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[PATCH] pi-futex: rt mutex docs 

Add rt-mutex documentation.

[rostedt@goodmis.org: Update rt-mutex-design.txt as per Randy Dunlap suggestions]
Signed-off-by: Ingo Molnar <mingo@elte.hu>
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
Signed-off-by: Arjan van de Ven <arjan@linux.intel.com>
Signed-off-by: Steven Rostedt <rostedt@goodmis.org>
Signed-off-by: Steven Rostedt <rostedt@goodmis.org>
Cc: "Randy.Dunlap" <rdunlap@xenotime.net>
Signed-off-by: Andrew Morton <akpm@osdl.org>
Signed-off-by: Linus Torvalds <torvalds@osdl.org>
This commit is contained in:
Steven Rostedt 2006-06-27 02:54:54 -07:00 committed by Linus Torvalds
parent 23f78d4a03
commit a6537be932
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Lightweight PI-futexes
----------------------
We are calling them lightweight for 3 reasons:
- in the user-space fastpath a PI-enabled futex involves no kernel work
(or any other PI complexity) at all. No registration, no extra kernel
calls - just pure fast atomic ops in userspace.
- even in the slowpath, the system call and scheduling pattern is very
similar to normal futexes.
- the in-kernel PI implementation is streamlined around the mutex
abstraction, with strict rules that keep the implementation
relatively simple: only a single owner may own a lock (i.e. no
read-write lock support), only the owner may unlock a lock, no
recursive locking, etc.
Priority Inheritance - why?
---------------------------
The short reply: user-space PI helps achieving/improving determinism for
user-space applications. In the best-case, it can help achieve
determinism and well-bound latencies. Even in the worst-case, PI will
improve the statistical distribution of locking related application
delays.
The longer reply:
-----------------
Firstly, sharing locks between multiple tasks is a common programming
technique that often cannot be replaced with lockless algorithms. As we
can see it in the kernel [which is a quite complex program in itself],
lockless structures are rather the exception than the norm - the current
ratio of lockless vs. locky code for shared data structures is somewhere
between 1:10 and 1:100. Lockless is hard, and the complexity of lockless
algorithms often endangers to ability to do robust reviews of said code.
I.e. critical RT apps often choose lock structures to protect critical
data structures, instead of lockless algorithms. Furthermore, there are
cases (like shared hardware, or other resource limits) where lockless
access is mathematically impossible.
Media players (such as Jack) are an example of reasonable application
design with multiple tasks (with multiple priority levels) sharing
short-held locks: for example, a highprio audio playback thread is
combined with medium-prio construct-audio-data threads and low-prio
display-colory-stuff threads. Add video and decoding to the mix and
we've got even more priority levels.
So once we accept that synchronization objects (locks) are an
unavoidable fact of life, and once we accept that multi-task userspace
apps have a very fair expectation of being able to use locks, we've got
to think about how to offer the option of a deterministic locking
implementation to user-space.
Most of the technical counter-arguments against doing priority
inheritance only apply to kernel-space locks. But user-space locks are
different, there we cannot disable interrupts or make the task
non-preemptible in a critical section, so the 'use spinlocks' argument
does not apply (user-space spinlocks have the same priority inversion
problems as other user-space locking constructs). Fact is, pretty much
the only technique that currently enables good determinism for userspace
locks (such as futex-based pthread mutexes) is priority inheritance:
Currently (without PI), if a high-prio and a low-prio task shares a lock
[this is a quite common scenario for most non-trivial RT applications],
even if all critical sections are coded carefully to be deterministic
(i.e. all critical sections are short in duration and only execute a
limited number of instructions), the kernel cannot guarantee any
deterministic execution of the high-prio task: any medium-priority task
could preempt the low-prio task while it holds the shared lock and
executes the critical section, and could delay it indefinitely.
Implementation:
---------------
As mentioned before, the userspace fastpath of PI-enabled pthread
mutexes involves no kernel work at all - they behave quite similarly to
normal futex-based locks: a 0 value means unlocked, and a value==TID
means locked. (This is the same method as used by list-based robust
futexes.) Userspace uses atomic ops to lock/unlock these mutexes without
entering the kernel.
