Most garbage collection work can proceed concurrently with the client or mutator Java threads. But in certain places, for example while tracing from thread stacks, the garbage collector needs to ensure that Java data processed by the collector is consistent and complete. At these points, the mutators should not hold references to the heap that are invisible to the garbage collector. And they should not be modifying the data that is visible to the collector.
Logically, the collector and mutator share a reader-writer lock on the Java heap and associated data structures. Mutators hold the lock in reader or shared mode while running Java code or touching heap-related data structures. The collector holds the lock in writer or exclusive mode while it needs the heap data structures to be stable. However, this reader-writer lock has a very customized implementation that also provides additional facilities, such as the ability to exclude only a single thread, so that we can specifically examine its heap references.
In order to ensure consistency of the Java data, the compiler inserts “suspend points”, sometimes also called “safe points” into the code. These allow a thread to respond to external requests.
Whenever a thread is runnable, i.e. whenever a thread logically holds the mutator lock in shared mode, it is expected to regularly execute such a suspend point, and check for pending requests. They are currently implemented by setting a flag in the thread structure[^1], which is then explicitly tested by the compiler-generated code.
A thread responds to suspend requests only when it is “runnable”, i.e. logically running Java code. When it runs native code, or is blocked in a kernel call, it logically releases the mutator lock. When the garbage collector needs mutator cooperation, and the thread is not runnable, it is assured that the mutator is not touching Java data, and hence the collector can safely perform the required action itself, on the mutator thread's behalf.
Normally, when a thread makes a JNI call, it is not considered runnable while executing native code. This makes the transitions to and from running native JNI code somewhat expensive (see below). But these transitions are necessary to ensure that such code, which does not execute “suspend points”, and can thus not cooperate with the GC, doesn't delay GC completion. @FastNative and @CriticalNative calls avoid these transitions, instead allowing the thread to remain “runnable”, at the expense of potentially delaying GC operations for the duration of the call.
Although we say that a thread is “suspended” when it is not running Java code, it may in fact still be running native code and touching data structures that are not considered “Java data”. This distinction can be a fine line. For example, a Java thread blocked on a Java monitor will normally be “suspended” and blocked on a mutex contained in the monitor data structure. But it may wake up for reasons beyond ARTs control, which will normally result in touching the mutex. The monitor code must be quite careful to ensure that this does not cause problems, especially if the ART runtime was shut down in the interim and the monitor data structure has been reclaimed.
When a thread changes between running Java and native code, it has to correspondingly change its state between “runnable” and one of several other states, all of which are considered to be “suspended” for our purposes. When a Java thread starts to execute native code, and may thus not respond promptly to suspend requests, it will normally create an object of type ScopedThreadSuspension. ScopedThreadSuspension's constructor changes state to the “suspended” state given as an argument, logically releasing the mutator lock and promising to no longer touch Java data structures. It also handles any pending suspension requests that slid in just before it changed state.
Conversely, ScopedThreadSuspension‘s destructor waits until the GC has finished any actions it is currently performing on the thread’s behalf and effectively released the mutator exclusive lock, and then returns to runnable state, re-acquiring the mutator lock.
Occasionally a thread running native code needs to temporarily again access Java data structures, performing the above transitions in the opposite order. ScopedObjectAccess is a similar RAII object whose constructor and destructor perform those transitions in the reverse order from ScopedThreadSuspension.
The mutator lock is not implemented as a conventional mutex. But it plays by the rules of our normal static thread-safety analysis. Thus a function that is expected to be called in runnable state, with the ability to access Java data, should be annotated with REQUIRES_SHARED(Locks::mutator_lock_).
There is an explicit mutator_lock_ object, of type MutatorMutex. MutatorMutex is seemingly a minor refinement of ReaderWriterMutex, but it is used entirely differently. It is acquired explicitly by clients that need to hold it exclusively, and in a small number of cases, it is acquired in shared mode, e.g. via SharedTryLock(), or by the GC itself. However, more commonly MutatorMutex::TransitionFromSuspendedToRunnable(), is used to logically acquire the mutator mutex, e.g. as part of ScopedObjectAccess construction.
