HAVE_ARCH_KSTACK_END
|
|
|
|
|
static inline int kstack_end(void *addr)
|
|
|
|
|
{
|
|
|
|
|
/* Reliable end of stack detection:
|
|
|
|
|
* Some APM bios versions misalign the stack
|
|
|
|
|
*/
|
|
|
|
|
return !(((unsigned long)addr+sizeof(void*)-1) & (THREAD_SIZE-sizeof(void*)));
|
|
|
|
|
}
|
|
|
|
|
#endif
|
|
|
|
|
|
|
|
|
|
extern union thread_union init_thread_union;
|
|
|
|
|
extern struct task_struct init_task;
|
|
|
|
|
|
|
|
|
|
extern struct mm_struct init_mm;
|
|
|
|
|
|
2007-10-19 14:40:06 +08:00
|
|
|
extern struct pid_namespace init_pid_ns;
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* find a task by one of its numerical ids
|
|
|
|
|
*
|
|
|
|
|
* find_task_by_pid_ns():
|
|
|
|
|
* finds a task by its pid in the specified namespace
|
2007-10-19 14:40:16 +08:00
|
|
|
* find_task_by_vpid():
|
|
|
|
|
* finds a task by its virtual pid
|
2007-10-19 14:40:06 +08:00
|
|
|
*
|
2008-07-25 16:48:36 +08:00
|
|
|
* see also find_vpid() etc in include/linux/pid.h
|
2007-10-19 14:40:06 +08:00
|
|
|
*/
|
|
|
|
|
|
2007-10-19 14:40:16 +08:00
|
|
|
extern struct task_struct *find_task_by_vpid(pid_t nr);
|
|
|
|
|
extern struct task_struct *find_task_by_pid_ns(pid_t nr,
|
|
|
|
|
struct pid_namespace *ns);
|
2007-10-19 14:40:06 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* per-UID process charging. */
|
2011-11-17 15:20:58 +08:00
|
|
|
extern struct user_struct * alloc_uid(kuid_t);
|
2005-04-17 06:20:36 +08:00
|
|
|
static inline struct user_struct *get_uid(struct user_struct *u)
|
|
|
|
|
{
|
|
|
|
|
atomic_inc(&u->__count);
|
|
|
|
|
return u;
|
|
|
|
|
}
|
|
|
|
|
extern void free_uid(struct user_struct *);
|
|
|
|
|
|
|
|
|
|
#include <asm/current.h>
|
|
|
|
|
|
2011-01-27 22:59:10 +08:00
|
|
|
extern void xtime_update(unsigned long ticks);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2008-02-14 07:03:15 +08:00
|
|
|
extern int wake_up_state(struct task_struct *tsk, unsigned int state);
|
|
|
|
|
extern int wake_up_process(struct task_struct *tsk);
|
2011-05-12 00:18:05 +08:00
|
|
|
extern void wake_up_new_task(struct task_struct *tsk);
|
2005-04-17 06:20:36 +08:00
|
|
|
#ifdef CONFIG_SMP
|
|
|
|
|
extern void kick_process(struct task_struct *tsk);
|
|
|
|
|
#else
|
|
|
|
|
static inline void kick_process(struct task_struct *tsk) { }
|
|
|
|
|
#endif
|
2011-05-12 00:18:05 +08:00
|
|
|
extern void sched_fork(struct task_struct *p);
|
2007-07-10 00:52:00 +08:00
|
|
|
extern void sched_dead(struct task_struct *p);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
extern void proc_caches_init(void);
|
|
|
|
|
extern void flush_signals(struct task_struct *);
|
2009-04-29 20:45:05 +08:00
|
|
|
extern void __flush_signals(struct task_struct *);
|
2007-05-09 17:34:37 +08:00
|
|
|
extern void ignore_signals(struct task_struct *);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void flush_signal_handlers(struct task_struct *, int force_default);
|
|
|
|
|
extern int dequeue_signal(struct task_struct *tsk, sigset_t *mask, siginfo_t *info);
|
|
|
|
|
|
|
|
|
|
static inline int dequeue_signal_lock(struct task_struct *tsk, sigset_t *mask, siginfo_t *info)
|
|
|
|
|
{
|
|
|
|
|
unsigned long flags;
|
|
|
|
|
int ret;
|
|
|
|
|
|
|
|
|
|
spin_lock_irqsave(&tsk->sighand->siglock, flags);
|
|
|
|
|
ret = dequeue_signal(tsk, mask, info);
|
|
|
|
|
spin_unlock_irqrestore(&tsk->sighand->siglock, flags);
|
|
|
|
|
|
|
|
|
|
return ret;
|
2011-06-23 05:08:18 +08:00
|
|
|
}
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
extern void block_all_signals(int (*notifier)(void *priv), void *priv,
|
|
|
|
|
sigset_t *mask);
|
|
|
|
|
extern void unblock_all_signals(void);
|
|
|
|
|
extern void release_task(struct task_struct * p);
|
|
|
|
|
extern int send_sig_info(int, struct siginfo *, struct task_struct *);
|
|
|
|
|
extern int force_sigsegv(int, struct task_struct *);
|
|
|
|
|
extern int force_sig_info(int, struct siginfo *, struct task_struct *);
|
2006-10-02 17:17:10 +08:00
|
|
|
extern int __kill_pgrp_info(int sig, struct siginfo *info, struct pid *pgrp);
|
|
|
|
|
extern int kill_pid_info(int sig, struct siginfo *info, struct pid *pid);
|
2011-09-26 23:45:18 +08:00
|
|
|
extern int kill_pid_info_as_cred(int, struct siginfo *, struct pid *,
|
|
|
|
|
const struct cred *, u32);
|
2006-10-02 17:17:10 +08:00
|
|
|
extern int kill_pgrp(struct pid *pid, int sig, int priv);
|
|
|
|
|
extern int kill_pid(struct pid *pid, int sig, int priv);
|
2007-02-09 23:11:47 +08:00
|
|
|
extern int kill_proc_info(int, struct siginfo *, pid_t);
|
2011-06-23 05:09:09 +08:00
|
|
|
extern __must_check bool do_notify_parent(struct task_struct *, int);
|
ptrace: __ptrace_detach: do __wake_up_parent() if we reap the tracee
The bug is old, it wasn't cause by recent changes.
Test case:
static void *tfunc(void *arg)
{
int pid = (long)arg;
assert(ptrace(PTRACE_ATTACH, pid, NULL, NULL) == 0);
kill(pid, SIGKILL);
sleep(1);
return NULL;
}
int main(void)
{
pthread_t th;
long pid = fork();
if (!pid)
pause();
signal(SIGCHLD, SIG_IGN);
assert(pthread_create(&th, NULL, tfunc, (void*)pid) == 0);
int r = waitpid(-1, NULL, __WNOTHREAD);
printf("waitpid: %d %m\n", r);
return 0;
}
Before the patch this program hangs, after this patch waitpid() correctly
fails with errno == -ECHILD.
The problem is, __ptrace_detach() reaps the EXIT_ZOMBIE tracee if its
->real_parent is our sub-thread and we ignore SIGCHLD. But in this case
we should wake up other threads which can sleep in do_wait().
Signed-off-by: Oleg Nesterov <oleg@redhat.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Vitaly Mayatskikh <vmayatsk@redhat.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-09-24 06:56:44 +08:00
|
|
|
extern void __wake_up_parent(struct task_struct *p, struct task_struct *parent);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void force_sig(int, struct task_struct *);
|
|
|
|
|
extern int send_sig(int, struct task_struct *, int);
|
2010-05-27 05:43:11 +08:00
|
|
|
extern int zap_other_threads(struct task_struct *p);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern struct sigqueue *sigqueue_alloc(void);
|
|
|
|
|
extern void sigqueue_free(struct sigqueue *);
|
2008-04-30 15:52:57 +08:00
|
|
|
extern int send_sigqueue(struct sigqueue *, struct task_struct *, int group);
|
2006-02-10 03:41:50 +08:00
|
|
|
extern int do_sigaction(int, struct k_sigaction *, struct k_sigaction *);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2012-05-22 11:33:55 +08:00
|
|
|
static inline void restore_saved_sigmask(void)
|
|
|
|
|
{
|
|
|
|
|
if (test_and_clear_restore_sigmask())
|
2012-04-28 01:58:59 +08:00
|
|
|
__set_current_blocked(¤t->saved_sigmask);
|
2012-05-22 11:33:55 +08:00
|
|
|
}
|
|
|
|
|
|
2012-05-02 21:59:21 +08:00
|
|
|
static inline sigset_t *sigmask_to_save(void)
|
|
|
|
|
{
|
|
|
|
|
sigset_t *res = ¤t->blocked;
|
|
|
|
|
if (unlikely(test_restore_sigmask()))
|
|
|
|
|
res = ¤t->saved_sigmask;
|
|
|
|
|
return res;
|
|
|
|
|
}
|
|
|
|
|
|
2006-10-02 17:19:00 +08:00
|
|
|
static inline int kill_cad_pid(int sig, int priv)
|
|
|
|
|
{
|
|
|
|
|
return kill_pid(cad_pid, sig, priv);
|
|
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* These can be the second arg to send_sig_info/send_group_sig_info. */
|
|
|
|
|
#define SEND_SIG_NOINFO ((struct siginfo *) 0)
|
|
|
|
|
#define SEND_SIG_PRIV ((struct siginfo *) 1)
|
|
|
|
|
#define SEND_SIG_FORCED ((struct siginfo *) 2)
|
|
|
|
|
|
2009-10-25 22:37:58 +08:00
|
|
|
/*
|
|
|
|
|
* True if we are on the alternate signal stack.