To handle the slowpath, we have added two new futex ops:
FUTEX_LOCK_PI
FUTEX_UNLOCK_PI
If the lock-acquire fastpath fails, [i.e. an atomic transition from 0 to
TID fails], then FUTEX_LOCK_PI is called. The kernel does all the
remaining work: if there is no futex-queue attached to the futex address
yet then the code looks up the task that owns the futex [it has put its
own TID into the futex value], and attaches a 'PI state' structure to
the futex-queue. The pi_state includes an rt-mutex, which is a PI-aware,
kernel-based synchronization object. The 'other' task is made the owner
of the rt-mutex, and the FUTEX_WAITERS bit is atomically set in the
futex value. Then this task tries to lock the rt-mutex, on which it
blocks. Once it returns, it has the mutex acquired, and it sets the
futex value to its own TID and returns. Userspace has no other work to
perform - it now owns the lock, and futex value contains
FUTEX_WAITERS|TID.
If the unlock side fastpath succeeds, [i.e. userspace manages to do a
TID -> 0 atomic transition of the futex value], then no kernel work is
triggered.
If the unlock fastpath fails (because the FUTEX_WAITERS bit is set),
then FUTEX_UNLOCK_PI is called, and the kernel unlocks the futex on the
behalf of userspace - and it also unlocks the attached
pi_state->rt_mutex and thus wakes up any potential waiters.
Note that under this approach, contrary to previous PI-futex approaches,
there is no prior 'registration' of a PI-futex. [which is not quite
possible anyway, due to existing ABI properties of pthread mutexes.]
Also, under this scheme, 'robustness' and 'PI' are two orthogonal
properties of futexes, and all four combinations are possible: futex,
robust-futex, PI-futex, robust+PI-futex.
More details about priority inheritance can be found in
Documentation/rtmutex.txt.

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#
# Copyright (c) 2006 Steven Rostedt
# Licensed under the GNU Free Documentation License, Version 1.2
#
RT-mutex implementation design
------------------------------
This document tries to describe the design of the rtmutex.c implementation.
It doesn't describe the reasons why rtmutex.c exists. For that please see
Documentation/rt-mutex.txt. Although this document does explain problems
that happen without this code, but that is in the concept to understand
what the code actually is doing.
The goal of this document is to help others understand the priority
inheritance (PI) algorithm that is used, as well as reasons for the
decisions that were made to implement PI in the manner that was done.
Unbounded Priority Inversion
----------------------------
Priority inversion is when a lower priority process executes while a higher
priority process wants to run. This happens for several reasons, and
most of the time it can't be helped. Anytime a high priority process wants
to use a resource that a lower priority process has (a mutex for example),
the high priority process must wait until the lower priority process is done
with the resource. This is a priority inversion. What we want to prevent
is something called unbounded priority inversion. That is when the high
priority process is prevented from running by a lower priority process for
an undetermined amount of time.
The classic example of unbounded priority inversion is were you have three
processes, let's call them processes A, B, and C, where A is the highest
priority process, C is the lowest, and B is in between. A tries to grab a lock
that C owns and must wait and lets C run to release the lock. But in the
meantime, B executes, and since B is of a higher priority than C, it preempts C,
but by doing so, it is in fact preempting A which is a higher priority process.
Now there's no way of knowing how long A will be sleeping waiting for C
to release the lock, because for all we know, B is a CPU hog and will
never give C a chance to release the lock. This is called unbounded priority
inversion.
Here's a little ASCII art to show the problem.
grab lock L1 (owned by C)
|
A ---+
C preempted by B
|
C +----+
B +-------->
B now keeps A from running.
Priority Inheritance (PI)
-------------------------
There are several ways to solve this issue, but other ways are out of scope
for this document. Here we only discuss PI.
PI is where a process inherits the priority of another process if the other
process blocks on a lock owned by the current process. To make this easier
to understand, let's use the previous example, with processes A, B, and C again.
This time, when A blocks on the lock owned by C, C would inherit the priority
of A. So now if B becomes runnable, it would not preempt C, since C now has
the high priority of A. As soon as C releases the lock, it loses its
inherited priority, and A then can continue with the resource that C had.
Terminology
-----------
Here I explain some terminology that is used in this document to help describe
the design that is used to implement PI.
PI chain - The PI chain is an ordered series of locks and processes that cause
processes to inherit priorities from a previous process that is
blocked on one of its locks. This is described in more detail
later in this document.
mutex - In this document, to differentiate from locks that implement
PI and spin locks that are used in the PI code, from now on
the PI locks will be called a mutex.
lock - In this document from now on, I will use the term lock when
referring to spin locks that are used to protect parts of the PI
algorithm. These locks disable preemption for UP (when
CONFIG_PREEMPT is enabled) and on SMP prevents multiple CPUs from
entering critical sections simultaneously.
spin lock - Same as lock above.
waiter - A waiter is a struct that is stored on the stack of a blocked
process. Since the scope of the waiter is within the code for
a process being blocked on the mutex, it is fine to allocate
the waiter on the process's stack (local variable). This
structure holds a pointer to the task, as well as the mutex that
the task is blocked on. It also has the plist node structures to
place the task in the waiter_list of a mutex as well as the
pi_list of a mutex owner task (described below).
waiter is sometimes used in reference to the task that is waiting
on a mutex. This is the same as waiter->task.
waiters - A list of processes that are blocked on a mutex.
top waiter - The highest priority process waiting on a specific mutex.
top pi waiter - The highest priority process waiting on one of the mutexes
that a specific process owns.