TransitionFromSuspendedToRunnable() does not physically acquire the ReaderWriterMutex in shared mode. Thus any thread acquiring the lock in exclusive mode must, in addition, explicitly arrange for mutator threads to be suspended via the thread suspension mechanism, and then make them runnable again on release.
Logically the mutator lock is held in shared/reader mode if either the underlying reader-writer lock is held in shared mode, or if a mutator is in runnable state. These two ways of holding the mutator mutex are not equivalent: In particular, we rely on the garbage collector never actually entering a “runnable” state while active (see below). However, it often runs with the explicit mutator mutex in shared mode, thus blocking others from acquiring it in exclusive mode.
Suspend point checks enable three kinds of communication with mutator threads:
Checkpoints : Checkpoint requests are used to get a thread to perform an action on our behalf. RequestCheckpoint() asks a specific thread to execute the closure supplied as an argument at its leisure. RequestSynchronousCheckpoint() in addition waits for the thread to complete running the closure, and handles suspended threads by running the closure on their behalf. In addition to these functions provided by Thread, ThreadList provides the RunCheckpoint() function that runs a checkpoint function on behalf of each thread, either by using RequestCheckpoint() to run it inside a running thread, or by ensuring that a suspended thread stays suspended, and then running the function on its behalf. RunCheckpoint() does not wait for completion of the function calls triggered by the resulting RequestCheckpoint() invocations.
Empty checkpoints : ThreadList provides RunEmptyCheckpoint(), which waits until all threads have either passed a suspend point, or have been suspended. This ensures that no thread is still executing Java code inside the same suspend-point-delimited code interval it was executing before the call. For example, a read-barrier started before a RunEmptyCheckpoint() call will have finished before the call returns.
Thread suspension : ThreadList provides a number of SuspendThread...() calls and a SuspendAll() call to suspend one or all threads until they are resumed by Resume() or ResumeAll(). The Suspend... calls guarantee that the target thread(s) are suspended (again, only in the sense of not running Java code) when the call returns.
It is easy to deadlock while attempting to run checkpoints, or suspending threads. In particular, we need to avoid situations in which we cannot suspend a thread because it is blocked, directly, or indirectly, on the GC completing its task. Deadlocks are avoided as follows:
Mutator lock ordering The mutator lock participates in the normal ART lock ordering hierarchy, as though it were a regular lock. See base/locks.h for the hierarchy. In particular, only locks at or below level kPostMutatorTopLockLevel may be acquired after acquiring the mutator lock, e.g. inside the scope of a ScopedObjectAccess. Similarly only locks at level strictly above kMutatatorLock may be held while acquiring the mutator lock, e.g. either by starting a ScopedObjectAccess, or ending a ScopedThreadSuspension.
This ensures that code that uses purely mutexes and threads state changes cannot deadlock: Since we always wait on a lower-level lock, the holder of the lowest-level lock can always progress. An attempt to initiate a checkpoint or to suspend another thread must also be treated as an acquisition of the mutator lock: A thread that is waiting for a lock before it can respond to the request is itself holding the mutator lock, and can only be blocked on lower-level locks. And acquisition of those can never depend on acquiring the mutator lock.
Checkpoints Running a checkpoint in a thread requires suspending that thread for the duration of the checkpoint, or running the checkpoint on the threads behalf while that thread is blocked from executing Java code. In the former case, the checkpoint code is run from CheckSuspend, which requires the mutator lock, so checkpoint code may only acquire mutexes at or below level kPostMutatorTopLockLevel. But that is not sufficient.
No matter whether the checkpoint is run in the target thread, or on its behalf, the target thread is effectively suspended and prevented from running Java code. However the target may hold arbitrary Java monitors, which it can no longer release. This may also prevent higher level mutexes from getting released. Thus checkpoint code should only acquire mutexes at level kPostMonitorLock or below.
