Mercurial > hg > graal-compiler
view src/share/vm/runtime/synchronizer.cpp @ 1120:9127aa69352e
6648438: 4/4 src/share/vm/prims/jvmtiEnv.cpp:457 assert(phase == JVMTI_PHASE_LIVE,"sanity check")
Summary: Return error on invalid JVMTI_PHASE instead of asserting.
Reviewed-by: dholmes, ohair
author | dcubed |
---|---|
date | Mon, 14 Dec 2009 09:51:09 -0700 |
parents | a7fac4381b50 |
children | 2328d1d3f8cf |
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line source
/* * Copyright 1998-2007 Sun Microsystems, Inc. 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 Sun Microsystems, Inc., 4150 Network Circle, Santa Clara, * CA 95054 USA or visit www.sun.com if you need additional information or * have any questions. * */ # include "incls/_precompiled.incl" # include "incls/_synchronizer.cpp.incl" #if defined(__GNUC__) && !defined(IA64) // Need to inhibit inlining for older versions of GCC to avoid build-time failures #define ATTR __attribute__((noinline)) #else #define ATTR #endif // Native markword accessors for synchronization and hashCode(). // // The "core" versions of monitor enter and exit reside in this file. // The interpreter and compilers contain specialized transliterated // variants of the enter-exit fast-path operations. See i486.ad fast_lock(), // for instance. If you make changes here, make sure to modify the // interpreter, and both C1 and C2 fast-path inline locking code emission. // // TODO: merge the objectMonitor and synchronizer classes. // // ----------------------------------------------------------------------------- #ifdef DTRACE_ENABLED // Only bother with this argument setup if dtrace is available // TODO-FIXME: probes should not fire when caller is _blocked. assert() accordingly. HS_DTRACE_PROBE_DECL5(hotspot, monitor__wait, jlong, uintptr_t, char*, int, long); HS_DTRACE_PROBE_DECL4(hotspot, monitor__waited, jlong, uintptr_t, char*, int); HS_DTRACE_PROBE_DECL4(hotspot, monitor__notify, jlong, uintptr_t, char*, int); HS_DTRACE_PROBE_DECL4(hotspot, monitor__notifyAll, jlong, uintptr_t, char*, int); HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__enter, jlong, uintptr_t, char*, int); HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__entered, jlong, uintptr_t, char*, int); HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__exit, jlong, uintptr_t, char*, int); #define DTRACE_MONITOR_PROBE_COMMON(klassOop, thread) \ char* bytes = NULL; \ int len = 0; \ jlong jtid = SharedRuntime::get_java_tid(thread); \ symbolOop klassname = ((oop)(klassOop))->klass()->klass_part()->name(); \ if (klassname != NULL) { \ bytes = (char*)klassname->bytes(); \ len = klassname->utf8_length(); \ } #define DTRACE_MONITOR_WAIT_PROBE(monitor, klassOop, thread, millis) \ { \ if (DTraceMonitorProbes) { \ DTRACE_MONITOR_PROBE_COMMON(klassOop, thread); \ HS_DTRACE_PROBE5(hotspot, monitor__wait, jtid, \ (monitor), bytes, len, (millis)); \ } \ } #define DTRACE_MONITOR_PROBE(probe, monitor, klassOop, thread) \ { \ if (DTraceMonitorProbes) { \ DTRACE_MONITOR_PROBE_COMMON(klassOop, thread); \ HS_DTRACE_PROBE4(hotspot, monitor__##probe, jtid, \ (uintptr_t)(monitor), bytes, len); \ } \ } #else // ndef DTRACE_ENABLED #define DTRACE_MONITOR_WAIT_PROBE(klassOop, thread, millis, mon) {;} #define DTRACE_MONITOR_PROBE(probe, klassOop, thread, mon) {;} #endif // ndef DTRACE_ENABLED // ObjectWaiter serves as a "proxy" or surrogate thread. // TODO-FIXME: Eliminate ObjectWaiter and use the thread-specific // ParkEvent instead. Beware, however, that the JVMTI code // knows about ObjectWaiters, so we'll have to reconcile that code. // See next_waiter(), first_waiter(), etc. class ObjectWaiter : public StackObj { public: enum TStates { TS_UNDEF, TS_READY, TS_RUN, TS_WAIT, TS_ENTER, TS_CXQ } ; enum Sorted { PREPEND, APPEND, SORTED } ; ObjectWaiter * volatile _next; ObjectWaiter * volatile _prev; Thread* _thread; ParkEvent * _event; volatile int _notified ; volatile TStates TState ; Sorted _Sorted ; // List placement disposition bool _active ; // Contention monitoring is enabled public: ObjectWaiter(Thread* thread) { _next = NULL; _prev = NULL; _notified = 0; TState = TS_RUN ; _thread = thread; _event = thread->_ParkEvent ; _active = false; assert (_event != NULL, "invariant") ; } void wait_reenter_begin(ObjectMonitor *mon) { JavaThread *jt = (JavaThread *)this->_thread; _active = JavaThreadBlockedOnMonitorEnterState::wait_reenter_begin(jt, mon); } void wait_reenter_end(ObjectMonitor *mon) { JavaThread *jt = (JavaThread *)this->_thread; JavaThreadBlockedOnMonitorEnterState::wait_reenter_end(jt, _active); } }; enum ManifestConstants { ClearResponsibleAtSTW = 0, MaximumRecheckInterval = 1000 } ; #undef TEVENT #define TEVENT(nom) {if (SyncVerbose) FEVENT(nom); } #define FEVENT(nom) { static volatile int ctr = 0 ; int v = ++ctr ; if ((v & (v-1)) == 0) { ::printf (#nom " : %d \n", v); ::fflush(stdout); }} #undef TEVENT #define TEVENT(nom) {;} // Performance concern: // OrderAccess::storestore() calls release() which STs 0 into the global volatile // OrderAccess::Dummy variable. This store is unnecessary for correctness. // Many threads STing into a common location causes considerable cache migration // or "sloshing" on large SMP system. As such, I avoid using OrderAccess::storestore() // until it's repaired. In some cases OrderAccess::fence() -- which incurs local // latency on the executing processor -- is a better choice as it scales on SMP // systems. See http://blogs.sun.com/dave/entry/biased_locking_in_hotspot for a // discussion of coherency costs. Note that all our current reference platforms // provide strong ST-ST order, so the issue is moot on IA32, x64, and SPARC. // // As a general policy we use "volatile" to control compiler-based reordering // and explicit fences (barriers) to control for architectural reordering performed // by the CPU(s) or platform. static int MBFence (int x) { OrderAccess::fence(); return x; } struct SharedGlobals { // These are highly shared mostly-read variables. // To avoid false-sharing they need to be the sole occupants of a $ line. double padPrefix [8]; volatile int stwRandom ; volatile int stwCycle ; // Hot RW variables -- Sequester to avoid false-sharing double padSuffix [16]; volatile int hcSequence ; double padFinal [8] ; } ; static SharedGlobals GVars ; // Tunables ... // The knob* variables are effectively final. Once set they should // never be modified hence. Consider using __read_mostly with GCC. static int Knob_LogSpins = 0 ; // enable jvmstat tally for spins static int Knob_HandOff = 0 ; static int Knob_Verbose = 0 ; static int Knob_ReportSettings = 0 ; static int Knob_SpinLimit = 5000 ; // derived by an external tool - static int Knob_SpinBase = 0 ; // Floor AKA SpinMin static int Knob_SpinBackOff = 0 ; // spin-loop backoff static int Knob_CASPenalty = -1 ; // Penalty for failed CAS static int Knob_OXPenalty = -1 ; // Penalty for observed _owner change static int Knob_SpinSetSucc = 1 ; // spinners set the _succ field static int Knob_SpinEarly = 1 ; static int Knob_SuccEnabled = 1 ; // futile wake throttling static int Knob_SuccRestrict = 0 ; // Limit successors + spinners to at-most-one static int Knob_MaxSpinners = -1 ; // Should be a function of # CPUs static int Knob_Bonus = 100 ; // spin success bonus static int Knob_BonusB = 100 ; // spin success bonus static int Knob_Penalty = 200 ; // spin failure penalty static int Knob_Poverty = 1000 ; static int Knob_SpinAfterFutile = 1 ; // Spin after returning from park() static int Knob_FixedSpin = 0 ; static int Knob_OState = 3 ; // Spinner checks thread state of _owner static int Knob_UsePause = 1 ; static int Knob_ExitPolicy = 0 ; static int Knob_PreSpin = 10 ; // 20-100 likely better static int Knob_ResetEvent = 0 ; static int BackOffMask = 0 ; static int Knob_FastHSSEC = 0 ; static int Knob_MoveNotifyee = 2 ; // notify() - disposition of notifyee static int Knob_QMode = 0 ; // EntryList-cxq policy - queue discipline static volatile int InitDone = 0 ; // hashCode() generation : // // Possibilities: // * MD5Digest of {obj,stwRandom} // * CRC32 of {obj,stwRandom} or any linear-feedback shift register function. // * A DES- or AES-style SBox[] mechanism // * One of the Phi-based schemes, such as: // 2654435761 = 2^32 * Phi (golden ratio) // HashCodeValue = ((uintptr_t(obj) >> 3) * 2654435761) ^ GVars.stwRandom ; // * A variation of Marsaglia's shift-xor RNG scheme. // * (obj ^ stwRandom) is appealing, but can result // in undesirable regularity in the hashCode values of adjacent objects // (objects allocated back-to-back, in particular). This could potentially // result