Mercurial > hg > truffle
view src/share/vm/runtime/mutex.cpp @ 15713:b68907673c8e
[inlining] shorter and equally informative, logNotInlined vs logNotInlinedMethod
author | Miguel Garcia <miguel.m.garcia@oracle.com> |
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date | Fri, 16 May 2014 15:28:00 +0200 |
parents | 51c6ce89d4dd |
children | 89152779163c |
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/* * Copyright (c) 1998, 2012, Oracle and/or its affiliates. All rights reserved. * DO NOT ALTER OR REMOVE COPYRIGHT NOTICES OR THIS FILE HEADER. * * This code is free software; you can redistribute it and/or modify it * under the terms of the GNU General Public License version 2 only, as * published by the Free Software Foundation. * * This code 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 * version 2 for more details (a copy is included in the LICENSE file that * accompanied this code). * * You should have received a copy of the GNU General Public License version * 2 along with this work; if not, write to the Free Software Foundation, * Inc., 51 Franklin St, Fifth Floor, Boston, MA 02110-1301 USA. * * Please contact Oracle, 500 Oracle Parkway, Redwood Shores, CA 94065 USA * or visit www.oracle.com if you need additional information or have any * questions. * */ #include "precompiled.hpp" #include "runtime/mutex.hpp" #include "runtime/osThread.hpp" #include "runtime/thread.inline.hpp" #include "utilities/events.hpp" #ifdef TARGET_OS_FAMILY_linux # include "mutex_linux.inline.hpp" #endif #ifdef TARGET_OS_FAMILY_solaris # include "mutex_solaris.inline.hpp" #endif #ifdef TARGET_OS_FAMILY_windows # include "mutex_windows.inline.hpp" #endif #ifdef TARGET_OS_FAMILY_bsd # include "mutex_bsd.inline.hpp" #endif // o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o // // Native Monitor-Mutex locking - theory of operations // // * Native Monitors are completely unrelated to Java-level monitors, // although the "back-end" slow-path implementations share a common lineage. // See objectMonitor:: in synchronizer.cpp. // Native Monitors do *not* support nesting or recursion but otherwise // they're basically Hoare-flavor monitors. // // * A thread acquires ownership of a Monitor/Mutex by CASing the LockByte // in the _LockWord from zero to non-zero. Note that the _Owner field // is advisory and is used only to verify that the thread calling unlock() // is indeed the last thread to have acquired the lock. // // * Contending threads "push" themselves onto the front of the contention // queue -- called the cxq -- with CAS and then spin/park. // The _LockWord contains the LockByte as well as the pointer to the head // of the cxq. Colocating the LockByte with the cxq precludes certain races. // // * Using a separately addressable LockByte allows for CAS:MEMBAR or CAS:0 // idioms. We currently use MEMBAR in the uncontended unlock() path, as // MEMBAR often has less latency than CAS. If warranted, we could switch to // a CAS:0 mode, using timers to close the resultant race, as is done // with Java Monitors in synchronizer.cpp. // // See the following for a discussion of the relative cost of atomics (CAS) // MEMBAR, and ways to eliminate such instructions from the common-case paths: // -- http://blogs.sun.com/dave/entry/biased_locking_in_hotspot // -- http://blogs.sun.com/dave/resource/MustangSync.pdf // -- http://blogs.sun.com/dave/resource/synchronization-public2.pdf // -- synchronizer.cpp // // * Overall goals - desiderata // 1. Minimize context switching // 2. Minimize lock migration // 3. Minimize CPI -- affinity and locality // 4. Minimize the execution of high-latency instructions such as CAS or MEMBAR // 5. Minimize outer lock hold times // 6. Behave gracefully on a loaded system // // * Thread flow and list residency: // // Contention queue --> EntryList --> OnDeck --> Owner --> !Owner // [..resident on monitor list..] // [...........contending..................] // // -- The contention queue (cxq) contains recently-arrived threads (RATs). // Threads on the cxq eventually drain into the EntryList. // -- Invariant: a thread appears on at most one list -- cxq, EntryList // or WaitSet -- at any one time. // -- For a given monitor there can be at most one "OnDeck" thread at any // given time but if needbe this particular invariant could be relaxed. // // * The WaitSet and EntryList linked lists are composed of ParkEvents. // I use ParkEvent instead of threads as ParkEvents are immortal and // type-stable, meaning we can safely unpark() a possibly stale // list element in the unlock()-path. (That's benign). // // * Succession policy - providing for progress: // // As necessary, the unlock()ing thread identifies, unlinks, and unparks // an "heir presumptive" tentative successor thread from the EntryList. // This becomes the so-called "OnDeck" thread, of which there can be only // one at any given time for a given monitor. The wakee will recontend // for ownership of monitor. // // Succession is provided for by a policy of competitive handoff. // The exiting thread does _not_ grant or pass ownership to the // successor thread. (This is also referred to as "handoff" succession"). // Instead the exiting thread releases ownership and possibly wakes // a successor, so the successor can (re)compete for ownership of the lock. // // Competitive handoff provides excellent overall throughput at the expense // of short-term fairness. If fairness is a concern then one remedy might // be to add an AcquireCounter field to the monitor. After a thread acquires // the lock it will decrement the AcquireCounter field. When the count // reaches 0 the thread would reset the AcquireCounter variable, abdicate // the lock directly to some thread on the EntryList, and then move itself to the // tail of the EntryList. // // But in practice most threads engage or otherwise participate in resource // bounded producer-consumer relationships, so lock domination is not usually // a practical concern. Recall too, that in general it's easier to construct // a fair lock from a fast lock, but not vice-versa. // // * The cxq can have multiple concurrent "pushers" but only one concurrent // detaching thread. This mechanism is immune from the ABA corruption. // More precisely, the CAS-based "push" onto cxq is ABA-oblivious. // We use OnDeck as a pseudo-lock to enforce the at-most-one detaching // thread constraint. // // * Taken together, the cxq and the EntryList constitute or form a // single logical queue of threads stalled trying to acquire the lock. // We use two distinct lists to reduce heat on the list ends. // Threads in lock() enqueue onto cxq while threads in unlock() will // dequeue from the EntryList. (c.f. Michael Scott's "2Q" algorithm). // A key desideratum is to minimize queue & monitor