Waiting This becomes much more problematic when we wait for something other than a lock. Waiting for something that may depend on the GC, while holding the mutator lock, can potentially lead to deadlock, since it will prevent the waiting thread from participating in GC checkpoints. Waiting while holding a lower-level lock like thread_list_lock_ is similarly unsafe in general, since a runnable thread may not respond to checkpoints until it acquires thread_list_lock_. In general, waiting for a condition variable while holding an unrelated lock is problematic, and these are specific instances of that general problem.
We do currently provide WaitHoldingLocks, and it is sometimes used with low-level locks held. But such code must somehow ensure that such waits eventually terminate without deadlock.
One common use of WaitHoldingLocks is to wait for weak reference processing. Special rules apply to avoid deadlocks in this case: Such waits must start after weak reference processing is disabled; the GC may not issue further nonempty checkpoints or suspend requests until weak reference processing has been reenabled, and threads have been notified. Thus the waiting thread's inability to respond to nonempty checkpoints and suspend requests cannot directly block the GC. Non-GC checkpoint or suspend requests that target a thread waiting on reference processing will block until reference processing completes.
Consider a case in which thread W1 waits on reference processing, while holding a low-level mutex M. Thread W2 holds the mutator lock and waits on M. We avoid a situation in which the GC needs to suspend or checkpoint W2 by briefly stopping the world to disable weak reference access. During the stop-the-world phase, W1 cannot yet be waiting for weak-reference access. Thus there is no danger of deadlock while entering this phase. After this phase, there is no need for W2 to suspend or execute a nonempty checkpoint. If we replaced the stop-the-world phase by a checkpoint, W2 could receive the checkpoint request too late, and be unable to respond.
Empty checkpoints can continue to occur during reference processing. Reference processing wait loops explicitly handle empty checkpoints, and an empty checkpoint request notifies the condition variable used to wait for reference processing, after acquiring reference_processor_lock_. This means that empty checkpoints do not preclude client threads from being in the middle of an operation that involves a weak reference access, while nonempty checkpoints do.
Suspending the GC Under unusual conditions, the GC can run on any thread. This means that when thread A suspends thread B for some other reason, Thread B might be running the garbage collector and conceivably thus cause it to block. This would be very deadlock prone. If Thread A allocates while Thread B is suspended in the GC, and the allocation requires the GC's help to complete, we deadlock.
Thus we ensure that the GC, together with anything else that can block GCs, cannot be blocked for thread suspension requests. This is accomplished by ensuring that it always appears to be in a suspended thread state. Since we only check for suspend requests when entering the runnable state, suspend requests go unnoticed until the GC completes. It may physically acquire and release the actual mutator_lock_ in either shared or exclusive mode.
Thread suspension is initiated by a registered thread, except that, for testing purposes, SuspendAll may be invoked with self == nullptr. We never suspend the initiating thread, explicitly exclusing it from SuspendAll(), and failing SuspendThreadBy...() requests to that effect.
The suspend calls invoke IncrementSuspendCount() to increment the thread suspend count for each thread. That adds a “suspend barrier” (atomic counter) to the per-thread list of such counters to decrement. It normally sets the kSuspendRequest (“should enter safepoint handler”) and kActiveSuspendBarrier (“need to notify us when suspended”) flags.
After setting these two flags, we check whether the thread is suspended and kSuspendRequest is still set. Since the thread is already suspended, it cannot be expected to respond to “pass the suspend barrier” (decrement the atomic counter) in a timely fashion. Hence we do so on its behalf. This decrements the “barrier” and removes it from the thread‘s list of barriers to decrement, and clears kActiveSuspendBarrier. kSuspendRequest remains to ensure the thread doesn’t prematurely return to runnable state.
