I know that we can use pthread_cond_init, pthread_cond_wait, and pthread_cond_broadcast to implement thread condition signaling in user space. But how can we implement it in a kernel module that is using kthread?
linux/mutex.h provides locking as discussed in this question. But it does not seem to be associated with any condition variables.
linex/wait.h seems to have condition waiting but it has no documentation so there are too many concerns around it. For example, wake_up_interruptable could have a race condition when used as suggested in examples with no mutex. It seems unlikely that the wait queues could be combined with mutex waits in a way that would implement semantics like pthread_cond_wait, because the thread scheduling functions implementing these different APIs need to be integrated to provide atomic semantics.
How can a kernel module get atomic sleep-and-unlock semantics, like in a normal user-space API, like pthread_cond_wait?
The specific semantics I need are (from the pthread_cond_wait* manpage):
These functions atomically release mutex and cause the calling thread
to block on the condition variable cond; atomically here means
"atomically with respect to access by another thread to the mutex and
then the condition variable". That is, if another thread is able to
acquire the mutex after the about-to-block thread has released it,
then a subsequent call to pthread_cond_broadcast() or
pthread_cond_signal() in that thread shall behave as if it were issued
after the about-to-block thread has blocked.
The important requirements are:
The condition is a shared structure, and can only be accessed when the mutex is held.
Atomicity prevents the condition from becoming true in between the mutex release and thread blocking. Without this, the thread could check the condition (when it's false), then release the mutex, then the condition becomes true (and a signal is generated and propagated), and then the thread falls asleep and misses the signal.
(Here is how pthread_cond_wait is used in a program).
Related
I'm reading APUE Chapter 12(3rd edition) and it says: we can set either PTHREAD_MUTEX_STALLED or PTHREAD_MUTEX_ROBUST to the mutex. But I don't think we need mutex with attribute PTHREAD_MUTEX_STALLED, mutex should always "robust" , so that we can be notified if the side which locked the mutex is dead. If mutex is "stalled", we will be suspending forever.
And I know that Windows' mutex is always be "robust" and we will be notified with error WAIT_ABANDONED if the side which locked the mutex is dead. So, in what kind of scenario, we have to use "stalled" mutex, not "robust" mutex?
Thanks for your attention.
I see the following reasons why stalled mutex exists:
If robust mutex is used then everytime you try to lock a mutex, you have to check for EOWNERDEAD. So it requires an additional check.
If EOWNERDEAD is returned by pthread_mutex_lock() then you probably need to check all the state of shared objects that are relevant to that mutex has to checked and the mutex state has to reinstated by calling pthread_mutex_lock().
It's the default mutex attribute. Hence, no need for the application to call:pthread_mutexattr_setrobust().
Historical: early pthread implementations didn't have robust mutexes.
So all the above mentioned additional checks are only required if an application thinks a thread might die unexpectedly while holding a mutex, which is not how most threaded applications are designed. So it's a decision for an application to make if default behaviour (stalled) is sufficient or robust mutexes are needed.
Scenario 1: release mutex then wait
Scenario 2: wait and then release mutex
Trying to understand conceptually what it does.
If the mutex were released before the calling thread is considered "blocked" on the condition variable, then another thread could lock the mutex, change the state that the predicate is based on, and call pthread_cond_signal without the waiting thread ever waking up (since it's not yet blocked). That's the problem.
Scenario 2, waiting then releasing the mutex, is internally how any real-world implementation has to work, since there's no such thing as an atomic implementation of the necessary behavior. But from the application's perspective, there's no way to observe the thread being part of the blocked set without the mutex also being released, so in the sense of the "abstract machine", it's atomic.
Edit: To go into more detail, the real-world implementation of a condition variable wait generally looks like:
Modify some internal state of the condition variable object such that the caller is considered to be part of the blocked set for it.
Unlock the mutex.
Perform a blocking wait operation, with the special property that it will return immediately if the state of the condition variable object from step 1 has changed due to a signal from any other thread.
Thus, the act of "blocking" is split between two steps, one of which happens before the mutex is unlocked (gaining membership in the blocked set) and the other of which happens after the mutex is unlocked (possibly sleeping and yielding control to other threads). It's this split that's able to make the "condition wait" operation "atomic" in the abstract machine.
Suppose a condition variable is used in a situation where the signaling thread modifies the state affecting the truth value of the predicate and calls pthread_cond_signal without holding the mutex associated with the condition variable? Is it true that this type of usage is always subject to race conditions where the signal may be missed?
To me, there seems to always be an obvious race:
Waiter evaluates the predicate as false, but before it can begin waiting...
Another thread changes state in a way that makes the predicate true.
That other thread calls pthread_cond_signal, which does nothing because there are no waiters yet.