in hashtable collisions and reduced hashtable efficiency. // There are simple ways to "diffuse" the middle address bits over the // generated hashCode values: // static inline intptr_t get_next_hash(Thread * Self, oop obj) { intptr_t value = 0 ; if (hashCode == 0) { // This form uses an unguarded global Park-Miller RNG, // so it's possible for two threads to race and generate the same RNG. // On MP system we'll have lots of RW access to a global, so the // mechanism induces lots of coherency traffic. value = os::random() ; } else if (hashCode == 1) { // This variation has the property of being stable (idempotent) // between STW operations. This can be useful in some of the 1-0 // synchronization schemes. intptr_t addrBits = intptr_t(obj) >> 3 ; value = addrBits ^ (addrBits >> 5) ^ GVars.stwRandom ; } else if (hashCode == 2) { value = 1 ; // for sensitivity testing } else if (hashCode == 3) { value = ++GVars.hcSequence ; } else if (hashCode == 4) { value = intptr_t(obj) ; } else { // Marsaglia's xor-shift scheme with thread-specific state // This is probably the best overall implementation -- we'll // likely make this the default in future releases. unsigned t = Self->_hashStateX ; t ^= (t << 11) ; Self->_hashStateX = Self->_hashStateY ; Self->_hashStateY = Self->_hashStateZ ; Self->_hashStateZ = Self->_hashStateW ; unsigned v = Self->_hashStateW ; v = (v ^ (v >> 19)) ^ (t ^ (t >> 8)) ; Self->_hashStateW = v ; value = v ; } value &= markOopDesc::hash_mask; if (value == 0) value = 0xBAD ; assert (value != markOopDesc::no_hash, "invariant") ; TEVENT (hashCode: GENERATE) ; return value; } void BasicLock::print_on(outputStream* st) const { st->print("monitor"); } void BasicLock::move_to(oop obj, BasicLock* dest) { // Check to see if we need to inflate the lock. This is only needed // if an object is locked using "this" lightweight monitor. In that // case, the displaced_header() is unlocked, because the // displaced_header() contains the header for the originally unlocked // object. However the object could have already been inflated. But it // does not matter, the inflation will just a no-op. For other cases, // the displaced header will be either 0x0 or 0x3, which are location // independent, therefore the BasicLock is free to move. // // During OSR we may need to relocate a BasicLock (which contains a // displaced word) from a location in an interpreter frame to a // new location in a compiled frame. "this" refers to the source // basiclock in the interpreter frame. "dest" refers to the destination // basiclock in the new compiled frame. We *always* inflate in move_to(). // The always-Inflate policy works properly, but in 1.5.0 it can sometimes // cause performance problems in code that makes heavy use of a small # of // uncontended locks. (We'd inflate during OSR, and then sync performance // would subsequently plummet because the thread would be forced thru the slow-path). // This problem has been made largely moot on IA32 by inlining the inflated fast-path // operations in Fast_Lock and Fast_Unlock in i486.ad. // // Note that there is a way to safely swing the object's markword from // one stack location to another. This avoids inflation. Obviously, // we need to ensure that both locations refer to the current thread's stack. // There are some subtle concurrency issues, however, and since the benefit is // is small (given the support for inflated fast-path locking in the fast_lock, etc) // we'll leave that optimization for another time. if (displaced_header()->is_neutral()) { ObjectSynchronizer::inflate_helper(obj); // WARNING: We can not put check here, because the inflation // will not update the displaced header. Once BasicLock is inflated, // no one should ever look at its content. } else { // Typically the displaced header will be 0 (recursive stack lock) or // unused_mark. Naively we'd like to assert that the displaced mark // value is either 0, neutral, or 3. But with the advent of the // store-before-CAS avoidance in fast_lock/compiler_lock_object // we can find any flavor mark in the displaced mark. } // [RGV] The next line appears to do nothing! intptr_t dh = (intptr_t) displaced_header(); dest->set_displaced_header(displaced_header()); } // ----------------------------------------------------------------------------- // standard