If SuspendAllInternal() does not immediately see a suspended state, then it is up to the target thread to decrement the suspend barrier. TransitionFromRunnableToSuspended() calls TransitionToSuspendedAndRunCheckpoints(), which changes the thread state and returns. TransitionFromRunnableToSuspended() then calls CheckActiveSuspendBarriers() to check for the kActiveSuspendBarrier flag and decrement the suspend barrier if set.
The suspend_count_lock_ is not consistently held in the target thread during this process. Thus correctness in resolving the race between a suspension-requesting thread and a target thread voluntarily suspending relies on first requesting suspension, and then checking whether the target is already suspended, The detailed correctness argument is given in a comment inside SuspendAllInternal(). This also ensures that the barrier cannot be decremented after the stack frame holding the barrier goes away.
This relies on the fact that the two stores in the two threads to the state and kActiveSuspendBarrier flag are ordered with respect to the later loads. That's guaranteed, since they are all stored in a single atomic<>. Thus even relaxed accesses are OK.
The actual suspend barrier representation still varies between SuspendAll() and SuspendThreadBy...(). The former relies on the fact that only one such barrier can be in use at a time, while the latter maintains a linked list of active suspend barriers for each target thread, relying on the fact that each one can appear on the list of only one thread, and we can thus use list nodes allocated in the stack frames of requesting threads.
Avoiding suspension cycles
Any thread can issue a SuspendThreadByPeer(), SuspendThreadByThreadId() or SuspendAll() request. But if Thread A increments Thread B‘s suspend count while Thread B increments Thread A’s suspend count, and they then both suspend during a subsequent thread transition, we're deadlocked.
For single-thread suspension requests, we refuse to initiate a suspend request from a registered thread that is also being asked to suspend (i.e. the suspend count is nonzero). Instead the requestor waits for that condition to change. This means that we cannot create a cycle in which each thread has asked to suspend the next one, and thus no thread can progress. The required atomicity of the requestor suspend count check with setting the suspend count of the target(s) target is ensured by holding suspend_count_lock_.
For SuspendAll(), we enforce a requirement that at most one SuspendAll() request is running at one time. We also set the kSuspensionImmune thread flag to prevent a single thread suspension of a thread currently between SuspendAll() and ResumeAll() calls. Thus once a SuspendAll() call starts, it will complete before it can be affected by suspension requests from other threads.
Logically when a state transitions to state kRunnable, it acquires the mutator lock (in shared mode). When it changes its state back to suspended, it releases that lock. These must enforce proper happens-before ordering with respect to a thread acquiring the mutator lock in exclusive mode. Thus changing the thread state to suspended normally requires a release operation, to ensure visibility to a thread wishing to “suspend”. Conversely, when we change the state back to kRunnable in TransitionFromSuspendedToRunnable we do so with an acquire compare-exchange operation to check that no suspension request remains.
Any suspending thread should clear kSuspendRequest with a release operation. This, together with the acquire compare-exchange when returning to runnable state, ensure that suspend-count decrements happen-before a thread becomes runnable again.
Flags are often set and tested with relaxed ordering, even when they may communicate a request to another thread. The reason this is safe varies. Generally the request is conformed by some other better synchronized interaction, which establishes the necessary ordering. In particular:
kCheckPointRequest : See below. The call happens-before checkpoint execution since checkpoint_function accesses are guarded by a lock, as are updates to the flag. Checkpoint completion ordering is ensured by waiting for both ThreadList::RunCheckpoint() to complete, and for a barrier indicating completion in “runnable” threads.
kEmptyCheckpointRequest : Currently (12/2024) in flux. See below and b/382722942 . We currently use acquire/release ordering to access this flag in some places to partially mitigate memory ordering issues here.
kActiveSuspendBarrier : Changes are protected by thread_suspend_count_lock_, and PassActiveSuspendBarriers rechecks it while holding that lock.
kPendingFlipFunction : Set while holding thread_list_lock_, but this lock is not consistently held while checking the flag, notably in EnsureFlipFunctionStarted(). We need acquire/release ordering to make sure that the requestor happens-before flip function execution, and the flip function is properly visible, among other things. We still commonly read the flag using a relaxed operation, but we confirm it with an acquire compare-exchange operation in EnsureFlipFunctionStarted() before actiong on it.
kRunningFlipFunction : Cleared with release ordering, and read with acquire ordering by WaitForFlipFunction and its callers, thus ensuring that flip function execution happens-before completion of WaitForFlipFunction.
kSuspensionImmune : Guarded by thread_suspend_count_lock_.