The waiter thread enters pthread_cond_wait, unaware that the predicate is now true, and waits indefinitely.
But does this same kind of race condition always exist if the situation is changed so that either (A) the mutex is held while calling pthread_cond_signal, just not while changing the state, or (B) so that the mutex is held while changing the state, just not while calling pthread_cond_signal?
I'm asking from a standpoint of wanting to know if there are any valid uses of the above not-best-practices usages, i.e. whether a correct condition-variable implementation needs to account for such usages in avoiding race conditions itself, or whether it can ignore them because they're already inherently racy.
The fundamental race here looks like this:
THREAD A THREAD B
Mutex lock
Check state
Change state
Signal
cvar wait
(never awakens)
If we take a lock EITHER on the state change OR the signal, OR both, then we avoid this; it's not possible for both the state-change and the signal to occur while thread A is in its critical section and holding the lock.
If we consider the reverse case, where thread A interleaves into thread B, there's no problem:
THREAD A THREAD B
Change state
Mutex lock
Check state
( no need to wait )
Mutex unlock
Signal (nobody cares)
So there's no particular need for thread B to hold a mutex over the entire operation; it just need to hold the mutex for some, possible infinitesimally small interval, between the state change and signal. Of course, if the state itself requires locking for safe manipulation, then the lock must be held over the state change as well.
Finally, note that dropping the mutex early is unlikely to be a performance improvement in most cases. Requiring the mutex to be held reduces contention over the internal locks in the condition variable, and in modern pthreads implementations, the system can 'move' the waiting thread from waiting on the cvar to waiting on the mutex without waking it up (thus avoiding it waking up only to immediately block on the mutex).
As pointed out in the comments, dropping the mutex may improve performance in some cases, by reducing the number of syscalls needed. Then again it could also lead to extra contention on the condition variable's internal mutex. Hard to say. It's probably not worth worrying about in any case.
Note that the applicable standards require that pthread_cond_signal be safely callable without holding the mutex:
The pthread_cond_signal() or pthread_cond_broadcast() functions may be called by a thread whether or not it currently owns the mutex that threads calling pthread_cond_wait() or pthread_cond_timedwait() have associated with the condition variable during their waits [...]
This usually means that condition variables have an internal lock over their internal data structures, or otherwise use some very careful lock-free algorithm.
The state must be modified inside a mutex, if for no other reason than the possibility of spurious wake-ups, which would lead to the reader reading the state while the writer is in the middle of writing it.
You can call pthread_cond_signal anytime after the state is changed. It doesn't have to be inside the mutex. POSIX guarantees that at least one waiter will awaken to check the new state. More to the point:
Calling pthread_cond_signal doesn't guarantee that a reader will acquire the mutex first. Another writer might get in before a reader gets a chance to check the new status. Condition variables don't guarantee that readers immediately follow writers (After all, what if there are no readers?)
Calling it after releasing the lock is actually better, since you don't risk having the just-awoken reader immediately going back to sleep trying to acquire the lock that the writer is still holding.
EDIT: #DietrichEpp makes a good point in the comments. The writer must change the state in such a way that the reader can never access an inconsistent state. It can do so either by acquiring the mutex used in the condition-variable, as I indicate above, or by ensuring that all state-changes are atomic.
The answer is, there is a race, and to eliminate that race, you must do this:
/* atomic op outside of mutex, and then: */
pthread_mutex_lock(&m);
pthread_mutex_unlock(&m);
pthread_cond_signal(&c);
The protection of the data doesn't matter, because you don't hold the mutex when calling pthread_cond_signal anyway.
See, by locking and unlocking the mutex, you have created a barrier. During that brief moment when the signaler has the mutex, there is a certainty: no other thread has the mutex. This means no other thread is executing any critical regions.
This means that all threads are either about to get the mutex to discover the change you have posted, or else they have already found that change and ran off with it (releasing the mutex), or else have not found they are looking for and have atomically given up the mutex to gone to sleep (and are guaranteed to be waiting nicely on the condition).
Without the mutex lock/unlock, you have no synchronization. The signal will sometimes fire as threads which didn't see the changed atomic value are transitioning to their atomic sleep to wait for it.
So this is what the mutex does from the point of view of a thread which is signaling. You can get the atomicity of access from something else, but not the synchronization.
P.S. I have implemented this logic before. The situation was in the Linux kernel (using my own mutexes and condition variables).
In my situation, it was impossible for the signaler to hold the mutex for the atomic operation on shared data. Why? Because the signaler did the operation in user space, inside a buffer shared between the kernel and user, and then (in some situations) made a system call into the kernel to wake up a thread. User space simply made some modifications to the buffer, and then if some conditions were satisfied, it would perform an ioctl.