Note: task and process are used interchangeably in this document, mostly to
differentiate between two processes that are being described together.
PI chain
--------
The PI chain is a list of processes and mutexes that may cause priority
inheritance to take place. Multiple chains may converge, but a chain
would never diverge, since a process can't be blocked on more than one
mutex at a time.
Example:
Process: A, B, C, D, E
Mutexes: L1, L2, L3, L4
A owns: L1
B blocked on L1
B owns L2
C blocked on L2
C owns L3
D blocked on L3
D owns L4
E blocked on L4
The chain would be:
E->L4->D->L3->C->L2->B->L1->A
To show where two chains merge, we could add another process F and
another mutex L5 where B owns L5 and F is blocked on mutex L5.
The chain for F would be:
F->L5->B->L1->A
Since a process may own more than one mutex, but never be blocked on more than
one, the chains merge.
Here we show both chains:
E->L4->D->L3->C->L2-+
|
+->B->L1->A
|
F->L5-+
For PI to work, the processes at the right end of these chains (or we may
also call it the Top of the chain) must be equal to or higher in priority
than the processes to the left or below in the chain.
Also since a mutex may have more than one process blocked on it, we can
have multiple chains merge at mutexes. If we add another process G that is
blocked on mutex L2:
G->L2->B->L1->A
And once again, to show how this can grow I will show the merging chains
again.
E->L4->D->L3->C-+
+->L2-+
| |
G-+ +->B->L1->A
|
F->L5-+
Plist
-----
Before I go further and talk about how the PI chain is stored through lists
on both mutexes and processes, I'll explain the plist. This is similar to
the struct list_head functionality that is already in the kernel.
The implementation of plist is out of scope for this document, but it is
very important to understand what it does.
There are a few differences between plist and list, the most important one
being that plist is a priority sorted linked list. This means that the
priorities of the plist are sorted, such that it takes O(1) to retrieve the
highest priority item in the list. Obviously this is useful to store processes
based on their priorities.
Another difference, which is important for implementation, is that, unlike
list, the head of the list is a different element than the nodes of a list.
So the head of the list is declared as struct plist_head and nodes that will
be added to the list are declared as struct plist_node.
Mutex Waiter List
-----------------
Every mutex keeps track of all the waiters that are blocked on itself. The mutex
has a plist to store these waiters by priority. This list is protected by
a spin lock that is located in the struct of the mutex. This lock is called
wait_lock. Since the modification of the waiter list is never done in
interrupt context, the wait_lock can be taken without disabling interrupts.
Task PI List
------------
To keep track of the PI chains, each process has its own PI list. This is
a list of all top waiters of the mutexes that are owned by the process.
Note that this list only holds the top waiters and not all waiters that are
blocked on mutexes owned by the process.
The top of the task's PI list is always the highest priority task that
is waiting on a mutex that is owned by the task. So if the task has
inherited a priority, it will always be the priority of the task that is
at the top of this list.
This list is stored in the task structure of a process as a plist called
pi_list. This list is protected by a spin lock also in the task structure,
called pi_lock. This lock may also be taken in interrupt context, so when
locking the pi_lock, interrupts must be disabled.
Depth of the PI Chain
---------------------
The maximum depth of the PI chain is not dynamic, and could actually be
defined. But is very complex to figure it out, since it depends on all
the nesting of mutexes. Let's look at the example where we have 3 mutexes,
L1, L2, and L3, and four separate functions func1, func2, func3 and func4.
The following shows a locking order of L1->L2->L3, but may not actually
be directly nested that way.
void func1(void)
{
mutex_lock(L1);
/* do anything */
mutex_unlock(L1);
}
void func2(void)
{
mutex_lock(L1);
mutex_lock(L2);
/* do something */
mutex_unlock(L2);
mutex_unlock(L1);
}
void func3(void)
{
mutex_lock(L2);
mutex_lock(L3);
/* do something else */
mutex_unlock(L3);
mutex_unlock(L2);
}
void func4(void)
{
mutex_lock(L3);
/* do something again */
mutex_unlock(L3);
}
Now we add 4 processes that run each of these functions separately.