|
|
|
|
|
*/
|
2005-04-17 06:20:36 +08:00
|
|
|
static inline int on_sig_stack(unsigned long sp)
|
|
|
|
|
{
|
2009-10-25 22:37:58 +08:00
|
|
|
#ifdef CONFIG_STACK_GROWSUP
|
|
|
|
|
return sp >= current->sas_ss_sp &&
|
|
|
|
|
sp - current->sas_ss_sp < current->sas_ss_size;
|
|
|
|
|
#else
|
|
|
|
|
return sp > current->sas_ss_sp &&
|
|
|
|
|
sp - current->sas_ss_sp <= current->sas_ss_size;
|
|
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline int sas_ss_flags(unsigned long sp)
|
|
|
|
|
{
|
|
|
|
|
return (current->sas_ss_size == 0 ? SS_DISABLE
|
|
|
|
|
: on_sig_stack(sp) ? SS_ONSTACK : 0);
|
|
|
|
|
}
|
|
|
|
|
|
2012-11-07 02:28:21 +08:00
|
|
|
static inline unsigned long sigsp(unsigned long sp, struct ksignal *ksig)
|
|
|
|
|
{
|
|
|
|
|
if (unlikely((ksig->ka.sa.sa_flags & SA_ONSTACK)) && ! sas_ss_flags(sp))
|
|
|
|
|
#ifdef CONFIG_STACK_GROWSUP
|
|
|
|
|
return current->sas_ss_sp;
|
|
|
|
|
#else
|
|
|
|
|
return current->sas_ss_sp + current->sas_ss_size;
|
|
|
|
|
#endif
|
|
|
|
|
return sp;
|
|
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
|
|
|
|
* Routines for handling mm_structs
|
|
|
|
|
*/
|
|
|
|
|
extern struct mm_struct * mm_alloc(void);
|
|
|
|
|
|
|
|
|
|
/* mmdrop drops the mm and the page tables */
|
2008-02-14 07:03:15 +08:00
|
|
|
extern void __mmdrop(struct mm_struct *);
|
2005-04-17 06:20:36 +08:00
|
|
|
static inline void mmdrop(struct mm_struct * mm)
|
|
|
|
|
{
|
2007-07-10 00:52:01 +08:00
|
|
|
if (unlikely(atomic_dec_and_test(&mm->mm_count)))
|
2005-04-17 06:20:36 +08:00
|
|
|
__mmdrop(mm);
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
/* mmput gets rid of the mappings and all user-space */
|
|
|
|
|
extern void mmput(struct mm_struct *);
|
|
|
|
|
/* Grab a reference to a task's mm, if it is not already going away */
|
|
|
|
|
extern struct mm_struct *get_task_mm(struct task_struct *task);
|
2012-02-02 09:04:09 +08:00
|
|
|
/*
|
|
|
|
|
* Grab a reference to a task's mm, if it is not already going away
|
|
|
|
|
* and ptrace_may_access with the mode parameter passed to it
|
|
|
|
|
* succeeds.
|
|
|
|
|
*/
|
|
|
|
|
extern struct mm_struct *mm_access(struct task_struct *task, unsigned int mode);
|
2005-04-17 06:20:36 +08:00
|
|
|
/* Remove the current tasks stale references to the old mm_struct */
|
|
|
|
|
extern void mm_release(struct task_struct *, struct mm_struct *);
|
2008-03-26 01:47:10 +08:00
|
|
|
/* Allocate a new mm structure and copy contents from tsk->mm */
|
|
|
|
|
extern struct mm_struct *dup_mm(struct task_struct *tsk);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2009-04-03 07:56:59 +08:00
|
|
|
extern int copy_thread(unsigned long, unsigned long, unsigned long,
|
2012-10-23 10:51:14 +08:00
|
|
|
struct task_struct *);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void flush_thread(void);
|
|
|
|
|
extern void exit_thread(void);
|
|
|
|
|
|
|
|
|
|
extern void exit_files(struct task_struct *);
|
2006-03-29 08:11:27 +08:00
|
|
|
extern void __cleanup_sighand(struct sighand_struct *);
|
2008-05-27 00:55:42 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void exit_itimers(struct signal_struct *);
|
2008-05-27 00:55:42 +08:00
|
|
|
extern void flush_itimer_signals(void);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2012-01-13 09:17:17 +08:00
|
|
|
extern void do_group_exit(int);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
extern int allow_signal(int);
|
|
|
|
|
extern int disallow_signal(int);
|
|
|
|
|
|
2010-08-18 06:52:56 +08:00
|
|
|
extern int do_execve(const char *,
|
|
|
|
|
const char __user * const __user *,
|
2012-10-21 09:49:33 +08:00
|
|
|
const char __user * const __user *);
|
2012-10-23 11:10:08 +08:00
|
|
|
extern long do_fork(unsigned long, unsigned long, unsigned long, int __user *, int __user *);
|
2006-07-03 15:25:41 +08:00
|
|
|
struct task_struct *fork_idle(int);
|
2012-09-22 07:55:31 +08:00
|
|
|
extern pid_t kernel_thread(int (*fn)(void *), void *arg, unsigned long flags);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
extern void set_task_comm(struct task_struct *tsk, char *from);
|
2008-02-05 14:27:21 +08:00
|
|
|
extern char *get_task_comm(char *to, struct task_struct *tsk);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
#ifdef CONFIG_SMP
|
2011-04-05 23:23:58 +08:00
|
|
|
void scheduler_ipi(void);
|
2008-07-26 10:45:58 +08:00
|
|
|
extern unsigned long wait_task_inactive(struct task_struct *, long match_state);
|
2005-04-17 06:20:36 +08:00
|
|
|
#else
|
2011-04-05 23:23:39 +08:00
|
|
|
static inline void scheduler_ipi(void) { }
|
2008-07-26 10:45:58 +08:00
|
|
|
static inline unsigned long wait_task_inactive(struct task_struct *p,
|
|
|
|
|
long match_state)
|
|
|
|
|
{
|
|
|
|
|
return 1;
|
|
|
|
|
}
|
2005-04-17 06:20:36 +08:00
|
|
|
#endif
|
|
|
|
|
|
2009-04-15 02:17:16 +08:00
|
|
|
#define next_task(p) \
|
|
|
|
|
list_entry_rcu((p)->tasks.next, struct task_struct, tasks)
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
#define for_each_process(p) \
|
|
|
|
|
for (p = &init_task ; (p = next_task(p)) != &init_task ; )
|
|
|
|
|
|
2009-07-10 09:48:23 +08:00
|
|
|
extern bool current_is_single_threaded(void);
|
CRED: Inaugurate COW credentials
Inaugurate copy-on-write credentials management. This uses RCU to manage the
credentials pointer in the task_struct with respect to accesses by other tasks.
A process may only modify its own credentials, and so does not need locking to
access or modify its own credentials.
A mutex (cred_replace_mutex) is added to the task_struct to control the effect
of PTRACE_ATTACHED on credential calculations, particularly with respect to
execve().
With this patch, the contents of an active credentials struct may not be
changed directly; rather a new set of credentials must be prepared, modified
and committed using something like the following sequence of events:
struct cred *new = prepare_creds();
int ret = blah(new);
if (ret < 0) {
abort_creds(new);
return ret;
}
return commit_creds(new);
There are some exceptions to this rule: the keyrings pointed to by the active
credentials may be instantiated - keyrings violate the COW rule as managing
COW keyrings is tricky, given that it is possible for a task to directly alter
the keys in a keyring in use by another task.
To help enforce this, various pointers to sets of credentials, such as those in
the task_struct, are declared const. The purpose of this is compile-time
discouragement of altering credentials through those pointers. Once a set of
credentials has been made public through one of these pointers, it may not be
modified, except under special circumstances:
(1) Its reference count may incremented and decremented.
(2) The keyrings to which it points may be modified, but not replaced.
The only safe way to modify anything else is to create a replacement and commit
using the functions described in Documentation/credentials.txt (which will be
added by a later patch).
This patch and the preceding patches have been tested with the LTP SELinux
testsuite.
This patch makes several logical sets of alteration:
(1) execve().
This now prepares and commits credentials in various places in the
security code rather than altering the current creds directly.
(2) Temporary credential overrides.
do_coredump() and sys_faccessat() now prepare their own credentials and
temporarily override the ones currently on the acting thread, whilst
preventing interference from other threads by holding cred_replace_mutex
on the thread being dumped.
This will be replaced in a future patch by something that hands down the
credentials directly to the functions being called, rather than altering
the task's objective credentials.
(3) LSM interface.
A number of functions have been changed, added or removed:
(*) security_capset_check(), ->capset_check()
(*) security_capset_set(), ->capset_set()
Removed in favour of security_capset().
(*) security_capset(), ->capset()
New. This is passed a pointer to the new creds, a pointer to the old
creds and the proposed capability sets. It should fill in the new
creds or return an error. All pointers, barring the pointer to the
new creds, are now const.
(*) security_bprm_apply_creds(), ->bprm_apply_creds()
Changed; now returns a value, which will cause the process to be
killed if it's an error.
(*) security_task_alloc(), ->task_alloc_security()
Removed in favour of security_prepare_creds().
(*) security_cred_free(), ->cred_free()
New. Free security data attached to cred->security.
(*) security_prepare_creds(), ->cred_prepare()
New. Duplicate any security data attached to cred->security.
(*) security_commit_creds(), ->cred_commit()
New. Apply any security effects for the upcoming installation of new
security by commit_creds().