metadata manipulation // that occurs while holding the "outer" monitor lock -- that is, we want to // minimize monitor lock holds times. // // The EntryList is ordered by the prevailing queue discipline and // can be organized in any convenient fashion, such as a doubly-linked list or // a circular doubly-linked list. If we need a priority queue then something akin // to Solaris' sleepq would work nicely. Viz., // -- http://agg.eng/ws/on10_nightly/source/usr/src/uts/common/os/sleepq.c. // -- http://cvs.opensolaris.org/source/xref/onnv/onnv-gate/usr/src/uts/common/os/sleepq.c // Queue discipline is enforced at ::unlock() time, when the unlocking thread // drains the cxq into the EntryList, and orders or reorders the threads on the // EntryList accordingly. // // Barring "lock barging", this mechanism provides fair cyclic ordering, // somewhat similar to an elevator-scan. // // * OnDeck // -- For a given monitor there can be at most one OnDeck thread at any given // instant. The OnDeck thread is contending for the lock, but has been // unlinked from the EntryList and cxq by some previous unlock() operations. // Once a thread has been designated the OnDeck thread it will remain so // until it manages to acquire the lock -- being OnDeck is a stable property. // -- Threads on the EntryList or cxq are _not allowed to attempt lock acquisition. // -- OnDeck also serves as an "inner lock" as follows. Threads in unlock() will, after // having cleared the LockByte and dropped the outer lock, attempt to "trylock" // OnDeck by CASing the field from null to non-null. If successful, that thread // is then responsible for progress and succession and can use CAS to detach and // drain the cxq into the EntryList. By convention, only this thread, the holder of // the OnDeck inner lock, can manipulate the EntryList or detach and drain the // RATs on the cxq into the EntryList. This avoids ABA corruption on the cxq as // we allow multiple concurrent "push" operations but restrict detach concurrency // to at most one thread. Having selected and detached a successor, the thread then // changes the OnDeck to refer to that successor, and then unparks the successor. // That successor will eventually acquire the lock and clear OnDeck. Beware // that the OnDeck usage as a lock is asymmetric. A thread in unlock() transiently // "acquires" OnDeck, performs queue manipulations, passes OnDeck to some successor, // and then the successor eventually "drops" OnDeck. Note that there's never // any sense of contention on the inner lock, however. Threads never contend // or wait for the inner lock. // -- OnDeck provides for futile wakeup throttling a described in section 3.3 of // See http://www.usenix.org/events/jvm01/full_papers/dice/dice.pdf // In a sense, OnDeck subsumes the ObjectMonitor _Succ and ObjectWaiter // TState fields found in Java-level objectMonitors. (See synchronizer.cpp). // // * Waiting threads reside on the WaitSet list -- wait() puts // the caller onto the WaitSet. Notify() or notifyAll() simply // transfers threads from the WaitSet to either the EntryList or cxq. // Subsequent unlock() operations will eventually unpark the notifyee. // Unparking a notifee in notify() proper is inefficient - if we were to do so // it's likely the notifyee would simply impale itself on the lock held // by the notifier. // // * The mechanism is obstruction-free in that if the holder of the transient // OnDeck lock in unlock() is preempted or otherwise stalls, other threads // can still acquire and release the outer lock and continue to make progress. // At worst, waking of already blocked contending threads may be delayed, // but nothing worse. (We only use "trylock" operations on the inner OnDeck // lock). // // * Note that thread-local storage must be initialized before a thread // uses Native monitors or mutexes. The native monitor-mutex subsystem // depends on Thread::current(). // // * The monitor synchronization subsystem avoids the use of native // synchronization primitives except for the narrow platform-specific // park-unpark abstraction. See the comments in os_solaris.cpp regarding // the semantics of park-unpark. Put another way, this monitor implementation // depends only on atomic operations and park-unpark. The monitor subsystem // manages all RUNNING->BLOCKED and BLOCKED->READY transitions while the // underlying OS manages the READY<->RUN transitions. // // * The memory consistency model provide by lock()-unlock() is at least as // strong or stronger than the Java Memory model defined by JSR-133. // That is, we guarantee at least entry consistency, if not stronger. // See http://g.oswego.edu/dl/jmm/cookbook.html. // // * Thread:: currently contains a set of purpose-specific ParkEvents: // _MutexEvent, _ParkEvent, etc. A better approach might be to do away with // the purpose-specific ParkEvents and instead implement a general per-thread // stack of available ParkEvents which we could provision on-demand. The // stack acts as a local cache to avoid excessive calls to ParkEvent::Allocate() // and ::Release(). A thread would simply pop an element from the local stack before it // enqueued or park()ed. When the contention was over the thread would // push the no-longer-needed ParkEvent back onto its stack. // // * A slightly reduced form of ILock() and IUnlock() have been partially // model-checked (Murphi) for safety and progress at T=1,2,3 and 4. // It'd be interesting to see if TLA/TLC could be useful as well. // // * Mutex-Monitor is a low-level "leaf" subsystem. That is, the monitor // code should never call other code in the JVM that might itself need to // acquire monitors or mutexes. That's true *except* in the case of the // ThreadBlockInVM state transition wrappers. The ThreadBlockInVM DTOR handles // mutator reentry (ingress) by checking for a pending safepoint in which case it will // call SafepointSynchronize::block(), which in turn may call Safepoint_lock->lock(), etc. // In that particular case a call to lock() for a given Monitor can end up recursively // calling lock() on another monitor. While distasteful, this is largely benign // as the calls come from jacket that wraps lock(), and not from deep within lock() itself. // // It's unfortunate that native mutexes and thread state transitions were convolved. // They're really separate concerns and should have remained that way. Melding // them together was facile -- a bit too facile. The current implementation badly // conflates the two concerns. // // * TODO-FIXME: // // -- Add DTRACE probes for contended acquire, contended acquired, contended unlock // We should also add DTRACE probes in the ParkEvent subsystem for // Park-entry, Park-exit, and Unpark. // // -- We have an excess of mutex-like constructs in the JVM, namely: // 1. objectMonitors for Java-level synchronization (synchronizer.cpp) // 2. low-level muxAcquire and muxRelease // 3. low-level spinAcquire and spinRelease // 4. native Mutex:: and Monitor:: // 5. jvm_raw_lock() and _unlock() // 6. JVMTI raw monitors -- distinct from (5) despite having a confusingly // similar name. // // o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o // CASPTR() uses the canonical argument order that dominates in the literature. // Our internal cmpxchg_ptr() uses a bastardized ordering to accommodate Sun .il templates. #define CASPTR(a,c,s) intptr_t(Atomic::cmpxchg_ptr ((void *)(s),(void *)(a),(void *)(c))) #define UNS(x) (uintptr_t(x)) #define TRACE(m) { static volatile int ctr = 0 ; int x = ++ctr ; if ((x & (x-1))==0) { ::printf ("%d:%s\n", x, #m); ::fflush(stdout); }} // Simplistic low-quality Marsaglia SHIFT-XOR RNG. // Bijective except for the trailing mask operation. // Useful for spin loops as the compiler can't optimize it away. static inline jint MarsagliaXORV (jint x) { if (x == 0) x = 1|os::random() ; x ^= x << 6; x ^= ((unsigned)x) >> 21; x ^= x << 7 ; return x & 0x7FFFFFFF ; } static int Stall (int its) { static volatile jint rv = 1 ; volatile int OnFrame = 0 ; jint v = rv ^ UNS(OnFrame) ; while (--its >= 0) { v = MarsagliaXORV (v) ; } // Make this impossible for the compiler to optimize away, // but (mostly) avoid W coherency sharing on MP systems. if (v == 0x12345) rv = v ; return v ; } int Monitor::TryLock () { intptr_t v = _LockWord.FullWord ; for (;;) { if ((v & _LBIT) != 0) return 0 ; const intptr_t u = CASPTR (&_LockWord, v, v|_LBIT) ; if (v == u) return 1 ; v = u ; } } int Monitor::TryFast () { // Optimistic fast-path form ... // Fast-path attempt for the common uncontended case. // Avoid RTS->RTO $ coherence upgrade on typical SMP systems. intptr_t v = CASPTR (&_LockWord, 0, _LBIT) ; // agro ... if (v == 0) return 1 ; for (;;) { if ((v & _LBIT) != 0) return 0 ; const intptr_t u = CASPTR (&_LockWord, v, v|_LBIT) ; if (v == u) return 1 ; v = u ; } } int Monitor::ILocked () { const intptr_t w = _LockWord.FullWord & 0xFF ; assert (w == 0 || w == _LBIT, "invariant") ; return w == _LBIT ; } // Polite TATAS spinlock with exponential backoff - bounded spin. // Ideally we'd use processor cycles, time or vtime to control // the loop, but we currently use iterations. // All the constants within were derived empirically but work over // over the spectrum of J2SE reference platforms. // On Niagara-class systems the back-off is unnecessary but // is relatively harmless. (At worst it'll slightly retard // acquisition times). The back-off is critical for older SMP systems // where constant fetching of the LockWord would otherwise impair // scalability. // // Clamp spinning at approximately 1/2 of a context-switch round-trip. // See synchronizer.cpp for details and rationale. int Monitor::TrySpin (Thread * const Self) { if (TryLock()) return 1 ; if (!os::is_MP()) return 0 ; int Probes = 0 ; int Delay = 0 ; int Steps = 0 ; int SpinMax = NativeMonitorSpinLimit ; int flgs = NativeMonitorFlags ; for (;;) { intptr_t v = _LockWord.FullWord; if ((v & _LBIT) == 0) { if (CASPTR (&_LockWord, v, v|_LBIT) == v) { return 1 ; } continue ; } if ((flgs & 8) == 0) { SpinPause () ; } // Periodically increase Delay -- variable Delay form // conceptually: delay *= 1 + 1/Exponent ++ Probes; if (Probes > SpinMax) return 0 ; if ((Probes & 0x7) == 0) { Delay = ((Delay << 1)|1) & 0x7FF ; // CONSIDER: Delay += 1 + (Delay/4); Delay &= 0x7FF ; } if (flgs & 2) continue ; // Consider checking _owner's schedctl state, if OFFPROC abort spin. // If the owner is OFFPROC then it's unlike that the lock will be dropped // in a timely fashion, which suggests that spinning would not be fruitful // or profitable. // Stall for "Delay" time units - iterations in the current implementation. // Avoid generating coherency traffic while stalled. // Possible ways to delay: // PAUSE, SLEEP, MEMBAR #sync, MEMBAR #halt, // wr %g0,%asi, gethrtime, rdstick, rdtick, rdtsc, etc. ... // Note that on Niagara-class systems we want to minimize STs in the // spin loop. N1 and brethren write-around the L1$ over the xbar into the L2$. // Furthermore, they don't have a W$ like traditional SPARC processors. // We currently use a Marsaglia Shift-Xor RNG loop. Steps += Delay ; if (Self != NULL) { jint rv = Self->rng[0] ; for (int k = Delay ; --k >= 0; ) { rv = MarsagliaXORV (rv) ; if ((flgs & 4) == 0 && SafepointSynchronize::do_call_back()) return 0 ; } Self->rng[0] = rv ; } else { Stall (Delay) ; } } } static int ParkCommon (ParkEvent * ev, jlong timo) { // Diagnostic support - periodically unwedge blocked threads intx nmt = NativeMonitorTimeout ; if (nmt > 0 && (nmt < timo || timo <= 0)) { timo = nmt ; } int err = OS_OK ; if (0 == timo) { ev->park() ; } else { err = ev->park(timo) ; } return err ; } inline int Monitor::AcquireOrPush (ParkEvent * ESelf) { intptr_t v = _LockWord.FullWord ; for (;;) { if ((v & _LBIT) == 0) { const intptr_t u = CASPTR (&_LockWord, v, v|_LBIT) ; if (u == v) return 1 ; // indicate acquired v = u ; } else { // Anticipate success ... ESelf->ListNext = (ParkEvent *) (v & ~_LBIT) ; const intptr_t u = CASPTR (&_LockWord, v, intptr_t(ESelf)|_LBIT) ; if (u == v) return 0 ; // indicate pushed onto cxq v = u ; } // Interference - LockWord change - just retry } } // ILock and IWait are the lowest level primitive internal blocking // synchronization functions. The callers of IWait and ILock must have // performed any needed state transitions beforehand. // IWait and ILock may directly call park() without any concern for thread state. // Note that ILock and IWait do *not* access _owner. // _owner is a higher-level logical concept. void Monitor::ILock (Thread * Self) { assert (_OnDeck != Self->_MutexEvent, "invariant") ; if (TryFast()) { Exeunt: assert (ILocked(), "invariant") ; return ; } ParkEvent * const ESelf = Self->_MutexEvent ; assert (_OnDeck != ESelf, "invariant") ; // As an optimization, spinners could conditionally try to set ONDECK to _LBIT // Synchronizer.cpp uses a similar optimization. if (TrySpin (Self)) goto Exeunt ; // Slow-path - the lock is contended. // Either Enqueue Self on cxq or acquire the outer lock. // LockWord encoding = (cxq,LOCKBYTE) ESelf->reset() ; OrderAccess::fence() ; // Optional optimization ... try barging on the inner lock if ((NativeMonitorFlags & 32) && CASPTR (&_OnDeck, NULL, UNS(Self)) == 0) { goto OnDeck_LOOP ; } if (AcquireOrPush (ESelf)) goto Exeunt ; // At any given time there is at most one ondeck thread. // ondeck implies not resident on cxq and not resident on EntryList // Only the OnDeck thread can try to acquire -- contended for -- the lock. // CONSIDER: use Self->OnDeck instead of m->OnDeck. // Deschedule Self so that others