Processes A, B, C, and D which run functions func1, func2, func3 and func4
respectively, and such that D runs first and A last. With D being preempted
in func4 in the "do something again" area, we have a locking that follows:
D owns L3
C blocked on L3
C owns L2
B blocked on L2
B owns L1
A blocked on L1
And thus we have the chain A->L1->B->L2->C->L3->D.
This gives us a PI depth of 4 (four processes), but looking at any of the
functions individually, it seems as though they only have at most a locking
depth of two. So, although the locking depth is defined at compile time,
it still is very difficult to find the possibilities of that depth.
Now since mutexes can be defined by user-land applications, we don't want a DOS
type of application that nests large amounts of mutexes to create a large
PI chain, and have the code holding spin locks while looking at a large
amount of data. So to prevent this, the implementation not only implements
a maximum lock depth, but also only holds at most two different locks at a
time, as it walks the PI chain. More about this below.
Mutex owner and flags
---------------------
The mutex structure contains a pointer to the owner of the mutex. If the
mutex is not owned, this owner is set to NULL. Since all architectures
have the task structure on at least a four byte alignment (and if this is
not true, the rtmutex.c code will be broken!), this allows for the two
least significant bits to be used as flags. This part is also described
in Documentation/rt-mutex.txt, but will also be briefly described here.
Bit 0 is used as the "Pending Owner" flag. This is described later.
Bit 1 is used as the "Has Waiters" flags. This is also described later
in more detail, but is set whenever there are waiters on a mutex.
cmpxchg Tricks
--------------
Some architectures implement an atomic cmpxchg (Compare and Exchange). This
is used (when applicable) to keep the fast path of grabbing and releasing
mutexes short.
cmpxchg is basically the following function performed atomically:
unsigned long _cmpxchg(unsigned long *A, unsigned long *B, unsigned long *C)
{
unsigned long T = *A;
if (*A == *B) {
*A = *C;
}
return T;
}
#define cmpxchg(a,b,c) _cmpxchg(&a,&b,&c)
This is really nice to have, since it allows you to only update a variable
if the variable is what you expect it to be. You know if it succeeded if
the return value (the old value of A) is equal to B.
The macro rt_mutex_cmpxchg is used to try to lock and unlock mutexes. If
the architecture does not support CMPXCHG, then this macro is simply set
to fail every time. But if CMPXCHG is supported, then this will
help out extremely to keep the fast path short.
The use of rt_mutex_cmpxchg with the flags in the owner field help optimize
the system for architectures that support it. This will also be explained
later in this document.
Priority adjustments
--------------------
The implementation of the PI code in rtmutex.c has several places that a
process must adjust its priority. With the help of the pi_list of a
process this is rather easy to know what needs to be adjusted.
The functions implementing the task adjustments are rt_mutex_adjust_prio,
__rt_mutex_adjust_prio (same as the former, but expects the task pi_lock
to already be taken), rt_mutex_get_prio, and rt_mutex_setprio.
rt_mutex_getprio and rt_mutex_setprio are only used in __rt_mutex_adjust_prio.
rt_mutex_getprio returns the priority that the task should have. Either the
task's own normal priority, or if a process of a higher priority is waiting on
a mutex owned by the task, then that higher priority should be returned.
Since the pi_list of a task holds an order by priority list of all the top
waiters of all the mutexes that the task owns, rt_mutex_getprio simply needs
to compare the top pi waiter to its own normal priority, and return the higher
priority back.
(Note: if looking at the code, you will notice that the lower number of
prio is returned. This is because the prio field in the task structure
is an inverse order of the actual priority. So a "prio" of 5 is
of higher priority than a "prio" of 10.)
__rt_mutex_adjust_prio examines the result of rt_mutex_getprio, and if the
result does not equal the task's current priority, then rt_mutex_setprio
is called to adjust the priority of the task to the new priority.
Note that rt_mutex_setprio is defined in kernel/sched.c to implement the
actual change in priority.
It is interesting to note that __rt_mutex_adjust_prio can either increase
or decrease the priority of the task. In the case that a higher priority
process has just blocked on a mutex owned by the task, __rt_mutex_adjust_prio
would increase/boost the task's priority. But if a higher priority task
were for some reason to leave the mutex (timeout or signal), this same function
would decrease/unboost the priority of the task. That is because the pi_list
always contains the highest priority task that is waiting on a mutex owned
by the task, so we only need to compare the priority of that top pi waiter
to the normal priority of the given task.