(*) security_task_post_setuid(), ->task_post_setuid()
Removed in favour of security_task_fix_setuid().
(*) security_task_fix_setuid(), ->task_fix_setuid()
Fix up the proposed new credentials for setuid(). This is used by
cap_set_fix_setuid() to implicitly adjust capabilities in line with
setuid() changes. Changes are made to the new credentials, rather
than the task itself as in security_task_post_setuid().
(*) security_task_reparent_to_init(), ->task_reparent_to_init()
Removed. Instead the task being reparented to init is referred
directly to init's credentials.
NOTE! This results in the loss of some state: SELinux's osid no
longer records the sid of the thread that forked it.
(*) security_key_alloc(), ->key_alloc()
(*) security_key_permission(), ->key_permission()
Changed. These now take cred pointers rather than task pointers to
refer to the security context.
(4) sys_capset().
This has been simplified and uses less locking. The LSM functions it
calls have been merged.
(5) reparent_to_kthreadd().
This gives the current thread the same credentials as init by simply using
commit_thread() to point that way.
(6) __sigqueue_alloc() and switch_uid()
__sigqueue_alloc() can't stop the target task from changing its creds
beneath it, so this function gets a reference to the currently applicable
user_struct which it then passes into the sigqueue struct it returns if
successful.
switch_uid() is now called from commit_creds(), and possibly should be
folded into that. commit_creds() should take care of protecting
__sigqueue_alloc().
(7) [sg]et[ug]id() and co and [sg]et_current_groups.
The set functions now all use prepare_creds(), commit_creds() and
abort_creds() to build and check a new set of credentials before applying
it.
security_task_set[ug]id() is called inside the prepared section. This
guarantees that nothing else will affect the creds until we've finished.
The calling of set_dumpable() has been moved into commit_creds().
Much of the functionality of set_user() has been moved into
commit_creds().
The get functions all simply access the data directly.
(8) security_task_prctl() and cap_task_prctl().
security_task_prctl() has been modified to return -ENOSYS if it doesn't
want to handle a function, or otherwise return the return value directly
rather than through an argument.
Additionally, cap_task_prctl() now prepares a new set of credentials, even
if it doesn't end up using it.
(9) Keyrings.
A number of changes have been made to the keyrings code:
(a) switch_uid_keyring(), copy_keys(), exit_keys() and suid_keys() have
all been dropped and built in to the credentials functions directly.
They may want separating out again later.
(b) key_alloc() and search_process_keyrings() now take a cred pointer
rather than a task pointer to specify the security context.
(c) copy_creds() gives a new thread within the same thread group a new
thread keyring if its parent had one, otherwise it discards the thread
keyring.
(d) The authorisation key now points directly to the credentials to extend
the search into rather pointing to the task that carries them.
(e) Installing thread, process or session keyrings causes a new set of
credentials to be created, even though it's not strictly necessary for
process or session keyrings (they're shared).
(10) Usermode helper.
The usermode helper code now carries a cred struct pointer in its
subprocess_info struct instead of a new session keyring pointer. This set
of credentials is derived from init_cred and installed on the new process
after it has been cloned.
call_usermodehelper_setup() allocates the new credentials and
call_usermodehelper_freeinfo() discards them if they haven't been used. A
special cred function (prepare_usermodeinfo_creds()) is provided
specifically for call_usermodehelper_setup() to call.
call_usermodehelper_setkeys() adjusts the credentials to sport the
supplied keyring as the new session keyring.
(11) SELinux.
SELinux has a number of changes, in addition to those to support the LSM
interface changes mentioned above:
(a) selinux_setprocattr() no longer does its check for whether the
current ptracer can access processes with the new SID inside the lock
that covers getting the ptracer's SID. Whilst this lock ensures that
the check is done with the ptracer pinned, the result is only valid
until the lock is released, so there's no point doing it inside the
lock.
(12) is_single_threaded().
This function has been extracted from selinux_setprocattr() and put into
a file of its own in the lib/ directory as join_session_keyring() now
wants to use it too.
The code in SELinux just checked to see whether a task shared mm_structs
with other tasks (CLONE_VM), but that isn't good enough. We really want
to know if they're part of the same thread group (CLONE_THREAD).
(13) nfsd.
The NFS server daemon now has to use the COW credentials to set the
credentials it is going to use. It really needs to pass the credentials
down to the functions it calls, but it can't do that until other patches
in this series have been applied.
Signed-off-by: David Howells <dhowells@redhat.com>
Acked-by: James Morris <jmorris@namei.org>
Signed-off-by: James Morris <jmorris@namei.org>
2008-11-14 07:39:23 +08:00
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
|
|
|
|
* Careful: do_each_thread/while_each_thread is a double loop so
|
|
|
|
|
* 'break' will not work as expected - use goto instead.
|
|
|
|
|
*/
|
|
|
|
|
#define do_each_thread(g, t) \
|
|
|
|
|
for (g = t = &init_task ; (g = t = next_task(g)) != &init_task ; ) do
|
|
|
|
|
|
|
|
|
|
#define while_each_thread(g, t) \
|
|
|
|
|
while ((t = next_thread(t)) != g)
|
|
|
|
|
|
2010-05-27 05:43:22 +08:00
|
|
|
static inline int get_nr_threads(struct task_struct *tsk)
|
|
|
|
|
{
|
2010-05-27 05:43:24 +08:00
|
|
|
return tsk->signal->nr_threads;
|
2010-05-27 05:43:22 +08:00
|
|
|
}
|
|
|
|
|
|
2011-06-23 05:10:26 +08:00
|
|
|
static inline bool thread_group_leader(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return p->exit_signal >= 0;
|
|
|
|
|
}
|
2005-04-17 06:20:36 +08:00
|
|
|
|
[PATCH] proc: readdir race fix (take 3)
The problem: An opendir, readdir, closedir sequence can fail to report
process ids that are continually in use throughout the sequence of system
calls. For this race to trigger the process that proc_pid_readdir stops at
must exit before readdir is called again.
This can cause ps to fail to report processes, and it is in violation of
posix guarantees and normal application expectations with respect to
readdir.
Currently there is no way to work around this problem in user space short
of providing a gargantuan buffer to user space so the directory read all
happens in on system call.
This patch implements the normal directory semantics for proc, that
guarantee that a directory entry that is neither created nor destroyed
while reading the directory entry will be returned. For directory that are
either created or destroyed during the readdir you may or may not see them.
Furthermore you may seek to a directory offset you have previously seen.
These are the guarantee that ext[23] provides and that posix requires, and
more importantly that user space expects. Plus it is a simple semantic to
implement reliable service. It is just a matter of calling readdir a
second time if you are wondering if something new has show up.
These better semantics are implemented by scanning through the pids in
numerical order and by making the file offset a pid plus a fixed offset.
The pid scan happens on the pid bitmap, which when you look at it is
remarkably efficient for a brute force algorithm. Given that a typical
cache line is 64 bytes and thus covers space for 64*8 == 200 pids. There
are only 40 cache lines for the entire 32K pid space. A typical system
will have 100 pids or more so this is actually fewer cache lines we have to
look at to scan a linked list, and the worst case of having to scan the
entire pid bitmap is pretty reasonable.
If we need something more efficient we can go to a more efficient data
structure for indexing the pids, but for now what we have should be
sufficient.
In addition this takes no additional locks and is actually less code than
what we are doing now.
Also another very subtle bug in this area has been fixed. It is possible
to catch a task in the middle of de_thread where a thread is assuming the
thread of it's thread group leader. This patch carefully handles that case
so if we hit it we don't fail to return the pid, that is undergoing the
de_thread dance.
Thanks to KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> for
providing the first fix, pointing this out and working on it.
[oleg@tv-sign.ru: fix it]
Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Jean Delvare <jdelvare@suse.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-10-02 17:17:04 +08:00
|
|
|
/* Do to the insanities of de_thread it is possible for a process
|
|
|
|
|
* to have the pid of the thread group leader without actually being
|
|
|
|
|
* the thread group leader. For iteration through the pids in proc
|
|
|
|
|
* all we care about is that we have a task with the appropriate
|
|
|
|
|
* pid, we don't actually care if we have the right task.
|
|
|
|
|
*/
|
2013-09-12 05:20:06 +08:00
|
|
|
static inline bool has_group_leader_pid(struct task_struct *p)
|
[PATCH] proc: readdir race fix (take 3)
The problem: An opendir, readdir, closedir sequence can fail to report
process ids that are continually in use throughout the sequence of system
calls. For this race to trigger the process that proc_pid_readdir stops at
must exit before readdir is called again.
This can cause ps to fail to report processes, and it is in violation of
posix guarantees and normal application expectations with respect to
readdir.
Currently there is no way to work around this problem in user space short
of providing a gargantuan buffer to user space so the directory read all
happens in on system call.
This patch implements the normal directory semantics for proc, that
guarantee that a directory entry that is neither created nor destroyed
while reading the directory entry will be returned. For directory that are
either created or destroyed during the readdir you may or may not see them.
Furthermore you may seek to a directory offset you have previously seen.
These are the guarantee that ext[23] provides and that posix requires, and
more importantly that user space expects. Plus it is a simple semantic to
implement reliable service. It is just a matter of calling readdir a
second time if you are wondering if something new has show up.
These better semantics are implemented by scanning through the pids in
numerical order and by making the file offset a pid plus a fixed offset.
The pid scan happens on the pid bitmap, which when you look at it is
remarkably efficient for a brute force algorithm. Given that a typical
cache line is 64 bytes and thus covers space for 64*8 == 200 pids. There
are only 40 cache lines for the entire 32K pid space. A typical system
will have 100 pids or more so this is actually fewer cache lines we have to
look at to scan a linked list, and the worst case of having to scan the
entire pid bitmap is pretty reasonable.