Checkpoints expose additional tricky issues, since they are also used to ensure that certain global state changes have propagated to all running Java threads.
ThreadList::RunCheckpoint() guarantees that “if at point X RunCheckpoint() has returned, and all checkpoints have been properly observed to have completed, then every thread has executed a code sequence S during which it remained in a suspended state, such that the call to RunCheckpoint happens-before the end of S, and the beginning of S happened before X.”
For each thread T, we attempt to atomically install the kCheckpointRequest flag while ensuring that T is runnable. If this succeeds, the fact that tlsPtr_.checkpoint_function is protected by thread_suspend_count_lock_, and checkpoint_function is written by the requestor, and read by T at a suspend point, ensures that the call happens-before S. Normally a barrier ensures that S happens-before X .
If we are unable to do so for a thread T because we find it suspended, we run checkpoint_function ourselves. Before running it we set kSuspendRequest (with a release store) while T is in S, preventing T from terminating S by becoming runnable, after an acquire load of the flags. This ensures that the RunCheckpoint() call happens-before the end of S. Since we learned of T's suspended state via an acquire load of its state, which it stored with a release store, we know that the beginning of S happens-before we return from RunCheckpoint(), and hence before X .
The case of empty checkpoints is worth highlighting. We have traditionally optimized that by avoiding ever suspending the affected threads. This appears correct if all operations are sequentially consistent. But once memory model issues are considered, it appears more expensive to do fully correctly than it is to forego the optimization, as we explain below:
We need to ensure that if Thread A performs some update U and then calls RunEmptyCheckpoint() then, when RunEmptyCheckpoint() returns, every Thread B will have passed a suspend point at which U became guaranteed visible to that thread. This could be ensured in different ways, depending on the observed state of Thread B:
Runnable : Use acquire/release ordering when setting and detecting the kEmptyCheckpointRequest flag, and then use a barrier to observe when thread B has done so. In this case, Thread B actually executes code at a suspend point. The acquire release ordering on kEmptyCheckpointRequest ensures that U happens-before the suspend point, and the barrier ensures that the suspend point happens-before the RunEmptyCheckpoint() return.
Suspended : In this case, we have no mechanism for the thread to execute synchronization operations, and things become trickier. Thread B is suspended, but it may restart at any time. Thread A wants to conclude that Thread B is already suspended at a suspend point, and thus it is safe to ignore Thread B. Effectively, it wants to perform the update U, read B's state, and continue, while B may change its state back to runnable, and then read the result of U. Both threads effectively perform a store (either U or to the state of B) and then read the other value. With only acquire/release ordering, this can result in A seeing the old suspended state, and B failing to see the update U, which is an incorrect outcome. : The canonical fix for this kind of store; load reordering problem is to make all operations sequentially consistent. There does not appear to be an easy way to limit this overhead to just empty checkpoint execution. Thus it appears to be better to forego this “optimization”.
[^1]: In the most recent versions of ART, compiler-generated code loads through the address at tlsPtr_.suspend_trigger. A thread suspension is requested by setting this to null, triggering a SIGSEGV, causing that thread to check for GC cooperation requests. The older mechanism instead sets an appropriate ThreadFlag entry to request suspension or a checkpoint. Note that the actual checkpoint function value is set, along with the flag, while holding suspend_count_lock_. If the target thread notices that a checkpoint is requested, it then acquires the suspend_count_lock_ to read the checkpoint function.