constructor, allows locking failures ObjectLocker::ObjectLocker(Handle obj, Thread* thread, bool doLock) { _dolock = doLock; _thread = thread; debug_only(if (StrictSafepointChecks) _thread->check_for_valid_safepoint_state(false);) _obj = obj; if (_dolock) { TEVENT (ObjectLocker) ; ObjectSynchronizer::fast_enter(_obj, &_lock, false, _thread); } } ObjectLocker::~ObjectLocker() { if (_dolock) { ObjectSynchronizer::fast_exit(_obj(), &_lock, _thread); } } // ----------------------------------------------------------------------------- PerfCounter * ObjectSynchronizer::_sync_Inflations = NULL ; PerfCounter * ObjectSynchronizer::_sync_Deflations = NULL ; PerfCounter * ObjectSynchronizer::_sync_ContendedLockAttempts = NULL ; PerfCounter * ObjectSynchronizer::_sync_FutileWakeups = NULL ; PerfCounter * ObjectSynchronizer::_sync_Parks = NULL ; PerfCounter * ObjectSynchronizer::_sync_EmptyNotifications = NULL ; PerfCounter * ObjectSynchronizer::_sync_Notifications = NULL ; PerfCounter * ObjectSynchronizer::_sync_PrivateA = NULL ; PerfCounter * ObjectSynchronizer::_sync_PrivateB = NULL ; PerfCounter * ObjectSynchronizer::_sync_SlowExit = NULL ; PerfCounter * ObjectSynchronizer::_sync_SlowEnter = NULL ; PerfCounter * ObjectSynchronizer::_sync_SlowNotify = NULL ; PerfCounter * ObjectSynchronizer::_sync_SlowNotifyAll = NULL ; PerfCounter * ObjectSynchronizer::_sync_FailedSpins = NULL ; PerfCounter * ObjectSynchronizer::_sync_SuccessfulSpins = NULL ; PerfCounter * ObjectSynchronizer::_sync_MonInCirculation = NULL ; PerfCounter * ObjectSynchronizer::_sync_MonScavenged = NULL ; PerfLongVariable * ObjectSynchronizer::_sync_MonExtant = NULL ; // One-shot global initialization for the sync subsystem. // We could also defer initialization and initialize on-demand // the first time we call inflate(). Initialization would // be protected - like so many things - by the MonitorCache_lock. void ObjectSynchronizer::Initialize () { static int InitializationCompleted = 0 ; assert (InitializationCompleted == 0, "invariant") ; InitializationCompleted = 1 ; if (UsePerfData) { EXCEPTION_MARK ; #define NEWPERFCOUNTER(n) {n = PerfDataManager::create_counter(SUN_RT, #n, PerfData::U_Events,CHECK); } #define NEWPERFVARIABLE(n) {n = PerfDataManager::create_variable(SUN_RT, #n, PerfData::U_Events,CHECK); } NEWPERFCOUNTER(_sync_Inflations) ; NEWPERFCOUNTER(_sync_Deflations) ; NEWPERFCOUNTER(_sync_ContendedLockAttempts) ; NEWPERFCOUNTER(_sync_FutileWakeups) ; NEWPERFCOUNTER(_sync_Parks) ; NEWPERFCOUNTER(_sync_EmptyNotifications) ; NEWPERFCOUNTER(_sync_Notifications) ; NEWPERFCOUNTER(_sync_SlowEnter) ; NEWPERFCOUNTER(_sync_SlowExit) ; NEWPERFCOUNTER(_sync_SlowNotify) ; NEWPERFCOUNTER(_sync_SlowNotifyAll) ; NEWPERFCOUNTER(_sync_FailedSpins) ; NEWPERFCOUNTER(_sync_SuccessfulSpins) ; NEWPERFCOUNTER(_sync_PrivateA) ; NEWPERFCOUNTER(_sync_PrivateB) ; NEWPERFCOUNTER(_sync_MonInCirculation) ; NEWPERFCOUNTER(_sync_MonScavenged) ; NEWPERFVARIABLE(_sync_MonExtant) ; #undef NEWPERFCOUNTER } } // Compile-time asserts // When possible, it's better to catch errors deterministically at // compile-time than at runtime. The down-side to using compile-time // asserts is that error message -- often something about negative array // indices -- is opaque. #define CTASSERT(x) { int tag[1-(2*!(x))]; printf ("Tag @%X\n", tag); } void ObjectMonitor::ctAsserts() { CTASSERT(offset_of (ObjectMonitor, _header) == 0); } static int Adjust (volatile int * adr, int dx) { int v ; for (v = *adr ; Atomic::cmpxchg (v + dx, adr, v) != v; v = *adr) ; return v ; } // Ad-hoc mutual exclusion primitives: SpinLock and Mux // // We employ SpinLocks _only for low-contention, fixed-length // short-duration critical sections where we're concerned // about native mutex_t or HotSpot Mutex:: latency. // The mux construct provides a spin-then-block mutual exclusion // mechanism. // // Testing has shown that contention on the ListLock guarding gFreeList // is common. If we implement ListLock as a simple SpinLock it's common // for the JVM to devolve to yielding with little progress. This is true // despite the fact that the critical sections protected by ListLock are // extremely short. // // TODO-FIXME: ListLock should be of type SpinLock. // We should make this a 1st-class type, integrated into the lock // hierarchy as leaf-locks. Critically, the SpinLock structure // should have sufficient padding to avoid false-sharing and excessive // cache-coherency traffic. typedef volatile int SpinLockT ; void Thread::SpinAcquire (volatile int * adr, const char * LockName) { if (Atomic::cmpxchg (1, adr, 0) == 0) { return ; // normal fast-path return } // Slow-path : We've encountered contention -- Spin/Yield/Block strategy. TEVENT (SpinAcquire - ctx) ; int ctr = 0 ; int Yields = 0 ; for (;;) { while (*adr != 0) { ++ctr ; if ((ctr & 0xFFF) == 0 || !os::is_MP()) { if (Yields > 5) { // Consider using a simple NakedSleep() instead. // Then SpinAcquire could be called by non-JVM threads Thread::current()->_ParkEvent->park(1) ; } else { os::NakedYield() ; ++Yields ; } } else { SpinPause() ; } } if (Atomic::cmpxchg (1, adr, 0) == 0) return ; } } void Thread::SpinRelease (volatile int * adr) { assert (*adr != 0, "invariant") ; OrderAccess::fence() ; // guarantee at least release consistency. // Roach-motel semantics. // It's safe if subsequent LDs and STs float "up" into the critical section, // but prior LDs and STs within the critical section can't be allowed // to reorder or float past the ST that releases the lock. *adr = 0 ; } // muxAcquire and muxRelease: // // * muxAcquire and muxRelease support a single-word lock-word construct. // The LSB of the word is set IFF the lock is held. // The remainder of the word points to the head of a singly-linked list // of threads blocked on the lock. // // * The current implementation of muxAcquire-muxRelease uses its own // dedicated Thread._MuxEvent instance. If we're interested in // minimizing the peak number of extant ParkEvent instances then // we could eliminate _MuxEvent and "borrow" _ParkEvent as long // as certain invariants were satisfied. Specifically, care would need // to be taken with regards to consuming unpark() "permits". // A safe rule of thumb is that a thread would never call muxAcquire() // if it's enqueued (cxq, EntryList, WaitList, etc) and will subsequently // park(). Otherwise the _ParkEvent park() operation in muxAcquire() could // consume an unpark() permit intended for monitorenter, for instance. // One way around this would be to widen the restricted-range semaphore // implemented in park(). Another alternative would be to provide // multiple instances of the PlatformEvent() for each thread. One // instance would be dedicated to muxAcquire-muxRelease, for instance. // // * Usage: // -- Only as leaf locks // -- for short-term locking only as muxAcquire does not perform // thread state transitions. // // Alternatives: // * We could implement muxAcquire and muxRelease with MCS or CLH locks // but with parking or spin-then-park instead of pure spinning. // * Use Taura-Oyama-Yonenzawa locks. // * It's possible to construct a 1-0 lock if we encode the lockword as // (List,LockByte). Acquire will CAS the full lockword while Release // will STB 0 into the LockByte. The 1-0 scheme admits stranding, so // acquiring threads use timers (ParkTimed) to detect and recover from // the stranding window. Thread/Node structures must be aligned on 256-byte // boundaries by using placement-new. // * Augment MCS with advisory back-link fields maintained with CAS(). // Pictorially: LockWord -> T1 <-> T2 <-> T3 <-> ... <-> Tn <-> Owner. // The validity of the backlinks must be ratified before we trust the value. // If the backlinks are invalid the exiting thread must back-track through the // the forward links, which are always trustworthy. // * Add a successor indication. The LockWord is currently encoded as // (List, LOCKBIT:1). We could also add a SUCCBIT or an explicit _succ variable // to provide the usual futile-wakeup optimization. // See RTStt for details. // * Consider schedctl.sc_nopreempt to cover the critical section. // typedef volatile intptr_t MutexT ; // Mux Lock-word enum MuxBits { LOCKBIT = 1 } ; void Thread::muxAcquire (volatile intptr_t * Lock, const char * LockName) { intptr_t w = Atomic::cmpxchg_ptr (LOCKBIT, Lock, 0) ; if (w == 0) return ; if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) { return ; } TEVENT (muxAcquire - Contention) ; ParkEvent * const Self = Thread::current()->_MuxEvent ; assert ((intptr_t(Self) & LOCKBIT) == 0, "invariant") ; for (;;) { int its = (os::is_MP() ? 