OpenCloudOS-Kernel/include/linux/memcontrol.h

865 lines
22 KiB
C
Raw Normal View History

/* memcontrol.h - Memory Controller
*
* Copyright IBM Corporation, 2007
* Author Balbir Singh <balbir@linux.vnet.ibm.com>
*
* Copyright 2007 OpenVZ SWsoft Inc
* Author: Pavel Emelianov <xemul@openvz.org>
*
* This program is free software; you can redistribute it and/or modify
* it under the terms of the GNU General Public License as published by
* the Free Software Foundation; either version 2 of the License, or
* (at your option) any later version.
*
* This program is distributed in the hope that it will be useful,
* but WITHOUT ANY WARRANTY; without even the implied warranty of
* MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE. See the
* GNU General Public License for more details.
*/
#ifndef _LINUX_MEMCONTROL_H
#define _LINUX_MEMCONTROL_H
#include <linux/cgroup.h>
memcg: add the pagefault count into memcg stats Two new stats in per-memcg memory.stat which tracks the number of page faults and number of major page faults. "pgfault" "pgmajfault" They are different from "pgpgin"/"pgpgout" stat which count number of pages charged/discharged to the cgroup and have no meaning of reading/ writing page to disk. It is valuable to track the two stats for both measuring application's performance as well as the efficiency of the kernel page reclaim path. Counting pagefaults per process is useful, but we also need the aggregated value since processes are monitored and controlled in cgroup basis in memcg. Functional test: check the total number of pgfault/pgmajfault of all memcgs and compare with global vmstat value: $ cat /proc/vmstat | grep fault pgfault 1070751 pgmajfault 553 $ cat /dev/cgroup/memory.stat | grep fault pgfault 1071138 pgmajfault 553 total_pgfault 1071142 total_pgmajfault 553 $ cat /dev/cgroup/A/memory.stat | grep fault pgfault 199 pgmajfault 0 total_pgfault 199 total_pgmajfault 0 Performance test: run page fault test(pft) wit 16 thread on faulting in 15G anon pages in 16G container. There is no regression noticed on the "flt/cpu/s" Sample output from pft: TAG pft:anon-sys-default: Gb Thr CLine User System Wall flt/cpu/s fault/wsec 15 16 1 0.67s 233.41s 14.76s 16798.546 266356.260 +-------------------------------------------------------------------------+ N Min Max Median Avg Stddev x 10 16682.962 17344.027 16913.524 16928.812 166.5362 + 10 16695.568 16923.896 16820.604 16824.652 84.816568 No difference proven at 95.0% confidence [akpm@linux-foundation.org: fix build] [hughd@google.com: shmem fix] Signed-off-by: Ying Han <yinghan@google.com> Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Acked-by: Balbir Singh <balbir@linux.vnet.ibm.com> Signed-off-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-05-27 07:25:38 +08:00
#include <linux/vm_event_item.h>
memcg: kmem controller infrastructure Introduce infrastructure for tracking kernel memory pages to a given memcg. This will happen whenever the caller includes the flag __GFP_KMEMCG flag, and the task belong to a memcg other than the root. In memcontrol.h those functions are wrapped in inline acessors. The idea is to later on, patch those with static branches, so we don't incur any overhead when no mem cgroups with limited kmem are being used. Users of this functionality shall interact with the memcg core code through the following functions: memcg_kmem_newpage_charge: will return true if the group can handle the allocation. At this point, struct page is not yet allocated. memcg_kmem_commit_charge: will either revert the charge, if struct page allocation failed, or embed memcg information into page_cgroup. memcg_kmem_uncharge_page: called at free time, will revert the charge. Signed-off-by: Glauber Costa <glommer@parallels.com> Acked-by: Michal Hocko <mhocko@suse.cz> Acked-by: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Tejun Heo <tj@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Frederic Weisbecker <fweisbec@redhat.com> Cc: Greg Thelen <gthelen@google.com> Cc: JoonSoo Kim <js1304@gmail.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik van Riel <riel@redhat.com> Cc: Suleiman Souhlal <suleiman@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-12-19 06:21:56 +08:00
#include <linux/hardirq.h>
memcg: use static branches when code not in use We can use static branches to patch the code in or out when not used. Because the _ACTIVE bit on kmem_accounted is only set after the increment is done, we guarantee that the root memcg will always be selected for kmem charges until all call sites are patched (see memcg_kmem_enabled). This guarantees that no mischarges are applied. Static branch decrement happens when the last reference count from the kmem accounting in memcg dies. This will only happen when the charges drop down to 0. When that happens, we need to disable the static branch only on those memcgs that enabled it. To achieve this, we would be forced to complicate the code by keeping track of which memcgs were the ones that actually enabled limits, and which ones got it from its parents. It is a lot simpler just to do static_key_slow_inc() on every child that is accounted. Signed-off-by: Glauber Costa <glommer@parallels.com> Acked-by: Michal Hocko <mhocko@suse.cz> Acked-by: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Suleiman Souhlal <suleiman@google.com> Cc: Tejun Heo <tj@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Frederic Weisbecker <fweisbec@redhat.com> Cc: Greg Thelen <gthelen@google.com> Cc: JoonSoo Kim <js1304@gmail.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik van Riel <riel@redhat.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-12-19 06:22:09 +08:00
#include <linux/jump_label.h>
#include <linux/page_counter.h>
#include <linux/vmpressure.h>
#include <linux/eventfd.h>
#include <linux/mmzone.h>
#include <linux/writeback.h>
#include <linux/page-flags.h>
memcg: add the pagefault count into memcg stats Two new stats in per-memcg memory.stat which tracks the number of page faults and number of major page faults. "pgfault" "pgmajfault" They are different from "pgpgin"/"pgpgout" stat which count number of pages charged/discharged to the cgroup and have no meaning of reading/ writing page to disk. It is valuable to track the two stats for both measuring application's performance as well as the efficiency of the kernel page reclaim path. Counting pagefaults per process is useful, but we also need the aggregated value since processes are monitored and controlled in cgroup basis in memcg. Functional test: check the total number of pgfault/pgmajfault of all memcgs and compare with global vmstat value: $ cat /proc/vmstat | grep fault pgfault 1070751 pgmajfault 553 $ cat /dev/cgroup/memory.stat | grep fault pgfault 1071138 pgmajfault 553 total_pgfault 1071142 total_pgmajfault 553 $ cat /dev/cgroup/A/memory.stat | grep fault pgfault 199 pgmajfault 0 total_pgfault 199 total_pgmajfault 0 Performance test: run page fault test(pft) wit 16 thread on faulting in 15G anon pages in 16G container. There is no regression noticed on the "flt/cpu/s" Sample output from pft: TAG pft:anon-sys-default: Gb Thr CLine User System Wall flt/cpu/s fault/wsec 15 16 1 0.67s 233.41s 14.76s 16798.546 266356.260 +-------------------------------------------------------------------------+ N Min Max Median Avg Stddev x 10 16682.962 17344.027 16913.524 16928.812 166.5362 + 10 16695.568 16923.896 16820.604 16824.652 84.816568 No difference proven at 95.0% confidence [akpm@linux-foundation.org: fix build] [hughd@google.com: shmem fix] Signed-off-by: Ying Han <yinghan@google.com> Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Acked-by: Balbir Singh <balbir@linux.vnet.ibm.com> Signed-off-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-05-27 07:25:38 +08:00
struct mem_cgroup;
struct page;
struct mm_struct;
struct kmem_cache;
/*
* The corresponding mem_cgroup_stat_names is defined in mm/memcontrol.c,
* These two lists should keep in accord with each other.
*/
enum mem_cgroup_stat_index {
/*
* For MEM_CONTAINER_TYPE_ALL, usage = pagecache + rss.
*/
MEM_CGROUP_STAT_CACHE, /* # of pages charged as cache */
MEM_CGROUP_STAT_RSS, /* # of pages charged as anon rss */
MEM_CGROUP_STAT_RSS_HUGE, /* # of pages charged as anon huge */
MEM_CGROUP_STAT_FILE_MAPPED, /* # of pages charged as file rss */
memcg: add per cgroup dirty page accounting When modifying PG_Dirty on cached file pages, update the new MEM_CGROUP_STAT_DIRTY counter. This is done in the same places where global NR_FILE_DIRTY is managed. The new memcg stat is visible in the per memcg memory.stat cgroupfs file. The most recent past attempt at this was http://thread.gmane.org/gmane.linux.kernel.cgroups/8632 The new accounting supports future efforts to add per cgroup dirty page throttling and writeback. It also helps an administrator break down a container's memory usage and provides evidence to understand memcg oom kills (the new dirty count is included in memcg oom kill messages). The ability to move page accounting between memcg (memory.move_charge_at_immigrate) makes this accounting more complicated than the global counter. The existing mem_cgroup_{begin,end}_page_stat() lock is used to serialize move accounting with stat updates. Typical update operation: memcg = mem_cgroup_begin_page_stat(page) if (TestSetPageDirty()) { [...] mem_cgroup_update_page_stat(memcg) } mem_cgroup_end_page_stat(memcg) Summary of mem_cgroup_end_page_stat() overhead: - Without CONFIG_MEMCG it's a no-op - With CONFIG_MEMCG and no inter memcg task movement, it's just rcu_read_lock() - With CONFIG_MEMCG and inter memcg task movement, it's rcu_read_lock() + spin_lock_irqsave() A memcg parameter is added to several routines because their callers now grab mem_cgroup_begin_page_stat() which returns the memcg later needed by for mem_cgroup_update_page_stat(). Because mem_cgroup_begin_page_stat() may disable interrupts, some adjustments are needed: - move __mark_inode_dirty() from __set_page_dirty() to its caller. __mark_inode_dirty() locking does not want interrupts disabled. - use spin_lock_irqsave(tree_lock) rather than spin_lock_irq() in __delete_from_page_cache(), replace_page_cache_page(), invalidate_complete_page2(), and __remove_mapping(). text data bss dec hex filename 8925147 1774832 1785856 12485835 be84cb vmlinux-!CONFIG_MEMCG-before 8925339 1774832 1785856 12486027 be858b vmlinux-!CONFIG_MEMCG-after +192 text bytes 8965977 1784992 1785856 12536825 bf4bf9 vmlinux-CONFIG_MEMCG-before 8966750 1784992 1785856 12537598 bf4efe vmlinux-CONFIG_MEMCG-after +773 text bytes Performance tests run on v4.0-rc1-36-g4f671fe2f952. Lower is better for all metrics, they're all wall clock or cycle counts. The read and write fault benchmarks just measure fault time, they do not include I/O time. * CONFIG_MEMCG not set: baseline patched kbuild 1m25.030000(+-0.088% 3 samples) 1m25.426667(+-0.120% 3 samples) dd write 100 MiB 0.859211561 +-15.10% 0.874162885 +-15.03% dd write 200 MiB 1.670653105 +-17.87% 1.669384764 +-11.99% dd write 1000 MiB 8.434691190 +-14.15% 8.474733215 +-14.77% read fault cycles 254.0(+-0.000% 10 samples) 253.0(+-0.000% 10 samples) write fault cycles 2021.2(+-3.070% 10 samples) 1984.5(+-1.036% 10 samples) * CONFIG_MEMCG=y root_memcg: baseline patched kbuild 1m25.716667(+-0.105% 3 samples) 1m25.686667(+-0.153% 3 samples) dd write 100 MiB 0.855650830 +-14.90% 0.887557919 +-14.90% dd write 200 MiB 1.688322953 +-12.72% 1.667682724 +-13.33% dd write 1000 MiB 8.418601605 +-14.30% 8.673532299 +-15.00% read fault cycles 266.0(+-0.000% 10 samples) 266.0(+-0.000% 10 samples) write fault cycles 2051.7(+-1.349% 10 samples) 2049.6(+-1.686% 10 samples) * CONFIG_MEMCG=y non-root_memcg: baseline patched kbuild 1m26.120000(+-0.273% 3 samples) 1m25.763333(+-0.127% 3 samples) dd write 100 MiB 0.861723964 +-15.25% 0.818129350 +-14.82% dd write 200 MiB 1.669887569 +-13.30% 1.698645885 +-13.27% dd write 1000 MiB 8.383191730 +-14.65% 8.351742280 +-14.52% read fault cycles 265.7(+-0.172% 10 samples) 267.0(+-0.000% 10 samples) write fault cycles 2070.6(+-1.512% 10 samples) 2084.4(+-2.148% 10 samples) As expected anon page faults are not affected by this patch. tj: Updated to apply on top of the recent cancel_dirty_page() changes. Signed-off-by: Sha Zhengju <handai.szj@gmail.com> Signed-off-by: Greg Thelen <gthelen@google.com> Signed-off-by: Tejun Heo <tj@kernel.org> Signed-off-by: Jens Axboe <axboe@fb.com>
2015-05-23 05:13:16 +08:00
MEM_CGROUP_STAT_DIRTY, /* # of dirty pages in page cache */
2013-09-13 06:13:53 +08:00
MEM_CGROUP_STAT_WRITEBACK, /* # of pages under writeback */
MEM_CGROUP_STAT_SWAP, /* # of pages, swapped out */
MEM_CGROUP_STAT_NSTATS,
/* default hierarchy stats */
MEMCG_KERNEL_STACK_KB = MEM_CGROUP_STAT_NSTATS,
MEMCG_SLAB_RECLAIMABLE,
MEMCG_SLAB_UNRECLAIMABLE,
MEMCG_SOCK,
MEMCG_NR_STAT,
};
struct mem_cgroup_reclaim_cookie {
pg_data_t *pgdat;
int priority;
unsigned int generation;
};
mm: memcontrol: default hierarchy interface for memory Introduce the basic control files to account, partition, and limit memory using cgroups in default hierarchy mode. This interface versioning allows us to address fundamental design issues in the existing memory cgroup interface, further explained below. The old interface will be maintained indefinitely, but a clearer model and improved workload performance should encourage existing users to switch over to the new one eventually. The control files are thus: - memory.current shows the current consumption of the cgroup and its descendants, in bytes. - memory.low configures the lower end of the cgroup's expected memory consumption range. The kernel considers memory below that boundary to be a reserve - the minimum that the workload needs in order to make forward progress - and generally avoids reclaiming it, unless there is an imminent risk of entering an OOM situation. - memory.high configures the upper end of the cgroup's expected memory consumption range. A cgroup whose consumption grows beyond this threshold is forced into direct reclaim, to work off the excess and to throttle new allocations heavily, but is generally allowed to continue and the OOM killer is not invoked. - memory.max configures the hard maximum amount of memory that the cgroup is allowed to consume before the OOM killer is invoked. - memory.events shows event counters that indicate how often the cgroup was reclaimed while below memory.low, how often it was forced to reclaim excess beyond memory.high, how often it hit memory.max, and how often it entered OOM due to memory.max. This allows users to identify configuration problems when observing a degradation in workload performance. An overcommitted system will have an increased rate of low boundary breaches, whereas increased rates of high limit breaches, maximum hits, or even OOM situations will indicate internally overcommitted cgroups. For existing users of memory cgroups, the following deviations from the current interface are worth pointing out and explaining: - The original lower boundary, the soft limit, is defined as a limit that is per default unset. As a result, the set of cgroups that global reclaim prefers is opt-in, rather than opt-out. The costs for optimizing these mostly negative lookups are so high that the implementation, despite its enormous size, does not even provide the basic desirable behavior. First off, the soft limit has no hierarchical meaning. All configured groups are organized in a global rbtree and treated like equal peers, regardless where they are located in the hierarchy. This makes subtree delegation impossible. Second, the soft limit reclaim pass is so aggressive that it not just introduces high allocation latencies into the system, but also impacts system performance due to overreclaim, to the point where the feature becomes self-defeating. The memory.low boundary on the other hand is a top-down allocated reserve. A cgroup enjoys reclaim protection when it and all its ancestors are below their low boundaries, which makes delegation of subtrees possible. Secondly, new cgroups have no reserve per default and in the common case most cgroups are eligible for the preferred reclaim pass. This allows the new low boundary to be efficiently implemented with just a minor addition to the generic reclaim code, without the need for out-of-band data structures and reclaim passes. Because the generic reclaim code considers all cgroups except for the ones running low in the preferred first reclaim pass, overreclaim of individual groups is eliminated as well, resulting in much better overall workload performance. - The original high boundary, the hard limit, is defined as a strict limit that can not budge, even if the OOM killer has to be called. But this generally goes against the goal of making the most out of the available memory. The memory consumption of workloads varies during runtime, and that requires users to overcommit. But doing that with a strict upper limit requires either a fairly accurate prediction of the working set size or adding slack to the limit. Since working set size estimation is hard and error prone, and getting it wrong results in OOM kills, most users tend to err on the side of a looser limit and end up wasting precious resources. The memory.high boundary on the other hand can be set much more conservatively. When hit, it throttles allocations by forcing them into direct reclaim to work off the excess, but it never invokes the OOM killer. As a result, a high boundary that is chosen too aggressively will not terminate the processes, but instead it will lead to gradual performance degradation. The user can monitor this and make corrections until the minimal memory footprint that still gives acceptable performance is found. In extreme cases, with many concurrent allocations and a complete breakdown of reclaim progress within the group, the high boundary can be exceeded. But even then it's mostly better to satisfy the allocation from the slack available in other groups or the rest of the system than killing the group. Otherwise, memory.max is there to limit this type of spillover and ultimately contain buggy or even malicious applications. - The original control file names are unwieldy and inconsistent in many different ways. For example, the upper boundary hit count is exported in the memory.failcnt file, but an OOM event count has to be manually counted by listening to memory.oom_control events, and lower boundary / soft limit events have to be counted by first setting a threshold for that value and then counting those events. Also, usage and limit files encode their units in the filename. That makes the filenames very long, even though this is not information that a user needs to be reminded of every time they type out those names. To address these naming issues, as well as to signal clearly that the new interface carries a new configuration model, the naming conventions in it necessarily differ from the old interface. - The original limit files indicate the state of an unset limit with a very high number, and a configured limit can be unset by echoing -1 into those files. But that very high number is implementation and architecture dependent and not very descriptive. And while -1 can be understood as an underflow into the highest possible value, -2 or -10M etc. do not work, so it's not inconsistent. memory.low, memory.high, and memory.max will use the string "infinity" to indicate and set the highest possible value. [akpm@linux-foundation.org: use seq_puts() for basic strings] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-12 07:26:06 +08:00
enum mem_cgroup_events_index {
MEM_CGROUP_EVENTS_PGPGIN, /* # of pages paged in */
MEM_CGROUP_EVENTS_PGPGOUT, /* # of pages paged out */
MEM_CGROUP_EVENTS_PGFAULT, /* # of page-faults */
MEM_CGROUP_EVENTS_PGMAJFAULT, /* # of major page-faults */
MEM_CGROUP_EVENTS_NSTATS,
/* default hierarchy events */
MEMCG_LOW = MEM_CGROUP_EVENTS_NSTATS,
MEMCG_HIGH,
MEMCG_MAX,
MEMCG_OOM,
MEMCG_NR_EVENTS,
};
/*
* Per memcg event counter is incremented at every pagein/pageout. With THP,
* it will be incremated by the number of pages. This counter is used for
* for trigger some periodic events. This is straightforward and better
* than using jiffies etc. to handle periodic memcg event.
*/
enum mem_cgroup_events_target {
MEM_CGROUP_TARGET_THRESH,
MEM_CGROUP_TARGET_SOFTLIMIT,
MEM_CGROUP_TARGET_NUMAINFO,
MEM_CGROUP_NTARGETS,
};
#ifdef CONFIG_MEMCG
#define MEM_CGROUP_ID_SHIFT 16
#define MEM_CGROUP_ID_MAX USHRT_MAX
mm: memcontrol: fix cgroup creation failure after many small jobs The memory controller has quite a bit of state that usually outlives the cgroup and pins its CSS until said state disappears. At the same time it imposes a 16-bit limit on the CSS ID space to economically store IDs in the wild. Consequently, when we use cgroups to contain frequent but small and short-lived jobs that leave behind some page cache, we quickly run into the 64k limitations of outstanding CSSs. Creating a new cgroup fails with -ENOSPC while there are only a few, or even no user-visible cgroups in existence. Although pinning CSSs past cgroup removal is common, there are only two instances that actually need an ID after a cgroup is deleted: cache shadow entries and swapout records. Cache shadow entries reference the ID weakly and can deal with the CSS having disappeared when it's looked up later. They pose no hurdle. Swap-out records do need to pin the css to hierarchically attribute swapins after the cgroup has been deleted; though the only pages that remain swapped out after offlining are tmpfs/shmem pages. And those references are under the user's control, so they are manageable. This patch introduces a private 16-bit memcg ID and switches swap and cache shadow entries over to using that. This ID can then be recycled after offlining when the CSS remains pinned only by objects that don't specifically need it. This script demonstrates the problem by faulting one cache page in a new cgroup and deleting it again: set -e mkdir -p pages for x in `seq 128000`; do [ $((x % 1000)) -eq 0 ] && echo $x mkdir /cgroup/foo echo $$ >/cgroup/foo/cgroup.procs echo trex >pages/$x echo $$ >/cgroup/cgroup.procs rmdir /cgroup/foo done When run on an unpatched kernel, we eventually run out of possible IDs even though there are no visible cgroups: [root@ham ~]# ./cssidstress.sh [...] 65000 mkdir: cannot create directory '/cgroup/foo': No space left on device After this patch, the IDs get released upon cgroup destruction and the cache and css objects get released once memory reclaim kicks in. [hannes@cmpxchg.org: init the IDR] Link: http://lkml.kernel.org/r/20160621154601.GA22431@cmpxchg.org Fixes: b2052564e66d ("mm: memcontrol: continue cache reclaim from offlined groups") Link: http://lkml.kernel.org/r/20160617162516.GD19084@cmpxchg.org Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reported-by: John Garcia <john.garcia@mesosphere.io> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: Tejun Heo <tj@kernel.org> Cc: Nikolay Borisov <kernel@kyup.com> Cc: <stable@vger.kernel.org> [3.19+] Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-21 06:44:57 +08:00
struct mem_cgroup_id {
int id;
atomic_t ref;
};
struct mem_cgroup_stat_cpu {
long count[MEMCG_NR_STAT];
unsigned long events[MEMCG_NR_EVENTS];
unsigned long nr_page_events;
unsigned long targets[MEM_CGROUP_NTARGETS];
};
struct mem_cgroup_reclaim_iter {
struct mem_cgroup *position;
/* scan generation, increased every round-trip */
unsigned int generation;
};
/*
* per-zone information in memory controller.