If we need something more efficient we can go to a more efficient data
structure for indexing the pids, but for now what we have should be
sufficient.
In addition this takes no additional locks and is actually less code than
what we are doing now.
Also another very subtle bug in this area has been fixed. It is possible
to catch a task in the middle of de_thread where a thread is assuming the
thread of it's thread group leader. This patch carefully handles that case
so if we hit it we don't fail to return the pid, that is undergoing the
de_thread dance.
Thanks to KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> for
providing the first fix, pointing this out and working on it.
[oleg@tv-sign.ru: fix it]
Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Jean Delvare <jdelvare@suse.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-10-02 17:17:04 +08:00
|
|
|
{
|
2013-09-12 05:20:06 +08:00
|
|
|
return task_pid(p) == p->signal->leader_pid;
|
[PATCH] proc: readdir race fix (take 3)
The problem: An opendir, readdir, closedir sequence can fail to report
process ids that are continually in use throughout the sequence of system
calls. For this race to trigger the process that proc_pid_readdir stops at
must exit before readdir is called again.
This can cause ps to fail to report processes, and it is in violation of
posix guarantees and normal application expectations with respect to
readdir.
Currently there is no way to work around this problem in user space short
of providing a gargantuan buffer to user space so the directory read all
happens in on system call.
This patch implements the normal directory semantics for proc, that
guarantee that a directory entry that is neither created nor destroyed
while reading the directory entry will be returned. For directory that are
either created or destroyed during the readdir you may or may not see them.
Furthermore you may seek to a directory offset you have previously seen.
These are the guarantee that ext[23] provides and that posix requires, and
more importantly that user space expects. Plus it is a simple semantic to
implement reliable service. It is just a matter of calling readdir a
second time if you are wondering if something new has show up.
These better semantics are implemented by scanning through the pids in
numerical order and by making the file offset a pid plus a fixed offset.
The pid scan happens on the pid bitmap, which when you look at it is
remarkably efficient for a brute force algorithm. Given that a typical
cache line is 64 bytes and thus covers space for 64*8 == 200 pids. There
are only 40 cache lines for the entire 32K pid space. A typical system
will have 100 pids or more so this is actually fewer cache lines we have to
look at to scan a linked list, and the worst case of having to scan the
entire pid bitmap is pretty reasonable.
If we need something more efficient we can go to a more efficient data
structure for indexing the pids, but for now what we have should be
sufficient.
In addition this takes no additional locks and is actually less code than
what we are doing now.
Also another very subtle bug in this area has been fixed. It is possible
to catch a task in the middle of de_thread where a thread is assuming the
thread of it's thread group leader. This patch carefully handles that case
so if we hit it we don't fail to return the pid, that is undergoing the
de_thread dance.
Thanks to KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> for
providing the first fix, pointing this out and working on it.
[oleg@tv-sign.ru: fix it]
Signed-off-by: Eric W. Biederman <ebiederm@xmission.com>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Cc: Jean Delvare <jdelvare@suse.de>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
2006-10-02 17:17:04 +08:00
|
|
|
}
|
|
|
|
|
|
2007-10-19 14:40:18 +08:00
|
|
|
static inline
|
2013-09-12 05:20:06 +08:00
|
|
|
bool same_thread_group(struct task_struct *p1, struct task_struct *p2)
|
2007-10-19 14:40:18 +08:00
|
|
|
{
|
2013-09-12 05:20:06 +08:00
|
|
|
return p1->signal == p2->signal;
|
2007-10-19 14:40:18 +08:00
|
|
|
}
|
|
|
|
|
|
2006-07-03 15:25:41 +08:00
|
|
|
static inline struct task_struct *next_thread(const struct task_struct *p)
|
2006-03-29 08:11:25 +08:00
|
|
|
{
|
2009-04-15 02:17:16 +08:00
|
|
|
return list_entry_rcu(p->thread_group.next,
|
|
|
|
|
struct task_struct, thread_group);
|
2006-03-29 08:11:25 +08:00
|
|
|
}
|
|
|
|
|
|
2007-10-26 16:17:22 +08:00
|
|
|
static inline int thread_group_empty(struct task_struct *p)
|
2005-04-17 06:20:36 +08:00
|
|
|
{
|
2006-03-29 08:11:25 +08:00
|
|
|
return list_empty(&p->thread_group);
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
#define delay_group_leader(p) \
|
|
|
|
|
(thread_group_leader(p) && !thread_group_empty(p))
|
|
|
|
|
|
|
|
|
|
/*
|
2006-06-23 17:05:18 +08:00
|
|
|
* Protects ->fs, ->files, ->mm, ->group_info, ->comm, keyring
|
2005-06-27 16:55:12 +08:00
|
|
|
* subscriptions and synchronises with wait4(). Also used in procfs. Also
|
Task Control Groups: basic task cgroup framework
Generic Process Control Groups
--------------------------
There have recently been various proposals floating around for
resource management/accounting and other task grouping subsystems in
the kernel, including ResGroups, User BeanCounters, NSProxy
cgroups, and others. These all need the basic abstraction of being
able to group together multiple processes in an aggregate, in order to
track/limit the resources permitted to those processes, or control
other behaviour of the processes, and all implement this grouping in
different ways.
This patchset provides a framework for tracking and grouping processes
into arbitrary "cgroups" and assigning arbitrary state to those
groupings, in order to control the behaviour of the cgroup as an
aggregate.
The intention is that the various resource management and
virtualization/cgroup efforts can also become task cgroup
clients, with the result that:
- the userspace APIs are (somewhat) normalised
- it's easier to test e.g. the ResGroups CPU controller in
conjunction with the BeanCounters memory controller, or use either of
them as the resource-control portion of a virtual server system.
- the additional kernel footprint of any of the competing resource
management systems is substantially reduced, since it doesn't need
to provide process grouping/containment, hence improving their
chances of getting into the kernel
This patch:
Add the main task cgroups framework - the cgroup filesystem, and the
basic structures for tracking membership and associating subsystem state
objects to tasks.
Signed-off-by: Paul Menage <menage@google.com>
Cc: Serge E. Hallyn <serue@us.ibm.com>
Cc: "Eric W. Biederman" <ebiederm@xmission.com>
Cc: Dave Hansen <haveblue@us.ibm.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Cc: Paul Jackson <pj@sgi.com>
Cc: Kirill Korotaev <dev@openvz.org>
Cc: Herbert Poetzl <herbert@13thfloor.at>
Cc: Srivatsa Vaddagiri <vatsa@in.ibm.com>
Cc: Cedric Le Goater <clg@fr.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-10-19 14:39:30 +08:00
|
|
|
* pins the final release of task.io_context. Also protects ->cpuset and
|
2012-03-06 06:59:13 +08:00
|
|
|
* ->cgroup.subsys[]. And ->vfork_done.
|
2005-04-17 06:20:36 +08:00
|
|
|
*
|
|
|
|
|
* Nests both inside and outside of read_lock(&tasklist_lock).
|
|
|
|
|
* It must not be nested with write_lock_irq(&tasklist_lock),
|
|
|
|
|
* neither inside nor outside.
|
|
|
|
|
*/
|
|
|
|
|
static inline void task_lock(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
spin_lock(&p->alloc_lock);
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void task_unlock(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
spin_unlock(&p->alloc_lock);
|
|
|
|
|
}
|
|
|
|
|
|
2010-10-28 06:34:06 +08:00
|
|
|
extern struct sighand_struct *__lock_task_sighand(struct task_struct *tsk,
|
2006-03-29 08:11:13 +08:00
|
|
|
unsigned long *flags);
|
|
|
|
|
|
2012-02-10 00:45:19 +08:00
|
|
|
static inline struct sighand_struct *lock_task_sighand(struct task_struct *tsk,
|
|
|
|
|
unsigned long *flags)
|
|
|
|
|
{
|
|
|
|
|
struct sighand_struct *ret;
|
|
|
|
|
|
|
|
|
|
ret = __lock_task_sighand(tsk, flags);
|
|
|
|
|
(void)__cond_lock(&tsk->sighand->siglock, ret);
|
|
|
|
|
return ret;
|
|
|
|
|
}
|
2010-10-28 06:34:06 +08:00
|
|
|
|
2006-03-29 08:11:13 +08:00
|
|
|
static inline void unlock_task_sighand(struct task_struct *tsk,
|
|
|
|
|
unsigned long *flags)
|
|
|
|
|
{
|
|
|
|
|
spin_unlock_irqrestore(&tsk->sighand->siglock, *flags);
|
|
|
|
|
}
|
|
|
|
|
|
2011-05-27 07:25:18 +08:00
|
|
|
#ifdef CONFIG_CGROUPS
|
2011-12-13 10:12:21 +08:00
|
|
|
static inline void threadgroup_change_begin(struct task_struct *tsk)
|
2011-05-27 07:25:18 +08:00
|
|
|
{
|
2011-12-13 10:12:21 +08:00
|
|
|
down_read(&tsk->signal->group_rwsem);
|
2011-05-27 07:25:18 +08:00
|
|
|
}
|
2011-12-13 10:12:21 +08:00
|
|
|
static inline void threadgroup_change_end(struct task_struct *tsk)
|
2011-05-27 07:25:18 +08:00
|
|
|
{
|
2011-12-13 10:12:21 +08:00
|
|
|
up_read(&tsk->signal->group_rwsem);
|
2011-05-27 07:25:18 +08:00
|
|
|
}
|
2011-12-13 10:12:21 +08:00
|
|
|
|
|
|
|
|
/**
|
|
|
|
|
* threadgroup_lock - lock threadgroup
|
|
|
|
|
* @tsk: member task of the threadgroup to lock
|
|
|
|
|
*
|
|
|
|
|
* Lock the threadgroup @tsk belongs to. No new task is allowed to enter
|
|
|
|
|
* and member tasks aren't allowed to exit (as indicated by PF_EXITING) or
|
2013-05-01 06:28:20 +08:00
|
|
|
* change ->group_leader/pid. This is useful for cases where the threadgroup
|
|
|
|
|
* needs to stay stable across blockable operations.