High level overview of the PI chain walk
----------------------------------------
The PI chain walk is implemented by the function rt_mutex_adjust_prio_chain.
The implementation has gone through several iterations, and has ended up
with what we believe is the best. It walks the PI chain by only grabbing
at most two locks at a time, and is very efficient.
The rt_mutex_adjust_prio_chain can be used either to boost or lower process
priorities.
rt_mutex_adjust_prio_chain is called with a task to be checked for PI
(de)boosting (the owner of a mutex that a process is blocking on), a flag to
check for deadlocking, the mutex that the task owns, and a pointer to a waiter
that is the process's waiter struct that is blocked on the mutex (although this
parameter may be NULL for deboosting).
For this explanation, I will not mention deadlock detection. This explanation
will try to stay at a high level.
When this function is called, there are no locks held. That also means
that the state of the owner and lock can change when entered into this function.
Before this function is called, the task has already had rt_mutex_adjust_prio
performed on it. This means that the task is set to the priority that it
should be at, but the plist nodes of the task's waiter have not been updated
with the new priorities, and that this task may not be in the proper locations
in the pi_lists and wait_lists that the task is blocked on. This function
solves all that.
A loop is entered, where task is the owner to be checked for PI changes that
was passed by parameter (for the first iteration). The pi_lock of this task is
taken to prevent any more changes to the pi_list of the task. This also
prevents new tasks from completing the blocking on a mutex that is owned by this
task.
If the task is not blocked on a mutex then the loop is exited. We are at
the top of the PI chain.
A check is now done to see if the original waiter (the process that is blocked
on the current mutex) is the top pi waiter of the task. That is, is this
waiter on the top of the task's pi_list. If it is not, it either means that
there is another process higher in priority that is blocked on one of the
mutexes that the task owns, or that the waiter has just woken up via a signal
or timeout and has left the PI chain. In either case, the loop is exited, since
we don't need to do any more changes to the priority of the current task, or any
task that owns a mutex that this current task is waiting on. A priority chain
walk is only needed when a new top pi waiter is made to a task.
The next check sees if the task's waiter plist node has the priority equal to
the priority the task is set at. If they are equal, then we are done with
the loop. Remember that the function started with the priority of the
task adjusted, but the plist nodes that hold the task in other processes
pi_lists have not been adjusted.
Next, we look at the mutex that the task is blocked on. The mutex's wait_lock
is taken. This is done by a spin_trylock, because the locking order of the
pi_lock and wait_lock goes in the opposite direction. If we fail to grab the
lock, the pi_lock is released, and we restart the loop.
Now that we have both the pi_lock of the task as well as the wait_lock of
the mutex the task is blocked on, we update the task's waiter's plist node
that is located on the mutex's wait_list.
Now we release the pi_lock of the task.
Next the owner of the mutex has its pi_lock taken, so we can update the
task's entry in the owner's pi_list. If the task is the highest priority
process on the mutex's wait_list, then we remove the previous top waiter
from the owner's pi_list, and replace it with the task.
Note: It is possible that the task was the current top waiter on the mutex,
in which case the task is not yet on the pi_list of the waiter. This
is OK, since plist_del does nothing if the plist node is not on any
list.
If the task was not the top waiter of the mutex, but it was before we
did the priority updates, that means we are deboosting/lowering the
task. In this case, the task is removed from the pi_list of the owner,
and the new top waiter is added.
Lastly, we unlock both the pi_lock of the task, as well as the mutex's
wait_lock, and continue the loop again. On the next iteration of the
loop, the previous owner of the mutex will be the task that will be
processed.
Note: One might think that the owner of this mutex might have changed
since we just grab the mutex's wait_lock. And one could be right.
The important thing to remember is that the owner could not have
become the task that is being processed in the PI chain, since
we have taken that task's pi_lock at the beginning of the loop.
So as long as there is an owner of this mutex that is not the same
process as the tasked being worked on, we are OK.
Looking closely at the code, one might be confused. The check for the
end of the PI chain is when the task isn't blocked on anything or the
task's waiter structure "task" element is NULL. This check is
protected only by the task's pi_lock. But the code to unlock the mutex
sets the task's waiter structure "task" element to NULL with only
the protection of the mutex's wait_lock, which was not taken yet.
Isn't this a race condition if the task becomes the new owner?
The answer is No! The trick is the spin_trylock of the mutex's
wait_lock. If we fail that lock, we release the pi_lock of the
task and continue the loop, doing the end of PI chain check again.