So in the ioctl call I did the mutex lock/unlock thing, and then hit the condition variable. This ensured that the thread would not miss the wake up related to that latest modification posted by user space.
At first I just had the condition variable signal, but it looked wrong without the involvement of the mutex, so I reasoned about the situation a little bit and realized that the mutex must simply be locked and unlocked to conform to the synchronization ritual which eliminates the lost wakeup.
I'm implementing pthread condition variables (based on Linux futexes) and I have an idea for avoiding the "stampede effect" on pthread_cond_broadcast with process-shared condition variables. For non-process-shared cond vars, futex requeue operations are traditionally (i.e. by NPTL) used to requeue waiters from the cond var's futex to the mutex's futex without waking them up, but this is in general impossible for process-shared cond vars, because pthread_cond_broadcast might not have a valid pointer to the associated mutex. In a worst case scenario, the mutex might not even be mapped in its memory space.
My idea for overcoming this issue is to have pthread_cond_broadcast only directly wake one waiter, and have that waiter perform the requeue operation when it wakes up, since it does have the needed pointer to the mutex.
Naturally there are a lot of ugly race conditions to consider if I pursue this approach, but if they can be overcome, are there any other reasons such an implementation would be invalid or undesirable? One potential issue I can think of that might not be able to be overcome is the race where the waiter (a separate process) responsible for the requeue gets killed before it can act, but it might be possible to overcome even this by putting the condvar futex in the robust mutex list so that the kernel performs a wake on it when the process dies.
There may be waiters belonging to multiple address spaces, each of which has mapped the mutex associated with the futex at a different address in memory. I'm not sure if FUTEX_REQUEUE is safe to use when the requeue point may not be mapped at the same address in all waiters; if it does then this isn't a problem.
There are other problems that won't be detected by robust futexes; for example, if your chosen waiter is busy in a signal handler, you could be kept waiting an arbitrarily long time. [As discussed in the comments, these are not an issue]
Note that with robust futexes, you must set the value of the futex & 0x3FFFFFFF to be the TID of the thread to be woken up; you must also set bit FUTEX_WAITERS on if you want a wakeup. This means that you must choose which thread to awaken from the broadcasting thread, or you will be unable to deal with thread death immediately after the FUTEX_WAKE. You'll also need to deal with the possibility of the thread dying immediately before the waker thread writes its TID into the state variable - perhaps having a 'pending master' field that is also registered in the robust mutex system would be a good idea.
I see no reason why this can't work, then, as long as you make sure to deal with the thread exit issues carefully. That said, it may be best to simply define in the kernel an extension to FUTEX_WAIT that takes a requeue point and comparison value as an argument, and let the kernel handle this in a simple, race-free manner.
I just don't see why you assume that the corresponding mutex might not be known. It is clearly stated
The effect of using more than one mutex for concurrent
pthread_cond_timedwait() or pthread_cond_wait() operations on the same
condition variable
is undefined; that is, a condition variable becomes bound to
a unique mutex when a thread waits on the condition variable, and this
(dynamic)
binding shall end when the wait returns.
So even for process shared mutexes and conditions this must hold, and any user space process must always have mapped the same and unique mutex that is associated to the condition.
Allowing users to associate different mutexes to a condition at the same time is nothing that I would support.
Suppose there are two threads, the main thread and say thread B(created by main). If B acquired a mutex(say pthread_mutex) and it has called pthread_exit without unlocking the lock. So what happens to the mutex? Does it become free?
nope. The mutex remaines locked. What actually happens to such a lock depends on its type, You can read about that here or here
If you created a robust mutex by setting up the right attributes before calling pthread_mutex_init, the mutex will enter a special state when the thread that holds the lock terminates, and the next thread to attempt to acquire the mutex will obtain an error of EOWNERDEAD. It is then responsible for cleaning up whatever state the mutex protects and calling pthread_mutex_consistent to make the mutex usable again, or calling pthread_mutex_unlock (which will make the mutex permanently unusable; further attempts to use it will return ENOTRECOVERABLE).
For non-robust mutexes, the mutex is permanently unusable if the thread that locked it terminates without unlocking it. Per the standard (see the resolution to issue 755 on the Austin Group tracker), the mutex remains locked and its formal ownership continues to belong to the thread that exited, and any thread that attempts to lock it will deadlock. If another thread attempts to unlock it, that's normally undefined behavior, unless the mutex was created with the PTHREAD_MUTEX_ERRORCHECK attribute, in which case an error will be returned.
On the other hand, many (most?) real-world implementations don't actually follow the requirements of the standard. An attempt to lock or unlock the mutex from another thread might spuriously succeed, since the thread id (used to track ownership) might have been reused and may now refer to a different thread (possibly the one making the new lock/unlock request). At least glibc's NPTL is known to exhibit this behavior.