|
2011-12-13 10:12:21 +08:00
|
|
|
*
|
|
|
|
|
* fork and exit paths explicitly call threadgroup_change_{begin|end}() for
|
|
|
|
|
* synchronization. While held, no new task will be added to threadgroup
|
|
|
|
|
* and no existing live task will have its PF_EXITING set.
|
|
|
|
|
*
|
2013-05-01 06:28:20 +08:00
|
|
|
* de_thread() does threadgroup_change_{begin|end}() when a non-leader
|
|
|
|
|
* sub-thread becomes a new leader.
|
2011-12-13 10:12:21 +08:00
|
|
|
*/
|
2011-12-13 10:12:21 +08:00
|
|
|
static inline void threadgroup_lock(struct task_struct *tsk)
|
2011-05-27 07:25:18 +08:00
|
|
|
{
|
2011-12-13 10:12:21 +08:00
|
|
|
down_write(&tsk->signal->group_rwsem);
|
2011-05-27 07:25:18 +08:00
|
|
|
}
|
2011-12-13 10:12:21 +08:00
|
|
|
|
|
|
|
|
/**
|
|
|
|
|
* threadgroup_unlock - unlock threadgroup
|
|
|
|
|
* @tsk: member task of the threadgroup to unlock
|
|
|
|
|
*
|
|
|
|
|
* Reverse threadgroup_lock().
|
|
|
|
|
*/
|
2011-12-13 10:12:21 +08:00
|
|
|
static inline void threadgroup_unlock(struct task_struct *tsk)
|
2011-05-27 07:25:18 +08:00
|
|
|
{
|
2011-12-13 10:12:21 +08:00
|
|
|
up_write(&tsk->signal->group_rwsem);
|
2011-05-27 07:25:18 +08:00
|
|
|
}
|
|
|
|
|
#else
|
2011-12-13 10:12:21 +08:00
|
|
|
static inline void threadgroup_change_begin(struct task_struct *tsk) {}
|
|
|
|
|
static inline void threadgroup_change_end(struct task_struct *tsk) {}
|
|
|
|
|
static inline void threadgroup_lock(struct task_struct *tsk) {}
|
|
|
|
|
static inline void threadgroup_unlock(struct task_struct *tsk) {}
|
2011-05-27 07:25:18 +08:00
|
|
|
#endif
|
|
|
|
|
|
2005-11-14 08:06:57 +08:00
|
|
|
#ifndef __HAVE_THREAD_FUNCTIONS
|
|
|
|
|
|
2007-05-09 17:35:17 +08:00
|
|
|
#define task_thread_info(task) ((struct thread_info *)(task)->stack)
|
|
|
|
|
#define task_stack_page(task) ((task)->stack)
|
2005-11-14 08:06:55 +08:00
|
|
|
|
2005-11-14 08:06:56 +08:00
|
|
|
static inline void setup_thread_stack(struct task_struct *p, struct task_struct *org)
|
|
|
|
|
{
|
|
|
|
|
*task_thread_info(p) = *task_thread_info(org);
|
|
|
|
|
task_thread_info(p)->task = p;
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline unsigned long *end_of_stack(struct task_struct *p)
|
|
|
|
|
{
|
2007-05-09 17:35:17 +08:00
|
|
|
return (unsigned long *)(task_thread_info(p) + 1);
|
2005-11-14 08:06:56 +08:00
|
|
|
}
|
|
|
|
|
|
2005-11-14 08:06:57 +08:00
|
|
|
#endif
|
|
|
|
|
|
2008-07-24 12:26:53 +08:00
|
|
|
static inline int object_is_on_stack(void *obj)
|
|
|
|
|
{
|
|
|
|
|
void *stack = task_stack_page(current);
|
|
|
|
|
|
|
|
|
|
return (obj >= stack) && (obj < (stack + THREAD_SIZE));
|
|
|
|
|
}
|
|
|
|
|
|
2008-04-18 14:56:15 +08:00
|
|
|
extern void thread_info_cache_init(void);
|
|
|
|
|
|
2008-04-23 05:38:23 +08:00
|
|
|
#ifdef CONFIG_DEBUG_STACK_USAGE
|
|
|
|
|
static inline unsigned long stack_not_used(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
unsigned long *n = end_of_stack(p);
|
|
|
|
|
|
|
|
|
|
do { /* Skip over canary */
|
|
|
|
|
n++;
|
|
|
|
|
} while (!*n);
|
|
|
|
|
|
|
|
|
|
return (unsigned long)n - (unsigned long)end_of_stack(p);
|
|
|
|
|
}
|
|
|
|
|
#endif
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/* set thread flags in other task's structures
|
|
|
|
|
* - see asm/thread_info.h for TIF_xxxx flags available
|
|
|
|
|
*/
|
|
|
|
|
static inline void set_tsk_thread_flag(struct task_struct *tsk, int flag)
|
|
|
|
|
{
|
2005-11-14 08:06:55 +08:00
|
|
|
set_ti_thread_flag(task_thread_info(tsk), flag);
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void clear_tsk_thread_flag(struct task_struct *tsk, int flag)
|
|
|
|
|
{
|
2005-11-14 08:06:55 +08:00
|
|
|
clear_ti_thread_flag(task_thread_info(tsk), flag);
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline int test_and_set_tsk_thread_flag(struct task_struct *tsk, int flag)
|
|
|
|
|
{
|
2005-11-14 08:06:55 +08:00
|
|
|
return test_and_set_ti_thread_flag(task_thread_info(tsk), flag);
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline int test_and_clear_tsk_thread_flag(struct task_struct *tsk, int flag)
|
|
|
|
|
{
|
2005-11-14 08:06:55 +08:00
|
|
|
return test_and_clear_ti_thread_flag(task_thread_info(tsk), flag);
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline int test_tsk_thread_flag(struct task_struct *tsk, int flag)
|
|
|
|
|
{
|
2005-11-14 08:06:55 +08:00
|
|
|
return test_ti_thread_flag(task_thread_info(tsk), flag);
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void set_tsk_need_resched(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
set_tsk_thread_flag(tsk,TIF_NEED_RESCHED);
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void clear_tsk_need_resched(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
clear_tsk_thread_flag(tsk,TIF_NEED_RESCHED);
|
|
|
|
|
}
|
|
|
|
|
|
2008-04-23 19:13:29 +08:00
|
|
|
static inline int test_tsk_need_resched(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
return unlikely(test_tsk_thread_flag(tsk,TIF_NEED_RESCHED));
|
|
|
|
|
}
|
|
|
|
|
|
2009-05-14 00:55:10 +08:00
|
|
|
static inline int restart_syscall(void)
|
|
|
|
|
{
|
|
|
|
|
set_tsk_thread_flag(current, TIF_SIGPENDING);
|
|
|
|
|
return -ERESTARTNOINTR;
|
|
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
static inline int signal_pending(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return unlikely(test_tsk_thread_flag(p,TIF_SIGPENDING));
|
|
|
|
|
}
|
2007-12-07 00:15:50 +08:00
|
|
|
|
2009-09-24 06:57:04 +08:00
|
|
|
static inline int __fatal_signal_pending(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return unlikely(sigismember(&p->pending.signal, SIGKILL));
|
|
|
|
|
}
|
2007-12-07 00:15:50 +08:00
|
|
|
|
|
|
|
|
static inline int fatal_signal_pending(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return signal_pending(p) && __fatal_signal_pending(p);
|
|
|
|
|
}
|
|
|
|
|
|
sched: fix TASK_WAKEKILL vs SIGKILL race
schedule() has the special "TASK_INTERRUPTIBLE && signal_pending()" case,
this allows us to do
current->state = TASK_INTERRUPTIBLE;
schedule();
without fear to sleep with pending signal.
However, the code like
current->state = TASK_KILLABLE;
schedule();
is not right, schedule() doesn't take TASK_WAKEKILL into account. This means
that mutex_lock_killable(), wait_for_completion_killable(), down_killable(),
schedule_timeout_killable() can miss SIGKILL (and btw the second SIGKILL has
no effect).
Introduce the new helper, signal_pending_state(), and change schedule() to
use it. Hopefully it will have more users, that is why the task's state is
passed separately.
Note this "__TASK_STOPPED | __TASK_TRACED" check in signal_pending_state().
This is needed to preserve the current behaviour (ptrace_notify). I hope
this check will be removed soon, but this (afaics good) change needs the
separate discussion.
The fast path is "(state & (INTERRUPTIBLE | WAKEKILL)) + signal_pending(p)",
basically the same that schedule() does now. However, this patch of course
bloats schedule().