*/
struct mem_cgroup_per_node {
struct lruvec lruvec;
unsigned long lru_size[NR_LRU_LISTS];
struct mem_cgroup_reclaim_iter iter[DEF_PRIORITY + 1];
struct rb_node tree_node; /* RB tree node */
unsigned long usage_in_excess;/* Set to the value by which */
/* the soft limit is exceeded*/
bool on_tree;
struct mem_cgroup *memcg; /* Back pointer, we cannot */
/* use container_of */
};
struct mem_cgroup_threshold {
struct eventfd_ctx *eventfd;
unsigned long threshold;
};
/* For threshold */
struct mem_cgroup_threshold_ary {
/* An array index points to threshold just below or equal to usage. */
int current_threshold;
/* Size of entries[] */
unsigned int size;
/* Array of thresholds */
struct mem_cgroup_threshold entries[0];
};
struct mem_cgroup_thresholds {
/* Primary thresholds array */
struct mem_cgroup_threshold_ary *primary;
/*
* Spare threshold array.
* This is needed to make mem_cgroup_unregister_event() "never fail".
* It must be able to store at least primary->size - 1 entries.
*/
struct mem_cgroup_threshold_ary *spare;
};
enum memcg_kmem_state {
KMEM_NONE,
KMEM_ALLOCATED,
KMEM_ONLINE,
};
/*
* The memory controller data structure. The memory controller controls both
* page cache and RSS per cgroup. We would eventually like to provide
* statistics based on the statistics developed by Rik Van Riel for clock-pro,
* to help the administrator determine what knobs to tune.
*/
struct mem_cgroup {
struct cgroup_subsys_state css;
mm: memcontrol: fix cgroup creation failure after many small jobs The memory controller has quite a bit of state that usually outlives the cgroup and pins its CSS until said state disappears. At the same time it imposes a 16-bit limit on the CSS ID space to economically store IDs in the wild. Consequently, when we use cgroups to contain frequent but small and short-lived jobs that leave behind some page cache, we quickly run into the 64k limitations of outstanding CSSs. Creating a new cgroup fails with -ENOSPC while there are only a few, or even no user-visible cgroups in existence. Although pinning CSSs past cgroup removal is common, there are only two instances that actually need an ID after a cgroup is deleted: cache shadow entries and swapout records. Cache shadow entries reference the ID weakly and can deal with the CSS having disappeared when it's looked up later. They pose no hurdle. Swap-out records do need to pin the css to hierarchically attribute swapins after the cgroup has been deleted; though the only pages that remain swapped out after offlining are tmpfs/shmem pages. And those references are under the user's control, so they are manageable. This patch introduces a private 16-bit memcg ID and switches swap and cache shadow entries over to using that. This ID can then be recycled after offlining when the CSS remains pinned only by objects that don't specifically need it. This script demonstrates the problem by faulting one cache page in a new cgroup and deleting it again: set -e mkdir -p pages for x in `seq 128000`; do [ $((x % 1000)) -eq 0 ] && echo $x mkdir /cgroup/foo echo $$ >/cgroup/foo/cgroup.procs echo trex >pages/$x echo $$ >/cgroup/cgroup.procs rmdir /cgroup/foo done When run on an unpatched kernel, we eventually run out of possible IDs even though there are no visible cgroups: [root@ham ~]# ./cssidstress.sh [...] 65000 mkdir: cannot create directory '/cgroup/foo': No space left on device After this patch, the IDs get released upon cgroup destruction and the cache and css objects get released once memory reclaim kicks in. [hannes@cmpxchg.org: init the IDR] Link: http://lkml.kernel.org/r/20160621154601.GA22431@cmpxchg.org Fixes: b2052564e66d ("mm: memcontrol: continue cache reclaim from offlined groups") Link: http://lkml.kernel.org/r/20160617162516.GD19084@cmpxchg.org Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reported-by: John Garcia <john.garcia@mesosphere.io> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: Tejun Heo <tj@kernel.org> Cc: Nikolay Borisov <kernel@kyup.com> Cc: <stable@vger.kernel.org> [3.19+] Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-21 06:44:57 +08:00
/* Private memcg ID. Used to ID objects that outlive the cgroup */
struct mem_cgroup_id id;
/* Accounted resources */
struct page_counter memory;
mm: memcontrol: charge swap to cgroup2 This patchset introduces swap accounting to cgroup2. This patch (of 7): In the legacy hierarchy we charge memsw, which is dubious, because: - memsw.limit must be >= memory.limit, so it is impossible to limit swap usage less than memory usage. Taking into account the fact that the primary limiting mechanism in the unified hierarchy is memory.high while memory.limit is either left unset or set to a very large value, moving memsw.limit knob to the unified hierarchy would effectively make it impossible to limit swap usage according to the user preference. - memsw.usage != memory.usage + swap.usage, because a page occupying both swap entry and a swap cache page is charged only once to memsw counter. As a result, it is possible to effectively eat up to memory.limit of memory pages *and* memsw.limit of swap entries, which looks unexpected. That said, we should provide a different swap limiting mechanism for cgroup2. This patch adds mem_cgroup->swap counter, which charges the actual number of swap entries used by a cgroup. It is only charged in the unified hierarchy, while the legacy hierarchy memsw logic is left intact. The swap usage can be monitored using new memory.swap.current file and limited using memory.swap.max. Note, to charge swap resource properly in the unified hierarchy, we have to make swap_entry_free uncharge swap only when ->usage reaches zero, not just ->count, i.e. when all references to a swap entry, including the one taken by swap cache, are gone. This is necessary, because otherwise swap-in could result in uncharging swap even if the page is still in swap cache and hence still occupies a swap entry. At the same time, this shouldn't break memsw counter logic, where a page is never charged twice for using both memory and swap, because in case of legacy hierarchy we uncharge swap on commit (see mem_cgroup_commit_charge). Signed-off-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@kernel.org> Cc: Tejun Heo <tj@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-21 07:02:56 +08:00
struct page_counter swap;
/* Legacy consumer-oriented counters */
struct page_counter memsw;
struct page_counter kmem;
struct page_counter tcpmem;
/* Normal memory consumption range */
unsigned long low;
unsigned long high;
/* Range enforcement for interrupt charges */
struct work_struct high_work;
unsigned long soft_limit;
/* vmpressure notifications */
struct vmpressure vmpressure;
/*
* Should the accounting and control be hierarchical, per subtree?
*/
bool use_hierarchy;
/* protected by memcg_oom_lock */
bool oom_lock;
int under_oom;
int swappiness;
/* OOM-Killer disable */
int oom_kill_disable;
/* handle for "memory.events" */
struct cgroup_file events_file;
/* protect arrays of thresholds */
struct mutex thresholds_lock;
/* thresholds for memory usage. RCU-protected */
struct mem_cgroup_thresholds thresholds;
/* thresholds for mem+swap usage. RCU-protected */
struct mem_cgroup_thresholds memsw_thresholds;
/* For oom notifier event fd */
struct list_head oom_notify;
/*
* Should we move charges of a task when a task is moved into this
* mem_cgroup ? And what type of charges should we move ?
*/
unsigned long move_charge_at_immigrate;
/*
* set > 0 if pages under this cgroup are moving to other cgroup.
*/
atomic_t moving_account;
/* taken only while moving_account > 0 */
spinlock_t move_lock;
struct task_struct *move_lock_task;
unsigned long move_lock_flags;
/*
* percpu counter.
*/
struct mem_cgroup_stat_cpu __percpu *stat;
unsigned long socket_pressure;
/* Legacy tcp memory accounting */
bool tcpmem_active;
int tcpmem_pressure;
#ifndef CONFIG_SLOB
/* Index in the kmem_cache->memcg_params.memcg_caches array */
int kmemcg_id;
enum memcg_kmem_state kmem_state;
#endif
int last_scanned_node;
#if MAX_NUMNODES > 1
nodemask_t scan_nodes;
atomic_t numainfo_events;
atomic_t numainfo_updating;
#endif
#ifdef CONFIG_CGROUP_WRITEBACK
struct list_head cgwb_list;
struct wb_domain cgwb_domain;
#endif
/* List of events which userspace want to receive */
struct list_head event_list;
spinlock_t event_list_lock;
struct mem_cgroup_per_node *nodeinfo[0];
/* WARNING: nodeinfo must be the last member here */
};
extern struct mem_cgroup *root_mem_cgroup;
static inline bool mem_cgroup_disabled(void)
{
return !cgroup_subsys_enabled(memory_cgrp_subsys);
}
/**
* mem_cgroup_events - count memory events against a cgroup
* @memcg: the memory cgroup
* @idx: the event index
* @nr: the number of events to account for
*/
static inline void mem_cgroup_events(struct mem_cgroup *memcg,
mm: memcontrol: default hierarchy interface for memory Introduce the basic control files to account, partition, and limit memory using cgroups in default hierarchy mode. This interface versioning allows us to address fundamental design issues in the existing memory cgroup interface, further explained below. The old interface will be maintained indefinitely, but a clearer model and improved workload performance should encourage existing users to switch over to the new one eventually. The control files are thus: - memory.current shows the current consumption of the cgroup and its descendants, in bytes. - memory.low configures the lower end of the cgroup's expected memory consumption range. The kernel considers memory below that boundary to be a reserve - the minimum that the workload needs in order to make forward progress - and generally avoids reclaiming it, unless there is an imminent risk of entering an OOM situation. - memory.high configures the upper end of the cgroup's expected memory consumption range. A cgroup whose consumption grows beyond this threshold is forced into direct reclaim, to work off the excess and to throttle new allocations heavily, but is generally allowed to continue and the OOM killer is not invoked. - memory.max configures the hard maximum amount of memory that the cgroup is allowed to consume before the OOM killer is invoked. - memory.events shows event counters that indicate how often the cgroup was reclaimed while below memory.low, how often it was forced to reclaim excess beyond memory.high, how often it hit memory.max, and how often it entered OOM due to memory.max. This allows users to identify configuration problems when observing a degradation in workload performance. An overcommitted system will have an increased rate of low boundary breaches, whereas increased rates of high limit breaches, maximum hits, or even OOM situations will indicate internally overcommitted cgroups. For existing users of memory cgroups, the following deviations from the current interface are worth pointing out and explaining: - The original lower boundary, the soft limit, is defined as a limit that is per default unset. As a result, the set of cgroups that global reclaim prefers is opt-in, rather than opt-out. The costs for optimizing these mostly negative lookups are so high that the implementation, despite its enormous size, does not even provide the basic desirable behavior. First off, the soft limit has no hierarchical meaning. All configured groups are organized in a global rbtree and treated like equal peers, regardless where they are located in the hierarchy. This makes subtree delegation impossible. Second, the soft limit reclaim pass is so aggressive that it not just introduces high allocation latencies into the system, but also impacts system performance due to overreclaim, to the point where the feature becomes self-defeating. The memory.low boundary on the other hand is a top-down allocated reserve. A cgroup enjoys reclaim protection when it and all its ancestors are below their low boundaries, which makes delegation of subtrees possible. Secondly, new cgroups have no reserve per default and in the common case most cgroups are eligible for the preferred reclaim pass. This allows the new low boundary to be efficiently implemented with just a minor addition to the generic reclaim code, without the need for out-of-band data structures and reclaim passes. Because the generic reclaim code considers all cgroups except for the ones running low in the preferred first reclaim pass, overreclaim of individual groups is eliminated as well, resulting in much better overall workload performance. - The original high boundary, the hard limit, is defined as a strict limit that can not budge, even if the OOM killer has to be called. But this generally goes against the goal of making the most out of the available memory. The memory consumption of workloads varies during runtime, and that requires users to overcommit. But doing that with a strict upper limit requires either a fairly accurate prediction of the working set size or adding slack to the limit. Since working set size estimation is hard and error prone, and getting it wrong results in OOM kills, most users tend to err on the side of a looser limit and end up wasting precious resources. The memory.high boundary on the other hand can be set much more conservatively. When hit, it throttles allocations by forcing them into direct reclaim to work off the excess, but it never invokes the OOM killer. As a result, a high boundary that is chosen too aggressively will not terminate the processes, but instead it will lead to gradual performance degradation. The user can monitor this and make corrections until the minimal memory footprint that still gives acceptable performance is found. In extreme cases, with many concurrent allocations and a complete breakdown of reclaim progress within the group, the high boundary can be exceeded. But even then it's mostly better to satisfy the allocation from the slack available in other groups or the rest of the system than killing the group. Otherwise, memory.max is there to limit this type of spillover and ultimately contain buggy or even malicious applications. - The original control file names are unwieldy and inconsistent in many different ways. For example, the upper boundary hit count is exported in the memory.failcnt file, but an OOM event count has to be manually counted by listening to memory.oom_control events, and lower boundary / soft limit events have to be counted by first setting a threshold for that value and then counting those events. Also, usage and limit files encode their units in the filename. That makes the filenames very long, even though this is not information that a user needs to be reminded of every time they type out those names. To address these naming issues, as well as to signal clearly that the new interface carries a new configuration model, the naming conventions in it necessarily differ from the old interface. - The original limit files indicate the state of an unset limit with a very high number, and a configured limit can be unset by echoing -1 into those files. But that very high number is implementation and architecture dependent and not very descriptive. And while -1 can be understood as an underflow into the highest possible value, -2 or -10M etc. do not work, so it's not inconsistent. memory.low, memory.high, and memory.max will use the string "infinity" to indicate and set the highest possible value. [akpm@linux-foundation.org: use seq_puts() for basic strings] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-12 07:26:06 +08:00
enum mem_cgroup_events_index idx,
unsigned int nr)
{
this_cpu_add(memcg->stat->events[idx], nr);
cgroup_file_notify(&memcg->events_file);
}
mm: memcontrol: default hierarchy interface for memory Introduce the basic control files to account, partition, and limit memory using cgroups in default hierarchy mode. This interface versioning allows us to address fundamental design issues in the existing memory cgroup interface, further explained below. The old interface will be maintained indefinitely, but a clearer model and improved workload performance should encourage existing users to switch over to the new one eventually. The control files are thus: - memory.current shows the current consumption of the cgroup and its descendants, in bytes. - memory.low configures the lower end of the cgroup's expected memory consumption range. The kernel considers memory below that boundary to be a reserve - the minimum that the workload needs in order to make forward progress - and generally avoids reclaiming it, unless there is an imminent risk of entering an OOM situation. - memory.high configures the upper end of the cgroup's expected memory consumption range. A cgroup whose consumption grows beyond this threshold is forced into direct reclaim, to work off the excess and to throttle new allocations heavily, but is generally allowed to continue and the OOM killer is not invoked. - memory.max configures the hard maximum amount of memory that the cgroup is allowed to consume before the OOM killer is invoked. - memory.events shows event counters that indicate how often the cgroup was reclaimed while below memory.low, how often it was forced to reclaim excess beyond memory.high, how often it hit memory.max, and how often it entered OOM due to memory.max. This allows users to identify configuration problems when observing a degradation in workload performance. An overcommitted system will have an increased rate of low boundary breaches, whereas increased rates of high limit breaches, maximum hits, or even OOM situations will indicate internally overcommitted cgroups. For existing users of memory cgroups, the following deviations from the current interface are worth pointing out and explaining: - The original lower boundary, the soft limit, is defined as a limit that is per default unset. As a result, the set of cgroups that global reclaim prefers is opt-in, rather than opt-out. The costs for optimizing these mostly negative lookups are so high that the implementation, despite its enormous size, does not even provide the basic desirable behavior. First off, the soft limit has no hierarchical meaning. All configured groups are organized in a global rbtree and treated like equal peers, regardless where they are located in the hierarchy. This makes subtree delegation impossible. Second, the soft limit reclaim pass is so aggressive that it not just introduces high allocation latencies into the system, but also impacts system performance due to overreclaim, to the point where the feature becomes self-defeating. The memory.low boundary on the other hand is a top-down allocated reserve. A cgroup enjoys reclaim protection when it and all its ancestors are below their low boundaries, which makes delegation of subtrees possible. Secondly, new cgroups have no reserve per default and in the common case most cgroups are eligible for the preferred reclaim pass. This allows the new low boundary to be efficiently implemented with just a minor addition to the generic reclaim code, without the need for out-of-band data structures and reclaim passes. Because the generic reclaim code considers all cgroups except for the ones running low in the preferred first reclaim pass, overreclaim of individual groups is eliminated as well, resulting in much better overall workload performance. - The original high boundary, the hard limit, is defined as a strict limit that can not budge, even if the OOM killer has to be called. But this generally goes against the goal of making the most out of the available memory. The memory consumption of workloads varies during runtime, and that requires users to overcommit. But doing that with a strict upper limit requires either a fairly accurate prediction of the working set size or adding slack to the limit. Since working set size estimation is hard and error prone, and getting it wrong results in OOM kills, most users tend to err on the side of a looser limit and end up wasting precious resources. The memory.high boundary on the other hand can be set much more conservatively. When hit, it throttles allocations by forcing them into direct reclaim to work off the excess, but it never invokes the OOM killer. As a result, a high boundary that is chosen too aggressively will not terminate the processes, but instead it will lead to gradual performance degradation. The user can monitor this and make corrections until the minimal memory footprint that still gives acceptable performance is found. In extreme cases, with many concurrent allocations and a complete breakdown of reclaim progress within the group, the high boundary can be exceeded. But even then it's mostly better to satisfy the allocation from the slack available in other groups or the rest of the system than killing the group. Otherwise, memory.max is there to limit this type of spillover and ultimately contain buggy or even malicious applications. - The original control file names are unwieldy and inconsistent in many different ways. For example, the upper boundary hit count is exported in the memory.failcnt file, but an OOM event count has to be manually counted by listening to memory.oom_control events, and lower boundary / soft limit events have to be counted by first setting a threshold for that value and then counting those events. Also, usage and limit files encode their units in the filename. That makes the filenames very long, even though this is not information that a user needs to be reminded of every time they type out those names. To address these naming issues, as well as to signal clearly that the new interface carries a new configuration model, the naming conventions in it necessarily differ from the old interface. - The original limit files indicate the state of an unset limit with a very high number, and a configured limit can be unset by echoing -1 into those files. But that very high number is implementation and architecture dependent and not very descriptive. And while -1 can be understood as an underflow into the highest possible value, -2 or -10M etc. do not work, so it's not inconsistent. memory.low, memory.high, and memory.max will use the string "infinity" to indicate and set the highest possible value. [akpm@linux-foundation.org: use seq_puts() for basic strings] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-12 07:26:06 +08:00
bool mem_cgroup_low(struct mem_cgroup *root, struct mem_cgroup *memcg);
mm: memcontrol: rewrite charge API These patches rework memcg charge lifetime to integrate more naturally with the lifetime of user pages. This drastically simplifies the code and reduces charging and uncharging overhead. The most expensive part of charging and uncharging is the page_cgroup bit spinlock, which is removed entirely after this series. Here are the top-10 profile entries of a stress test that reads a 128G sparse file on a freshly booted box, without even a dedicated cgroup (i.e. executing in the root memcg). Before: 15.36% cat [kernel.kallsyms] [k] copy_user_generic_string 13.31% cat [kernel.kallsyms] [k] memset 11.48% cat [kernel.kallsyms] [k] do_mpage_readpage 4.23% cat [kernel.kallsyms] [k] get_page_from_freelist 2.38% cat [kernel.kallsyms] [k] put_page 2.32% cat [kernel.kallsyms] [k] __mem_cgroup_commit_charge 2.18% kswapd0 [kernel.kallsyms] [k] __mem_cgroup_uncharge_common 1.92% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.86% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.62% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn After: 15.67% cat [kernel.kallsyms] [k] copy_user_generic_string 13.48% cat [kernel.kallsyms] [k] memset 11.42% cat [kernel.kallsyms] [k] do_mpage_readpage 3.98% cat [kernel.kallsyms] [k] get_page_from_freelist 2.46% cat [kernel.kallsyms] [k] put_page 2.13% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.88% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.67% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn 1.39% kswapd0 [kernel.kallsyms] [k] free_pcppages_bulk 1.30% cat [kernel.kallsyms] [k] kfree As you can see, the memcg footprint has shrunk quite a bit. text data bss dec hex filename 37970 9892 400 48262 bc86 mm/memcontrol.o.old 35239 9892 400 45531 b1db mm/memcontrol.o This patch (of 4): The memcg charge API charges pages before they are rmapped - i.e. have an actual "type" - and so every callsite needs its own set of charge and uncharge functions to know what type is being operated on. Worse, uncharge has to happen from a context that is still type-specific, rather than at the end of the page's lifetime with exclusive access, and so requires a lot of synchronization. Rewrite the charge API to provide a generic set of try_charge(), commit_charge() and cancel_charge() transaction operations, much like what's currently done for swap-in: mem_cgroup_try_charge() attempts to reserve a charge, reclaiming pages from the memcg if necessary. mem_cgroup_commit_charge() commits the page to the charge once it has a valid page->mapping and PageAnon() reliably tells the type. mem_cgroup_cancel_charge() aborts the transaction. This reduces the charge API and enables subsequent patches to drastically simplify uncharging. As pages need to be committed after rmap is established but before they are added to the LRU, page_add_new_anon_rmap() must stop doing LRU additions again. Revive lru_cache_add_active_or_unevictable(). [hughd@google.com: fix shmem_unuse] [hughd@google.com: Add comments on the private use of -EAGAIN] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Hugh Dickins <hughd@google.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:20 +08:00
int mem_cgroup_try_charge(struct page *page, struct mm_struct *mm,
memcg: adjust to support new THP refcounting As with rmap, with new refcounting we cannot rely on PageTransHuge() to check if we need to charge size of huge page form the cgroup. We need to get information from caller to know whether it was mapped with PMD or PTE. We do uncharge when last reference on the page gone. At that point if we see PageTransHuge() it means we need to unchange whole huge page. The tricky part is partial unmap -- when we try to unmap part of huge page. We don't do a special handing of this situation, meaning we don't uncharge the part of huge page unless last user is gone or split_huge_page() is triggered. In case of cgroup memory pressure happens the partial unmapped page will be split through shrinker. This should be good enough. Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Acked-by: Jerome Marchand <jmarchan@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hughd@google.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Steve Capper <steve.capper@linaro.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Christoph Lameter <cl@linux.com> Cc: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-16 08:52:20 +08:00
gfp_t gfp_mask, struct mem_cgroup **memcgp,
bool compound);