In the code to release the lock, the wait_lock of the mutex is held
the entire time, and it is not let go when we grab the pi_lock of the
new owner of the mutex. So if the switch of a new owner were to happen
after the check for end of the PI chain and the grabbing of the
wait_lock, the unlocking code would spin on the new owner's pi_lock
but never give up the wait_lock. So the PI chain loop is guaranteed to
fail the spin_trylock on the wait_lock, release the pi_lock, and
try again.
If you don't quite understand the above, that's OK. You don't have to,
unless you really want to make a proof out of it ;)
Pending Owners and Lock stealing
--------------------------------
One of the flags in the owner field of the mutex structure is "Pending Owner".
What this means is that an owner was chosen by the process releasing the
mutex, but that owner has yet to wake up and actually take the mutex.
Why is this important? Why can't we just give the mutex to another process
and be done with it?
The PI code is to help with real-time processes, and to let the highest
priority process run as long as possible with little latencies and delays.
If a high priority process owns a mutex that a lower priority process is
blocked on, when the mutex is released it would be given to the lower priority
process. What if the higher priority process wants to take that mutex again.
The high priority process would fail to take that mutex that it just gave up
and it would need to boost the lower priority process to run with full
latency of that critical section (since the low priority process just entered
it).
There's no reason a high priority process that gives up a mutex should be
penalized if it tries to take that mutex again. If the new owner of the
mutex has not woken up yet, there's no reason that the higher priority process
could not take that mutex away.
To solve this, we introduced Pending Ownership and Lock Stealing. When a
new process is given a mutex that it was blocked on, it is only given
pending ownership. This means that it's the new owner, unless a higher
priority process comes in and tries to grab that mutex. If a higher priority
process does come along and wants that mutex, we let the higher priority
process "steal" the mutex from the pending owner (only if it is still pending)
and continue with the mutex.
Taking of a mutex (The walk through)
------------------------------------
OK, now let's take a look at the detailed walk through of what happens when
taking a mutex.
The first thing that is tried is the fast taking of the mutex. This is
done when we have CMPXCHG enabled (otherwise the fast taking automatically
fails). Only when the owner field of the mutex is NULL can the lock be
taken with the CMPXCHG and nothing else needs to be done.
If there is contention on the lock, whether it is owned or pending owner
we go about the slow path (rt_mutex_slowlock).
The slow path function is where the task's waiter structure is created on
the stack. This is because the waiter structure is only needed for the
scope of this function. The waiter structure holds the nodes to store
the task on the wait_list of the mutex, and if need be, the pi_list of
the owner.
The wait_lock of the mutex is taken since the slow path of unlocking the
mutex also takes this lock.
We then call try_to_take_rt_mutex. This is where the architecture that
does not implement CMPXCHG would always grab the lock (if there's no
contention).
try_to_take_rt_mutex is used every time the task tries to grab a mutex in the
slow path. The first thing that is done here is an atomic setting of
the "Has Waiters" flag of the mutex's owner field. Yes, this could really
be false, because if the the mutex has no owner, there are no waiters and
the current task also won't have any waiters. But we don't have the lock
yet, so we assume we are going to be a waiter. The reason for this is to
play nice for those architectures that do have CMPXCHG. By setting this flag
now, the owner of the mutex can't release the mutex without going into the
slow unlock path, and it would then need to grab the wait_lock, which this
code currently holds. So setting the "Has Waiters" flag forces the owner
to synchronize with this code.
Now that we know that we can't have any races with the owner releasing the
mutex, we check to see if we can take the ownership. This is done if the
mutex doesn't have a owner, or if we can steal the mutex from a pending
owner. Let's look at the situations we have here.
1) Has owner that is pending
----------------------------
The mutex has a owner, but it hasn't woken up and the mutex flag
"Pending Owner" is set. The first check is to see if the owner isn't the
current task. This is because this function is also used for the pending
owner to grab the mutex. When a pending owner wakes up, it checks to see
if it can take the mutex, and this is done if the owner is already set to
itself. If so, we succeed and leave the function, clearing the "Pending
Owner" bit.
If the pending owner is not current, we check to see if the current priority is
higher than the pending owner. If not, we fail the function and return.
There's also something special about a pending owner. That is a pending owner
is never blocked on a mutex. So there is no PI chain to worry about. It also
means that if the mutex doesn't have any waiters, there's no accounting needed
to update the pending owner's pi_list, since we only worry about processes
blocked on the current mutex.
If there are waiters on this mutex, and we just stole the ownership, we need
to take the top waiter, remove it from the pi_list of the pending owner, and
add it to the current pi_list. Note that at this moment, the pending owner
is no longer on the list of waiters. This is fine, since the pending owner
would add itself back when it realizes that it had the ownership stolen
from itself. When the pending owner tries to grab the mutex, it will fail
in try_to_take_rt_mutex if the owner field points to another process.