Signed-off-by: Oleg Nesterov <oleg@tv-sign.ru>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-06-09 01:20:41 +08:00
|
|
|
static inline int signal_pending_state(long state, struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
if (!(state & (TASK_INTERRUPTIBLE | TASK_WAKEKILL)))
|
|
|
|
|
return 0;
|
|
|
|
|
if (!signal_pending(p))
|
|
|
|
|
return 0;
|
|
|
|
|
|
|
|
|
|
return (state & TASK_INTERRUPTIBLE) || __fatal_signal_pending(p);
|
|
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
|
|
|
|
* cond_resched() and cond_resched_lock(): latency reduction via
|
|
|
|
|
* explicit rescheduling in places that are safe. The return
|
|
|
|
|
* value indicates whether a reschedule was done in fact.
|
|
|
|
|
* cond_resched_lock() will drop the spinlock before scheduling,
|
|
|
|
|
* cond_resched_softirq() will enable bhs before scheduling.
|
|
|
|
|
*/
|
2008-05-12 07:04:48 +08:00
|
|
|
extern int _cond_resched(void);
|
2009-07-16 21:44:29 +08:00
|
|
|
|
2009-07-16 21:44:29 +08:00
|
|
|
#define cond_resched() ({ \
|
|
|
|
|
__might_sleep(__FILE__, __LINE__, 0); \
|
|
|
|
|
_cond_resched(); \
|
|
|
|
|
})
|
2009-07-16 21:44:29 +08:00
|
|
|
|
2009-07-16 21:44:29 +08:00
|
|
|
extern int __cond_resched_lock(spinlock_t *lock);
|
|
|
|
|
|
2011-06-08 07:13:27 +08:00
|
|
|
#ifdef CONFIG_PREEMPT_COUNT
|
2009-07-25 02:05:23 +08:00
|
|
|
#define PREEMPT_LOCK_OFFSET PREEMPT_OFFSET
|
2008-01-26 04:08:28 +08:00
|
|
|
#else
|
2009-07-25 02:05:23 +08:00
|
|
|
#define PREEMPT_LOCK_OFFSET 0
|
2008-01-26 04:08:28 +08:00
|
|
|
#endif
|
2009-07-25 02:05:23 +08:00
|
|
|
|
2009-07-16 21:44:29 +08:00
|
|
|
#define cond_resched_lock(lock) ({ \
|
2009-07-25 02:05:23 +08:00
|
|
|
__might_sleep(__FILE__, __LINE__, PREEMPT_LOCK_OFFSET); \
|
2009-07-16 21:44:29 +08:00
|
|
|
__cond_resched_lock(lock); \
|
|
|
|
|
})
|
|
|
|
|
|
|
|
|
|
extern int __cond_resched_softirq(void);
|
|
|
|
|
|
2010-10-05 08:03:16 +08:00
|
|
|
#define cond_resched_softirq() ({ \
|
|
|
|
|
__might_sleep(__FILE__, __LINE__, SOFTIRQ_DISABLE_OFFSET); \
|
|
|
|
|
__cond_resched_softirq(); \
|
2009-07-16 21:44:29 +08:00
|
|
|
})
|
2005-04-17 06:20:36 +08:00
|
|
|
|
2013-05-22 13:50:31 +08:00
|
|
|
static inline void cond_resched_rcu(void)
|
|
|
|
|
{
|
|
|
|
|
#if defined(CONFIG_DEBUG_ATOMIC_SLEEP) || !defined(CONFIG_PREEMPT_RCU)
|
|
|
|
|
rcu_read_unlock();
|
|
|
|
|
cond_resched();
|
|
|
|
|
rcu_read_lock();
|
|
|
|
|
#endif
|
|
|
|
|
}
|
|
|
|
|
|
2005-04-17 06:20:36 +08:00
|
|
|
/*
|
|
|
|
|
* Does a critical section need to be broken due to another
|
2008-01-30 20:31:20 +08:00
|
|
|
* task waiting?: (technically does not depend on CONFIG_PREEMPT,
|
|
|
|
|
* but a general need for low latency)
|
2005-04-17 06:20:36 +08:00
|
|
|
*/
|
2008-01-30 20:31:20 +08:00
|
|
|
static inline int spin_needbreak(spinlock_t *lock)
|
2005-04-17 06:20:36 +08:00
|
|
|
{
|
2008-01-30 20:31:20 +08:00
|
|
|
#ifdef CONFIG_PREEMPT
|
|
|
|
|
return spin_is_contended(lock);
|
|
|
|
|
#else
|
2005-04-17 06:20:36 +08:00
|
|
|
return 0;
|
2008-01-30 20:31:20 +08:00
|
|
|
#endif
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
2013-03-22 05:49:32 +08:00
|
|
|
/*
|
|
|
|
|
* Idle thread specific functions to determine the need_resched
|
|
|
|
|
* polling state. We have two versions, one based on TS_POLLING in
|
|
|
|
|
* thread_info.status and one based on TIF_POLLING_NRFLAG in
|
|
|
|
|
* thread_info.flags
|
|
|
|
|
*/
|
|
|
|
|
#ifdef TS_POLLING
|
|
|
|
|
static inline int tsk_is_polling(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return task_thread_info(p)->status & TS_POLLING;
|
|
|
|
|
}
|
2013-09-11 18:43:13 +08:00
|
|
|
static inline void __current_set_polling(void)
|
2013-03-22 05:49:33 +08:00
|
|
|
{
|
|
|
|
|
current_thread_info()->status |= TS_POLLING;
|
|
|
|
|
}
|
|
|
|
|
|
2013-09-11 18:43:13 +08:00
|
|
|
static inline bool __must_check current_set_polling_and_test(void)
|
|
|
|
|
{
|
|
|
|
|
__current_set_polling();
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Polling state must be visible before we test NEED_RESCHED,
|
|
|
|
|
* paired by resched_task()
|
|
|
|
|
*/
|
|
|
|
|
smp_mb();
|
|
|
|
|
|
|
|
|
|
return unlikely(tif_need_resched());
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void __current_clr_polling(void)
|
2013-03-22 05:49:33 +08:00
|
|
|
{
|
|
|
|
|
current_thread_info()->status &= ~TS_POLLING;
|
2013-09-11 18:43:13 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline bool __must_check current_clr_polling_and_test(void)
|
|
|
|
|
{
|
|
|
|
|
__current_clr_polling();
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Polling state must be visible before we test NEED_RESCHED,
|
|
|
|
|
* paired by resched_task()
|
|
|
|
|
*/
|
|
|
|
|
smp_mb();
|
|
|
|
|
|
|
|
|
|
return unlikely(tif_need_resched());
|
2013-03-22 05:49:33 +08:00
|
|
|
}
|
2013-03-22 05:49:32 +08:00
|
|
|
#elif defined(TIF_POLLING_NRFLAG)
|
|
|
|
|
static inline int tsk_is_polling(struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return test_tsk_thread_flag(p, TIF_POLLING_NRFLAG);
|
|
|
|
|
}
|
2013-09-11 18:43:13 +08:00
|
|
|
|
|
|
|
|
static inline void __current_set_polling(void)
|
2013-03-22 05:49:33 +08:00
|
|
|
{
|
|
|
|
|
set_thread_flag(TIF_POLLING_NRFLAG);
|
|
|
|
|
}
|
|
|
|
|
|
2013-09-11 18:43:13 +08:00
|
|
|
static inline bool __must_check current_set_polling_and_test(void)
|
|
|
|
|
{
|
|
|
|
|
__current_set_polling();
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Polling state must be visible before we test NEED_RESCHED,
|
|
|
|
|
* paired by resched_task()
|
|
|
|
|
*
|
|
|
|
|
* XXX: assumes set/clear bit are identical barrier wise.
|
|
|
|
|
*/
|
|
|
|
|
smp_mb__after_clear_bit();
|
|
|
|
|
|
|
|
|
|
return unlikely(tif_need_resched());
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void __current_clr_polling(void)
|
2013-03-22 05:49:33 +08:00
|
|
|
{
|
|
|
|
|
clear_thread_flag(TIF_POLLING_NRFLAG);
|
|
|
|
|
}
|
2013-09-11 18:43:13 +08:00
|
|
|
|
|
|
|
|
static inline bool __must_check current_clr_polling_and_test(void)
|
|
|
|
|
{
|
|
|
|
|
__current_clr_polling();
|
|
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Polling state must be visible before we test NEED_RESCHED,
|
|
|
|
|
* paired by resched_task()
|
|
|
|
|
*/
|
|
|
|
|
smp_mb__after_clear_bit();
|
|
|
|
|
|
|
|
|
|
return unlikely(tif_need_resched());
|
|
|
|
|
}
|
|
|
|
|
|
2013-03-22 05:49:32 +08:00
|
|
|
#else
|
|
|
|
|
static inline int tsk_is_polling(struct task_struct *p) { return 0; }
|
2013-09-11 18:43:13 +08:00
|
|
|
static inline void __current_set_polling(void) { }
|
|
|
|
|
static inline void __current_clr_polling(void) { }
|
|
|
|
|
|
|
|
|
|
static inline bool __must_check current_set_polling_and_test(void)
|
|
|
|
|
{
|
|
|
|
|
return unlikely(tif_need_resched());
|
|
|
|
|
}
|
|
|
|
|
static inline bool __must_check current_clr_polling_and_test(void)
|
|
|
|
|
{
|
|
|
|
|
return unlikely(tif_need_resched());
|
|
|
|
|
}
|
2013-03-22 05:49:32 +08:00
|
|
|
#endif
|
|
|
|
|
|
2013-09-27 23:30:03 +08:00
|
|
|
static __always_inline bool need_resched(void)
|
|
|
|
|
{
|
|
|
|
|
return unlikely(tif_need_resched());
|
|
|
|
|
}
|
|
|
|
|
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
/*
|
|
|
|
|
* Thread group CPU time accounting.
|
|
|
|
|
*/
|
2009-02-05 19:24:16 +08:00
|
|
|
void thread_group_cputime(struct task_struct *tsk, struct task_cputime *times);
|
2009-02-11 18:30:27 +08:00
|
|
|
void thread_group_cputimer(struct task_struct *tsk, struct task_cputime *times);
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
|
2008-11-25 00:06:57 +08:00
|
|
|
static inline void thread_group_cputime_init(struct signal_struct *sig)
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
{
|
2009-07-26 00:56:56 +08:00
|
|
|
raw_spin_lock_init(&sig->cputimer.lock);
|
timers: fix itimer/many thread hang
Overview
This patch reworks the handling of POSIX CPU timers, including the
ITIMER_PROF, ITIMER_VIRT timers and rlimit handling. It was put together
with the help of Roland McGrath, the owner and original writer of this code.