mm: memcontrol: rewrite charge API These patches rework memcg charge lifetime to integrate more naturally with the lifetime of user pages. This drastically simplifies the code and reduces charging and uncharging overhead. The most expensive part of charging and uncharging is the page_cgroup bit spinlock, which is removed entirely after this series. Here are the top-10 profile entries of a stress test that reads a 128G sparse file on a freshly booted box, without even a dedicated cgroup (i.e. executing in the root memcg). Before: 15.36% cat [kernel.kallsyms] [k] copy_user_generic_string 13.31% cat [kernel.kallsyms] [k] memset 11.48% cat [kernel.kallsyms] [k] do_mpage_readpage 4.23% cat [kernel.kallsyms] [k] get_page_from_freelist 2.38% cat [kernel.kallsyms] [k] put_page 2.32% cat [kernel.kallsyms] [k] __mem_cgroup_commit_charge 2.18% kswapd0 [kernel.kallsyms] [k] __mem_cgroup_uncharge_common 1.92% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.86% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.62% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn After: 15.67% cat [kernel.kallsyms] [k] copy_user_generic_string 13.48% cat [kernel.kallsyms] [k] memset 11.42% cat [kernel.kallsyms] [k] do_mpage_readpage 3.98% cat [kernel.kallsyms] [k] get_page_from_freelist 2.46% cat [kernel.kallsyms] [k] put_page 2.13% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.88% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.67% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn 1.39% kswapd0 [kernel.kallsyms] [k] free_pcppages_bulk 1.30% cat [kernel.kallsyms] [k] kfree As you can see, the memcg footprint has shrunk quite a bit. text data bss dec hex filename 37970 9892 400 48262 bc86 mm/memcontrol.o.old 35239 9892 400 45531 b1db mm/memcontrol.o This patch (of 4): The memcg charge API charges pages before they are rmapped - i.e. have an actual "type" - and so every callsite needs its own set of charge and uncharge functions to know what type is being operated on. Worse, uncharge has to happen from a context that is still type-specific, rather than at the end of the page's lifetime with exclusive access, and so requires a lot of synchronization. Rewrite the charge API to provide a generic set of try_charge(), commit_charge() and cancel_charge() transaction operations, much like what's currently done for swap-in: mem_cgroup_try_charge() attempts to reserve a charge, reclaiming pages from the memcg if necessary. mem_cgroup_commit_charge() commits the page to the charge once it has a valid page->mapping and PageAnon() reliably tells the type. mem_cgroup_cancel_charge() aborts the transaction. This reduces the charge API and enables subsequent patches to drastically simplify uncharging. As pages need to be committed after rmap is established but before they are added to the LRU, page_add_new_anon_rmap() must stop doing LRU additions again. Revive lru_cache_add_active_or_unevictable(). [hughd@google.com: fix shmem_unuse] [hughd@google.com: Add comments on the private use of -EAGAIN] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Hugh Dickins <hughd@google.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:20 +08:00
void mem_cgroup_commit_charge(struct page *page, struct mem_cgroup *memcg,
memcg: adjust to support new THP refcounting As with rmap, with new refcounting we cannot rely on PageTransHuge() to check if we need to charge size of huge page form the cgroup. We need to get information from caller to know whether it was mapped with PMD or PTE. We do uncharge when last reference on the page gone. At that point if we see PageTransHuge() it means we need to unchange whole huge page. The tricky part is partial unmap -- when we try to unmap part of huge page. We don't do a special handing of this situation, meaning we don't uncharge the part of huge page unless last user is gone or split_huge_page() is triggered. In case of cgroup memory pressure happens the partial unmapped page will be split through shrinker. This should be good enough. Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Acked-by: Jerome Marchand <jmarchan@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hughd@google.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Steve Capper <steve.capper@linaro.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Christoph Lameter <cl@linux.com> Cc: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-16 08:52:20 +08:00
bool lrucare, bool compound);
void mem_cgroup_cancel_charge(struct page *page, struct mem_cgroup *memcg,
bool compound);
mm: memcontrol: rewrite uncharge API The memcg uncharging code that is involved towards the end of a page's lifetime - truncation, reclaim, swapout, migration - is impressively complicated and fragile. Because anonymous and file pages were always charged before they had their page->mapping established, uncharges had to happen when the page type could still be known from the context; as in unmap for anonymous, page cache removal for file and shmem pages, and swap cache truncation for swap pages. However, these operations happen well before the page is actually freed, and so a lot of synchronization is necessary: - Charging, uncharging, page migration, and charge migration all need to take a per-page bit spinlock as they could race with uncharging. - Swap cache truncation happens during both swap-in and swap-out, and possibly repeatedly before the page is actually freed. This means that the memcg swapout code is called from many contexts that make no sense and it has to figure out the direction from page state to make sure memory and memory+swap are always correctly charged. - On page migration, the old page might be unmapped but then reused, so memcg code has to prevent untimely uncharging in that case. Because this code - which should be a simple charge transfer - is so special-cased, it is not reusable for replace_page_cache(). But now that charged pages always have a page->mapping, introduce mem_cgroup_uncharge(), which is called after the final put_page(), when we know for sure that nobody is looking at the page anymore. For page migration, introduce mem_cgroup_migrate(), which is called after the migration is successful and the new page is fully rmapped. Because the old page is no longer uncharged after migration, prevent double charges by decoupling the page's memcg association (PCG_USED and pc->mem_cgroup) from the page holding an actual charge. The new bits PCG_MEM and PCG_MEMSW represent the respective charges and are transferred to the new page during migration. mem_cgroup_migrate() is suitable for replace_page_cache() as well, which gets rid of mem_cgroup_replace_page_cache(). However, care needs to be taken because both the source and the target page can already be charged and on the LRU when fuse is splicing: grab the page lock on the charge moving side to prevent changing pc->mem_cgroup of a page under migration. Also, the lruvecs of both pages change as we uncharge the old and charge the new during migration, and putback may race with us, so grab the lru lock and isolate the pages iff on LRU to prevent races and ensure the pages are on the right lruvec afterward. Swap accounting is massively simplified: because the page is no longer uncharged as early as swap cache deletion, a new mem_cgroup_swapout() can transfer the page's memory+swap charge (PCG_MEMSW) to the swap entry before the final put_page() in page reclaim. Finally, page_cgroup changes are now protected by whatever protection the page itself offers: anonymous pages are charged under the page table lock, whereas page cache insertions, swapin, and migration hold the page lock. Uncharging happens under full exclusion with no outstanding references. Charging and uncharging also ensure that the page is off-LRU, which serializes against charge migration. Remove the very costly page_cgroup lock and set pc->flags non-atomically. [mhocko@suse.cz: mem_cgroup_charge_statistics needs preempt_disable] [vdavydov@parallels.com: fix flags definition] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Cc: Hugh Dickins <hughd@google.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Tested-by: Jet Chen <jet.chen@intel.com> Acked-by: Michal Hocko <mhocko@suse.cz> Tested-by: Felipe Balbi <balbi@ti.com> Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:22 +08:00
void mem_cgroup_uncharge(struct page *page);
void mem_cgroup_uncharge_list(struct list_head *page_list);
memcg: coalesce uncharge during unmap/truncate In massive parallel enviroment, res_counter can be a performance bottleneck. One strong techinque to reduce lock contention is reducing calls by coalescing some amount of calls into one. Considering charge/uncharge chatacteristic, - charge is done one by one via demand-paging. - uncharge is done by - in chunk at munmap, truncate, exit, execve... - one by one via vmscan/paging. It seems we have a chance to coalesce uncharges for improving scalability at unmap/truncation. This patch is a for coalescing uncharge. For avoiding scattering memcg's structure to functions under /mm, this patch adds memcg batch uncharge information to the task. A reason for per-task batching is for making use of caller's context information. We do batched uncharge (deleyed uncharge) when truncation/unmap occurs but do direct uncharge when uncharge is called by memory reclaim (vmscan.c). The degree of coalescing depends on callers - at invalidate/trucate... pagevec size - at unmap ....ZAP_BLOCK_SIZE (memory itself will be freed in this degree.) Then, we'll not coalescing too much. On x86-64 8cpu server, I tested overheads of memcg at page fault by running a program which does map/fault/unmap in a loop. Running a task per a cpu by taskset and see sum of the number of page faults in 60secs. [without memcg config] 40156968 page-faults # 0.085 M/sec ( +- 0.046% ) 27.67 cache-miss/faults [root cgroup] 36659599 page-faults # 0.077 M/sec ( +- 0.247% ) 31.58 miss/faults [in a child cgroup] 18444157 page-faults # 0.039 M/sec ( +- 0.133% ) 69.96 miss/faults [child with this patch] 27133719 page-faults # 0.057 M/sec ( +- 0.155% ) 47.16 miss/faults We can see some amounts of improvement. (root cgroup doesn't affected by this patch) Another patch for "charge" will follow this and above will be improved more. Changelog(since 2009/10/02): - renamed filed of memcg_batch (as pages to bytes, memsw to memsw_bytes) - some clean up and commentary/description updates. - added initialize code to copy_process(). (possible bug fix) Changelog(old): - fixed !CONFIG_MEM_CGROUP case. - rebased onto the latest mmotm + softlimit fix patches. - unified patch for callers - added commetns. - make ->do_batch as bool. - removed css_get() at el. We don't need it. Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-12-16 08:47:03 +08:00
void mem_cgroup_migrate(struct page *oldpage, struct page *newpage);
memcg: coalesce uncharge during unmap/truncate In massive parallel enviroment, res_counter can be a performance bottleneck. One strong techinque to reduce lock contention is reducing calls by coalescing some amount of calls into one. Considering charge/uncharge chatacteristic, - charge is done one by one via demand-paging. - uncharge is done by - in chunk at munmap, truncate, exit, execve... - one by one via vmscan/paging. It seems we have a chance to coalesce uncharges for improving scalability at unmap/truncation. This patch is a for coalescing uncharge. For avoiding scattering memcg's structure to functions under /mm, this patch adds memcg batch uncharge information to the task. A reason for per-task batching is for making use of caller's context information. We do batched uncharge (deleyed uncharge) when truncation/unmap occurs but do direct uncharge when uncharge is called by memory reclaim (vmscan.c). The degree of coalescing depends on callers - at invalidate/trucate... pagevec size - at unmap ....ZAP_BLOCK_SIZE (memory itself will be freed in this degree.) Then, we'll not coalescing too much. On x86-64 8cpu server, I tested overheads of memcg at page fault by running a program which does map/fault/unmap in a loop. Running a task per a cpu by taskset and see sum of the number of page faults in 60secs. [without memcg config] 40156968 page-faults # 0.085 M/sec ( +- 0.046% ) 27.67 cache-miss/faults [root cgroup] 36659599 page-faults # 0.077 M/sec ( +- 0.247% ) 31.58 miss/faults [in a child cgroup] 18444157 page-faults # 0.039 M/sec ( +- 0.133% ) 69.96 miss/faults [child with this patch] 27133719 page-faults # 0.057 M/sec ( +- 0.155% ) 47.16 miss/faults We can see some amounts of improvement. (root cgroup doesn't affected by this patch) Another patch for "charge" will follow this and above will be improved more. Changelog(since 2009/10/02): - renamed filed of memcg_batch (as pages to bytes, memsw to memsw_bytes) - some clean up and commentary/description updates. - added initialize code to copy_process(). (possible bug fix) Changelog(old): - fixed !CONFIG_MEM_CGROUP case. - rebased onto the latest mmotm + softlimit fix patches. - unified patch for callers - added commetns. - make ->do_batch as bool. - removed css_get() at el. We don't need it. Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-12-16 08:47:03 +08:00
static struct mem_cgroup_per_node *
mem_cgroup_nodeinfo(struct mem_cgroup *memcg, int nid)
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
{
return memcg->nodeinfo[nid];
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
}
/**
* mem_cgroup_lruvec - get the lru list vector for a node or a memcg zone
* @node: node of the wanted lruvec
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
* @memcg: memcg of the wanted lruvec
*
* Returns the lru list vector holding pages for a given @node or a given
* @memcg and @zone. This can be the node lruvec, if the memory controller
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
* is disabled.
*/
static inline struct lruvec *mem_cgroup_lruvec(struct pglist_data *pgdat,
struct mem_cgroup *memcg)
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
{
struct mem_cgroup_per_node *mz;
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
struct lruvec *lruvec;
if (mem_cgroup_disabled()) {
lruvec = node_lruvec(pgdat);
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
goto out;
}
mz = mem_cgroup_nodeinfo(memcg, pgdat->node_id);
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
lruvec = &mz->lruvec;
out:
/*
* Since a node can be onlined after the mem_cgroup was created,
mm, vmscan: move LRU lists to node This moves the LRU lists from the zone to the node and related data such as counters, tracing, congestion tracking and writeback tracking. Unfortunately, due to reclaim and compaction retry logic, it is necessary to account for the number of LRU pages on both zone and node logic. Most reclaim logic is based on the node counters but the retry logic uses the zone counters which do not distinguish inactive and active sizes. It would be possible to leave the LRU counters on a per-zone basis but it's a heavier calculation across multiple cache lines that is much more frequent than the retry checks. Other than the LRU counters, this is mostly a mechanical patch but note that it introduces a number of anomalies. For example, the scans are per-zone but using per-node counters. We also mark a node as congested when a zone is congested. This causes weird problems that are fixed later but is easier to review. In the event that there is excessive overhead on 32-bit systems due to the nodes being on LRU then there are two potential solutions 1. Long-term isolation of highmem pages when reclaim is lowmem When pages are skipped, they are immediately added back onto the LRU list. If lowmem reclaim persisted for long periods of time, the same highmem pages get continually scanned. The idea would be that lowmem keeps those pages on a separate list until a reclaim for highmem pages arrives that splices the highmem pages back onto the LRU. It potentially could be implemented similar to the UNEVICTABLE list. That would reduce the skip rate with the potential corner case is that highmem pages have to be scanned and reclaimed to free lowmem slab pages. 2. Linear scan lowmem pages if the initial LRU shrink fails This will break LRU ordering but may be preferable and faster during memory pressure than skipping LRU pages. Link: http://lkml.kernel.org/r/1467970510-21195-4-git-send-email-mgorman@techsingularity.net Signed-off-by: Mel Gorman <mgorman@techsingularity.net> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vlastimil Babka <vbabka@suse.cz> Cc: Hillf Danton <hillf.zj@alibaba-inc.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@surriel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:31 +08:00
* we have to be prepared to initialize lruvec->pgdat here;
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
* and if offlined then reonlined, we need to reinitialize it.
*/
if (unlikely(lruvec->pgdat != pgdat))
lruvec->pgdat = pgdat;
mm: fix vm-scalability regression in cgroup-aware workingset code Commit 23047a96d7cf ("mm: workingset: per-cgroup cache thrash detection") added a page->mem_cgroup lookup to the cache eviction, refault, and activation paths, as well as locking to the activation path, and the vm-scalability tests showed a regression of -23%. While the test in question is an artificial worst-case scenario that doesn't occur in real workloads - reading two sparse files in parallel at full CPU speed just to hammer the LRU paths - there is still some optimizations that can be done in those paths. Inline the lookup functions to eliminate calls. Also, page->mem_cgroup doesn't need to be stabilized when counting an activation; we merely need to hold the RCU lock to prevent the memcg from being freed. This cuts down on overhead quite a bit: 23047a96d7cfcfca 063f6715e77a7be5770d6081fe ---------------- -------------------------- %stddev %change %stddev \ | \ 21621405 +- 0% +11.3% 24069657 +- 2% vm-scalability.throughput [linux@roeck-us.net: drop unnecessary include file] [hannes@cmpxchg.org: add WARN_ON_ONCE()s] Link: http://lkml.kernel.org/r/20160707194024.GA26580@cmpxchg.org Link: http://lkml.kernel.org/r/20160624175101.GA3024@cmpxchg.org Reported-by: Ye Xiaolong <xiaolong.ye@intel.com> Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.com> Cc: Vladimir Davydov <vdavydov@virtuozzo.com> Signed-off-by: Guenter Roeck <linux@roeck-us.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:10 +08:00
return lruvec;
}
mm, vmscan: move LRU lists to node This moves the LRU lists from the zone to the node and related data such as counters, tracing, congestion tracking and writeback tracking. Unfortunately, due to reclaim and compaction retry logic, it is necessary to account for the number of LRU pages on both zone and node logic. Most reclaim logic is based on the node counters but the retry logic uses the zone counters which do not distinguish inactive and active sizes. It would be possible to leave the LRU counters on a per-zone basis but it's a heavier calculation across multiple cache lines that is much more frequent than the retry checks. Other than the LRU counters, this is mostly a mechanical patch but note that it introduces a number of anomalies. For example, the scans are per-zone but using per-node counters. We also mark a node as congested when a zone is congested. This causes weird problems that are fixed later but is easier to review. In the event that there is excessive overhead on 32-bit systems due to the nodes being on LRU then there are two potential solutions 1. Long-term isolation of highmem pages when reclaim is lowmem When pages are skipped, they are immediately added back onto the LRU list. If lowmem reclaim persisted for long periods of time, the same highmem pages get continually scanned. The idea would be that lowmem keeps those pages on a separate list until a reclaim for highmem pages arrives that splices the highmem pages back onto the LRU. It potentially could be implemented similar to the UNEVICTABLE list. That would reduce the skip rate with the potential corner case is that highmem pages have to be scanned and reclaimed to free lowmem slab pages. 2. Linear scan lowmem pages if the initial LRU shrink fails This will break LRU ordering but may be preferable and faster during memory pressure than skipping LRU pages. Link: http://lkml.kernel.org/r/1467970510-21195-4-git-send-email-mgorman@techsingularity.net Signed-off-by: Mel Gorman <mgorman@techsingularity.net> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vlastimil Babka <vbabka@suse.cz> Cc: Hillf Danton <hillf.zj@alibaba-inc.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@surriel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:31 +08:00
struct lruvec *mem_cgroup_page_lruvec(struct page *, struct pglist_data *);
bool task_in_mem_cgroup(struct task_struct *task, struct mem_cgroup *memcg);
struct mem_cgroup *mem_cgroup_from_task(struct task_struct *p);
static inline
struct mem_cgroup *mem_cgroup_from_css(struct cgroup_subsys_state *css){
return css ? container_of(css, struct mem_cgroup, css) : NULL;
}
#define mem_cgroup_from_counter(counter, member) \
container_of(counter, struct mem_cgroup, member)
struct mem_cgroup *mem_cgroup_iter(struct mem_cgroup *,
struct mem_cgroup *,
struct mem_cgroup_reclaim_cookie *);
void mem_cgroup_iter_break(struct mem_cgroup *, struct mem_cgroup *);
static inline unsigned short mem_cgroup_id(struct mem_cgroup *memcg)
{
if (mem_cgroup_disabled())
return 0;
mm: memcontrol: fix cgroup creation failure after many small jobs The memory controller has quite a bit of state that usually outlives the cgroup and pins its CSS until said state disappears. At the same time it imposes a 16-bit limit on the CSS ID space to economically store IDs in the wild. Consequently, when we use cgroups to contain frequent but small and short-lived jobs that leave behind some page cache, we quickly run into the 64k limitations of outstanding CSSs. Creating a new cgroup fails with -ENOSPC while there are only a few, or even no user-visible cgroups in existence. Although pinning CSSs past cgroup removal is common, there are only two instances that actually need an ID after a cgroup is deleted: cache shadow entries and swapout records. Cache shadow entries reference the ID weakly and can deal with the CSS having disappeared when it's looked up later. They pose no hurdle. Swap-out records do need to pin the css to hierarchically attribute swapins after the cgroup has been deleted; though the only pages that remain swapped out after offlining are tmpfs/shmem pages. And those references are under the user's control, so they are manageable. This patch introduces a private 16-bit memcg ID and switches swap and cache shadow entries over to using that. This ID can then be recycled after offlining when the CSS remains pinned only by objects that don't specifically need it. This script demonstrates the problem by faulting one cache page in a new cgroup and deleting it again: set -e mkdir -p pages for x in `seq 128000`; do [ $((x % 1000)) -eq 0 ] && echo $x mkdir /cgroup/foo echo $$ >/cgroup/foo/cgroup.procs echo trex >pages/$x echo $$ >/cgroup/cgroup.procs rmdir /cgroup/foo done When run on an unpatched kernel, we eventually run out of possible IDs even though there are no visible cgroups: [root@ham ~]# ./cssidstress.sh [...] 65000 mkdir: cannot create directory '/cgroup/foo': No space left on device After this patch, the IDs get released upon cgroup destruction and the cache and css objects get released once memory reclaim kicks in. [hannes@cmpxchg.org: init the IDR] Link: http://lkml.kernel.org/r/20160621154601.GA22431@cmpxchg.org Fixes: b2052564e66d ("mm: memcontrol: continue cache reclaim from offlined groups") Link: http://lkml.kernel.org/r/20160617162516.GD19084@cmpxchg.org Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reported-by: John Garcia <john.garcia@mesosphere.io> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: Tejun Heo <tj@kernel.org> Cc: Nikolay Borisov <kernel@kyup.com> Cc: <stable@vger.kernel.org> [3.19+] Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-21 06:44:57 +08:00
return memcg->id.id;
}
mm: memcontrol: fix cgroup creation failure after many small jobs The memory controller has quite a bit of state that usually outlives the cgroup and pins its CSS until said state disappears. At the same time it imposes a 16-bit limit on the CSS ID space to economically store IDs in the wild. Consequently, when we use cgroups to contain frequent but small and short-lived jobs that leave behind some page cache, we quickly run into the 64k limitations of outstanding CSSs. Creating a new cgroup fails with -ENOSPC while there are only a few, or even no user-visible cgroups in existence. Although pinning CSSs past cgroup removal is common, there are only two instances that actually need an ID after a cgroup is deleted: cache shadow entries and swapout records. Cache shadow entries reference the ID weakly and can deal with the CSS having disappeared when it's looked up later. They pose no hurdle. Swap-out records do need to pin the css to hierarchically attribute swapins after the cgroup has been deleted; though the only pages that remain swapped out after offlining are tmpfs/shmem pages. And those references are under the user's control, so they are manageable. This patch introduces a private 16-bit memcg ID and switches swap and cache shadow entries over to using that. This ID can then be recycled after offlining when the CSS remains pinned only by objects that don't specifically need it. This script demonstrates the problem by faulting one cache page in a new cgroup and deleting it again: set -e mkdir -p pages for x in `seq 128000`; do [ $((x % 1000)) -eq 0 ] && echo $x mkdir /cgroup/foo echo $$ >/cgroup/foo/cgroup.procs echo trex >pages/$x echo $$ >/cgroup/cgroup.procs rmdir /cgroup/foo done When run on an unpatched kernel, we eventually run out of possible IDs even though there are no visible cgroups: [root@ham ~]# ./cssidstress.sh [...] 65000 mkdir: cannot create directory '/cgroup/foo': No space left on device After this patch, the IDs get released upon cgroup destruction and the cache and css objects get released once memory reclaim kicks in. [hannes@cmpxchg.org: init the IDR] Link: http://lkml.kernel.org/r/20160621154601.GA22431@cmpxchg.org Fixes: b2052564e66d ("mm: memcontrol: continue cache reclaim from offlined groups") Link: http://lkml.kernel.org/r/20160617162516.GD19084@cmpxchg.org Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reported-by: John Garcia <john.garcia@mesosphere.io> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: Tejun Heo <tj@kernel.org> Cc: Nikolay Borisov <kernel@kyup.com> Cc: <stable@vger.kernel.org> [3.19+] Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-21 06:44:57 +08:00
struct mem_cgroup *mem_cgroup_from_id(unsigned short id);
/**
* parent_mem_cgroup - find the accounting parent of a memcg
* @memcg: memcg whose parent to find
*
* Returns the parent memcg, or NULL if this is the root or the memory
* controller is in legacy no-hierarchy mode.