2) No owner
-----------
If there is no owner (or we successfully stole the lock), we set the owner
of the mutex to current, and set the flag of "Has Waiters" if the current
mutex actually has waiters, or we clear the flag if it doesn't. See, it was
OK that we set that flag early, since now it is cleared.
3) Failed to grab ownership
---------------------------
The most interesting case is when we fail to take ownership. This means that
there exists an owner, or there's a pending owner with equal or higher
priority than the current task.
We'll continue on the failed case.
If the mutex has a timeout, we set up a timer to go off to break us out
of this mutex if we failed to get it after a specified amount of time.
Now we enter a loop that will continue to try to take ownership of the mutex, or
fail from a timeout or signal.
Once again we try to take the mutex. This will usually fail the first time
in the loop, since it had just failed to get the mutex. But the second time
in the loop, this would likely succeed, since the task would likely be
the pending owner.
If the mutex is TASK_INTERRUPTIBLE a check for signals and timeout is done
here.
The waiter structure has a "task" field that points to the task that is blocked
on the mutex. This field can be NULL the first time it goes through the loop
or if the task is a pending owner and had it's mutex stolen. If the "task"
field is NULL then we need to set up the accounting for it.
Task blocks on mutex
--------------------
The accounting of a mutex and process is done with the waiter structure of
the process. The "task" field is set to the process, and the "lock" field
to the mutex. The plist nodes are initialized to the processes current
priority.
Since the wait_lock was taken at the entry of the slow lock, we can safely
add the waiter to the wait_list. If the current process is the highest
priority process currently waiting on this mutex, then we remove the
previous top waiter process (if it exists) from the pi_list of the owner,
and add the current process to that list. Since the pi_list of the owner
has changed, we call rt_mutex_adjust_prio on the owner to see if the owner
should adjust its priority accordingly.
If the owner is also blocked on a lock, and had its pi_list changed
(or deadlock checking is on), we unlock the wait_lock of the mutex and go ahead
and run rt_mutex_adjust_prio_chain on the owner, as described earlier.
Now all locks are released, and if the current process is still blocked on a
mutex (waiter "task" field is not NULL), then we go to sleep (call schedule).
Waking up in the loop
---------------------
The schedule can then wake up for a few reasons.
1) we were given pending ownership of the mutex.
2) we received a signal and was TASK_INTERRUPTIBLE
3) we had a timeout and was TASK_INTERRUPTIBLE
In any of these cases, we continue the loop and once again try to grab the
ownership of the mutex. If we succeed, we exit the loop, otherwise we continue
and on signal and timeout, will exit the loop, or if we had the mutex stolen
we just simply add ourselves back on the lists and go back to sleep.
Note: For various reasons, because of timeout and signals, the steal mutex
algorithm needs to be careful. This is because the current process is
still on the wait_list. And because of dynamic changing of priorities,
especially on SCHED_OTHER tasks, the current process can be the
highest priority task on the wait_list.
Failed to get mutex on Timeout or Signal
----------------------------------------
If a timeout or signal occurred, the waiter's "task" field would not be
NULL and the task needs to be taken off the wait_list of the mutex and perhaps
pi_list of the owner. If this process was a high priority process, then
the rt_mutex_adjust_prio_chain needs to be executed again on the owner,
but this time it will be lowering the priorities.
Unlocking the Mutex
-------------------
The unlocking of a mutex also has a fast path for those architectures with
CMPXCHG. Since the taking of a mutex on contention always sets the
"Has Waiters" flag of the mutex's owner, we use this to know if we need to
take the slow path when unlocking the mutex. If the mutex doesn't have any
waiters, the owner field of the mutex would equal the current process and
the mutex can be unlocked by just replacing the owner field with NULL.
If the owner field has the "Has Waiters" bit set (or CMPXCHG is not available),
the slow unlock path is taken.
The first thing done in the slow unlock path is to take the wait_lock of the
mutex. This synchronizes the locking and unlocking of the mutex.
A check is made to see if the mutex has waiters or not. On architectures that
do not have CMPXCHG, this is the location that the owner of the mutex will
determine if a waiter needs to be awoken or not. On architectures that
do have CMPXCHG, that check is done in the fast path, but it is still needed
in the slow path too. If a waiter of a mutex woke up because of a signal
or timeout between the time the owner failed the fast path CMPXCHG check and
the grabbing of the wait_lock, the mutex may not have any waiters, thus the
owner still needs to make this check. If there are no waiters than the mutex
owner field is set to NULL, the wait_lock is released and nothing more is
needed.