The problem we ran into, and the reason for this rework, has to do with using
a profiling timer in a process with a large number of threads. It appears
that the performance of the old implementation of run_posix_cpu_timers() was
at least O(n*3) (where "n" is the number of threads in a process) or worse.
Everything is fine with an increasing number of threads until the time taken
for that routine to run becomes the same as or greater than the tick time, at
which point things degrade rather quickly.
This patch fixes bug 9906, "Weird hang with NPTL and SIGPROF."
Code Changes
This rework corrects the implementation of run_posix_cpu_timers() to make it
run in constant time for a particular machine. (Performance may vary between
one machine and another depending upon whether the kernel is built as single-
or multiprocessor and, in the latter case, depending upon the number of
running processors.) To do this, at each tick we now update fields in
signal_struct as well as task_struct. The run_posix_cpu_timers() function
uses those fields to make its decisions.
We define a new structure, "task_cputime," to contain user, system and
scheduler times and use these in appropriate places:
struct task_cputime {
cputime_t utime;
cputime_t stime;
unsigned long long sum_exec_runtime;
};
This is included in the structure "thread_group_cputime," which is a new
substructure of signal_struct and which varies for uniprocessor versus
multiprocessor kernels. For uniprocessor kernels, it uses "task_cputime" as
a simple substructure, while for multiprocessor kernels it is a pointer:
struct thread_group_cputime {
struct task_cputime totals;
};
struct thread_group_cputime {
struct task_cputime *totals;
};
We also add a new task_cputime substructure directly to signal_struct, to
cache the earliest expiration of process-wide timers, and task_cputime also
replaces the it_*_expires fields of task_struct (used for earliest expiration
of thread timers). The "thread_group_cputime" structure contains process-wide
timers that are updated via account_user_time() and friends. In the non-SMP
case the structure is a simple aggregator; unfortunately in the SMP case that
simplicity was not achievable due to cache-line contention between CPUs (in
one measured case performance was actually _worse_ on a 16-cpu system than
the same test on a 4-cpu system, due to this contention). For SMP, the
thread_group_cputime counters are maintained as a per-cpu structure allocated
using alloc_percpu(). The timer functions update only the timer field in
the structure corresponding to the running CPU, obtained using per_cpu_ptr().
We define a set of inline functions in sched.h that we use to maintain the
thread_group_cputime structure and hide the differences between UP and SMP
implementations from the rest of the kernel. The thread_group_cputime_init()
function initializes the thread_group_cputime structure for the given task.
The thread_group_cputime_alloc() is a no-op for UP; for SMP it calls the
out-of-line function thread_group_cputime_alloc_smp() to allocate and fill
in the per-cpu structures and fields. The thread_group_cputime_free()
function, also a no-op for UP, in SMP frees the per-cpu structures. The
thread_group_cputime_clone_thread() function (also a UP no-op) for SMP calls
thread_group_cputime_alloc() if the per-cpu structures haven't yet been
allocated. The thread_group_cputime() function fills the task_cputime
structure it is passed with the contents of the thread_group_cputime fields;
in UP it's that simple but in SMP it must also safely check that tsk->signal
is non-NULL (if it is it just uses the appropriate fields of task_struct) and,
if so, sums the per-cpu values for each online CPU. Finally, the three
functions account_group_user_time(), account_group_system_time() and
account_group_exec_runtime() are used by timer functions to update the
respective fields of the thread_group_cputime structure.
Non-SMP operation is trivial and will not be mentioned further.
The per-cpu structure is always allocated when a task creates its first new
thread, via a call to thread_group_cputime_clone_thread() from copy_signal().
It is freed at process exit via a call to thread_group_cputime_free() from
cleanup_signal().
All functions that formerly summed utime/stime/sum_sched_runtime values from
from all threads in the thread group now use thread_group_cputime() to
snapshot the values in the thread_group_cputime structure or the values in
the task structure itself if the per-cpu structure hasn't been allocated.
Finally, the code in kernel/posix-cpu-timers.c has changed quite a bit.
The run_posix_cpu_timers() function has been split into a fast path and a
slow path; the former safely checks whether there are any expired thread
timers and, if not, just returns, while the slow path does the heavy lifting.
With the dedicated thread group fields, timers are no longer "rebalanced" and
the process_timer_rebalance() function and related code has gone away. All
summing loops are gone and all code that used them now uses the
thread_group_cputime() inline. When process-wide timers are set, the new
task_cputime structure in signal_struct is used to cache the earliest
expiration; this is checked in the fast path.
Performance
The fix appears not to add significant overhead to existing operations. It
generally performs the same as the current code except in two cases, one in
which it performs slightly worse (Case 5 below) and one in which it performs
very significantly better (Case 2 below). Overall it's a wash except in those
two cases.
I've since done somewhat more involved testing on a dual-core Opteron system.
Case 1: With no itimer running, for a test with 100,000 threads, the fixed
kernel took 1428.5 seconds, 513 seconds more than the unfixed system,
all of which was spent in the system. There were twice as many
voluntary context switches with the fix as without it.
Case 2: With an itimer running at .01 second ticks and 4000 threads (the most
an unmodified kernel can handle), the fixed kernel ran the test in
eight percent of the time (5.8 seconds as opposed to 70 seconds) and
had better tick accuracy (.012 seconds per tick as opposed to .023
seconds per tick).
Case 3: A 4000-thread test with an initial timer tick of .01 second and an
interval of 10,000 seconds (i.e. a timer that ticks only once) had
very nearly the same performance in both cases: 6.3 seconds elapsed
for the fixed kernel versus 5.5 seconds for the unfixed kernel.
With fewer threads (eight in these tests), the Case 1 test ran in essentially
the same time on both the modified and unmodified kernels (5.2 seconds versus
5.8 seconds). The Case 2 test ran in about the same time as well, 5.9 seconds
versus 5.4 seconds but again with much better tick accuracy, .013 seconds per
tick versus .025 seconds per tick for the unmodified kernel.
Since the fix affected the rlimit code, I also tested soft and hard CPU limits.
Case 4: With a hard CPU limit of 20 seconds and eight threads (and an itimer
running), the modified kernel was very slightly favored in that while
it killed the process in 19.997 seconds of CPU time (5.002 seconds of
wall time), only .003 seconds of that was system time, the rest was
user time. The unmodified kernel killed the process in 20.001 seconds
of CPU (5.014 seconds of wall time) of which .016 seconds was system
time. Really, though, the results were too close to call. The results
were essentially the same with no itimer running.
Case 5: With a soft limit of 20 seconds and a hard limit of 2000 seconds
(where the hard limit would never be reached) and an itimer running,
the modified kernel exhibited worse tick accuracy than the unmodified
kernel: .050 seconds/tick versus .028 seconds/tick. Otherwise,
performance was almost indistinguishable. With no itimer running this
test exhibited virtually identical behavior and times in both cases.
In times past I did some limited performance testing. those results are below.
On a four-cpu Opteron system without this fix, a sixteen-thread test executed
in 3569.991 seconds, of which user was 3568.435s and system was 1.556s. On
the same system with the fix, user and elapsed time were about the same, but
system time dropped to 0.007 seconds. Performance with eight, four and one
thread were comparable. Interestingly, the timer ticks with the fix seemed
more accurate: The sixteen-thread test with the fix received 149543 ticks
for 0.024 seconds per tick, while the same test without the fix received 58720
for 0.061 seconds per tick. Both cases were configured for an interval of
0.01 seconds. Again, the other tests were comparable. Each thread in this
test computed the primes up to 25,000,000.
I also did a test with a large number of threads, 100,000 threads, which is
impossible without the fix. In this case each thread computed the primes only
up to 10,000 (to make the runtime manageable). System time dominated, at
1546.968 seconds out of a total 2176.906 seconds (giving a user time of
629.938s). It received 147651 ticks for 0.015 seconds per tick, still quite
accurate. There is obviously no comparable test without the fix.