*/
static inline struct mem_cgroup *parent_mem_cgroup(struct mem_cgroup *memcg)
{
if (!memcg->memory.parent)
return NULL;
return mem_cgroup_from_counter(memcg->memory.parent, memory);
}
static inline bool mem_cgroup_is_descendant(struct mem_cgroup *memcg,
struct mem_cgroup *root)
{
if (root == memcg)
return true;
if (!root->use_hierarchy)
return false;
return cgroup_is_descendant(memcg->css.cgroup, root->css.cgroup);
}
static inline bool mm_match_cgroup(struct mm_struct *mm,
struct mem_cgroup *memcg)
{
struct mem_cgroup *task_memcg;
bool match = false;
mm: memcg: count pte references from every member of the reclaimed hierarchy The rmap walker checking page table references has historically ignored references from VMAs that were not part of the memcg that was being reclaimed during memcg hard limit reclaim. When transitioning global reclaim to memcg hierarchy reclaim, I missed that bit and now references from outside a memcg are ignored even during global reclaim. Reverting back to traditional behaviour - count all references during global reclaim and only mind references of the memcg being reclaimed during limit reclaim would be one option. However, the more generic idea is to ignore references exactly then when they are outside the hierarchy that is currently under reclaim; because only then will their reclamation be of any use to help the pressure situation. It makes no sense to ignore references from a sibling memcg and then evict a page that will be immediately refaulted by that sibling which contributes to the same usage of the common ancestor under reclaim. The solution: make the rmap walker ignore references from VMAs that are not part of the hierarchy that is being reclaimed. Flat limit reclaim will stay the same, hierarchical limit reclaim will mind the references only to pages that the hierarchy owns. Global reclaim, since it reclaims from all memcgs, will be fixed to regard all references. [akpm@linux-foundation.org: name the args in the declaration] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reported-by: Konstantin Khlebnikov <khlebnikov@openvz.org> Acked-by: Konstantin Khlebnikov<khlebnikov@openvz.org> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Tejun Heo <tj@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-05-30 06:06:25 +08:00
rcu_read_lock();
task_memcg = mem_cgroup_from_task(rcu_dereference(mm->owner));
if (task_memcg)
match = mem_cgroup_is_descendant(task_memcg, memcg);
rcu_read_unlock();
mm: memcg: count pte references from every member of the reclaimed hierarchy The rmap walker checking page table references has historically ignored references from VMAs that were not part of the memcg that was being reclaimed during memcg hard limit reclaim. When transitioning global reclaim to memcg hierarchy reclaim, I missed that bit and now references from outside a memcg are ignored even during global reclaim. Reverting back to traditional behaviour - count all references during global reclaim and only mind references of the memcg being reclaimed during limit reclaim would be one option. However, the more generic idea is to ignore references exactly then when they are outside the hierarchy that is currently under reclaim; because only then will their reclamation be of any use to help the pressure situation. It makes no sense to ignore references from a sibling memcg and then evict a page that will be immediately refaulted by that sibling which contributes to the same usage of the common ancestor under reclaim. The solution: make the rmap walker ignore references from VMAs that are not part of the hierarchy that is being reclaimed. Flat limit reclaim will stay the same, hierarchical limit reclaim will mind the references only to pages that the hierarchy owns. Global reclaim, since it reclaims from all memcgs, will be fixed to regard all references. [akpm@linux-foundation.org: name the args in the declaration] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reported-by: Konstantin Khlebnikov <khlebnikov@openvz.org> Acked-by: Konstantin Khlebnikov<khlebnikov@openvz.org> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Tejun Heo <tj@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-05-30 06:06:25 +08:00
return match;
}
Memory controller: memory accounting Add the accounting hooks. The accounting is carried out for RSS and Page Cache (unmapped) pages. There is now a common limit and accounting for both. The RSS accounting is accounted at page_add_*_rmap() and page_remove_rmap() time. Page cache is accounted at add_to_page_cache(), __delete_from_page_cache(). Swap cache is also accounted for. Each page's page_cgroup is protected with the last bit of the page_cgroup pointer, this makes handling of race conditions involving simultaneous mappings of a page easier. A reference count is kept in the page_cgroup to deal with cases where a page might be unmapped from the RSS of all tasks, but still lives in the page cache. Credits go to Vaidyanathan Srinivasan for helping with reference counting work of the page cgroup. Almost all of the page cache accounting code has help from Vaidyanathan Srinivasan. [hugh@veritas.com: fix swapoff breakage] [akpm@linux-foundation.org: fix locking] Signed-off-by: Vaidyanathan Srinivasan <svaidy@linux.vnet.ibm.com> Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com> Cc: Pavel Emelianov <xemul@openvz.org> Cc: Paul Menage <menage@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Kirill Korotaev <dev@sw.ru> Cc: Herbert Poetzl <herbert@13thfloor.at> Cc: David Rientjes <rientjes@google.com> Cc: <Valdis.Kletnieks@vt.edu> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-02-07 16:13:53 +08:00
struct cgroup_subsys_state *mem_cgroup_css_from_page(struct page *page);
memcg: add page_cgroup_ino helper This patchset introduces a new user API for tracking user memory pages that have not been used for a given period of time. The purpose of this is to provide the userspace with the means of tracking a workload's working set, i.e. the set of pages that are actively used by the workload. Knowing the working set size can be useful for partitioning the system more efficiently, e.g. by tuning memory cgroup limits appropriately, or for job placement within a compute cluster. ==== USE CASES ==== The unified cgroup hierarchy has memory.low and memory.high knobs, which are defined as the low and high boundaries for the workload working set size. However, the working set size of a workload may be unknown or change in time. With this patch set, one can periodically estimate the amount of memory unused by each cgroup and tune their memory.low and memory.high parameters accordingly, therefore optimizing the overall memory utilization. Another use case is balancing workloads within a compute cluster. Knowing how much memory is not really used by a workload unit may help take a more optimal decision when considering migrating the unit to another node within the cluster. Also, as noted by Minchan, this would be useful for per-process reclaim (https://lwn.net/Articles/545668/). With idle tracking, we could reclaim idle pages only by smart user memory manager. ==== USER API ==== The user API consists of two new files: * /sys/kernel/mm/page_idle/bitmap. This file implements a bitmap where each bit corresponds to a page, indexed by PFN. When the bit is set, the corresponding page is idle. A page is considered idle if it has not been accessed since it was marked idle. To mark a page idle one should set the bit corresponding to the page by writing to the file. A value written to the file is OR-ed with the current bitmap value. Only user memory pages can be marked idle, for other page types input is silently ignored. Writing to this file beyond max PFN results in the ENXIO error. Only available when CONFIG_IDLE_PAGE_TRACKING is set. This file can be used to estimate the amount of pages that are not used by a particular workload as follows: 1. mark all pages of interest idle by setting corresponding bits in the /sys/kernel/mm/page_idle/bitmap 2. wait until the workload accesses its working set 3. read /sys/kernel/mm/page_idle/bitmap and count the number of bits set * /proc/kpagecgroup. This file contains a 64-bit inode number of the memory cgroup each page is charged to, indexed by PFN. Only available when CONFIG_MEMCG is set. This file can be used to find all pages (including unmapped file pages) accounted to a particular cgroup. Using /sys/kernel/mm/page_idle/bitmap, one can then estimate the cgroup working set size. For an example of using these files for estimating the amount of unused memory pages per each memory cgroup, please see the script attached below. ==== REASONING ==== The reason to introduce the new user API instead of using /proc/PID/{clear_refs,smaps} is that the latter has two serious drawbacks: - it does not count unmapped file pages - it affects the reclaimer logic The new API attempts to overcome them both. For more details on how it is achieved, please see the comment to patch 6. ==== PATCHSET STRUCTURE ==== The patch set is organized as follows: - patch 1 adds page_cgroup_ino() helper for the sake of /proc/kpagecgroup and patches 2-3 do related cleanup - patch 4 adds /proc/kpagecgroup, which reports cgroup ino each page is charged to - patch 5 introduces a new mmu notifier callback, clear_young, which is a lightweight version of clear_flush_young; it is used in patch 6 - patch 6 implements the idle page tracking feature, including the userspace API, /sys/kernel/mm/page_idle/bitmap - patch 7 exports idle flag via /proc/kpageflags ==== SIMILAR WORKS ==== Originally, the patch for tracking idle memory was proposed back in 2011 by Michel Lespinasse (see http://lwn.net/Articles/459269/). The main difference between Michel's patch and this one is that Michel implemented a kernel space daemon for estimating idle memory size per cgroup while this patch only provides the userspace with the minimal API for doing the job, leaving the rest up to the userspace. However, they both share the same idea of Idle/Young page flags to avoid affecting the reclaimer logic. ==== PERFORMANCE EVALUATION ==== SPECjvm2008 (https://www.spec.org/jvm2008/) was used to evaluate the performance impact introduced by this patch set. Three runs were carried out: - base: kernel without the patch - patched: patched kernel, the feature is not used - patched-active: patched kernel, 1 minute-period daemon is used for tracking idle memory For tracking idle memory, idlememstat utility was used: https://github.com/locker/idlememstat testcase base patched patched-active compiler 537.40 ( 0.00)% 532.26 (-0.96)% 538.31 ( 0.17)% compress 305.47 ( 0.00)% 301.08 (-1.44)% 300.71 (-1.56)% crypto 284.32 ( 0.00)% 282.21 (-0.74)% 284.87 ( 0.19)% derby 411.05 ( 0.00)% 413.44 ( 0.58)% 412.07 ( 0.25)% mpegaudio 189.96 ( 0.00)% 190.87 ( 0.48)% 189.42 (-0.28)% scimark.large 46.85 ( 0.00)% 46.41 (-0.94)% 47.83 ( 2.09)% scimark.small 412.91 ( 0.00)% 415.41 ( 0.61)% 421.17 ( 2.00)% serial 204.23 ( 0.00)% 213.46 ( 4.52)% 203.17 (-0.52)% startup 36.76 ( 0.00)% 35.49 (-3.45)% 35.64 (-3.05)% sunflow 115.34 ( 0.00)% 115.08 (-0.23)% 117.37 ( 1.76)% xml 620.55 ( 0.00)% 619.95 (-0.10)% 620.39 (-0.03)% composite 211.50 ( 0.00)% 211.15 (-0.17)% 211.67 ( 0.08)% time idlememstat: 17.20user 65.16system 2:15:23elapsed 1%CPU (0avgtext+0avgdata 8476maxresident)k 448inputs+40outputs (1major+36052minor)pagefaults 0swaps ==== SCRIPT FOR COUNTING IDLE PAGES PER CGROUP ==== #! /usr/bin/python # import os import stat import errno import struct CGROUP_MOUNT = "/sys/fs/cgroup/memory" BUFSIZE = 8 * 1024 # must be multiple of 8 def get_hugepage_size(): with open("/proc/meminfo", "r") as f: for s in f: k, v = s.split(":") if k == "Hugepagesize": return int(v.split()[0]) * 1024 PAGE_SIZE = os.sysconf("SC_PAGE_SIZE") HUGEPAGE_SIZE = get_hugepage_size() def set_idle(): f = open("/sys/kernel/mm/page_idle/bitmap", "wb", BUFSIZE) while True: try: f.write(struct.pack("Q", pow(2, 64) - 1)) except IOError as err: if err.errno == errno.ENXIO: break raise f.close() def count_idle(): f_flags = open("/proc/kpageflags", "rb", BUFSIZE) f_cgroup = open("/proc/kpagecgroup", "rb", BUFSIZE) with open("/sys/kernel/mm/page_idle/bitmap", "rb", BUFSIZE) as f: while f.read(BUFSIZE): pass # update idle flag idlememsz = {} while True: s1, s2 = f_flags.read(8), f_cgroup.read(8) if not s1 or not s2: break flags, = struct.unpack('Q', s1) cgino, = struct.unpack('Q', s2) unevictable = (flags >> 18) & 1 huge = (flags >> 22) & 1 idle = (flags >> 25) & 1 if idle and not unevictable: idlememsz[cgino] = idlememsz.get(cgino, 0) + \ (HUGEPAGE_SIZE if huge else PAGE_SIZE) f_flags.close() f_cgroup.close() return idlememsz if __name__ == "__main__": print "Setting the idle flag for each page..." set_idle() raw_input("Wait until the workload accesses its working set, " "then press Enter") print "Counting idle pages..." idlememsz = count_idle() for dir, subdirs, files in os.walk(CGROUP_MOUNT): ino = os.stat(dir)[stat.ST_INO] print dir + ": " + str(idlememsz.get(ino, 0) / 1024) + " kB" ==== END SCRIPT ==== This patch (of 8): Add page_cgroup_ino() helper to memcg. This function returns the inode number of the closest online ancestor of the memory cgroup a page is charged to. It is required for exporting information about which page is charged to which cgroup to userspace, which will be introduced by a following patch. Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Reviewed-by: Andres Lagar-Cavilla <andreslc@google.com> Cc: Minchan Kim <minchan@kernel.org> Cc: Raghavendra K T <raghavendra.kt@linux.vnet.ibm.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Greg Thelen <gthelen@google.com> Cc: Michel Lespinasse <walken@google.com> Cc: David Rientjes <rientjes@google.com> Cc: Pavel Emelyanov <xemul@parallels.com> Cc: Cyrill Gorcunov <gorcunov@openvz.org> Cc: Jonathan Corbet <corbet@lwn.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-09-10 06:35:28 +08:00
ino_t page_cgroup_ino(struct page *page);
static inline bool mem_cgroup_online(struct mem_cgroup *memcg)
{
if (mem_cgroup_disabled())
return true;
return !!(memcg->css.flags & CSS_ONLINE);
}
/*
* For memory reclaim.
*/
int mem_cgroup_select_victim_node(struct mem_cgroup *memcg);
void mem_cgroup_update_lru_size(struct lruvec *lruvec, enum lru_list lru,
mm, vmscan: Update all zone LRU sizes before updating memcg Minchan Kim reported setting the following warning on a 32-bit system although it can affect 64-bit systems. WARNING: CPU: 4 PID: 1322 at mm/memcontrol.c:998 mem_cgroup_update_lru_size+0x103/0x110 mem_cgroup_update_lru_size(f44b4000, 1, -7): zid 1 lru_size 1 but empty Modules linked in: CPU: 4 PID: 1322 Comm: cp Not tainted 4.7.0-rc4-mm1+ #143 Hardware name: QEMU Standard PC (i440FX + PIIX, 1996), BIOS Bochs 01/01/2011 Call Trace: dump_stack+0x76/0xaf __warn+0xea/0x110 ? mem_cgroup_update_lru_size+0x103/0x110 warn_slowpath_fmt+0x3b/0x40 mem_cgroup_update_lru_size+0x103/0x110 isolate_lru_pages.isra.61+0x2e2/0x360 shrink_active_list+0xac/0x2a0 ? __delay+0xe/0x10 shrink_node_memcg+0x53c/0x7a0 shrink_node+0xab/0x2a0 do_try_to_free_pages+0xc6/0x390 try_to_free_pages+0x245/0x590 LRU list contents and counts are updated separately. Counts are updated before pages are added to the LRU and updated after pages are removed. The warning above is from a check in mem_cgroup_update_lru_size that ensures that list sizes of zero are empty. The problem is that node-lru needs to account for highmem pages if CONFIG_HIGHMEM is set. One impact of the implementation is that the sizes are updated in multiple passes when pages from multiple zones were isolated. This happens whether HIGHMEM is set or not. When multiple zones are isolated, it's possible for a debugging check in memcg to be tripped. This patch forces all the zone counts to be updated before the memcg function is called. Link: http://lkml.kernel.org/r/1468588165-12461-6-git-send-email-mgorman@techsingularity.net Signed-off-by: Mel Gorman <mgorman@techsingularity.net> Tested-by: Minchan Kim <minchan@kernel.org> Reported-by: Minchan Kim <minchan@kernel.org> Acked-by: Minchan Kim <minchan@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:47:17 +08:00
int nr_pages);
unsigned long mem_cgroup_node_nr_lru_pages(struct mem_cgroup *memcg,
int nid, unsigned int lru_mask);
static inline
unsigned long mem_cgroup_get_lru_size(struct lruvec *lruvec, enum lru_list lru)
{
struct mem_cgroup_per_node *mz;
mz = container_of(lruvec, struct mem_cgroup_per_node, lruvec);
return mz->lru_size[lru];
}
memcg: punt high overage reclaim to return-to-userland path Currently, try_charge() tries to reclaim memory synchronously when the high limit is breached; however, if the allocation doesn't have __GFP_WAIT, synchronous reclaim is skipped. If a process performs only speculative allocations, it can blow way past the high limit. This is actually easily reproducible by simply doing "find /". slab/slub allocator tries speculative allocations first, so as long as there's memory which can be consumed without blocking, it can keep allocating memory regardless of the high limit. This patch makes try_charge() always punt the over-high reclaim to the return-to-userland path. If try_charge() detects that high limit is breached, it adds the overage to current->memcg_nr_pages_over_high and schedules execution of mem_cgroup_handle_over_high() which performs synchronous reclaim from the return-to-userland path. As long as kernel doesn't have a run-away allocation spree, this should provide enough protection while making kmemcg behave more consistently. It also has the following benefits. - All over-high reclaims can use GFP_KERNEL regardless of the specific gfp mask in use, e.g. GFP_NOFS, when the limit was breached. - It copes with prio inversion. Previously, a low-prio task with small memory.high might perform over-high reclaim with a bunch of locks held. If a higher prio task needed any of these locks, it would have to wait until the low prio task finished reclaim and released the locks. By handing over-high reclaim to the task exit path this issue can be avoided. Signed-off-by: Tejun Heo <tj@kernel.org> Acked-by: Michal Hocko <mhocko@kernel.org> Reviewed-by: Vladimir Davydov <vdavydov@parallels.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-11-06 10:46:11 +08:00
void mem_cgroup_handle_over_high(void);
void mem_cgroup_print_oom_info(struct mem_cgroup *memcg,
struct task_struct *p);
static inline void mem_cgroup_oom_enable(void)
{
WARN_ON(current->memcg_may_oom);
current->memcg_may_oom = 1;
}
static inline void mem_cgroup_oom_disable(void)
{
WARN_ON(!current->memcg_may_oom);
current->memcg_may_oom = 0;
}
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
static inline bool task_in_memcg_oom(struct task_struct *p)
{
return p->memcg_in_oom;
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
}
bool mem_cgroup_oom_synchronize(bool wait);
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
#ifdef CONFIG_MEMCG_SWAP
extern int do_swap_account;
#endif
void lock_page_memcg(struct page *page);
void unlock_page_memcg(struct page *page);
mm: memcontrol: fix missed end-writeback page accounting Commit 0a31bc97c80c ("mm: memcontrol: rewrite uncharge API") changed page migration to uncharge the old page right away. The page is locked, unmapped, truncated, and off the LRU, but it could race with writeback ending, which then doesn't unaccount the page properly: test_clear_page_writeback() migration wait_on_page_writeback() TestClearPageWriteback() mem_cgroup_migrate() clear PCG_USED mem_cgroup_update_page_stat() if (PageCgroupUsed(pc)) decrease memcg pages under writeback release pc->mem_cgroup->move_lock The per-page statistics interface is heavily optimized to avoid a function call and a lookup_page_cgroup() in the file unmap fast path, which means it doesn't verify whether a page is still charged before clearing PageWriteback() and it has to do it in the stat update later. Rework it so that it looks up the page's memcg once at the beginning of the transaction and then uses it throughout. The charge will be verified before clearing PageWriteback() and migration can't uncharge the page as long as that is still set. The RCU lock will protect the memcg past uncharge. As far as losing the optimization goes, the following test results are from a microbenchmark that maps, faults, and unmaps a 4GB sparse file three times in a nested fashion, so that there are two negative passes that don't account but still go through the new transaction overhead. There is no actual difference: old: 33.195102545 seconds time elapsed ( +- 0.01% ) new: 33.199231369 seconds time elapsed ( +- 0.03% ) The time spent in page_remove_rmap()'s callees still adds up to the same, but the time spent in the function itself seems reduced: # Children Self Command Shared Object Symbol old: 0.12% 0.11% filemapstress [kernel.kallsyms] [k] page_remove_rmap new: 0.12% 0.08% filemapstress [kernel.kallsyms] [k] page_remove_rmap Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: <stable@vger.kernel.org> [3.17.x] Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-10-30 05:50:48 +08:00
/**
* mem_cgroup_update_page_stat - update page state statistics
* @page: the page
* @idx: page state item to account
* @val: number of pages (positive or negative)
*
* The @page must be locked or the caller must use lock_page_memcg()
* to prevent double accounting when the page is concurrently being
* moved to another memcg:
*
* lock_page(page) or lock_page_memcg(page)
* if (TestClearPageState(page))
* mem_cgroup_update_page_stat(page, state, -1);
* unlock_page(page) or unlock_page_memcg(page)
*/
static inline void mem_cgroup_update_page_stat(struct page *page,
enum mem_cgroup_stat_index idx, int val)
{
VM_BUG_ON(!(rcu_read_lock_held() || PageLocked(page)));
if (page->mem_cgroup)
this_cpu_add(page->mem_cgroup->stat->count[idx], val);
}
static inline void mem_cgroup_inc_page_stat(struct page *page,
enum mem_cgroup_stat_index idx)
{
mem_cgroup_update_page_stat(page, idx, 1);
}
static inline void mem_cgroup_dec_page_stat(struct page *page,
enum mem_cgroup_stat_index idx)
{
mem_cgroup_update_page_stat(page, idx, -1);
}
unsigned long mem_cgroup_soft_limit_reclaim(pg_data_t *pgdat, int order,
gfp_t gfp_mask,
unsigned long *total_scanned);
oom: badness heuristic rewrite This a complete rewrite of the oom killer's badness() heuristic which is used to determine which task to kill in oom conditions. The goal is to make it as simple and predictable as possible so the results are better understood and we end up killing the task which will lead to the most memory freeing while still respecting the fine-tuning from userspace. Instead of basing the heuristic on mm->total_vm for each task, the task's rss and swap space is used instead. This is a better indication of the amount of memory that will be freeable if the oom killed task is chosen and subsequently exits. This helps specifically in cases where KDE or GNOME is chosen for oom kill on desktop systems instead of a memory hogging task. The baseline for the heuristic is a proportion of memory that each task is currently using in memory plus swap compared to the amount of "allowable" memory. "Allowable," in this sense, means the system-wide resources for unconstrained oom conditions, the set of mempolicy nodes, the mems attached to current's cpuset, or a memory controller's limit. The proportion is given on a scale of 0 (never kill) to 1000 (always kill), roughly meaning that if a task has a badness() score of 500 that the task consumes approximately 50% of allowable memory resident in RAM or in swap space. The proportion is always relative to the amount of "allowable" memory and not the total amount of RAM systemwide so that mempolicies and cpusets may operate in isolation; they shall not need to know the true size of the machine on which they are running if they are bound to a specific set of nodes or mems, respectively. Root tasks are given 3% extra memory just like __vm_enough_memory() provides in LSMs. In the event of two tasks consuming similar amounts of memory, it is generally better to save root's task. Because of the change in the badness() heuristic's baseline, it is also necessary to introduce a new user interface to tune it. It's not possible to redefine the meaning of /proc/pid/oom_adj with a new scale since the ABI cannot be changed for backward compatability. Instead, a new tunable, /proc/pid/oom_score_adj, is added that ranges from -1000 to +1000. It may be used to polarize the heuristic such that certain tasks are never considered for oom kill while others may always be considered. The value is added directly into the badness() score so a value of -500, for example, means to discount 50% of its memory consumption in comparison to other tasks either on the system, bound to the mempolicy, in the cpuset, or sharing the same memory controller. /proc/pid/oom_adj is changed so that its meaning is rescaled into the units used by /proc/pid/oom_score_adj, and vice versa. Changing one of these per-task tunables will rescale the value of the other to an equivalent meaning. Although /proc/pid/oom_adj was originally defined as a bitshift on the badness score, it now shares the same linear growth as /proc/pid/oom_score_adj but with different granularity. This is required so the ABI is not broken with userspace applications and allows oom_adj to be deprecated for future removal. Signed-off-by: David Rientjes <rientjes@google.com> Cc: Nick Piggin <npiggin@suse.de> Cc: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Oleg Nesterov <oleg@redhat.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>