may run. while (_OnDeck != ESelf) { ParkCommon (ESelf, 0) ; } // Self is now in the ONDECK position and will remain so until it // manages to acquire the lock. OnDeck_LOOP: for (;;) { assert (_OnDeck == ESelf, "invariant") ; if (TrySpin (Self)) break ; // CONSIDER: if ESelf->TryPark() && TryLock() break ... // It's probably wise to spin only if we *actually* blocked // CONSIDER: check the lockbyte, if it remains set then // preemptively drain the cxq into the EntryList. // The best place and time to perform queue operations -- lock metadata -- // is _before having acquired the outer lock, while waiting for the lock to drop. ParkCommon (ESelf, 0) ; } assert (_OnDeck == ESelf, "invariant") ; _OnDeck = NULL ; // Note that we current drop the inner lock (clear OnDeck) in the slow-path // epilog immediately after having acquired the outer lock. // But instead we could consider the following optimizations: // A. Shift or defer dropping the inner lock until the subsequent IUnlock() operation. // This might avoid potential reacquisition of the inner lock in IUlock(). // B. While still holding the inner lock, attempt to opportunistically select // and unlink the next ONDECK thread from the EntryList. // If successful, set ONDECK to refer to that thread, otherwise clear ONDECK. // It's critical that the select-and-unlink operation run in constant-time as // it executes when holding the outer lock and may artificially increase the // effective length of the critical section. // Note that (A) and (B) are tantamount to succession by direct handoff for // the inner lock. goto Exeunt ; } void Monitor::IUnlock (bool RelaxAssert) { assert (ILocked(), "invariant") ; // Conceptually we need a MEMBAR #storestore|#loadstore barrier or fence immediately // before the store that releases the lock. Crucially, all the stores and loads in the // critical section must be globally visible before the store of 0 into the lock-word // that releases the lock becomes globally visible. That is, memory accesses in the // critical section should not be allowed to bypass or overtake the following ST that // releases the lock. As such, to prevent accesses within the critical section // from "leaking" out, we need a release fence between the critical section and the // store that releases the lock. In practice that release barrier is elided on // platforms with strong memory models such as TSO. // // Note that the OrderAccess::storeload() fence that appears after unlock store // provides for progress conditions and succession and is _not related to exclusion // safety or lock release consistency. OrderAccess::release_store(&_LockWord.Bytes[_LSBINDEX], 0); // drop outer lock OrderAccess::storeload (); ParkEvent * const w = _OnDeck ; assert (RelaxAssert || w != Thread::current()->_MutexEvent, "invariant") ; if (w != NULL) { // Either we have a valid ondeck thread or ondeck is transiently "locked" // by some exiting thread as it arranges for succession. The LSBit of // OnDeck allows us to discriminate two cases. If the latter, the // responsibility for progress and succession lies with that other thread. // For good performance, we also depend on the fact that redundant unpark() // operations are cheap. That is, repeated Unpark()ing of the ONDECK thread // is inexpensive. This approach provides implicit futile wakeup throttling. // Note that the referent "w" might be stale with respect to the lock. // In that case the following unpark() is harmless and the worst that'll happen // is a spurious return from a park() operation. Critically, if "w" _is stale, // then progress is known to have occurred as that means the thread associated // with "w" acquired the lock. In that case this thread need take no further // action to guarantee progress. if ((UNS(w) & _LBIT) == 0) w->unpark() ; return ; } intptr_t cxq = _LockWord.FullWord ; if (((cxq & ~_LBIT)|UNS(_EntryList)) == 0) { return ; // normal fast-path exit - cxq and EntryList both empty } if (cxq & _LBIT) { // Optional optimization ... // Some other thread acquired the lock in the window since this // thread released it. Succession is now that thread's responsibility. return ; } Succession: // Slow-path exit - this thread must ensure succession and progress. // OnDeck serves as lock to protect cxq and EntryList. // Only the holder of OnDeck can manipulate EntryList or detach the RATs from cxq. // Avoid ABA - allow multiple concurrent producers (enqueue via push-CAS) // but only one concurrent consumer (detacher of RATs). // Consider protecting this critical section with schedctl on Solaris. // Unlike a normal lock, however, the exiting thread "locks" OnDeck, // picks a successor and marks that thread as OnDeck. That successor // thread will then clear OnDeck once it eventually acquires the outer lock. if (CASPTR (&_OnDeck, NULL, _LBIT) != UNS(NULL)) { return ; } ParkEvent * List = _EntryList ; if (List != NULL) { // Transfer the head of the EntryList to the OnDeck position. // Once OnDeck, a thread stays OnDeck until it acquires the lock. // For a given lock there is at most OnDeck thread at any one instant. WakeOne: assert (List == _EntryList, "invariant") ; ParkEvent * const w = List ; assert (RelaxAssert || w != Thread::current()->_MutexEvent, "invariant") ; _EntryList = w->ListNext ; // as a diagnostic measure consider setting w->_ListNext = BAD assert (UNS(_OnDeck) == _LBIT, "invariant") ; _OnDeck = w ; // pass OnDeck to w. // w will clear OnDeck once it acquires the outer lock // Another optional optimization ... // For heavily contended locks it's not uncommon that some other // thread acquired the lock while this thread was arranging succession. // Try to defer the unpark() operation - Delegate the responsibility // for unpark()ing the OnDeck thread to the current or subsequent owners // That is, the new owner is responsible for unparking the OnDeck thread. OrderAccess::storeload() ; cxq = _LockWord.FullWord ; if (cxq & _LBIT) return ; w->unpark() ; return ; } cxq = _LockWord.FullWord ; if ((cxq & ~_LBIT) != 0) { // The EntryList is empty but the cxq is populated. // drain RATs from cxq into EntryList // Detach RATs segment with CAS and then merge into EntryList for (;;) { // optional optimization - if locked, the owner is responsible for succession if (cxq & _LBIT) goto Punt ; const intptr_t vfy = CASPTR (&_LockWord, cxq, cxq & _LBIT) ; if (vfy == cxq) break ; cxq = vfy ; // Interference - LockWord changed - Just retry // We can see concurrent interference from contending threads // pushing themselves onto the cxq or from lock-unlock operations. // From the perspective of this thread, EntryList is stable and // the cxq is prepend-only -- the head is volatile but the interior // of the cxq is stable. In theory if we encounter interference from threads // pushing onto cxq we could simply break off the original cxq suffix and // move that segment to the EntryList, avoiding a 2nd or multiple CAS attempts // on the high-traffic LockWord variable. For instance lets say the cxq is "ABCD" // when we first fetch cxq above. Between the fetch -- where we observed "A" // -- and CAS -- where we attempt to CAS null over A -- "PQR" arrive, // yielding cxq = "PQRABCD". In this case we could simply set A.ListNext // null, leaving cxq = "PQRA" and transfer the "BCD" segment to the EntryList. // Note too, that it's safe for this thread to traverse the cxq // without taking any special concurrency precautions. } // We don't currently reorder the cxq segment as we move it onto // the EntryList, but it might make sense to reverse the order // or perhaps sort by thread priority. See the comments in // synchronizer.cpp objectMonitor::exit(). assert (_EntryList == NULL, "invariant") ; _EntryList = List = (ParkEvent *)(cxq & ~_LBIT) ; assert (List != NULL, "invariant") ; goto WakeOne ; } // cxq|EntryList is empty. // w == NULL implies that cxq|EntryList == NULL in the past. // Possible race - rare inopportune interleaving. // A thread could have added itself to cxq since this thread previously checked. // Detect and recover by refetching cxq. Punt: assert (UNS(_OnDeck) == _LBIT, "invariant") ; _OnDeck = NULL ; // Release inner lock. OrderAccess::storeload(); // Dekker duality - pivot point // Resample LockWord/cxq to recover from possible race. // For instance, while this thread T1 held OnDeck, some other thread T2 might // acquire the outer lock. Another thread T3 might try to acquire the outer // lock, but encounter contention and enqueue itself on cxq. T2 then drops the // outer lock, but skips succession as this thread T1 still holds OnDeck. // T1 is and remains responsible for ensuring succession of T3. // // Note that we don't need to recheck EntryList, just cxq. // If threads moved onto EntryList since we dropped OnDeck // that implies some other thread forced succession. cxq = _LockWord.FullWord ; if ((cxq & ~_LBIT) != 0 && (cxq & _LBIT) == 0) { goto Succession ; // potential race -- re-run succession } return ; } bool Monitor::notify() { assert (_owner == Thread::current(), "invariant") ; assert (ILocked(), "invariant") ; if (_WaitSet == NULL) return true ; NotifyCount ++ ; // Transfer one thread from the WaitSet to the EntryList or cxq. // Currently we just unlink the head of the WaitSet and prepend to the cxq. // And of course we could just unlink it and unpark it, too, but // in that case it'd likely impale itself on the reentry. Thread::muxAcquire (_WaitLock, "notify:WaitLock") ; ParkEvent * nfy = _WaitSet ; if (nfy != NULL) { // DCL idiom _WaitSet = nfy->ListNext ; assert (nfy->Notified == 0, "invariant") ; // push nfy onto the cxq for (;;) { const intptr_t v = _LockWord.FullWord ; assert ((v & 0xFF) == _LBIT, "invariant") ; nfy->ListNext = (ParkEvent *)(v & ~_LBIT); if (CASPTR (&_LockWord, v, UNS(nfy)|_LBIT) == v) break; // interference - _LockWord changed -- just retry } // Note that setting Notified before pushing nfy onto the cxq is // also legal and safe, but the safety properties are much more // subtle, so for the sake of code stewardship ... OrderAccess::fence() ; nfy->Notified = 1; } Thread::muxRelease (_WaitLock) ; if (nfy != NULL && (NativeMonitorFlags & 16)) { // Experimental code ... light up the wakee in the hope that this thread (the owner) // will drop the lock just about the time the wakee comes ONPROC. nfy->unpark() ; } assert (ILocked(), "invariant") ; return true ; } // Currently notifyAll() transfers the waiters one-at-a-time from the waitset // to the cxq. This could be done more efficiently with a single bulk en-mass transfer, // but in practice notifyAll() for large #s of threads is rare and not time-critical. // Beware too, that we invert the order of the waiters. Lets say that the // waitset is "ABCD" and the cxq is "XYZ". After a notifyAll() the waitset // will be empty and the cxq will be "DCBAXYZ". This is benign, of course. bool Monitor::notify_all() { assert (_owner == Thread::current(), "invariant") ; assert (ILocked(), "invariant") ; while (_WaitSet != NULL) notify() ; return true ; } int Monitor::IWait (Thread * Self, jlong timo) { assert (ILocked(), "invariant") ; // Phases: // 1. Enqueue Self on WaitSet - currently prepend // 2. unlock - drop the outer lock // 3. wait for either notification or timeout // 4. lock - reentry - reacquire the outer lock ParkEvent * const ESelf = Self->_MutexEvent ; ESelf->Notified = 0 ; ESelf->reset() ; OrderAccess::fence() ; // Add Self to WaitSet // Ideally only the holder of the outer lock would manipulate the WaitSet - // That is, the outer lock would implicitly protect the WaitSet. // But if a thread in wait() encounters a timeout it will need to dequeue itself // from the WaitSet _before it becomes the owner of the lock. We need to dequeue // as the ParkEvent -- which serves as a proxy for the thread -- can't reside // on both the WaitSet and the EntryList|cxq at the same time.. That is, a thread // on the WaitSet can't be allowed to compete for the lock until it has managed to // unlink its ParkEvent from WaitSet. Thus the need for WaitLock. // Contention on the WaitLock is minimal. // // Another viable approach would be add another ParkEvent, "WaitEvent" to the // thread class. The WaitSet would be composed of WaitEvents. Only the // owner of the outer lock would manipulate the WaitSet. A thread in wait() // could then compete for the outer lock, and then, if necessary, unlink itself // from the WaitSet only after having acquired the outer lock. More precisely, // there would be no WaitLock. A thread in in wait() would enqueue its WaitEvent // on the WaitSet; release the outer lock; wait for either notification or timeout; // reacquire the inner lock; and then, if needed, unlink itself from the WaitSet. // // Alternatively, a 2nd set of list link fields in the ParkEvent might suffice. // One set would be for the WaitSet and one for the EntryList. // We could also deconstruct the ParkEvent into a "pure" event and add a // new immortal/TSM "ListElement" class that referred to ParkEvents. // In that case we could have one ListElement on the WaitSet and another // on the EntryList, with both referring to the same pure Event. Thread::muxAcquire (_WaitLock, "wait:WaitLock:Add") ; ESelf->ListNext = _WaitSet ; _WaitSet = ESelf ; Thread::muxRelease (_WaitLock) ; // Release the outer lock // We call IUnlock (RelaxAssert=true) as a thread T1 might // enqueue itself on the WaitSet, call IUnlock(), drop the lock, // and then stall before it can attempt to wake a successor. // Some other thread T2 acquires the lock, and calls notify(), moving // T1 from the WaitSet to