Signed-off-by: Frank Mayhar <fmayhar@google.com>
Cc: Roland McGrath <roland@redhat.com>
Cc: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Ingo Molnar <mingo@elte.hu>
2008-09-13 00:54:39 +08:00
|
|
|
}
|
|
|
|
|
|
2007-05-24 04:57:44 +08:00
|
|
|
/*
|
|
|
|
|
* Reevaluate whether the task has signals pending delivery.
|
|
|
|
|
* Wake the task if so.
|
|
|
|
|
* This is required every time the blocked sigset_t changes.
|
|
|
|
|
* callers must hold sighand->siglock.
|
|
|
|
|
*/
|
|
|
|
|
extern void recalc_sigpending_and_wake(struct task_struct *t);
|
2005-04-17 06:20:36 +08:00
|
|
|
extern void recalc_sigpending(void);
|
|
|
|
|
|
2013-01-22 03:47:41 +08:00
|
|
|
extern void signal_wake_up_state(struct task_struct *t, unsigned int state);
|
|
|
|
|
|
|
|
|
|
static inline void signal_wake_up(struct task_struct *t, bool resume)
|
|
|
|
|
{
|
|
|
|
|
signal_wake_up_state(t, resume ? TASK_WAKEKILL : 0);
|
|
|
|
|
}
|
|
|
|
|
static inline void ptrace_signal_wake_up(struct task_struct *t, bool resume)
|
|
|
|
|
{
|
|
|
|
|
signal_wake_up_state(t, resume ? __TASK_TRACED : 0);
|
|
|
|
|
}
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
/*
|
|
|
|
|
* Wrappers for p->thread_info->cpu access. No-op on UP.
|
|
|
|
|
*/
|
|
|
|
|
#ifdef CONFIG_SMP
|
|
|
|
|
|
|
|
|
|
static inline unsigned int task_cpu(const struct task_struct *p)
|
|
|
|
|
{
|
2005-11-14 08:06:55 +08:00
|
|
|
return task_thread_info(p)->cpu;
|
2005-04-17 06:20:36 +08:00
|
|
|
}
|
|
|
|
|
|
2007-07-10 00:51:58 +08:00
|
|
|
extern void set_task_cpu(struct task_struct *p, unsigned int cpu);
|
2005-04-17 06:20:36 +08:00
|
|
|
|
|
|
|
|
#else
|
|
|
|
|
|
|
|
|
|
static inline unsigned int task_cpu(const struct task_struct *p)
|
|
|
|
|
{
|
|
|
|
|
return 0;
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void set_task_cpu(struct task_struct *p, unsigned int cpu)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
#endif /* CONFIG_SMP */
|
|
|
|
|
|
2008-11-25 00:05:14 +08:00
|
|
|
extern long sched_setaffinity(pid_t pid, const struct cpumask *new_mask);
|
|
|
|
|
extern long sched_getaffinity(pid_t pid, struct cpumask *mask);
|
2006-06-27 17:54:42 +08:00
|
|
|
|
2010-01-20 20:26:18 +08:00
|
|
|
#ifdef CONFIG_CGROUP_SCHED
|
2011-01-07 15:17:36 +08:00
|
|
|
extern struct task_group root_task_group;
|
2012-06-22 19:36:05 +08:00
|
|
|
#endif /* CONFIG_CGROUP_SCHED */
|
2007-10-15 23:00:09 +08:00
|
|
|
|
2009-02-27 17:43:54 +08:00
|
|
|
extern int task_can_switch_user(struct user_struct *up,
|
|
|
|
|
struct task_struct *tsk);
|
|
|
|
|
|
[PATCH] ifdef ->rchar, ->wchar, ->syscr, ->syscw from task_struct
They are fat: 4x8 bytes in task_struct.
They are uncoditionally updated in every fork, read, write and sendfile.
They are used only if you have some "extended acct fields feature".
And please, please, please, read(2) knows about bytes, not characters,
why it is called "rchar"?
Signed-off-by: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Jay Lan <jlan@engr.sgi.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-02-10 17:46:45 +08:00
|
|
|
#ifdef CONFIG_TASK_XACCT
|
|
|
|
|
static inline void add_rchar(struct task_struct *tsk, ssize_t amt)
|
|
|
|
|
{
|
2008-07-28 06:48:12 +08:00
|
|
|
tsk->ioac.rchar += amt;
|
[PATCH] ifdef ->rchar, ->wchar, ->syscr, ->syscw from task_struct
They are fat: 4x8 bytes in task_struct.
They are uncoditionally updated in every fork, read, write and sendfile.
They are used only if you have some "extended acct fields feature".
And please, please, please, read(2) knows about bytes, not characters,
why it is called "rchar"?
Signed-off-by: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Jay Lan <jlan@engr.sgi.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-02-10 17:46:45 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void add_wchar(struct task_struct *tsk, ssize_t amt)
|
|
|
|
|
{
|
2008-07-28 06:48:12 +08:00
|
|
|
tsk->ioac.wchar += amt;
|
[PATCH] ifdef ->rchar, ->wchar, ->syscr, ->syscw from task_struct
They are fat: 4x8 bytes in task_struct.
They are uncoditionally updated in every fork, read, write and sendfile.
They are used only if you have some "extended acct fields feature".
And please, please, please, read(2) knows about bytes, not characters,
why it is called "rchar"?
Signed-off-by: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Jay Lan <jlan@engr.sgi.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-02-10 17:46:45 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void inc_syscr(struct task_struct *tsk)
|
|
|
|
|
{
|
2008-07-28 06:48:12 +08:00
|
|
|
tsk->ioac.syscr++;
|
[PATCH] ifdef ->rchar, ->wchar, ->syscr, ->syscw from task_struct
They are fat: 4x8 bytes in task_struct.
They are uncoditionally updated in every fork, read, write and sendfile.
They are used only if you have some "extended acct fields feature".
And please, please, please, read(2) knows about bytes, not characters,
why it is called "rchar"?
Signed-off-by: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Jay Lan <jlan@engr.sgi.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-02-10 17:46:45 +08:00
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void inc_syscw(struct task_struct *tsk)
|
|
|
|
|
{
|
2008-07-28 06:48:12 +08:00
|
|
|
tsk->ioac.syscw++;
|
[PATCH] ifdef ->rchar, ->wchar, ->syscr, ->syscw from task_struct
They are fat: 4x8 bytes in task_struct.
They are uncoditionally updated in every fork, read, write and sendfile.
They are used only if you have some "extended acct fields feature".
And please, please, please, read(2) knows about bytes, not characters,
why it is called "rchar"?
Signed-off-by: Alexey Dobriyan <adobriyan@gmail.com>
Cc: Jay Lan <jlan@engr.sgi.com>
Cc: Balbir Singh <balbir@in.ibm.com>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2007-02-10 17:46:45 +08:00
|
|
|
}
|
|
|
|
|
#else
|
|
|
|
|
static inline void add_rchar(struct task_struct *tsk, ssize_t amt)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void add_wchar(struct task_struct *tsk, ssize_t amt)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void inc_syscr(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
|
|
|
|
|
static inline void inc_syscw(struct task_struct *tsk)
|
|
|
|
|
{
|
|
|
|
|
}
|
|
|
|
|
#endif
|
|
|
|
|
|
2008-02-05 14:28:59 +08:00
|
|
|
#ifndef TASK_SIZE_OF
|
|
|
|
|
#define TASK_SIZE_OF(tsk) TASK_SIZE
|
|
|
|
|
#endif
|
|
|
|
|
|
cgroups: add an owner to the mm_struct
Remove the mem_cgroup member from mm_struct and instead adds an owner.
This approach was suggested by Paul Menage. The advantage of this approach
is that, once the mm->owner is known, using the subsystem id, the cgroup
can be determined. It also allows several control groups that are
virtually grouped by mm_struct, to exist independent of the memory
controller i.e., without adding mem_cgroup's for each controller, to
mm_struct.
A new config option CONFIG_MM_OWNER is added and the memory resource
controller selects this config option.
This patch also adds cgroup callbacks to notify subsystems when mm->owner
changes. The mm_cgroup_changed callback is called with the task_lock() of
the new task held and is called just prior to changing the mm->owner.
I am indebted to Paul Menage for the several reviews of this patchset and
helping me make it lighter and simpler.
This patch was tested on a powerpc box, it was compiled with both the
MM_OWNER config turned on and off.
After the thread group leader exits, it's moved to init_css_state by
cgroup_exit(), thus all future charges from runnings threads would be
redirected to the init_css_set's subsystem.
Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com>
Cc: Pavel Emelianov <xemul@openvz.org>
Cc: Hugh Dickins <hugh@veritas.com>
Cc: Sudhir Kumar <skumar@linux.vnet.ibm.com>
Cc: YAMAMOTO Takashi <yamamoto@valinux.co.jp>
Cc: Hirokazu Takahashi <taka@valinux.co.jp>
Cc: David Rientjes <rientjes@google.com>,
Cc: Balbir Singh <balbir@linux.vnet.ibm.com>
Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com>
Acked-by: Pekka Enberg <penberg@cs.helsinki.fi>
Reviewed-by: Paul Menage <menage@google.com>
Cc: Oleg Nesterov <oleg@tv-sign.ru>
Signed-off-by: Andrew Morton <akpm@linux-foundation.org>
Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-04-29 16:00:16 +08:00
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#ifdef CONFIG_MM_OWNER
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extern void mm_update_next_owner(struct mm_struct *mm);
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extern void mm_init_owner(struct mm_struct *mm, struct task_struct *p);
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#else
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static inline void mm_update_next_owner(struct mm_struct *mm)
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{
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}
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static inline void mm_init_owner(struct mm_struct *mm, struct task_struct *p)
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{
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}
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#endif /* CONFIG_MM_OWNER */
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2009-11-20 00:16:37 +08:00
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static inline unsigned long task_rlimit(const struct task_struct *tsk,
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unsigned int limit)
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{
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return ACCESS_ONCE(tsk->signal->rlim[limit].rlim_cur);
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}
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static inline unsigned long task_rlimit_max(const struct task_struct *tsk,
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unsigned int limit)
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{
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return ACCESS_ONCE(tsk->signal->rlim[limit].rlim_max);
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}
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static inline unsigned long rlimit(unsigned int limit)
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{
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return task_rlimit(current, limit);
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}
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static inline unsigned long rlimit_max(unsigned int limit)
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{
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return task_rlimit_max(current, limit);
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}
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2005-04-17 06:20:36 +08:00
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#endif
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