2010-08-10 08:19:46 +08:00
static inline void mem_cgroup_count_vm_event(struct mm_struct *mm,
enum vm_event_item idx)
{
struct mem_cgroup *memcg;
if (mem_cgroup_disabled())
return;
rcu_read_lock();
memcg = mem_cgroup_from_task(rcu_dereference(mm->owner));
if (unlikely(!memcg))
goto out;
switch (idx) {
case PGFAULT:
this_cpu_inc(memcg->stat->events[MEM_CGROUP_EVENTS_PGFAULT]);
break;
case PGMAJFAULT:
this_cpu_inc(memcg->stat->events[MEM_CGROUP_EVENTS_PGMAJFAULT]);
break;
default:
BUG();
}
out:
rcu_read_unlock();
}
#ifdef CONFIG_TRANSPARENT_HUGEPAGE
void mem_cgroup_split_huge_fixup(struct page *head);
#endif
#else /* CONFIG_MEMCG */
#define MEM_CGROUP_ID_SHIFT 0
#define MEM_CGROUP_ID_MAX 0
memcg: introduce charge-commit-cancel style of functions There is a small race in do_swap_page(). When the page swapped-in is charged, the mapcount can be greater than 0. But, at the same time some process (shares it ) call unmap and make mapcount 1->0 and the page is uncharged. CPUA CPUB mapcount == 1. (1) charge if mapcount==0 zap_pte_range() (2) mapcount 1 => 0. (3) uncharge(). (success) (4) set page's rmap() mapcount 0=>1 Then, this swap page's account is leaked. For fixing this, I added a new interface. - charge account to res_counter by PAGE_SIZE and try to free pages if necessary. - commit register page_cgroup and add to LRU if necessary. - cancel uncharge PAGE_SIZE because of do_swap_page failure. CPUA (1) charge (always) (2) set page's rmap (mapcount > 0) (3) commit charge was necessary or not after set_pte(). This protocol uses PCG_USED bit on page_cgroup for avoiding over accounting. Usual mem_cgroup_charge_common() does charge -> commit at a time. And this patch also adds following function to clarify all charges. - mem_cgroup_newpage_charge() ....replacement for mem_cgroup_charge() called against newly allocated anon pages. - mem_cgroup_charge_migrate_fixup() called only from remove_migration_ptes(). we'll have to rewrite this later.(this patch just keeps old behavior) This function will be removed by additional patch to make migration clearer. Good for clarifying "what we do" Then, we have 4 following charge points. - newpage - swap-in - add-to-cache. - migration. [akpm@linux-foundation.org: add missing inline directives to stubs] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> 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>
2009-01-08 10:07:48 +08:00
struct mem_cgroup;
static inline bool mem_cgroup_disabled(void)
{
return true;
}
mm: memcontrol: default hierarchy interface for memory Introduce the basic control files to account, partition, and limit memory using cgroups in default hierarchy mode. This interface versioning allows us to address fundamental design issues in the existing memory cgroup interface, further explained below. The old interface will be maintained indefinitely, but a clearer model and improved workload performance should encourage existing users to switch over to the new one eventually. The control files are thus: - memory.current shows the current consumption of the cgroup and its descendants, in bytes. - memory.low configures the lower end of the cgroup's expected memory consumption range. The kernel considers memory below that boundary to be a reserve - the minimum that the workload needs in order to make forward progress - and generally avoids reclaiming it, unless there is an imminent risk of entering an OOM situation. - memory.high configures the upper end of the cgroup's expected memory consumption range. A cgroup whose consumption grows beyond this threshold is forced into direct reclaim, to work off the excess and to throttle new allocations heavily, but is generally allowed to continue and the OOM killer is not invoked. - memory.max configures the hard maximum amount of memory that the cgroup is allowed to consume before the OOM killer is invoked. - memory.events shows event counters that indicate how often the cgroup was reclaimed while below memory.low, how often it was forced to reclaim excess beyond memory.high, how often it hit memory.max, and how often it entered OOM due to memory.max. This allows users to identify configuration problems when observing a degradation in workload performance. An overcommitted system will have an increased rate of low boundary breaches, whereas increased rates of high limit breaches, maximum hits, or even OOM situations will indicate internally overcommitted cgroups. For existing users of memory cgroups, the following deviations from the current interface are worth pointing out and explaining: - The original lower boundary, the soft limit, is defined as a limit that is per default unset. As a result, the set of cgroups that global reclaim prefers is opt-in, rather than opt-out. The costs for optimizing these mostly negative lookups are so high that the implementation, despite its enormous size, does not even provide the basic desirable behavior. First off, the soft limit has no hierarchical meaning. All configured groups are organized in a global rbtree and treated like equal peers, regardless where they are located in the hierarchy. This makes subtree delegation impossible. Second, the soft limit reclaim pass is so aggressive that it not just introduces high allocation latencies into the system, but also impacts system performance due to overreclaim, to the point where the feature becomes self-defeating. The memory.low boundary on the other hand is a top-down allocated reserve. A cgroup enjoys reclaim protection when it and all its ancestors are below their low boundaries, which makes delegation of subtrees possible. Secondly, new cgroups have no reserve per default and in the common case most cgroups are eligible for the preferred reclaim pass. This allows the new low boundary to be efficiently implemented with just a minor addition to the generic reclaim code, without the need for out-of-band data structures and reclaim passes. Because the generic reclaim code considers all cgroups except for the ones running low in the preferred first reclaim pass, overreclaim of individual groups is eliminated as well, resulting in much better overall workload performance. - The original high boundary, the hard limit, is defined as a strict limit that can not budge, even if the OOM killer has to be called. But this generally goes against the goal of making the most out of the available memory. The memory consumption of workloads varies during runtime, and that requires users to overcommit. But doing that with a strict upper limit requires either a fairly accurate prediction of the working set size or adding slack to the limit. Since working set size estimation is hard and error prone, and getting it wrong results in OOM kills, most users tend to err on the side of a looser limit and end up wasting precious resources. The memory.high boundary on the other hand can be set much more conservatively. When hit, it throttles allocations by forcing them into direct reclaim to work off the excess, but it never invokes the OOM killer. As a result, a high boundary that is chosen too aggressively will not terminate the processes, but instead it will lead to gradual performance degradation. The user can monitor this and make corrections until the minimal memory footprint that still gives acceptable performance is found. In extreme cases, with many concurrent allocations and a complete breakdown of reclaim progress within the group, the high boundary can be exceeded. But even then it's mostly better to satisfy the allocation from the slack available in other groups or the rest of the system than killing the group. Otherwise, memory.max is there to limit this type of spillover and ultimately contain buggy or even malicious applications. - The original control file names are unwieldy and inconsistent in many different ways. For example, the upper boundary hit count is exported in the memory.failcnt file, but an OOM event count has to be manually counted by listening to memory.oom_control events, and lower boundary / soft limit events have to be counted by first setting a threshold for that value and then counting those events. Also, usage and limit files encode their units in the filename. That makes the filenames very long, even though this is not information that a user needs to be reminded of every time they type out those names. To address these naming issues, as well as to signal clearly that the new interface carries a new configuration model, the naming conventions in it necessarily differ from the old interface. - The original limit files indicate the state of an unset limit with a very high number, and a configured limit can be unset by echoing -1 into those files. But that very high number is implementation and architecture dependent and not very descriptive. And while -1 can be understood as an underflow into the highest possible value, -2 or -10M etc. do not work, so it's not inconsistent. memory.low, memory.high, and memory.max will use the string "infinity" to indicate and set the highest possible value. [akpm@linux-foundation.org: use seq_puts() for basic strings] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Vladimir Davydov <vdavydov@parallels.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-12 07:26:06 +08:00
static inline void mem_cgroup_events(struct mem_cgroup *memcg,
enum mem_cgroup_events_index idx,
unsigned int nr)
{
}
static inline bool mem_cgroup_low(struct mem_cgroup *root,
struct mem_cgroup *memcg)
{
return false;
}
mm: memcontrol: rewrite charge API These patches rework memcg charge lifetime to integrate more naturally with the lifetime of user pages. This drastically simplifies the code and reduces charging and uncharging overhead. The most expensive part of charging and uncharging is the page_cgroup bit spinlock, which is removed entirely after this series. Here are the top-10 profile entries of a stress test that reads a 128G sparse file on a freshly booted box, without even a dedicated cgroup (i.e. executing in the root memcg). Before: 15.36% cat [kernel.kallsyms] [k] copy_user_generic_string 13.31% cat [kernel.kallsyms] [k] memset 11.48% cat [kernel.kallsyms] [k] do_mpage_readpage 4.23% cat [kernel.kallsyms] [k] get_page_from_freelist 2.38% cat [kernel.kallsyms] [k] put_page 2.32% cat [kernel.kallsyms] [k] __mem_cgroup_commit_charge 2.18% kswapd0 [kernel.kallsyms] [k] __mem_cgroup_uncharge_common 1.92% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.86% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.62% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn After: 15.67% cat [kernel.kallsyms] [k] copy_user_generic_string 13.48% cat [kernel.kallsyms] [k] memset 11.42% cat [kernel.kallsyms] [k] do_mpage_readpage 3.98% cat [kernel.kallsyms] [k] get_page_from_freelist 2.46% cat [kernel.kallsyms] [k] put_page 2.13% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.88% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.67% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn 1.39% kswapd0 [kernel.kallsyms] [k] free_pcppages_bulk 1.30% cat [kernel.kallsyms] [k] kfree As you can see, the memcg footprint has shrunk quite a bit. text data bss dec hex filename 37970 9892 400 48262 bc86 mm/memcontrol.o.old 35239 9892 400 45531 b1db mm/memcontrol.o This patch (of 4): The memcg charge API charges pages before they are rmapped - i.e. have an actual "type" - and so every callsite needs its own set of charge and uncharge functions to know what type is being operated on. Worse, uncharge has to happen from a context that is still type-specific, rather than at the end of the page's lifetime with exclusive access, and so requires a lot of synchronization. Rewrite the charge API to provide a generic set of try_charge(), commit_charge() and cancel_charge() transaction operations, much like what's currently done for swap-in: mem_cgroup_try_charge() attempts to reserve a charge, reclaiming pages from the memcg if necessary. mem_cgroup_commit_charge() commits the page to the charge once it has a valid page->mapping and PageAnon() reliably tells the type. mem_cgroup_cancel_charge() aborts the transaction. This reduces the charge API and enables subsequent patches to drastically simplify uncharging. As pages need to be committed after rmap is established but before they are added to the LRU, page_add_new_anon_rmap() must stop doing LRU additions again. Revive lru_cache_add_active_or_unevictable(). [hughd@google.com: fix shmem_unuse] [hughd@google.com: Add comments on the private use of -EAGAIN] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Hugh Dickins <hughd@google.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:20 +08:00
static inline int mem_cgroup_try_charge(struct page *page, struct mm_struct *mm,
gfp_t gfp_mask,
memcg: adjust to support new THP refcounting As with rmap, with new refcounting we cannot rely on PageTransHuge() to check if we need to charge size of huge page form the cgroup. We need to get information from caller to know whether it was mapped with PMD or PTE. We do uncharge when last reference on the page gone. At that point if we see PageTransHuge() it means we need to unchange whole huge page. The tricky part is partial unmap -- when we try to unmap part of huge page. We don't do a special handing of this situation, meaning we don't uncharge the part of huge page unless last user is gone or split_huge_page() is triggered. In case of cgroup memory pressure happens the partial unmapped page will be split through shrinker. This should be good enough. Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Acked-by: Jerome Marchand <jmarchan@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hughd@google.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Steve Capper <steve.capper@linaro.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Christoph Lameter <cl@linux.com> Cc: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-16 08:52:20 +08:00
struct mem_cgroup **memcgp,
bool compound)
memcg: introduce charge-commit-cancel style of functions There is a small race in do_swap_page(). When the page swapped-in is charged, the mapcount can be greater than 0. But, at the same time some process (shares it ) call unmap and make mapcount 1->0 and the page is uncharged. CPUA CPUB mapcount == 1. (1) charge if mapcount==0 zap_pte_range() (2) mapcount 1 => 0. (3) uncharge(). (success) (4) set page's rmap() mapcount 0=>1 Then, this swap page's account is leaked. For fixing this, I added a new interface. - charge account to res_counter by PAGE_SIZE and try to free pages if necessary. - commit register page_cgroup and add to LRU if necessary. - cancel uncharge PAGE_SIZE because of do_swap_page failure. CPUA (1) charge (always) (2) set page's rmap (mapcount > 0) (3) commit charge was necessary or not after set_pte(). This protocol uses PCG_USED bit on page_cgroup for avoiding over accounting. Usual mem_cgroup_charge_common() does charge -> commit at a time. And this patch also adds following function to clarify all charges. - mem_cgroup_newpage_charge() ....replacement for mem_cgroup_charge() called against newly allocated anon pages. - mem_cgroup_charge_migrate_fixup() called only from remove_migration_ptes(). we'll have to rewrite this later.(this patch just keeps old behavior) This function will be removed by additional patch to make migration clearer. Good for clarifying "what we do" Then, we have 4 following charge points. - newpage - swap-in - add-to-cache. - migration. [akpm@linux-foundation.org: add missing inline directives to stubs] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> 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>
2009-01-08 10:07:48 +08:00
{
mm: memcontrol: rewrite charge API These patches rework memcg charge lifetime to integrate more naturally with the lifetime of user pages. This drastically simplifies the code and reduces charging and uncharging overhead. The most expensive part of charging and uncharging is the page_cgroup bit spinlock, which is removed entirely after this series. Here are the top-10 profile entries of a stress test that reads a 128G sparse file on a freshly booted box, without even a dedicated cgroup (i.e. executing in the root memcg). Before: 15.36% cat [kernel.kallsyms] [k] copy_user_generic_string 13.31% cat [kernel.kallsyms] [k] memset 11.48% cat [kernel.kallsyms] [k] do_mpage_readpage 4.23% cat [kernel.kallsyms] [k] get_page_from_freelist 2.38% cat [kernel.kallsyms] [k] put_page 2.32% cat [kernel.kallsyms] [k] __mem_cgroup_commit_charge 2.18% kswapd0 [kernel.kallsyms] [k] __mem_cgroup_uncharge_common 1.92% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.86% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.62% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn After: 15.67% cat [kernel.kallsyms] [k] copy_user_generic_string 13.48% cat [kernel.kallsyms] [k] memset 11.42% cat [kernel.kallsyms] [k] do_mpage_readpage 3.98% cat [kernel.kallsyms] [k] get_page_from_freelist 2.46% cat [kernel.kallsyms] [k] put_page 2.13% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.88% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.67% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn 1.39% kswapd0 [kernel.kallsyms] [k] free_pcppages_bulk 1.30% cat [kernel.kallsyms] [k] kfree As you can see, the memcg footprint has shrunk quite a bit. text data bss dec hex filename 37970 9892 400 48262 bc86 mm/memcontrol.o.old 35239 9892 400 45531 b1db mm/memcontrol.o This patch (of 4): The memcg charge API charges pages before they are rmapped - i.e. have an actual "type" - and so every callsite needs its own set of charge and uncharge functions to know what type is being operated on. Worse, uncharge has to happen from a context that is still type-specific, rather than at the end of the page's lifetime with exclusive access, and so requires a lot of synchronization. Rewrite the charge API to provide a generic set of try_charge(), commit_charge() and cancel_charge() transaction operations, much like what's currently done for swap-in: mem_cgroup_try_charge() attempts to reserve a charge, reclaiming pages from the memcg if necessary. mem_cgroup_commit_charge() commits the page to the charge once it has a valid page->mapping and PageAnon() reliably tells the type. mem_cgroup_cancel_charge() aborts the transaction. This reduces the charge API and enables subsequent patches to drastically simplify uncharging. As pages need to be committed after rmap is established but before they are added to the LRU, page_add_new_anon_rmap() must stop doing LRU additions again. Revive lru_cache_add_active_or_unevictable(). [hughd@google.com: fix shmem_unuse] [hughd@google.com: Add comments on the private use of -EAGAIN] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Hugh Dickins <hughd@google.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:20 +08:00
*memcgp = NULL;
memcg: introduce charge-commit-cancel style of functions There is a small race in do_swap_page(). When the page swapped-in is charged, the mapcount can be greater than 0. But, at the same time some process (shares it ) call unmap and make mapcount 1->0 and the page is uncharged. CPUA CPUB mapcount == 1. (1) charge if mapcount==0 zap_pte_range() (2) mapcount 1 => 0. (3) uncharge(). (success) (4) set page's rmap() mapcount 0=>1 Then, this swap page's account is leaked. For fixing this, I added a new interface. - charge account to res_counter by PAGE_SIZE and try to free pages if necessary. - commit register page_cgroup and add to LRU if necessary. - cancel uncharge PAGE_SIZE because of do_swap_page failure. CPUA (1) charge (always) (2) set page's rmap (mapcount > 0) (3) commit charge was necessary or not after set_pte(). This protocol uses PCG_USED bit on page_cgroup for avoiding over accounting. Usual mem_cgroup_charge_common() does charge -> commit at a time. And this patch also adds following function to clarify all charges. - mem_cgroup_newpage_charge() ....replacement for mem_cgroup_charge() called against newly allocated anon pages. - mem_cgroup_charge_migrate_fixup() called only from remove_migration_ptes(). we'll have to rewrite this later.(this patch just keeps old behavior) This function will be removed by additional patch to make migration clearer. Good for clarifying "what we do" Then, we have 4 following charge points. - newpage - swap-in - add-to-cache. - migration. [akpm@linux-foundation.org: add missing inline directives to stubs] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> 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>
2009-01-08 10:07:48 +08:00
return 0;
}
mm: memcontrol: rewrite charge API These patches rework memcg charge lifetime to integrate more naturally with the lifetime of user pages. This drastically simplifies the code and reduces charging and uncharging overhead. The most expensive part of charging and uncharging is the page_cgroup bit spinlock, which is removed entirely after this series. Here are the top-10 profile entries of a stress test that reads a 128G sparse file on a freshly booted box, without even a dedicated cgroup (i.e. executing in the root memcg). Before: 15.36% cat [kernel.kallsyms] [k] copy_user_generic_string 13.31% cat [kernel.kallsyms] [k] memset 11.48% cat [kernel.kallsyms] [k] do_mpage_readpage 4.23% cat [kernel.kallsyms] [k] get_page_from_freelist 2.38% cat [kernel.kallsyms] [k] put_page 2.32% cat [kernel.kallsyms] [k] __mem_cgroup_commit_charge 2.18% kswapd0 [kernel.kallsyms] [k] __mem_cgroup_uncharge_common 1.92% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.86% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.62% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn After: 15.67% cat [kernel.kallsyms] [k] copy_user_generic_string 13.48% cat [kernel.kallsyms] [k] memset 11.42% cat [kernel.kallsyms] [k] do_mpage_readpage 3.98% cat [kernel.kallsyms] [k] get_page_from_freelist 2.46% cat [kernel.kallsyms] [k] put_page 2.13% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.88% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.67% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn 1.39% kswapd0 [kernel.kallsyms] [k] free_pcppages_bulk 1.30% cat [kernel.kallsyms] [k] kfree As you can see, the memcg footprint has shrunk quite a bit. text data bss dec hex filename 37970 9892 400 48262 bc86 mm/memcontrol.o.old 35239 9892 400 45531 b1db mm/memcontrol.o This patch (of 4): The memcg charge API charges pages before they are rmapped - i.e. have an actual "type" - and so every callsite needs its own set of charge and uncharge functions to know what type is being operated on. Worse, uncharge has to happen from a context that is still type-specific, rather than at the end of the page's lifetime with exclusive access, and so requires a lot of synchronization. Rewrite the charge API to provide a generic set of try_charge(), commit_charge() and cancel_charge() transaction operations, much like what's currently done for swap-in: mem_cgroup_try_charge() attempts to reserve a charge, reclaiming pages from the memcg if necessary. mem_cgroup_commit_charge() commits the page to the charge once it has a valid page->mapping and PageAnon() reliably tells the type. mem_cgroup_cancel_charge() aborts the transaction. This reduces the charge API and enables subsequent patches to drastically simplify uncharging. As pages need to be committed after rmap is established but before they are added to the LRU, page_add_new_anon_rmap() must stop doing LRU additions again. Revive lru_cache_add_active_or_unevictable(). [hughd@google.com: fix shmem_unuse] [hughd@google.com: Add comments on the private use of -EAGAIN] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Hugh Dickins <hughd@google.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:20 +08:00