If there are waiters, then we need to wake one up and give that waiter
pending ownership.
On the wake up code, the pi_lock of the current owner is taken. The top
waiter of the lock is found and removed from the wait_list of the mutex
as well as the pi_list of the current owner. The task field of the new
pending owner's waiter structure is set to NULL, and the owner field of the
mutex is set to the new owner with the "Pending Owner" bit set, as well
as the "Has Waiters" bit if there still are other processes blocked on the
mutex.
The pi_lock of the previous owner is released, and the new pending owner's
pi_lock is taken. Remember that this is the trick to prevent the race
condition in rt_mutex_adjust_prio_chain from adding itself as a waiter
on the mutex.
We now clear the "pi_blocked_on" field of the new pending owner, and if
the mutex still has waiters pending, we add the new top waiter to the pi_list
of the pending owner.
Finally we unlock the pi_lock of the pending owner and wake it up.
Contact
-------
For updates on this document, please email Steven Rostedt <rostedt@goodmis.org>
Credits
-------
Author: Steven Rostedt <rostedt@goodmis.org>
Reviewers: Ingo Molnar, Thomas Gleixner, Thomas Duetsch, and Randy Dunlap
Updates
-------
This document was originally written for 2.6.17-rc3-mm1

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RT-mutex subsystem with PI support
----------------------------------
RT-mutexes with priority inheritance are used to support PI-futexes,
which enable pthread_mutex_t priority inheritance attributes
(PTHREAD_PRIO_INHERIT). [See Documentation/pi-futex.txt for more details
about PI-futexes.]
This technology was developed in the -rt tree and streamlined for
pthread_mutex support.
Basic principles:
-----------------
RT-mutexes extend the semantics of simple mutexes by the priority
inheritance protocol.
A low priority owner of a rt-mutex inherits the priority of a higher
priority waiter until the rt-mutex is released. If the temporarily
boosted owner blocks on a rt-mutex itself it propagates the priority
boosting to the owner of the other rt_mutex it gets blocked on. The
priority boosting is immediately removed once the rt_mutex has been
unlocked.
This approach allows us to shorten the block of high-prio tasks on
mutexes which protect shared resources. Priority inheritance is not a
magic bullet for poorly designed applications, but it allows
well-designed applications to use userspace locks in critical parts of
an high priority thread, without losing determinism.
The enqueueing of the waiters into the rtmutex waiter list is done in
priority order. For same priorities FIFO order is chosen. For each
rtmutex, only the top priority waiter is enqueued into the owner's
priority waiters list. This list too queues in priority order. Whenever
the top priority waiter of a task changes (for example it timed out or
got a signal), the priority of the owner task is readjusted. [The
priority enqueueing is handled by "plists", see include/linux/plist.h
for more details.]
RT-mutexes are optimized for fastpath operations and have no internal
locking overhead when locking an uncontended mutex or unlocking a mutex
without waiters. The optimized fastpath operations require cmpxchg
support. [If that is not available then the rt-mutex internal spinlock
is used]
The state of the rt-mutex is tracked via the owner field of the rt-mutex
structure:
rt_mutex->owner holds the task_struct pointer of the owner. Bit 0 and 1
are used to keep track of the "owner is pending" and "rtmutex has
waiters" state.
owner bit1 bit0
NULL 0 0 mutex is free (fast acquire possible)
NULL 0 1 invalid state
NULL 1 0 Transitional state*
NULL 1 1 invalid state
taskpointer 0 0 mutex is held (fast release possible)
taskpointer 0 1 task is pending owner
taskpointer 1 0 mutex is held and has waiters
taskpointer 1 1 task is pending owner and mutex has waiters
Pending-ownership handling is a performance optimization:
pending-ownership is assigned to the first (highest priority) waiter of
the mutex, when the mutex is released. The thread is woken up and once
it starts executing it can acquire the mutex. Until the mutex is taken
by it (bit 0 is cleared) a competing higher priority thread can "steal"
the mutex which puts the woken up thread back on the waiters list.
The pending-ownership optimization is especially important for the
uninterrupted workflow of high-prio tasks which repeatedly
takes/releases locks that have lower-prio waiters. Without this
optimization the higher-prio thread would ping-pong to the lower-prio
task [because at unlock time we always assign a new owner].
(*) The "mutex has waiters" bit gets set to take the lock. If the lock
doesn't already have an owner, this bit is quickly cleared if there are
no waiters. So this is a transitional state to synchronize with looking
at the owner field of the mutex and the mutex owner releasing the lock.