the cxq. T2 then drops the lock. T1 resumes, // and then finds *itself* on the cxq. During the course of a normal // IUnlock() call a thread should _never find itself on the EntryList // or cxq, but in the case of wait() it's possible. // See synchronizer.cpp objectMonitor::wait(). IUnlock (true) ; // Wait for either notification or timeout // Beware that in some circumstances we might propagate // spurious wakeups back to the caller. for (;;) { if (ESelf->Notified) break ; int err = ParkCommon (ESelf, timo) ; if (err == OS_TIMEOUT || (NativeMonitorFlags & 1)) break ; } // Prepare for reentry - if necessary, remove ESelf from WaitSet // ESelf can be: // 1. Still on the WaitSet. This can happen if we exited the loop by timeout. // 2. On the cxq or EntryList // 3. Not resident on cxq, EntryList or WaitSet, but in the OnDeck position. OrderAccess::fence() ; int WasOnWaitSet = 0 ; if (ESelf->Notified == 0) { Thread::muxAcquire (_WaitLock, "wait:WaitLock:remove") ; if (ESelf->Notified == 0) { // DCL idiom assert (_OnDeck != ESelf, "invariant") ; // can't be both OnDeck and on WaitSet // ESelf is resident on the WaitSet -- unlink it. // A doubly-linked list would be better here so we can unlink in constant-time. // We have to unlink before we potentially recontend as ESelf might otherwise // end up on the cxq|EntryList -- it can't be on two lists at once. ParkEvent * p = _WaitSet ; ParkEvent * q = NULL ; // classic q chases p while (p != NULL && p != ESelf) { q = p ; p = p->ListNext ; } assert (p == ESelf, "invariant") ; if (p == _WaitSet) { // found at head assert (q == NULL, "invariant") ; _WaitSet = p->ListNext ; } else { // found in interior assert (q->ListNext == p, "invariant") ; q->ListNext = p->ListNext ; } WasOnWaitSet = 1 ; // We were *not* notified but instead encountered timeout } Thread::muxRelease (_WaitLock) ; } // Reentry phase - reacquire the lock if (WasOnWaitSet) { // ESelf was previously on the WaitSet but we just unlinked it above // because of a timeout. ESelf is not resident on any list and is not OnDeck assert (_OnDeck != ESelf, "invariant") ; ILock (Self) ; } else { // A prior notify() operation moved ESelf from the WaitSet to the cxq. // ESelf is now on the cxq, EntryList or at the OnDeck position. // The following fragment is extracted from Monitor::ILock() for (;;) { if (_OnDeck == ESelf && TrySpin(Self)) break ; ParkCommon (ESelf, 0) ; } assert (_OnDeck == ESelf, "invariant") ; _OnDeck = NULL ; } assert (ILocked(), "invariant") ; return WasOnWaitSet != 0 ; // return true IFF timeout } // ON THE VMTHREAD SNEAKING PAST HELD LOCKS: // In particular, there are certain types of global lock that may be held // by a Java thread while it is blocked at a safepoint but before it has // written the _owner field. These locks may be sneakily acquired by the // VM thread during a safepoint to avoid deadlocks. Alternatively, one should // identify all such locks, and ensure that Java threads never block at // safepoints while holding them (_no_safepoint_check_flag). While it // seems as though this could increase the time to reach a safepoint // (or at least increase the mean, if not the variance), the latter // approach might make for a cleaner, more maintainable JVM design. // // Sneaking is vile and reprehensible and should be excised at the 1st // opportunity. It's possible that the need for sneaking could be obviated // as follows. Currently, a thread might (a) while TBIVM, call pthread_mutex_lock // or ILock() thus acquiring the "physical" lock underlying Monitor/Mutex. // (b) stall at the TBIVM exit point as a safepoint is in effect. Critically, // it'll stall at the TBIVM reentry state transition after having acquired the // underlying lock, but before having set _owner and having entered the actual // critical section. The lock-sneaking facility leverages that fact and allowed the // VM thread to logically acquire locks that had already be physically locked by mutators // but where mutators were known blocked by the reentry thread state transition. // // If we were to modify the Monitor-Mutex so that TBIVM state transitions tightly // wrapped calls to park(), then we could likely do away with sneaking. We'd // decouple lock acquisition and parking. The critical invariant to eliminating // sneaking is to ensure that we never "physically" acquire the lock while TBIVM. // An easy way to accomplish this is to wrap the park calls in a narrow TBIVM jacket. // One difficulty with this approach is that the TBIVM wrapper could recurse and // call lock() deep from within a lock() call, while the MutexEvent was already enqueued. // Using a stack (N=2 at minimum) of ParkEvents would take care of that problem. // // But of course the proper ultimate approach is to avoid schemes that require explicit // sneaking or dependence on any any clever invariants or subtle implementation properties // of Mutex-Monitor and instead directly address the underlying design flaw. void Monitor::lock (Thread * Self) { #ifdef CHECK_UNHANDLED_OOPS // Clear unhandled oops so we get a crash right away. Only clear for non-vm // or GC threads. if (Self->is_Java_thread()) { Self->clear_unhandled_oops(); } #endif // CHECK_UNHANDLED_OOPS debug_only(check_prelock_state(Self)); assert (_owner != Self , "invariant") ; assert (_OnDeck != Self->_MutexEvent, "invariant") ; if (TryFast()) { Exeunt: assert (ILocked(), "invariant") ; assert (owner() == NULL, "invariant"); set_owner (Self); return ; } // The lock is contended ... bool can_sneak = Self->is_VM_thread() && SafepointSynchronize::is_at_safepoint(); if (can_sneak && _owner == NULL) { // a java thread has locked the lock but has not entered the // critical region -- let's just pretend we've locked the lock // and go on. we note this with _snuck so we can also // pretend to unlock when the time comes. _snuck = true; goto Exeunt ; } // Try a brief spin to avoid passing thru thread state transition ... if (TrySpin (Self)) goto Exeunt ; check_block_state(Self); if (Self->is_Java_thread()) { // Horribile dictu - we suffer through a state transition assert(rank() > Mutex::special, "Potential deadlock with special or lesser rank mutex"); ThreadBlockInVM tbivm ((JavaThread *) Self) ; ILock (Self) ; } else { // Mirabile dictu ILock (Self) ; } goto Exeunt ; } void Monitor::lock() { this->lock(Thread::current()); } // Lock without safepoint check - a degenerate variant of lock(). // Should ONLY be used by safepoint code and other code // that is guaranteed not to block while running inside the VM. If this is called with // thread state set to be in VM, the safepoint synchronization code will deadlock! void Monitor::lock_without_safepoint_check (Thread * Self) { assert (_owner != Self, "invariant") ; ILock (Self) ; assert (_owner == NULL, "invariant"); set_owner (Self); } void Monitor::lock_without_safepoint_check () { lock_without_safepoint_check (Thread::current()) ; } // Returns true if thread succeceed [sic] in grabbing the lock, otherwise false. bool Monitor::try_lock() { Thread * const Self = Thread::current(); debug_only(check_prelock_state(Self)); // assert(!thread->is_inside_signal_handler(), "don't lock inside signal handler"); // Special case, where all Java threads are stopped. // The lock may have been acquired but _owner is not yet set. // In that case the VM thread can safely grab the lock. // It strikes me this should appear _after the TryLock() fails, below. bool can_sneak = Self->is_VM_thread() && SafepointSynchronize::is_at_safepoint(); if (can_sneak && _owner == NULL) { set_owner(Self); // Do not need to be atomic, since we are at a safepoint _snuck = true; return true; } if (TryLock()) { // We got the lock assert (_owner == NULL, "invariant"); set_owner (Self); return true; } return false; } void Monitor::unlock() { assert (_owner == Thread::current(), "invariant") ; assert (_OnDeck != Thread::current()->_MutexEvent , "invariant") ; set_owner (NULL) ; if (_snuck) { assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak"); _snuck = false; return ; } IUnlock (false) ; } // Yet another degenerate version of Monitor::lock() or lock_without_safepoint_check() // jvm_raw_lock() and _unlock() can be called by non-Java threads via JVM_RawMonitorEnter. // // There's no expectation that JVM_RawMonitors will interoperate properly with the native // Mutex-Monitor constructs. We happen to implement JVM_RawMonitors in terms of // native Mutex-Monitors simply as a matter of convenience. A simple abstraction layer // over a pthread_mutex_t would work equally as well, but require more platform-specific // code -- a "PlatformMutex". Alternatively, a simply layer over muxAcquire-muxRelease // would work too. // // Since the caller might be a foreign thread, we don't necessarily have a Thread.MutexEvent // instance available. Instead, we transiently allocate a ParkEvent on-demand if // we encounter contention. That ParkEvent remains associated with the thread // until it manages to acquire the lock, at which time we return the ParkEvent // to the global ParkEvent free list. This is correct and suffices for our purposes. // // Beware that the original jvm_raw_unlock() had a "_snuck" test but that // jvm_raw_lock() didn't have the corresponding test. I suspect that's an // oversight, but I've replicated the original suspect logic in the new code ... void Monitor::jvm_raw_lock() { assert(rank() == native, "invariant"); if (TryLock()) { Exeunt: assert (ILocked(), "invariant") ; assert (_owner == NULL, "invariant"); // This can potentially be called by non-java Threads. Thus, the ThreadLocalStorage // might return NULL. Don't call set_owner since it will break on an NULL owner // Consider installing a non-null "ANON" distinguished value instead of just NULL. _owner = ThreadLocalStorage::thread(); return ; } if (TrySpin(NULL)) goto Exeunt ; // slow-path - apparent contention // Allocate a ParkEvent for transient use. // The ParkEvent remains associated with this thread until // the time the thread manages to acquire the lock. ParkEvent * const ESelf = ParkEvent::Allocate(NULL) ; ESelf->reset() ; OrderAccess::storeload() ; // Either Enqueue Self on cxq or acquire the outer lock. if (AcquireOrPush (ESelf)) { ParkEvent::Release (ESelf) ; // surrender the ParkEvent goto Exeunt ; } // At any given time there is at most one ondeck thread. // ondeck implies not resident on cxq and not resident on EntryList // Only the OnDeck thread can try to acquire -- contended for -- the lock. // CONSIDER: use Self->OnDeck instead of m->OnDeck. for (;;) { if (_OnDeck == ESelf && TrySpin(NULL)) break ; ParkCommon (ESelf, 0) ; } assert (_OnDeck == ESelf, "invariant") ; _OnDeck = NULL ; ParkEvent::Release (ESelf) ; // surrender the ParkEvent goto Exeunt ; } void Monitor::jvm_raw_unlock() { // Nearly the same as Monitor::unlock() ... // directly set _owner instead of using set_owner(null) _owner = NULL ; if (_snuck) { // ??? assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak"); _snuck = false; return ; } IUnlock(false) ; } bool Monitor::wait(bool no_safepoint_check, long timeout, bool as_suspend_equivalent) { Thread * const Self = Thread::current() ; assert (_owner == Self, "invariant") ; assert (ILocked(), "invariant") ; // as_suspend_equivalent logically implies !no_safepoint_check guarantee (!as_suspend_equivalent || !no_safepoint_check, "invariant") ; // !no_safepoint_check logically implies java_thread guarantee (no_safepoint_check || Self->is_Java_thread(), "invariant") ; #ifdef ASSERT Monitor * least = get_least_ranked_lock_besides_this(Self->owned_locks()); assert(least != this, "Specification of get_least_... call above"); if (least != NULL && least->rank() <= special) { tty->print("Attempting to wait on monitor %s/%d while holding" " lock %s/%d -- possible deadlock", name(), rank(), least->name(), least->rank()); assert(false, "Shouldn't block(wait) while holding a lock of rank special"); } #endif // ASSERT int wait_status ; // conceptually set the owner to NULL in anticipation of // abdicating the lock in wait set_owner(NULL); if (no_safepoint_check) { wait_status = IWait (Self, timeout) ; } else { assert (Self->is_Java_thread(), "invariant") ; JavaThread *jt = (JavaThread *)Self; // Enter safepoint region - ornate and Rococo ... ThreadBlockInVM tbivm(jt); OSThreadWaitState osts(Self->osthread(), false /* not Object.wait() */); if (as_suspend_equivalent) { jt->set_suspend_equivalent(); // cleared by handle_special_suspend_equivalent_condition() or // java_suspend_self() } wait_status = IWait (Self, timeout) ; // were we externally suspended while we were waiting? if (as_suspend_equivalent && jt->handle_special_suspend_equivalent_condition()) { // Our event wait has finished and we own the lock, but // while we were waiting another thread suspended us. We don't // want to hold the lock while suspended because that // would surprise the thread that suspended us. assert (ILocked(), "invariant") ; IUnlock (true) ; jt->java_suspend_self(); ILock (Self) ; assert (ILocked(), "invariant") ; } } // Conceptually reestablish ownership of the lock. // The "real" lock -- the LockByte -- was reacquired by IWait(). assert (ILocked(), "invariant") ; assert (_owner == NULL, "invariant") ; set_owner (Self) ; return wait_status != 0 ; // return true IFF timeout } Monitor::~Monitor() { assert ((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "") ; } void Monitor::ClearMonitor (Monitor * m, const char *name) { m->_owner = NULL ; m->_snuck = false ; if (name == NULL) { strcpy(m->_name, "UNKNOWN") ; } else { strncpy(m->_name, name, MONITOR_NAME_LEN - 1); m->_name[MONITOR_NAME_LEN - 1] = '\0'; } m->_LockWord.FullWord = 0 ; m->_EntryList = NULL ; m->_OnDeck = NULL ; m->_WaitSet = NULL ; m->_WaitLock[0] = 0 ; } Monitor::Monitor() { ClearMonitor(this); } Monitor::Monitor (int Rank, const char * name, bool allow_vm_block) { ClearMonitor (this, name) ; #ifdef ASSERT _allow_vm_block = allow_vm_block; _rank = Rank ; #endif } Mutex::~Mutex() { assert ((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "") ; } Mutex::Mutex (int Rank, const char * name, bool allow_vm_block) { ClearMonitor ((Monitor *) this, name) ; #ifdef ASSERT _allow_vm_block = allow_vm_block; _rank = Rank ; #endif } bool Monitor::owned_by_self() const { bool ret = _owner == Thread::current(); assert (!ret || _LockWord.Bytes[_LSBINDEX] != 0, "invariant") ; return ret; } void Monitor::print_on_error(outputStream* st) const { st->print("[" PTR_FORMAT, this); st->print("] %s", _name); st->print(" - owner thread: " PTR_FORMAT, _owner); } // ---------------------------------------------------------------------------------- // Non-product code #ifndef PRODUCT void Monitor::print_on(outputStream* st) const { st->print_cr("Mutex: [0x%lx/0x%lx] %s - owner: 0x%lx", this, _LockWord.FullWord, _name, _owner); } #endif #ifndef PRODUCT #ifdef ASSERT Monitor * Monitor::get_least_ranked_lock(Monitor * locks) { Monitor *res, *tmp; for (res = tmp = locks; tmp != NULL; tmp = tmp->next()) { if (tmp->rank() < res->rank()) { res = tmp; } } if (!SafepointSynchronize::is_at_safepoint()) { // In this case, we expect the held locks to be // in increasing rank order (modulo any native ranks) for (tmp = locks; tmp != NULL; tmp = tmp->next()) { if (tmp->next() != NULL) { assert(tmp->rank() == Mutex::native || tmp->rank() <= tmp->next()->rank(), "mutex rank anomaly?"); } } } return res; } Monitor* Monitor::get_least_ranked_lock_besides_this(Monitor* locks) { Monitor *res, *tmp; for (res = NULL, tmp = locks; tmp != NULL; tmp = tmp->next()) { if (tmp != this && (res == NULL || tmp->rank() < res->rank())) { res = tmp; } } if (!SafepointSynchronize::is_at_safepoint()) { // In this case, we expect the held locks to be // in increasing rank order (modulo any native ranks) for (tmp = locks; tmp != NULL; tmp = tmp->next()) { if (tmp->next() != NULL) { assert(tmp->rank() == Mutex::native || tmp->rank() <= tmp->next()->rank(), "mutex rank anomaly?"); } } } return res; } bool Monitor::contains(Monitor* locks, Monitor * lock) { for (; locks != NULL; locks = locks->next()) { if (locks == lock) return true; } return false; } #endif // Called immediately after lock acquisition or release as a diagnostic // to track the lock-set of the thread and test for rank violations that // might indicate exposure to deadlock. // Rather like an EventListener for _owner (:>). void Monitor::set_owner_implementation(Thread *new_owner) { // This function is solely responsible for maintaining // and checking the invariant that threads and locks // are in a 1/N relation, with some some locks unowned. // It uses the Mutex::_owner, Mutex::_next, and // Thread::_owned_locks fields, and no other function // changes those fields. // It is illegal to set the mutex from one non-NULL // owner to another--it must be owned by NULL as an // intermediate state. if (new_owner != NULL) { // the thread is acquiring this lock assert(new_owner == Thread::current(), "Should I be doing this?"); assert(_owner == NULL, "setting the owner thread of an already owned mutex"); _owner = new_owner; // set the owner // link "this" into the owned locks list #ifdef ASSERT // Thread::_owned_locks is under the same ifdef Monitor* locks = get_least_ranked_lock(new_owner->owned_locks()); // Mutex::set_owner_implementation is a friend of Thread assert(this->rank() >= 0, "bad lock rank"); // Deadlock avoidance rules require us to acquire Mutexes only in // a global total order. For example m1 is the lowest ranked mutex // that the thread holds and m2 is the mutex the thread is trying // to acquire, then deadlock avoidance rules require that the rank // of m2 be less than the rank of m1. // The rank Mutex::native is an exception in that it is not subject // to the verification rules. // Here are some further notes relating to mutex acquisition anomalies: // . under Solaris, the interrupt lock gets acquired when doing // profiling, so any lock could be held. // . it is also ok to acquire Safepoint_lock at the very end while we // already hold Terminator_lock - may happen because of periodic safepoints if (this->rank() != Mutex::native && this->rank() != Mutex::suspend_resume && locks != NULL && locks->rank() <= this->rank() && !SafepointSynchronize::is_at_safepoint() && this != Interrupt_lock && this != ProfileVM_lock && !(this == Safepoint_lock && contains(locks, Terminator_lock) && SafepointSynchronize::is_synchronizing())) { new_owner->print_owned_locks(); fatal(err_msg("acquiring lock %s/%d out of order with lock %s/%d -- " "possible deadlock", this->name(), this->rank(), locks->name(), locks->rank())); } this->_next = new_owner->_owned_locks; new_owner->_owned_locks = this; #endif } else { // the thread is releasing this lock Thread* old_owner = _owner; debug_only(_last_owner = old_owner); assert(old_owner != NULL, "removing the owner thread of an unowned mutex"); assert(old_owner == Thread::current(), "removing the owner thread of an unowned mutex"); _owner = NULL; // set the owner #ifdef ASSERT Monitor *locks = old_owner->owned_locks(); // remove "this" from the owned locks list Monitor *prev = NULL; bool found = false; for (; locks != NULL; prev = locks, locks = locks->next()) { if (locks == this) { found = true; break; } } assert(found, "Removing a lock not owned"); if (prev == NULL) { old_owner->_owned_locks = _next; } else { prev->_next = _next; } _next = NULL; #endif } } // Factored out common sanity checks for locking mutex'es. Used by lock() and try_lock() void Monitor::check_prelock_state(Thread *thread) { assert((!thread->is_Java_thread() || ((JavaThread *)thread)->thread_state() == _thread_in_vm) || rank() == Mutex::special, "wrong thread state for using locks"); if (StrictSafepointChecks) { if (thread->is_VM_thread() && !allow_vm_block()) { fatal(err_msg("VM thread using lock %s (not allowed to block on)", name())); } debug_only(if (rank() != Mutex::special) \ thread->check_for_valid_safepoint_state(false);) } if (thread->is_Watcher_thread()) { assert(!WatcherThread::watcher_thread()->has_crash_protection(), "locking not allowed when crash protection is set"); } } void Monitor::check_block_state(Thread *thread) { if (!_allow_vm_block && thread->is_VM_thread()) { warning("VM thread blocked on lock"); print(); BREAKPOINT; } assert(_owner != thread, "deadlock: blocking on monitor owned by current thread"); } #endif // PRODUCT