static inline void mem_cgroup_commit_charge(struct page *page,
struct mem_cgroup *memcg,
memcg: adjust to support new THP refcounting As with rmap, with new refcounting we cannot rely on PageTransHuge() to check if we need to charge size of huge page form the cgroup. We need to get information from caller to know whether it was mapped with PMD or PTE. We do uncharge when last reference on the page gone. At that point if we see PageTransHuge() it means we need to unchange whole huge page. The tricky part is partial unmap -- when we try to unmap part of huge page. We don't do a special handing of this situation, meaning we don't uncharge the part of huge page unless last user is gone or split_huge_page() is triggered. In case of cgroup memory pressure happens the partial unmapped page will be split through shrinker. This should be good enough. Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Acked-by: Jerome Marchand <jmarchan@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hughd@google.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Steve Capper <steve.capper@linaro.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Christoph Lameter <cl@linux.com> Cc: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-16 08:52:20 +08:00
bool lrucare, bool compound)
memcg: introduce charge-commit-cancel style of functions There is a small race in do_swap_page(). When the page swapped-in is charged, the mapcount can be greater than 0. But, at the same time some process (shares it ) call unmap and make mapcount 1->0 and the page is uncharged. CPUA CPUB mapcount == 1. (1) charge if mapcount==0 zap_pte_range() (2) mapcount 1 => 0. (3) uncharge(). (success) (4) set page's rmap() mapcount 0=>1 Then, this swap page's account is leaked. For fixing this, I added a new interface. - charge account to res_counter by PAGE_SIZE and try to free pages if necessary. - commit register page_cgroup and add to LRU if necessary. - cancel uncharge PAGE_SIZE because of do_swap_page failure. CPUA (1) charge (always) (2) set page's rmap (mapcount > 0) (3) commit charge was necessary or not after set_pte(). This protocol uses PCG_USED bit on page_cgroup for avoiding over accounting. Usual mem_cgroup_charge_common() does charge -> commit at a time. And this patch also adds following function to clarify all charges. - mem_cgroup_newpage_charge() ....replacement for mem_cgroup_charge() called against newly allocated anon pages. - mem_cgroup_charge_migrate_fixup() called only from remove_migration_ptes(). we'll have to rewrite this later.(this patch just keeps old behavior) This function will be removed by additional patch to make migration clearer. Good for clarifying "what we do" Then, we have 4 following charge points. - newpage - swap-in - add-to-cache. - migration. [akpm@linux-foundation.org: add missing inline directives to stubs] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> 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>
2009-01-08 10:07:48 +08:00
{
}
mm: memcontrol: rewrite charge API These patches rework memcg charge lifetime to integrate more naturally with the lifetime of user pages. This drastically simplifies the code and reduces charging and uncharging overhead. The most expensive part of charging and uncharging is the page_cgroup bit spinlock, which is removed entirely after this series. Here are the top-10 profile entries of a stress test that reads a 128G sparse file on a freshly booted box, without even a dedicated cgroup (i.e. executing in the root memcg). Before: 15.36% cat [kernel.kallsyms] [k] copy_user_generic_string 13.31% cat [kernel.kallsyms] [k] memset 11.48% cat [kernel.kallsyms] [k] do_mpage_readpage 4.23% cat [kernel.kallsyms] [k] get_page_from_freelist 2.38% cat [kernel.kallsyms] [k] put_page 2.32% cat [kernel.kallsyms] [k] __mem_cgroup_commit_charge 2.18% kswapd0 [kernel.kallsyms] [k] __mem_cgroup_uncharge_common 1.92% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.86% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.62% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn After: 15.67% cat [kernel.kallsyms] [k] copy_user_generic_string 13.48% cat [kernel.kallsyms] [k] memset 11.42% cat [kernel.kallsyms] [k] do_mpage_readpage 3.98% cat [kernel.kallsyms] [k] get_page_from_freelist 2.46% cat [kernel.kallsyms] [k] put_page 2.13% kswapd0 [kernel.kallsyms] [k] shrink_page_list 1.88% cat [kernel.kallsyms] [k] __radix_tree_lookup 1.67% cat [kernel.kallsyms] [k] __pagevec_lru_add_fn 1.39% kswapd0 [kernel.kallsyms] [k] free_pcppages_bulk 1.30% cat [kernel.kallsyms] [k] kfree As you can see, the memcg footprint has shrunk quite a bit. text data bss dec hex filename 37970 9892 400 48262 bc86 mm/memcontrol.o.old 35239 9892 400 45531 b1db mm/memcontrol.o This patch (of 4): The memcg charge API charges pages before they are rmapped - i.e. have an actual "type" - and so every callsite needs its own set of charge and uncharge functions to know what type is being operated on. Worse, uncharge has to happen from a context that is still type-specific, rather than at the end of the page's lifetime with exclusive access, and so requires a lot of synchronization. Rewrite the charge API to provide a generic set of try_charge(), commit_charge() and cancel_charge() transaction operations, much like what's currently done for swap-in: mem_cgroup_try_charge() attempts to reserve a charge, reclaiming pages from the memcg if necessary. mem_cgroup_commit_charge() commits the page to the charge once it has a valid page->mapping and PageAnon() reliably tells the type. mem_cgroup_cancel_charge() aborts the transaction. This reduces the charge API and enables subsequent patches to drastically simplify uncharging. As pages need to be committed after rmap is established but before they are added to the LRU, page_add_new_anon_rmap() must stop doing LRU additions again. Revive lru_cache_add_active_or_unevictable(). [hughd@google.com: fix shmem_unuse] [hughd@google.com: Add comments on the private use of -EAGAIN] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Michal Hocko <mhocko@suse.cz> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Hugh Dickins <hughd@google.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:20 +08:00
static inline void mem_cgroup_cancel_charge(struct page *page,
memcg: adjust to support new THP refcounting As with rmap, with new refcounting we cannot rely on PageTransHuge() to check if we need to charge size of huge page form the cgroup. We need to get information from caller to know whether it was mapped with PMD or PTE. We do uncharge when last reference on the page gone. At that point if we see PageTransHuge() it means we need to unchange whole huge page. The tricky part is partial unmap -- when we try to unmap part of huge page. We don't do a special handing of this situation, meaning we don't uncharge the part of huge page unless last user is gone or split_huge_page() is triggered. In case of cgroup memory pressure happens the partial unmapped page will be split through shrinker. This should be good enough. Signed-off-by: Kirill A. Shutemov <kirill.shutemov@linux.intel.com> Tested-by: Sasha Levin <sasha.levin@oracle.com> Tested-by: Aneesh Kumar K.V <aneesh.kumar@linux.vnet.ibm.com> Acked-by: Vlastimil Babka <vbabka@suse.cz> Acked-by: Jerome Marchand <jmarchan@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Hugh Dickins <hughd@google.com> Cc: Dave Hansen <dave.hansen@intel.com> Cc: Mel Gorman <mgorman@suse.de> Cc: Rik van Riel <riel@redhat.com> Cc: Naoya Horiguchi <n-horiguchi@ah.jp.nec.com> Cc: Steve Capper <steve.capper@linaro.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Christoph Lameter <cl@linux.com> Cc: David Rientjes <rientjes@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-16 08:52:20 +08:00
struct mem_cgroup *memcg,
bool compound)
memcg: introduce charge-commit-cancel style of functions There is a small race in do_swap_page(). When the page swapped-in is charged, the mapcount can be greater than 0. But, at the same time some process (shares it ) call unmap and make mapcount 1->0 and the page is uncharged. CPUA CPUB mapcount == 1. (1) charge if mapcount==0 zap_pte_range() (2) mapcount 1 => 0. (3) uncharge(). (success) (4) set page's rmap() mapcount 0=>1 Then, this swap page's account is leaked. For fixing this, I added a new interface. - charge account to res_counter by PAGE_SIZE and try to free pages if necessary. - commit register page_cgroup and add to LRU if necessary. - cancel uncharge PAGE_SIZE because of do_swap_page failure. CPUA (1) charge (always) (2) set page's rmap (mapcount > 0) (3) commit charge was necessary or not after set_pte(). This protocol uses PCG_USED bit on page_cgroup for avoiding over accounting. Usual mem_cgroup_charge_common() does charge -> commit at a time. And this patch also adds following function to clarify all charges. - mem_cgroup_newpage_charge() ....replacement for mem_cgroup_charge() called against newly allocated anon pages. - mem_cgroup_charge_migrate_fixup() called only from remove_migration_ptes(). we'll have to rewrite this later.(this patch just keeps old behavior) This function will be removed by additional patch to make migration clearer. Good for clarifying "what we do" Then, we have 4 following charge points. - newpage - swap-in - add-to-cache. - migration. [akpm@linux-foundation.org: add missing inline directives to stubs] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Reviewed-by: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> 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>
2009-01-08 10:07:48 +08:00
{
}
mm: memcontrol: rewrite uncharge API The memcg uncharging code that is involved towards the end of a page's lifetime - truncation, reclaim, swapout, migration - is impressively complicated and fragile. Because anonymous and file pages were always charged before they had their page->mapping established, uncharges had to happen when the page type could still be known from the context; as in unmap for anonymous, page cache removal for file and shmem pages, and swap cache truncation for swap pages. However, these operations happen well before the page is actually freed, and so a lot of synchronization is necessary: - Charging, uncharging, page migration, and charge migration all need to take a per-page bit spinlock as they could race with uncharging. - Swap cache truncation happens during both swap-in and swap-out, and possibly repeatedly before the page is actually freed. This means that the memcg swapout code is called from many contexts that make no sense and it has to figure out the direction from page state to make sure memory and memory+swap are always correctly charged. - On page migration, the old page might be unmapped but then reused, so memcg code has to prevent untimely uncharging in that case. Because this code - which should be a simple charge transfer - is so special-cased, it is not reusable for replace_page_cache(). But now that charged pages always have a page->mapping, introduce mem_cgroup_uncharge(), which is called after the final put_page(), when we know for sure that nobody is looking at the page anymore. For page migration, introduce mem_cgroup_migrate(), which is called after the migration is successful and the new page is fully rmapped. Because the old page is no longer uncharged after migration, prevent double charges by decoupling the page's memcg association (PCG_USED and pc->mem_cgroup) from the page holding an actual charge. The new bits PCG_MEM and PCG_MEMSW represent the respective charges and are transferred to the new page during migration. mem_cgroup_migrate() is suitable for replace_page_cache() as well, which gets rid of mem_cgroup_replace_page_cache(). However, care needs to be taken because both the source and the target page can already be charged and on the LRU when fuse is splicing: grab the page lock on the charge moving side to prevent changing pc->mem_cgroup of a page under migration. Also, the lruvecs of both pages change as we uncharge the old and charge the new during migration, and putback may race with us, so grab the lru lock and isolate the pages iff on LRU to prevent races and ensure the pages are on the right lruvec afterward. Swap accounting is massively simplified: because the page is no longer uncharged as early as swap cache deletion, a new mem_cgroup_swapout() can transfer the page's memory+swap charge (PCG_MEMSW) to the swap entry before the final put_page() in page reclaim. Finally, page_cgroup changes are now protected by whatever protection the page itself offers: anonymous pages are charged under the page table lock, whereas page cache insertions, swapin, and migration hold the page lock. Uncharging happens under full exclusion with no outstanding references. Charging and uncharging also ensure that the page is off-LRU, which serializes against charge migration. Remove the very costly page_cgroup lock and set pc->flags non-atomically. [mhocko@suse.cz: mem_cgroup_charge_statistics needs preempt_disable] [vdavydov@parallels.com: fix flags definition] Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Cc: Hugh Dickins <hughd@google.com> Cc: Tejun Heo <tj@kernel.org> Cc: Vladimir Davydov <vdavydov@parallels.com> Tested-by: Jet Chen <jet.chen@intel.com> Acked-by: Michal Hocko <mhocko@suse.cz> Tested-by: Felipe Balbi <balbi@ti.com> Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2014-08-09 05:19:22 +08:00
static inline void mem_cgroup_uncharge(struct page *page)
memcg: coalesce uncharge during unmap/truncate In massive parallel enviroment, res_counter can be a performance bottleneck. One strong techinque to reduce lock contention is reducing calls by coalescing some amount of calls into one. Considering charge/uncharge chatacteristic, - charge is done one by one via demand-paging. - uncharge is done by - in chunk at munmap, truncate, exit, execve... - one by one via vmscan/paging. It seems we have a chance to coalesce uncharges for improving scalability at unmap/truncation. This patch is a for coalescing uncharge. For avoiding scattering memcg's structure to functions under /mm, this patch adds memcg batch uncharge information to the task. A reason for per-task batching is for making use of caller's context information. We do batched uncharge (deleyed uncharge) when truncation/unmap occurs but do direct uncharge when uncharge is called by memory reclaim (vmscan.c). The degree of coalescing depends on callers - at invalidate/trucate... pagevec size - at unmap ....ZAP_BLOCK_SIZE (memory itself will be freed in this degree.) Then, we'll not coalescing too much. On x86-64 8cpu server, I tested overheads of memcg at page fault by running a program which does map/fault/unmap in a loop. Running a task per a cpu by taskset and see sum of the number of page faults in 60secs. [without memcg config] 40156968 page-faults # 0.085 M/sec ( +- 0.046% ) 27.67 cache-miss/faults [root cgroup] 36659599 page-faults # 0.077 M/sec ( +- 0.247% ) 31.58 miss/faults [in a child cgroup] 18444157 page-faults # 0.039 M/sec ( +- 0.133% ) 69.96 miss/faults [child with this patch] 27133719 page-faults # 0.057 M/sec ( +- 0.155% ) 47.16 miss/faults We can see some amounts of improvement. (root cgroup doesn't affected by this patch) Another patch for "charge" will follow this and above will be improved more. Changelog(since 2009/10/02): - renamed filed of memcg_batch (as pages to bytes, memsw to memsw_bytes) - some clean up and commentary/description updates. - added initialize code to copy_process(). (possible bug fix) Changelog(old): - fixed !CONFIG_MEM_CGROUP case. - rebased onto the latest mmotm + softlimit fix patches. - unified patch for callers - added commetns. - make ->do_batch as bool. - removed css_get() at el. We don't need it. Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-12-16 08:47:03 +08:00
{
}
static inline void mem_cgroup_uncharge_list(struct list_head *page_list)
Memory controller: memory accounting Add the accounting hooks. The accounting is carried out for RSS and Page Cache (unmapped) pages. There is now a common limit and accounting for both. The RSS accounting is accounted at page_add_*_rmap() and page_remove_rmap() time. Page cache is accounted at add_to_page_cache(), __delete_from_page_cache(). Swap cache is also accounted for. Each page's page_cgroup is protected with the last bit of the page_cgroup pointer, this makes handling of race conditions involving simultaneous mappings of a page easier. A reference count is kept in the page_cgroup to deal with cases where a page might be unmapped from the RSS of all tasks, but still lives in the page cache. Credits go to Vaidyanathan Srinivasan for helping with reference counting work of the page cgroup. Almost all of the page cache accounting code has help from Vaidyanathan Srinivasan. [hugh@veritas.com: fix swapoff breakage] [akpm@linux-foundation.org: fix locking] Signed-off-by: Vaidyanathan Srinivasan <svaidy@linux.vnet.ibm.com> Signed-off-by: Balbir Singh <balbir@linux.vnet.ibm.com> Cc: Pavel Emelianov <xemul@openvz.org> Cc: Paul Menage <menage@google.com> Cc: Peter Zijlstra <a.p.zijlstra@chello.nl> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Nick Piggin <nickpiggin@yahoo.com.au> Cc: Kirill Korotaev <dev@sw.ru> Cc: Herbert Poetzl <herbert@13thfloor.at> Cc: David Rientjes <rientjes@google.com> Cc: <Valdis.Kletnieks@vt.edu> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-02-07 16:13:53 +08:00
{
}
static inline void mem_cgroup_migrate(struct page *old, struct page *new)
memcg: remove refcnt from page_cgroup memcg: performance improvements Patch Description 1/5 ... remove refcnt fron page_cgroup patch (shmem handling is fixed) 2/5 ... swapcache handling patch 3/5 ... add helper function for shmem's memory reclaim patch 4/5 ... optimize by likely/unlikely ppatch 5/5 ... remove redundunt check patch (shmem handling is fixed.) Unix bench result. == 2.6.26-rc2-mm1 + memory resource controller Execl Throughput 2915.4 lps (29.6 secs, 3 samples) C Compiler Throughput 1019.3 lpm (60.0 secs, 3 samples) Shell Scripts (1 concurrent) 5796.0 lpm (60.0 secs, 3 samples) Shell Scripts (8 concurrent) 1097.7 lpm (60.0 secs, 3 samples) Shell Scripts (16 concurrent) 565.3 lpm (60.0 secs, 3 samples) File Read 1024 bufsize 2000 maxblocks 1022128.0 KBps (30.0 secs, 3 samples) File Write 1024 bufsize 2000 maxblocks 544057.0 KBps (30.0 secs, 3 samples) File Copy 1024 bufsize 2000 maxblocks 346481.0 KBps (30.0 secs, 3 samples) File Read 256 bufsize 500 maxblocks 319325.0 KBps (30.0 secs, 3 samples) File Write 256 bufsize 500 maxblocks 148788.0 KBps (30.0 secs, 3 samples) File Copy 256 bufsize 500 maxblocks 99051.0 KBps (30.0 secs, 3 samples) File Read 4096 bufsize 8000 maxblocks 2058917.0 KBps (30.0 secs, 3 samples) File Write 4096 bufsize 8000 maxblocks 1606109.0 KBps (30.0 secs, 3 samples) File Copy 4096 bufsize 8000 maxblocks 854789.0 KBps (30.0 secs, 3 samples) Dc: sqrt(2) to 99 decimal places 126145.2 lpm (30.0 secs, 3 samples) INDEX VALUES TEST BASELINE RESULT INDEX Execl Throughput 43.0 2915.4 678.0 File Copy 1024 bufsize 2000 maxblocks 3960.0 346481.0 875.0 File Copy 256 bufsize 500 maxblocks 1655.0 99051.0 598.5 File Copy 4096 bufsize 8000 maxblocks 5800.0 854789.0 1473.8 Shell Scripts (8 concurrent) 6.0 1097.7 1829.5 ========= FINAL SCORE 991.3 == 2.6.26-rc2-mm1 + this set == Execl Throughput 3012.9 lps (29.9 secs, 3 samples) C Compiler Throughput 981.0 lpm (60.0 secs, 3 samples) Shell Scripts (1 concurrent) 5872.0 lpm (60.0 secs, 3 samples) Shell Scripts (8 concurrent) 1120.3 lpm (60.0 secs, 3 samples) Shell Scripts (16 concurrent) 578.0 lpm (60.0 secs, 3 samples) File Read 1024 bufsize 2000 maxblocks 1003993.0 KBps (30.0 secs, 3 samples) File Write 1024 bufsize 2000 maxblocks 550452.0 KBps (30.0 secs, 3 samples) File Copy 1024 bufsize 2000 maxblocks 347159.0 KBps (30.0 secs, 3 samples) File Read 256 bufsize 500 maxblocks 314644.0 KBps (30.0 secs, 3 samples) File Write 256 bufsize 500 maxblocks 151852.0 KBps (30.0 secs, 3 samples) File Copy 256 bufsize 500 maxblocks 101000.0 KBps (30.0 secs, 3 samples) File Read 4096 bufsize 8000 maxblocks 2033256.0 KBps (30.0 secs, 3 samples) File Write 4096 bufsize 8000 maxblocks 1611814.0 KBps (30.0 secs, 3 samples) File Copy 4096 bufsize 8000 maxblocks 847979.0 KBps (30.0 secs, 3 samples) Dc: sqrt(2) to 99 decimal places 128148.7 lpm (30.0 secs, 3 samples) INDEX VALUES TEST BASELINE RESULT INDEX Execl Throughput 43.0 3012.9 700.7 File Copy 1024 bufsize 2000 maxblocks 3960.0 347159.0 876.7 File Copy 256 bufsize 500 maxblocks 1655.0 101000.0 610.3 File Copy 4096 bufsize 8000 maxblocks 5800.0 847979.0 1462.0 Shell Scripts (8 concurrent) 6.0 1120.3 1867.2 ========= FINAL SCORE 1004.6 This patch: Remove refcnt from page_cgroup(). After this, * A page is charged only when !page_mapped() && no page_cgroup is assigned. * Anon page is newly mapped. * File page is added to mapping->tree. * A page is uncharged only when * Anon page is fully unmapped. * File page is removed from LRU. There is no change in behavior from user's view. This patch also removes unnecessary calls in rmap.c which was used only for refcnt mangement. [akpm@linux-foundation.org: fix warning] [hugh@veritas.com: fix shmem_unuse_inode charging] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: "Eric W. Biederman" <ebiederm@xmission.com> Cc: Pavel Emelyanov <xemul@openvz.org> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Hugh Dickins <hugh@veritas.com> Cc: YAMAMOTO Takashi <yamamoto@valinux.co.jp> Cc: Paul Menage <menage@google.com> Cc: David Rientjes <rientjes@google.com> Signed-off-by: Hugh Dickins <hugh@veritas.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2008-07-25 16:47:14 +08:00
{
}
static inline struct lruvec *mem_cgroup_lruvec(struct pglist_data *pgdat,
struct mem_cgroup *memcg)
memcg: synchronized LRU A big patch for changing memcg's LRU semantics. Now, - page_cgroup is linked to mem_cgroup's its own LRU (per zone). - LRU of page_cgroup is not synchronous with global LRU. - page and page_cgroup is one-to-one and statically allocated. - To find page_cgroup is on what LRU, you have to check pc->mem_cgroup as - lru = page_cgroup_zoneinfo(pc, nid_of_pc, zid_of_pc); - SwapCache is handled. And, when we handle LRU list of page_cgroup, we do following. pc = lookup_page_cgroup(page); lock_page_cgroup(pc); .....................(1) mz = page_cgroup_zoneinfo(pc); spin_lock(&mz->lru_lock); .....add to LRU spin_unlock(&mz->lru_lock); unlock_page_cgroup(pc); But (1) is spin_lock and we have to be afraid of dead-lock with zone->lru_lock. So, trylock() is used at (1), now. Without (1), we can't trust "mz" is correct. This is a trial to remove this dirty nesting of locks. This patch changes mz->lru_lock to be zone->lru_lock. Then, above sequence will be written as spin_lock(&zone->lru_lock); # in vmscan.c or swap.c via global LRU mem_cgroup_add/remove/etc_lru() { pc = lookup_page_cgroup(page); mz = page_cgroup_zoneinfo(pc); if (PageCgroupUsed(pc)) { ....add to LRU } spin_lock(&zone->lru_lock); # in vmscan.c or swap.c via global LRU This is much simpler. (*) We're safe even if we don't take lock_page_cgroup(pc). Because.. 1. When pc->mem_cgroup can be modified. - at charge. - at account_move(). 2. at charge the PCG_USED bit is not set before pc->mem_cgroup is fixed. 3. at account_move() the page is isolated and not on LRU. Pros. - easy for maintenance. - memcg can make use of laziness of pagevec. - we don't have to duplicated LRU/Active/Unevictable bit in page_cgroup. - LRU status of memcg will be synchronized with global LRU's one. - # of locks are reduced. - account_move() is simplified very much. Cons. - may increase cost of LRU rotation. (no impact if memcg is not configured.) Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Pavel Emelyanov <xemul@openvz.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-01-08 10:08:01 +08:00
{
return node_lruvec(pgdat);
memcg: synchronized LRU A big patch for changing memcg's LRU semantics. Now, - page_cgroup is linked to mem_cgroup's its own LRU (per zone). - LRU of page_cgroup is not synchronous with global LRU. - page and page_cgroup is one-to-one and statically allocated. - To find page_cgroup is on what LRU, you have to check pc->mem_cgroup as - lru = page_cgroup_zoneinfo(pc, nid_of_pc, zid_of_pc); - SwapCache is handled. And, when we handle LRU list of page_cgroup, we do following. pc = lookup_page_cgroup(page); lock_page_cgroup(pc); .....................(1) mz = page_cgroup_zoneinfo(pc); spin_lock(&mz->lru_lock); .....add to LRU spin_unlock(&mz->lru_lock); unlock_page_cgroup(pc); But (1) is spin_lock and we have to be afraid of dead-lock with zone->lru_lock. So, trylock() is used at (1), now. Without (1), we can't trust "mz" is correct. This is a trial to remove this dirty nesting of locks. This patch changes mz->lru_lock to be zone->lru_lock. Then, above sequence will be written as spin_lock(&zone->lru_lock); # in vmscan.c or swap.c via global LRU mem_cgroup_add/remove/etc_lru() { pc = lookup_page_cgroup(page); mz = page_cgroup_zoneinfo(pc); if (PageCgroupUsed(pc)) { ....add to LRU } spin_lock(&zone->lru_lock); # in vmscan.c or swap.c via global LRU This is much simpler. (*) We're safe even if we don't take lock_page_cgroup(pc). Because.. 1. When pc->mem_cgroup can be modified. - at charge. - at account_move(). 2. at charge the PCG_USED bit is not set before pc->mem_cgroup is fixed. 3. at account_move() the page is isolated and not on LRU. Pros. - easy for maintenance. - memcg can make use of laziness of pagevec. - we don't have to duplicated LRU/Active/Unevictable bit in page_cgroup. - LRU status of memcg will be synchronized with global LRU's one. - # of locks are reduced. - account_move() is simplified very much. Cons. - may increase cost of LRU rotation. (no impact if memcg is not configured.) Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Li Zefan <lizf@cn.fujitsu.com> Cc: Balbir Singh <balbir@in.ibm.com> Cc: Pavel Emelyanov <xemul@openvz.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2009-01-08 10:08:01 +08:00
}
static inline struct lruvec *mem_cgroup_page_lruvec(struct page *page,
mm, vmscan: move LRU lists to node This moves the LRU lists from the zone to the node and related data such as counters, tracing, congestion tracking and writeback tracking. Unfortunately, due to reclaim and compaction retry logic, it is necessary to account for the number of LRU pages on both zone and node logic. Most reclaim logic is based on the node counters but the retry logic uses the zone counters which do not distinguish inactive and active sizes. It would be possible to leave the LRU counters on a per-zone basis but it's a heavier calculation across multiple cache lines that is much more frequent than the retry checks. Other than the LRU counters, this is mostly a mechanical patch but note that it introduces a number of anomalies. For example, the scans are per-zone but using per-node counters. We also mark a node as congested when a zone is congested. This causes weird problems that are fixed later but is easier to review. In the event that there is excessive overhead on 32-bit systems due to the nodes being on LRU then there are two potential solutions 1. Long-term isolation of highmem pages when reclaim is lowmem When pages are skipped, they are immediately added back onto the LRU list. If lowmem reclaim persisted for long periods of time, the same highmem pages get continually scanned. The idea would be that lowmem keeps those pages on a separate list until a reclaim for highmem pages arrives that splices the highmem pages back onto the LRU. It potentially could be implemented similar to the UNEVICTABLE list. That would reduce the skip rate with the potential corner case is that highmem pages have to be scanned and reclaimed to free lowmem slab pages. 2. Linear scan lowmem pages if the initial LRU shrink fails This will break LRU ordering but may be preferable and faster during memory pressure than skipping LRU pages. Link: http://lkml.kernel.org/r/1467970510-21195-4-git-send-email-mgorman@techsingularity.net Signed-off-by: Mel Gorman <mgorman@techsingularity.net> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vlastimil Babka <vbabka@suse.cz> Cc: Hillf Danton <hillf.zj@alibaba-inc.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@surriel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:31 +08:00
struct pglist_data *pgdat)
{
mm, vmscan: move LRU lists to node This moves the LRU lists from the zone to the node and related data such as counters, tracing, congestion tracking and writeback tracking. Unfortunately, due to reclaim and compaction retry logic, it is necessary to account for the number of LRU pages on both zone and node logic. Most reclaim logic is based on the node counters but the retry logic uses the zone counters which do not distinguish inactive and active sizes. It would be possible to leave the LRU counters on a per-zone basis but it's a heavier calculation across multiple cache lines that is much more frequent than the retry checks. Other than the LRU counters, this is mostly a mechanical patch but note that it introduces a number of anomalies. For example, the scans are per-zone but using per-node counters. We also mark a node as congested when a zone is congested. This causes weird problems that are fixed later but is easier to review. In the event that there is excessive overhead on 32-bit systems due to the nodes being on LRU then there are two potential solutions 1. Long-term isolation of highmem pages when reclaim is lowmem When pages are skipped, they are immediately added back onto the LRU list. If lowmem reclaim persisted for long periods of time, the same highmem pages get continually scanned. The idea would be that lowmem keeps those pages on a separate list until a reclaim for highmem pages arrives that splices the highmem pages back onto the LRU. It potentially could be implemented similar to the UNEVICTABLE list. That would reduce the skip rate with the potential corner case is that highmem pages have to be scanned and reclaimed to free lowmem slab pages. 2. Linear scan lowmem pages if the initial LRU shrink fails This will break LRU ordering but may be preferable and faster during memory pressure than skipping LRU pages. Link: http://lkml.kernel.org/r/1467970510-21195-4-git-send-email-mgorman@techsingularity.net Signed-off-by: Mel Gorman <mgorman@techsingularity.net> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Acked-by: Vlastimil Babka <vbabka@suse.cz> Cc: Hillf Danton <hillf.zj@alibaba-inc.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Minchan Kim <minchan@kernel.org> Cc: Rik van Riel <riel@surriel.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-07-29 06:45:31 +08:00
return &pgdat->lruvec;
}
static inline bool mm_match_cgroup(struct mm_struct *mm,
struct mem_cgroup *memcg)
{
return true;
}
static inline bool task_in_mem_cgroup(struct task_struct *task,
const struct mem_cgroup *memcg)
{
return true;
}
static inline struct mem_cgroup *
mem_cgroup_iter(struct mem_cgroup *root,
struct mem_cgroup *prev,
struct mem_cgroup_reclaim_cookie *reclaim)
{
return NULL;
}
static inline void mem_cgroup_iter_break(struct mem_cgroup *root,
struct mem_cgroup *prev)
{
}
static inline unsigned short mem_cgroup_id(struct mem_cgroup *memcg)
{
return 0;
}
static inline struct mem_cgroup *mem_cgroup_from_id(unsigned short id)
{
WARN_ON_ONCE(id);
/* XXX: This should always return root_mem_cgroup */
return NULL;
}
static inline bool mem_cgroup_online(struct mem_cgroup *memcg)
{
return true;
}
static inline unsigned long
mem_cgroup_get_lru_size(struct lruvec *lruvec, enum lru_list lru)
{
return 0;
}
static inline unsigned long
mem_cgroup_node_nr_lru_pages(struct mem_cgroup *memcg,
int nid, unsigned int lru_mask)
{
return 0;
}
static inline void
mem_cgroup_print_oom_info(struct mem_cgroup *memcg, struct task_struct *p)
{
}
static inline void lock_page_memcg(struct page *page)
memcg: use new logic for page stat accounting Now, page-stat-per-memcg is recorded into per page_cgroup flag by duplicating page's status into the flag. The reason is that memcg has a feature to move a page from a group to another group and we have race between "move" and "page stat accounting", Under current logic, assume CPU-A and CPU-B. CPU-A does "move" and CPU-B does "page stat accounting". When CPU-A goes 1st, CPU-A CPU-B update "struct page" info. move_lock_mem_cgroup(memcg) see pc->flags copy page stat to new group overwrite pc->mem_cgroup. move_unlock_mem_cgroup(memcg) move_lock_mem_cgroup(mem) set pc->flags update page stat accounting move_unlock_mem_cgroup(mem) stat accounting is guarded by move_lock_mem_cgroup() and "move" logic (CPU-A) doesn't see changes in "struct page" information. But it's costly to have the same information both in 'struct page' and 'struct page_cgroup'. And, there is a potential problem. For example, assume we have PG_dirty accounting in memcg. PG_..is a flag for struct page. PCG_ is a flag for struct page_cgroup. (This is just an example. The same problem can be found in any kind of page stat accounting.) CPU-A CPU-B TestSet PG_dirty (delay) TestClear PG_dirty if (TestClear(PCG_dirty)) memcg->nr_dirty-- if (TestSet(PCG_dirty)) memcg->nr_dirty++ Here, memcg->nr_dirty = +1, this is wrong. This race was reported by Greg Thelen <gthelen@google.com>. Now, only FILE_MAPPED is supported but fortunately, it's serialized by page table lock and this is not real bug, _now_, If this potential problem is caused by having duplicated information in struct page and struct page_cgroup, we may be able to fix this by using original 'struct page' information. But we'll have a problem in "move account" Assume we use only PG_dirty. CPU-A CPU-B TestSet PG_dirty (delay) move_lock_mem_cgroup() if (PageDirty(page)) new_memcg->nr_dirty++ pc->mem_cgroup = new_memcg; move_unlock_mem_cgroup() move_lock_mem_cgroup() memcg = pc->mem_cgroup new_memcg->nr_dirty++ accounting information may be double-counted. This was original reason to have PCG_xxx flags but it seems PCG_xxx has another problem. I think we need a bigger lock as move_lock_mem_cgroup(page) TestSetPageDirty(page) update page stats (without any checks) move_unlock_mem_cgroup(page) This fixes both of problems and we don't have to duplicate page flag into page_cgroup. Please note: move_lock_mem_cgroup() is held only when there are possibility of "account move" under the system. So, in most path, status update will go without atomic locks. This patch introduces mem_cgroup_begin_update_page_stat() and mem_cgroup_end_update_page_stat() both should be called at modifying 'struct page' information if memcg takes care of it. as mem_cgroup_begin_update_page_stat() modify page information mem_cgroup_update_page_stat() => never check any 'struct page' info, just update counters. mem_cgroup_end_update_page_stat(). This patch is slow because we need to call begin_update_page_stat()/ end_update_page_stat() regardless of accounted will be changed or not. A following patch adds an easy optimization and reduces the cost. [akpm@linux-foundation.org: s/lock/locked/] [hughd@google.com: fix deadlock by avoiding stat lock when anon] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Greg Thelen <gthelen@google.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Ying Han <yinghan@google.com> Signed-off-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-03-22 07:34:25 +08:00
{
}
static inline void unlock_page_memcg(struct page *page)
memcg: use new logic for page stat accounting Now, page-stat-per-memcg is recorded into per page_cgroup flag by duplicating page's status into the flag. The reason is that memcg has a feature to move a page from a group to another group and we have race between "move" and "page stat accounting", Under current logic, assume CPU-A and CPU-B. CPU-A does "move" and CPU-B does "page stat accounting". When CPU-A goes 1st, CPU-A CPU-B update "struct page" info. move_lock_mem_cgroup(memcg) see pc->flags copy page stat to new group overwrite pc->mem_cgroup. move_unlock_mem_cgroup(memcg) move_lock_mem_cgroup(mem) set pc->flags update page stat accounting move_unlock_mem_cgroup(mem) stat accounting is guarded by move_lock_mem_cgroup() and "move" logic (CPU-A) doesn't see changes in "struct page" information. But it's costly to have the same information both in 'struct page' and 'struct page_cgroup'. And, there is a potential problem. For example, assume we have PG_dirty accounting in memcg. PG_..is a flag for struct page. PCG_ is a flag for struct page_cgroup. (This is just an example. The same problem can be found in any kind of page stat accounting.) CPU-A CPU-B TestSet PG_dirty (delay) TestClear PG_dirty if (TestClear(PCG_dirty)) memcg->nr_dirty-- if (TestSet(PCG_dirty)) memcg->nr_dirty++ Here, memcg->nr_dirty = +1, this is wrong. This race was reported by Greg Thelen <gthelen@google.com>. Now, only FILE_MAPPED is supported but fortunately, it's serialized by page table lock and this is not real bug, _now_, If this potential problem is caused by having duplicated information in struct page and struct page_cgroup, we may be able to fix this by using original 'struct page' information. But we'll have a problem in "move account" Assume we use only PG_dirty. CPU-A CPU-B TestSet PG_dirty (delay) move_lock_mem_cgroup() if (PageDirty(page)) new_memcg->nr_dirty++ pc->mem_cgroup = new_memcg; move_unlock_mem_cgroup() move_lock_mem_cgroup() memcg = pc->mem_cgroup new_memcg->nr_dirty++ accounting information may be double-counted. This was original reason to have PCG_xxx flags but it seems PCG_xxx has another problem. I think we need a bigger lock as move_lock_mem_cgroup(page) TestSetPageDirty(page) update page stats (without any checks) move_unlock_mem_cgroup(page) This fixes both of problems and we don't have to duplicate page flag into page_cgroup. Please note: move_lock_mem_cgroup() is held only when there are possibility of "account move" under the system. So, in most path, status update will go without atomic locks. This patch introduces mem_cgroup_begin_update_page_stat() and mem_cgroup_end_update_page_stat() both should be called at modifying 'struct page' information if memcg takes care of it. as mem_cgroup_begin_update_page_stat() modify page information mem_cgroup_update_page_stat() => never check any 'struct page' info, just update counters. mem_cgroup_end_update_page_stat(). This patch is slow because we need to call begin_update_page_stat()/ end_update_page_stat() regardless of accounted will be changed or not. A following patch adds an easy optimization and reduces the cost. [akpm@linux-foundation.org: s/lock/locked/] [hughd@google.com: fix deadlock by avoiding stat lock when anon] Signed-off-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Greg Thelen <gthelen@google.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Cc: Ying Han <yinghan@google.com> Signed-off-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-03-22 07:34:25 +08:00
{
}
memcg: punt high overage reclaim to return-to-userland path Currently, try_charge() tries to reclaim memory synchronously when the high limit is breached; however, if the allocation doesn't have __GFP_WAIT, synchronous reclaim is skipped. If a process performs only speculative allocations, it can blow way past the high limit. This is actually easily reproducible by simply doing "find /". slab/slub allocator tries speculative allocations first, so as long as there's memory which can be consumed without blocking, it can keep allocating memory regardless of the high limit. This patch makes try_charge() always punt the over-high reclaim to the return-to-userland path. If try_charge() detects that high limit is breached, it adds the overage to current->memcg_nr_pages_over_high and schedules execution of mem_cgroup_handle_over_high() which performs synchronous reclaim from the return-to-userland path. As long as kernel doesn't have a run-away allocation spree, this should provide enough protection while making kmemcg behave more consistently. It also has the following benefits. - All over-high reclaims can use GFP_KERNEL regardless of the specific gfp mask in use, e.g. GFP_NOFS, when the limit was breached. - It copes with prio inversion. Previously, a low-prio task with small memory.high might perform over-high reclaim with a bunch of locks held. If a higher prio task needed any of these locks, it would have to wait until the low prio task finished reclaim and released the locks. By handing over-high reclaim to the task exit path this issue can be avoided. Signed-off-by: Tejun Heo <tj@kernel.org> Acked-by: Michal Hocko <mhocko@kernel.org> Reviewed-by: Vladimir Davydov <vdavydov@parallels.com> Acked-by: Johannes Weiner <hannes@cmpxchg.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-11-06 10:46:11 +08:00
static inline void mem_cgroup_handle_over_high(void)
{
}
static inline void mem_cgroup_oom_enable(void)
{
}
static inline void mem_cgroup_oom_disable(void)
{
}
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
static inline bool task_in_memcg_oom(struct task_struct *p)
{
return false;
}
static inline bool mem_cgroup_oom_synchronize(bool wait)
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
{
return false;
}
static inline void mem_cgroup_inc_page_stat(struct page *page,
enum mem_cgroup_stat_index idx)
{
}
static inline void mem_cgroup_dec_page_stat(struct page *page,
enum mem_cgroup_stat_index idx)
{
}
2009-09-24 06:56:39 +08:00
static inline
unsigned long mem_cgroup_soft_limit_reclaim(pg_data_t *pgdat, int order,
gfp_t gfp_mask,
unsigned long *total_scanned)
2009-09-24 06:56:39 +08:00
{
return 0;
2009-09-24 06:56:39 +08:00
}
static inline void mem_cgroup_split_huge_fixup(struct page *head)
{
}
memcg: add the pagefault count into memcg stats Two new stats in per-memcg memory.stat which tracks the number of page faults and number of major page faults. "pgfault" "pgmajfault" They are different from "pgpgin"/"pgpgout" stat which count number of pages charged/discharged to the cgroup and have no meaning of reading/ writing page to disk. It is valuable to track the two stats for both measuring application's performance as well as the efficiency of the kernel page reclaim path. Counting pagefaults per process is useful, but we also need the aggregated value since processes are monitored and controlled in cgroup basis in memcg. Functional test: check the total number of pgfault/pgmajfault of all memcgs and compare with global vmstat value: $ cat /proc/vmstat | grep fault pgfault 1070751 pgmajfault 553 $ cat /dev/cgroup/memory.stat | grep fault pgfault 1071138 pgmajfault 553 total_pgfault 1071142 total_pgmajfault 553 $ cat /dev/cgroup/A/memory.stat | grep fault pgfault 199 pgmajfault 0 total_pgfault 199 total_pgmajfault 0 Performance test: run page fault test(pft) wit 16 thread on faulting in 15G anon pages in 16G container. There is no regression noticed on the "flt/cpu/s" Sample output from pft: TAG pft:anon-sys-default: Gb Thr CLine User System Wall flt/cpu/s fault/wsec 15 16 1 0.67s 233.41s 14.76s 16798.546 266356.260 +-------------------------------------------------------------------------+ N Min Max Median Avg Stddev x 10 16682.962 17344.027 16913.524 16928.812 166.5362 + 10 16695.568 16923.896 16820.604 16824.652 84.816568 No difference proven at 95.0% confidence [akpm@linux-foundation.org: fix build] [hughd@google.com: shmem fix] Signed-off-by: Ying Han <yinghan@google.com> Acked-by: KAMEZAWA Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: KOSAKI Motohiro <kosaki.motohiro@jp.fujitsu.com> Reviewed-by: Minchan Kim <minchan.kim@gmail.com> Cc: Daisuke Nishimura <nishimura@mxp.nes.nec.co.jp> Acked-by: Balbir Singh <balbir@linux.vnet.ibm.com> Signed-off-by: Hugh Dickins <hughd@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2011-05-27 07:25:38 +08:00
static inline
void mem_cgroup_count_vm_event(struct mm_struct *mm, enum vm_event_item idx)
{
}
#endif /* CONFIG_MEMCG */
writeback: make backing_dev_info host cgroup-specific bdi_writebacks For the planned cgroup writeback support, on each bdi (backing_dev_info), each memcg will be served by a separate wb (bdi_writeback). This patch updates bdi so that a bdi can host multiple wbs (bdi_writebacks). On the default hierarchy, blkcg implicitly enables memcg. This allows using memcg's page ownership for attributing writeback IOs, and every memcg - blkcg combination can be served by its own wb by assigning a dedicated wb to each memcg. This means that there may be multiple wb's of a bdi mapped to the same blkcg. As congested state is per blkcg - bdi combination, those wb's should share the same congested state. This is achieved by tracking congested state via bdi_writeback_congested structs which are keyed by blkcg. bdi->wb remains unchanged and will keep serving the root cgroup. cgwb's (cgroup wb's) for non-root cgroups are created on-demand or looked up while dirtying an inode according to the memcg of the page being dirtied or current task. Each cgwb is indexed on bdi->cgwb_tree by its memcg id. Once an inode is associated with its wb, it can be retrieved using inode_to_wb(). Currently, none of the filesystems has FS_CGROUP_WRITEBACK and all pages will keep being associated with bdi->wb. v3: inode_attach_wb() in account_page_dirtied() moved inside mapping_cap_account_dirty() block where it's known to be !NULL. Also, an unnecessary NULL check before kfree() removed. Both detected by the kbuild bot. v2: Updated so that wb association is per inode and wb is per memcg rather than blkcg. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: kbuild test robot <fengguang.wu@intel.com> Cc: Dan Carpenter <dan.carpenter@oracle.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Jan Kara <jack@suse.cz> Signed-off-by: Jens Axboe <axboe@fb.com>
2015-05-23 05:13:37 +08:00
#ifdef CONFIG_CGROUP_WRITEBACK
writeback: make backing_dev_info host cgroup-specific bdi_writebacks For the planned cgroup writeback support, on each bdi (backing_dev_info), each memcg will be served by a separate wb (bdi_writeback). This patch updates bdi so that a bdi can host multiple wbs (bdi_writebacks). On the default hierarchy, blkcg implicitly enables memcg. This allows using memcg's page ownership for attributing writeback IOs, and every memcg - blkcg combination can be served by its own wb by assigning a dedicated wb to each memcg. This means that there may be multiple wb's of a bdi mapped to the same blkcg. As congested state is per blkcg - bdi combination, those wb's should share the same congested state. This is achieved by tracking congested state via bdi_writeback_congested structs which are keyed by blkcg. bdi->wb remains unchanged and will keep serving the root cgroup. cgwb's (cgroup wb's) for non-root cgroups are created on-demand or looked up while dirtying an inode according to the memcg of the page being dirtied or current task. Each cgwb is indexed on bdi->cgwb_tree by its memcg id. Once an inode is associated with its wb, it can be retrieved using inode_to_wb(). Currently, none of the filesystems has FS_CGROUP_WRITEBACK and all pages will keep being associated with bdi->wb. v3: inode_attach_wb() in account_page_dirtied() moved inside mapping_cap_account_dirty() block where it's known to be !NULL. Also, an unnecessary NULL check before kfree() removed. Both detected by the kbuild bot. v2: Updated so that wb association is per inode and wb is per memcg rather than blkcg. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: kbuild test robot <fengguang.wu@intel.com> Cc: Dan Carpenter <dan.carpenter@oracle.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Jan Kara <jack@suse.cz> Signed-off-by: Jens Axboe <axboe@fb.com>
2015-05-23 05:13:37 +08:00
struct list_head *mem_cgroup_cgwb_list(struct mem_cgroup *memcg);
struct wb_domain *mem_cgroup_wb_domain(struct bdi_writeback *wb);
2015-09-30 01:04:26 +08:00
void mem_cgroup_wb_stats(struct bdi_writeback *wb, unsigned long *pfilepages,
unsigned long *pheadroom, unsigned long *pdirty,
unsigned long *pwriteback);
#else /* CONFIG_CGROUP_WRITEBACK */
static inline struct wb_domain *mem_cgroup_wb_domain(struct bdi_writeback *wb)
{
return NULL;
}
writeback: implement memcg writeback domain based throttling While cgroup writeback support now connects memcg and blkcg so that writeback IOs are properly attributed and controlled, the IO back pressure propagation mechanism implemented in balance_dirty_pages() and its subroutines wasn't aware of cgroup writeback. Processes belonging to a memcg may have access to only subset of total memory available in the system and not factoring this into dirty throttling rendered it completely ineffective for processes under memcg limits and memcg ended up building a separate ad-hoc degenerate mechanism directly into vmscan code to limit page dirtying. The previous patches updated balance_dirty_pages() and its subroutines so that they can deal with multiple wb_domain's (writeback domains) and defined per-memcg wb_domain. Processes belonging to a non-root memcg are bound to two wb_domains, global wb_domain and memcg wb_domain, and should be throttled according to IO pressures from both domains. This patch updates dirty throttling code so that it repeats similar calculations for the two domains - the differences between the two are few and minor - and applies the lower of the two sets of resulting constraints. wb_over_bg_thresh(), which controls when background writeback terminates, is also updated to consider both global and memcg wb_domains. It returns true if dirty is over bg_thresh for either domain. This makes the dirty throttling mechanism operational for memcg domains including writeback-bandwidth-proportional dirty page distribution inside them but the ad-hoc memcg throttling mechanism in vmscan is still in place. The next patch will rip it out. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: Jens Axboe <axboe@kernel.dk> Cc: Jan Kara <jack@suse.cz> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by: Jens Axboe <axboe@fb.com>
2015-05-23 06:23:35 +08:00
static inline void mem_cgroup_wb_stats(struct bdi_writeback *wb,
2015-09-30 01:04:26 +08:00
unsigned long *pfilepages,
unsigned long *pheadroom,
writeback: implement memcg writeback domain based throttling While cgroup writeback support now connects memcg and blkcg so that writeback IOs are properly attributed and controlled, the IO back pressure propagation mechanism implemented in balance_dirty_pages() and its subroutines wasn't aware of cgroup writeback. Processes belonging to a memcg may have access to only subset of total memory available in the system and not factoring this into dirty throttling rendered it completely ineffective for processes under memcg limits and memcg ended up building a separate ad-hoc degenerate mechanism directly into vmscan code to limit page dirtying. The previous patches updated balance_dirty_pages() and its subroutines so that they can deal with multiple wb_domain's (writeback domains) and defined per-memcg wb_domain. Processes belonging to a non-root memcg are bound to two wb_domains, global wb_domain and memcg wb_domain, and should be throttled according to IO pressures from both domains. This patch updates dirty throttling code so that it repeats similar calculations for the two domains - the differences between the two are few and minor - and applies the lower of the two sets of resulting constraints. wb_over_bg_thresh(), which controls when background writeback terminates, is also updated to consider both global and memcg wb_domains. It returns true if dirty is over bg_thresh for either domain. This makes the dirty throttling mechanism operational for memcg domains including writeback-bandwidth-proportional dirty page distribution inside them but the ad-hoc memcg throttling mechanism in vmscan is still in place. The next patch will rip it out. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: Jens Axboe <axboe@kernel.dk> Cc: Jan Kara <jack@suse.cz> Cc: Wu Fengguang <fengguang.wu@intel.com> Cc: Greg Thelen <gthelen@google.com> Signed-off-by: Jens Axboe <axboe@fb.com>
2015-05-23 06:23:35 +08:00
unsigned long *pdirty,
unsigned long *pwriteback)
{
}
#endif /* CONFIG_CGROUP_WRITEBACK */
writeback: make backing_dev_info host cgroup-specific bdi_writebacks For the planned cgroup writeback support, on each bdi (backing_dev_info), each memcg will be served by a separate wb (bdi_writeback). This patch updates bdi so that a bdi can host multiple wbs (bdi_writebacks). On the default hierarchy, blkcg implicitly enables memcg. This allows using memcg's page ownership for attributing writeback IOs, and every memcg - blkcg combination can be served by its own wb by assigning a dedicated wb to each memcg. This means that there may be multiple wb's of a bdi mapped to the same blkcg. As congested state is per blkcg - bdi combination, those wb's should share the same congested state. This is achieved by tracking congested state via bdi_writeback_congested structs which are keyed by blkcg. bdi->wb remains unchanged and will keep serving the root cgroup. cgwb's (cgroup wb's) for non-root cgroups are created on-demand or looked up while dirtying an inode according to the memcg of the page being dirtied or current task. Each cgwb is indexed on bdi->cgwb_tree by its memcg id. Once an inode is associated with its wb, it can be retrieved using inode_to_wb(). Currently, none of the filesystems has FS_CGROUP_WRITEBACK and all pages will keep being associated with bdi->wb. v3: inode_attach_wb() in account_page_dirtied() moved inside mapping_cap_account_dirty() block where it's known to be !NULL. Also, an unnecessary NULL check before kfree() removed. Both detected by the kbuild bot. v2: Updated so that wb association is per inode and wb is per memcg rather than blkcg. Signed-off-by: Tejun Heo <tj@kernel.org> Cc: kbuild test robot <fengguang.wu@intel.com> Cc: Dan Carpenter <dan.carpenter@oracle.com> Cc: Jens Axboe <axboe@kernel.dk> Cc: Jan Kara <jack@suse.cz> Signed-off-by: Jens Axboe <axboe@fb.com>
2015-05-23 05:13:37 +08:00
struct sock;
void sock_update_memcg(struct sock *sk);
void sock_release_memcg(struct sock *sk);
bool mem_cgroup_charge_skmem(struct mem_cgroup *memcg, unsigned int nr_pages);
void mem_cgroup_uncharge_skmem(struct mem_cgroup *memcg, unsigned int nr_pages);
#ifdef CONFIG_MEMCG
extern struct static_key_false memcg_sockets_enabled_key;
#define mem_cgroup_sockets_enabled static_branch_unlikely(&memcg_sockets_enabled_key)
static inline bool mem_cgroup_under_socket_pressure(struct mem_cgroup *memcg)
net: tcp_memcontrol: sanitize tcp memory accounting callbacks There won't be a tcp control soft limit, so integrating the memcg code into the global skmem limiting scheme complicates things unnecessarily. Replace this with simple and clear charge and uncharge calls--hidden behind a jump label--to account skb memory. Note that this is not purely aesthetic: as a result of shoehorning the per-memcg code into the same memory accounting functions that handle the global level, the old code would compare the per-memcg consumption against the smaller of the per-memcg limit and the global limit. This allowed the total consumption of multiple sockets to exceed the global limit, as long as the individual sockets stayed within bounds. After this change, the code will always compare the per-memcg consumption to the per-memcg limit, and the global consumption to the global limit, and thus close this loophole. Without a soft limit, the per-memcg memory pressure state in sockets is generally questionable. However, we did it until now, so we continue to enter it when the hard limit is hit, and packets are dropped, to let other sockets in the cgroup know that they shouldn't grow their transmit windows, either. However, keep it simple in the new callback model and leave memory pressure lazily when the next packet is accepted (as opposed to doing it synchroneously when packets are processed). When packets are dropped, network performance will already be in the toilet, so that should be a reasonable trade-off. As described above, consumption is now checked on the per-memcg level and the global level separately. Likewise, memory pressure states are maintained on both the per-memcg level and the global level, and a socket is considered under pressure when either level asserts as much. Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: David S. Miller <davem@davemloft.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-15 07:21:14 +08:00
{
if (!cgroup_subsys_on_dfl(memory_cgrp_subsys) && memcg->tcpmem_pressure)
return true;
do {
if (time_before(jiffies, memcg->socket_pressure))
return true;
} while ((memcg = parent_mem_cgroup(memcg)));
return false;
net: tcp_memcontrol: sanitize tcp memory accounting callbacks There won't be a tcp control soft limit, so integrating the memcg code into the global skmem limiting scheme complicates things unnecessarily. Replace this with simple and clear charge and uncharge calls--hidden behind a jump label--to account skb memory. Note that this is not purely aesthetic: as a result of shoehorning the per-memcg code into the same memory accounting functions that handle the global level, the old code would compare the per-memcg consumption against the smaller of the per-memcg limit and the global limit. This allowed the total consumption of multiple sockets to exceed the global limit, as long as the individual sockets stayed within bounds. After this change, the code will always compare the per-memcg consumption to the per-memcg limit, and the global consumption to the global limit, and thus close this loophole. Without a soft limit, the per-memcg memory pressure state in sockets is generally questionable. However, we did it until now, so we continue to enter it when the hard limit is hit, and packets are dropped, to let other sockets in the cgroup know that they shouldn't grow their transmit windows, either. However, keep it simple in the new callback model and leave memory pressure lazily when the next packet is accepted (as opposed to doing it synchroneously when packets are processed). When packets are dropped, network performance will already be in the toilet, so that should be a reasonable trade-off. As described above, consumption is now checked on the per-memcg level and the global level separately. Likewise, memory pressure states are maintained on both the per-memcg level and the global level, and a socket is considered under pressure when either level asserts as much. Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: David S. Miller <davem@davemloft.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-15 07:21:14 +08:00
}
#else
#define mem_cgroup_sockets_enabled 0
static inline bool mem_cgroup_under_socket_pressure(struct mem_cgroup *memcg)
net: tcp_memcontrol: sanitize tcp memory accounting callbacks There won't be a tcp control soft limit, so integrating the memcg code into the global skmem limiting scheme complicates things unnecessarily. Replace this with simple and clear charge and uncharge calls--hidden behind a jump label--to account skb memory. Note that this is not purely aesthetic: as a result of shoehorning the per-memcg code into the same memory accounting functions that handle the global level, the old code would compare the per-memcg consumption against the smaller of the per-memcg limit and the global limit. This allowed the total consumption of multiple sockets to exceed the global limit, as long as the individual sockets stayed within bounds. After this change, the code will always compare the per-memcg consumption to the per-memcg limit, and the global consumption to the global limit, and thus close this loophole. Without a soft limit, the per-memcg memory pressure state in sockets is generally questionable. However, we did it until now, so we continue to enter it when the hard limit is hit, and packets are dropped, to let other sockets in the cgroup know that they shouldn't grow their transmit windows, either. However, keep it simple in the new callback model and leave memory pressure lazily when the next packet is accepted (as opposed to doing it synchroneously when packets are processed). When packets are dropped, network performance will already be in the toilet, so that should be a reasonable trade-off. As described above, consumption is now checked on the per-memcg level and the global level separately. Likewise, memory pressure states are maintained on both the per-memcg level and the global level, and a socket is considered under pressure when either level asserts as much. Signed-off-by: Johannes Weiner <hannes@cmpxchg.org> Reviewed-by: Vladimir Davydov <vdavydov@virtuozzo.com> Acked-by: David S. Miller <davem@davemloft.net> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2016-01-15 07:21:14 +08:00
{
return false;
}
#endif
memcg: kmem controller infrastructure Introduce infrastructure for tracking kernel memory pages to a given memcg. This will happen whenever the caller includes the flag __GFP_KMEMCG flag, and the task belong to a memcg other than the root. In memcontrol.h those functions are wrapped in inline acessors. The idea is to later on, patch those with static branches, so we don't incur any overhead when no mem cgroups with limited kmem are being used. Users of this functionality shall interact with the memcg core code through the following functions: memcg_kmem_newpage_charge: will return true if the group can handle the allocation. At this point, struct page is not yet allocated. memcg_kmem_commit_charge: will either revert the charge, if struct page allocation failed, or embed memcg information into page_cgroup. memcg_kmem_uncharge_page: called at free time, will revert the charge. Signed-off-by: Glauber Costa <glommer@parallels.com> Acked-by: Michal Hocko <mhocko@suse.cz> Acked-by: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Tejun Heo <tj@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Frederic Weisbecker <fweisbec@redhat.com> Cc: Greg Thelen <gthelen@google.com> Cc: JoonSoo Kim <js1304@gmail.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik van Riel <riel@redhat.com> Cc: Suleiman Souhlal <suleiman@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-12-19 06:21:56 +08:00
struct kmem_cache *memcg_kmem_get_cache(struct kmem_cache *cachep);
void memcg_kmem_put_cache(struct kmem_cache *cachep);
int memcg_kmem_charge_memcg(struct page *page, gfp_t gfp, int order,
struct mem_cgroup *memcg);
int memcg_kmem_charge(struct page *page, gfp_t gfp, int order);
void memcg_kmem_uncharge(struct page *page, int order);
#if defined(CONFIG_MEMCG) && !defined(CONFIG_SLOB)
extern struct static_key_false memcg_kmem_enabled_key;
extern int memcg_nr_cache_ids;
void memcg_get_cache_ids(void);
void memcg_put_cache_ids(void);
/*
* Helper macro to loop through all memcg-specific caches. Callers must still
* check if the cache is valid (it is either valid or NULL).
* the slab_mutex must be held when looping through those caches
*/
#define for_each_memcg_cache_index(_idx) \
for ((_idx) = 0; (_idx) < memcg_nr_cache_ids; (_idx)++)
memcg: kmem controller infrastructure Introduce infrastructure for tracking kernel memory pages to a given memcg. This will happen whenever the caller includes the flag __GFP_KMEMCG flag, and the task belong to a memcg other than the root. In memcontrol.h those functions are wrapped in inline acessors. The idea is to later on, patch those with static branches, so we don't incur any overhead when no mem cgroups with limited kmem are being used. Users of this functionality shall interact with the memcg core code through the following functions: memcg_kmem_newpage_charge: will return true if the group can handle the allocation. At this point, struct page is not yet allocated. memcg_kmem_commit_charge: will either revert the charge, if struct page allocation failed, or embed memcg information into page_cgroup. memcg_kmem_uncharge_page: called at free time, will revert the charge. Signed-off-by: Glauber Costa <glommer@parallels.com> Acked-by: Michal Hocko <mhocko@suse.cz> Acked-by: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Tejun Heo <tj@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Frederic Weisbecker <fweisbec@redhat.com> Cc: Greg Thelen <gthelen@google.com> Cc: JoonSoo Kim <js1304@gmail.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik van Riel <riel@redhat.com> Cc: Suleiman Souhlal <suleiman@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-12-19 06:21:56 +08:00
static inline bool memcg_kmem_enabled(void)
{
return static_branch_unlikely(&memcg_kmem_enabled_key);
memcg: kmem controller infrastructure Introduce infrastructure for tracking kernel memory pages to a given memcg. This will happen whenever the caller includes the flag __GFP_KMEMCG flag, and the task belong to a memcg other than the root. In memcontrol.h those functions are wrapped in inline acessors. The idea is to later on, patch those with static branches, so we don't incur any overhead when no mem cgroups with limited kmem are being used. Users of this functionality shall interact with the memcg core code through the following functions: memcg_kmem_newpage_charge: will return true if the group can handle the allocation. At this point, struct page is not yet allocated. memcg_kmem_commit_charge: will either revert the charge, if struct page allocation failed, or embed memcg information into page_cgroup. memcg_kmem_uncharge_page: called at free time, will revert the charge. Signed-off-by: Glauber Costa <glommer@parallels.com> Acked-by: Michal Hocko <mhocko@suse.cz> Acked-by: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Tejun Heo <tj@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Frederic Weisbecker <fweisbec@redhat.com> Cc: Greg Thelen <gthelen@google.com> Cc: JoonSoo Kim <js1304@gmail.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik van Riel <riel@redhat.com> Cc: Suleiman Souhlal <suleiman@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-12-19 06:21:56 +08:00
}
/*
* helper for accessing a memcg's index. It will be used as an index in the
* child cache array in kmem_cache, and also to derive its name. This function
* will return -1 when this is not a kmem-limited memcg.
*/
static inline int memcg_cache_id(struct mem_cgroup *memcg)
{
return memcg ? memcg->kmemcg_id : -1;
}
/**
* memcg_kmem_update_page_stat - update kmem page state statistics
* @page: the page
* @idx: page state item to account
* @val: number of pages (positive or negative)
*/
static inline void memcg_kmem_update_page_stat(struct page *page,
enum mem_cgroup_stat_index idx, int val)
{
if (memcg_kmem_enabled() && page->mem_cgroup)
this_cpu_add(page->mem_cgroup->stat->count[idx], val);
}
memcg: kmem controller infrastructure Introduce infrastructure for tracking kernel memory pages to a given memcg. This will happen whenever the caller includes the flag __GFP_KMEMCG flag, and the task belong to a memcg other than the root. In memcontrol.h those functions are wrapped in inline acessors. The idea is to later on, patch those with static branches, so we don't incur any overhead when no mem cgroups with limited kmem are being used. Users of this functionality shall interact with the memcg core code through the following functions: memcg_kmem_newpage_charge: will return true if the group can handle the allocation. At this point, struct page is not yet allocated. memcg_kmem_commit_charge: will either revert the charge, if struct page allocation failed, or embed memcg information into page_cgroup. memcg_kmem_uncharge_page: called at free time, will revert the charge. Signed-off-by: Glauber Costa <glommer@parallels.com> Acked-by: Michal Hocko <mhocko@suse.cz> Acked-by: Kamezawa Hiroyuki <kamezawa.hiroyu@jp.fujitsu.com> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@cs.helsinki.fi> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Tejun Heo <tj@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Frederic Weisbecker <fweisbec@redhat.com> Cc: Greg Thelen <gthelen@google.com> Cc: JoonSoo Kim <js1304@gmail.com> Cc: Mel Gorman <mel@csn.ul.ie> Cc: Rik van Riel <riel@redhat.com> Cc: Suleiman Souhlal <suleiman@google.com> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2012-12-19 06:21:56 +08:00
#else
#define for_each_memcg_cache_index(_idx) \
for (; NULL; )
static inline bool memcg_kmem_enabled(void)
{
return false;
}
static inline int memcg_cache_id(struct mem_cgroup *memcg)
{
return -1;
}
memcg: add rwsem to synchronize against memcg_caches arrays relocation We need a stable value of memcg_nr_cache_ids in kmem_cache_create() (memcg_alloc_cache_params() wants it for root caches), where we only hold the slab_mutex and no memcg-related locks. As a result, we have to update memcg_nr_cache_ids under the slab_mutex, which we can only take on the slab's side (see memcg_update_array_size). This looks awkward and will become even worse when per-memcg list_lru is introduced, which also wants stable access to memcg_nr_cache_ids. To get rid of this dependency between the memcg_nr_cache_ids and the slab_mutex, this patch introduces a special rwsem. The rwsem is held for writing during memcg_caches arrays relocation and memcg_nr_cache_ids updates. Therefore one can take it for reading to get a stable access to memcg_caches arrays and/or memcg_nr_cache_ids. Currently the semaphore is taken for reading only from kmem_cache_create, right before taking the slab_mutex, so right now there's no much point in using rwsem instead of mutex. However, once list_lru is made per-memcg it will allow list_lru initializations to proceed concurrently. Signed-off-by: Vladimir Davydov <vdavydov@parallels.com> Cc: Dave Chinner <david@fromorbit.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Michal Hocko <mhocko@suse.cz> Cc: Greg Thelen <gthelen@google.com> Cc: Glauber Costa <glommer@gmail.com> Cc: Alexander Viro <viro@zeniv.linux.org.uk> Cc: Christoph Lameter <cl@linux.com> Cc: Pekka Enberg <penberg@kernel.org> Cc: David Rientjes <rientjes@google.com> Cc: Joonsoo Kim <iamjoonsoo.kim@lge.com> Cc: Tejun Heo <tj@kernel.org> Signed-off-by: Andrew Morton <akpm@linux-foundation.org> Signed-off-by: Linus Torvalds <torvalds@linux-foundation.org>
2015-02-13 06:59:01 +08:00
static inline void memcg_get_cache_ids(void)
{
}
static inline void memcg_put_cache_ids(void)
{
}
static inline void memcg_kmem_update_page_stat(struct page *page,
enum mem_cgroup_stat_index idx, int val)
{
}
#endif /* CONFIG_MEMCG && !CONFIG_SLOB */
#endif /* _LINUX_MEMCONTROL_H */