100 : 0) + 1 ; // Optional spin phase: spin-then-park strategy while (--its >= 0) { w = *Lock ; if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) { return ; } } Self->reset() ; Self->OnList = intptr_t(Lock) ; // The following fence() isn't _strictly necessary as the subsequent // CAS() both serializes execution and ratifies the fetched *Lock value. OrderAccess::fence(); for (;;) { w = *Lock ; if ((w & LOCKBIT) == 0) { if (Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) { Self->OnList = 0 ; // hygiene - allows stronger asserts return ; } continue ; // Interference -- *Lock changed -- Just retry } assert (w & LOCKBIT, "invariant") ; Self->ListNext = (ParkEvent *) (w & ~LOCKBIT ); if (Atomic::cmpxchg_ptr (intptr_t(Self)|LOCKBIT, Lock, w) == w) break ; } while (Self->OnList != 0) { Self->park() ; } } } void Thread::muxAcquireW (volatile intptr_t * Lock, ParkEvent * ev) { intptr_t w = Atomic::cmpxchg_ptr (LOCKBIT, Lock, 0) ; if (w == 0) return ; if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) { return ; } TEVENT (muxAcquire - Contention) ; ParkEvent * ReleaseAfter = NULL ; if (ev == NULL) { ev = ReleaseAfter = ParkEvent::Allocate (NULL) ; } assert ((intptr_t(ev) & LOCKBIT) == 0, "invariant") ; for (;;) { guarantee (ev->OnList == 0, "invariant") ; int its = (os::is_MP() ? 100 : 0) + 1 ; // Optional spin phase: spin-then-park strategy while (--its >= 0) { w = *Lock ; if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) { if (ReleaseAfter != NULL) { ParkEvent::Release (ReleaseAfter) ; } return ; } } ev->reset() ; ev->OnList = intptr_t(Lock) ; // The following fence() isn't _strictly necessary as the subsequent // CAS() both serializes execution and ratifies the fetched *Lock value. OrderAccess::fence(); for (;;) { w = *Lock ; if ((w & LOCKBIT) == 0) { if (Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) { ev->OnList = 0 ; // We call ::Release while holding the outer lock, thus // artificially lengthening the critical section. // Consider deferring the ::Release() until the subsequent unlock(), // after we've dropped the outer lock. if (ReleaseAfter != NULL) { ParkEvent::Release (ReleaseAfter) ; } return ; } continue ; // Interference -- *Lock changed -- Just retry } assert (w & LOCKBIT, "invariant") ; ev->ListNext = (ParkEvent *) (w & ~LOCKBIT ); if (Atomic::cmpxchg_ptr (intptr_t(ev)|LOCKBIT, Lock, w) == w) break ; } while (ev->OnList != 0) { ev->park() ; } } } // Release() must extract a successor from the list and then wake that thread. // It can "pop" the front of the list or use a detach-modify-reattach (DMR) scheme // similar to that used by ParkEvent::Allocate() and ::Release(). DMR-based // Release() would : // (A) CAS() or swap() null to *Lock, releasing the lock and detaching the list. // (B) Extract a successor from the private list "in-hand" // (C) attempt to CAS() the residual back into *Lock over null. // If there were any newly arrived threads and the CAS() would fail. // In that case Release() would detach the RATs, re-merge the list in-hand // with the RATs and repeat as needed. Alternately, Release() might // detach and extract a successor, but then pass the residual list to the wakee. // The wakee would be responsible for reattaching and remerging before it // competed for the lock. // // Both "pop" and DMR are immune from ABA corruption -- there can be // multiple concurrent pushers, but only one popper or detacher. // This implementation pops from the head of the list. This is unfair, // but tends to provide excellent throughput as hot threads remain hot. // (We wake recently run threads first). void Thread::muxRelease (volatile intptr_t * Lock) { for (;;) { const intptr_t w = Atomic::cmpxchg_ptr (0, Lock, LOCKBIT) ; assert (w & LOCKBIT, "invariant") ; if (w == LOCKBIT) return ; ParkEvent * List = (ParkEvent *) (w & ~LOCKBIT) ; assert (List != NULL, "invariant") ; assert (List->OnList == intptr_t(Lock), "invariant") ; ParkEvent * nxt = List->ListNext ; // The following CAS() releases the lock and pops the head element. if (Atomic::cmpxchg_ptr (intptr_t(nxt), Lock, w) != w) { continue ; } List->OnList = 0 ; OrderAccess::fence() ; List->unpark () ; return ; } } // ObjectMonitor Lifecycle // ----------------------- // Inflation unlinks monitors from the global gFreeList and // associates them with objects. Deflation -- which occurs at // STW-time -- disassociates idle monitors from objects. Such // scavenged monitors are returned to the gFreeList. // // The global list is protected by ListLock. All the critical sections // are short and operate in constant-time. // // ObjectMonitors reside in type-stable memory (TSM) and are immortal. // // Lifecycle: // -- unassigned and on the global free list // -- unassigned and on a thread's private omFreeList // -- assigned to an object. The object is inflated and the mark refers // to the objectmonitor. // // TODO-FIXME: // // * We currently protect the gFreeList with a simple lock. // An alternate lock-free scheme would be to pop elements from the gFreeList // with CAS. This would be safe from ABA corruption as long we only // recycled previously appearing elements onto the list in deflate_idle_monitors() // at STW-time. Completely new elements could always be pushed onto the gFreeList // with CAS. Elements that appeared previously on the list could only // be installed at STW-time. // // * For efficiency and to help reduce the store-before-CAS penalty // the objectmonitors on gFreeList or local free lists should be ready to install // with the exception of _header and _object. _object can be set after inflation. // In particular, keep all objectMonitors on a thread's private list in ready-to-install // state with m.Owner set properly. // // * We could all diffuse contention by using multiple global (FreeList, Lock) // pairs -- threads could use trylock() and a cyclic-scan strategy to search for // an unlocked free list. // // * Add lifecycle tags and assert()s. // // * Be more consistent about when we clear an objectmonitor's fields: // A. After extracting the objectmonitor from a free list. // B. After adding an objectmonitor to a free list. // ObjectMonitor * ObjectSynchronizer::gBlockList = NULL ; ObjectMonitor * volatile ObjectSynchronizer::gFreeList = NULL ; static volatile intptr_t ListLock = 0 ; // protects global monitor free-list cache #define CHAINMARKER ((oop)-1) ObjectMonitor * ATTR ObjectSynchronizer::omAlloc (Thread * Self) { // A large MAXPRIVATE value reduces both list lock contention // and list coherency traffic, but also tends to increase the // number of objectMonitors in circulation as well as the STW // scavenge costs. As usual, we lean toward time in space-time // tradeoffs. const int MAXPRIVATE = 1024 ; for (;;) { ObjectMonitor * m ; // 1: try to allocate from the thread's local omFreeList. // Threads will attempt to allocate first from their local list, then // from the global list, and only after those attempts fail will the thread // attempt to instantiate new monitors. Thread-local free lists take // heat off the ListLock and improve allocation latency, as well as reducing // coherency traffic on the shared global list. m = Self->omFreeList ; if (m != NULL) { Self->omFreeList = m->FreeNext ; Self->omFreeCount -- ; // CONSIDER: set m->FreeNext = BAD -- diagnostic hygiene guarantee (m->object() == NULL, "invariant") ; return m ; } // 2: try to allocate from the global gFreeList // CONSIDER: use muxTry() instead of muxAcquire(). // If the muxTry() fails then drop immediately into case 3. // If we're using thread-local free lists then try // to reprovision the caller's free list. if (gFreeList != NULL) { // Reprovision the thread's omFreeList. // Use bulk transfers to reduce the allocation rate and heat // on various locks. Thread::muxAcquire (&ListLock, "omAlloc") ; for (int i = Self->omFreeProvision; --i >= 0 && gFreeList != NULL; ) { ObjectMonitor * take = gFreeList ; gFreeList = take->FreeNext ; guarantee (take->object() == NULL, "invariant") ; guarantee (!take->is_busy(), "invariant") ; take->Recycle() ; omRelease (Self, take) ; } Thread::muxRelease (&ListLock) ; Self->omFreeProvision += 1 + (Self->omFreeProvision/2) ; if (Self->omFreeProvision > MAXPRIVATE ) Self->omFreeProvision = MAXPRIVATE ; TEVENT (omFirst - reprovision) ; continue ; } // 3: allocate a block of new ObjectMonitors // Both the local and global free lists are empty -- resort to malloc(). // In the current implementation objectMonitors are TSM - immortal. assert (_BLOCKSIZE > 1, "invariant") ; ObjectMonitor * temp = new ObjectMonitor[_BLOCKSIZE]; // NOTE: (almost) no way to recover if allocation failed. // We might be able to induce a STW safepoint and scavenge enough // objectMonitors to permit progress. if (temp == NULL) { vm_exit_out_of_memory (sizeof (ObjectMonitor[_BLOCKSIZE]), "Allocate ObjectMonitors") ; } // Format the block. // initialize the linked list, each monitor points to its next // forming the single linked free list, the very first monitor // will points to next block, which forms the block list. // The trick of using the 1st element in the block as gBlockList // linkage should be reconsidered. A better implementation would // look like: class Block { Block * next; int N; ObjectMonitor Body [N] ; } for (int i = 1; i < _BLOCKSIZE ; i++) { temp[i].FreeNext = &temp[i+1]; } // terminate the last monitor as the end of list temp[_BLOCKSIZE - 1].FreeNext = NULL ; // Element [0] is reserved for global list linkage temp[0].set_object(CHAINMARKER); // Consider carving out this thread's current request from the // block in hand. This avoids some lock traffic and redundant // list activity. // Acquire the ListLock to manipulate BlockList and FreeList. // An Oyama-Taura-Yonezawa scheme might be more efficient. Thread::muxAcquire (&ListLock, "omAlloc [2]") ; // Add the new block to the list of extant blocks (gBlockList). // The very first objectMonitor in a block is reserved and dedicated. // It serves as blocklist "next" linkage. temp[0].FreeNext = gBlockList; gBlockList = temp; // Add the new string of objectMonitors to the global free list temp[_BLOCKSIZE - 1].FreeNext = gFreeList ; gFreeList = temp + 1; Thread::muxRelease (&ListLock) ; TEVENT (Allocate block of monitors) ; } } // Place "m" on the caller's private per-thread omFreeList. // In practice there's no need to clamp or limit the number of // monitors on a thread's omFreeList as the only time we'll call // omRelease is to return a monitor to the free list after a CAS // attempt failed. This doesn't allow unbounded #s of monitors to // accumulate on a thread's free list. // // In the future the usage of omRelease() might change and monitors // could migrate between free lists. In that case to avoid excessive // accumulation we could limit omCount to (omProvision*2), otherwise return // the objectMonitor to the global list. We should drain (return) in reasonable chunks. // That is, *not* one-at-a-time. void ObjectSynchronizer::omRelease (Thread * Self, ObjectMonitor * m) { guarantee (m->object() == NULL, "invariant") ; m->FreeNext = Self->omFreeList ; Self->omFreeList = m ; Self->omFreeCount ++ ; } // Return the monitors of a moribund thread's local free list to // the global free list. Typically a thread calls omFlush() when // it's dying. We could also consider having the VM thread steal // monitors from threads that have not run java code over a few // consecutive STW safepoints. Relatedly, we might decay // omFreeProvision at STW safepoints. // // We currently call omFlush() from the Thread:: dtor _after the thread // has been excised from the thread list and is no longer a mutator. // That means that omFlush() can run concurrently with a safepoint and // the scavenge operator. Calling omFlush() from JavaThread::exit() might // be a better choice as we could safely reason that that the JVM is // not at a safepoint at the time of the call, and thus there could // be not inopportune interleavings between omFlush() and the scavenge // operator. void ObjectSynchronizer::omFlush (Thread * Self) { ObjectMonitor * List = Self->omFreeList ; // Null-terminated SLL Self->omFreeList = NULL ; if (List == NULL) return ; ObjectMonitor * Tail = NULL ; ObjectMonitor * s ; for (s = List ; s != NULL ; s = s->FreeNext) { Tail = s ; guarantee (s->object() == NULL, "invariant") ; guarantee (!s->is_busy(), "invariant") ; s->set_owner (NULL) ; // redundant but good hygiene TEVENT (omFlush - Move one) ; } guarantee (Tail != NULL && List != NULL, "invariant") ; Thread::muxAcquire (&ListLock, "omFlush") ; Tail->FreeNext = gFreeList ; gFreeList = List ; Thread::muxRelease (&ListLock) ; TEVENT (omFlush) ; } // Get the next block in the block list. static inline ObjectMonitor* next(ObjectMonitor* block) { assert(block->object() == CHAINMARKER, "must be a block header"); block = block->FreeNext ; assert(block == NULL || block->object() == CHAINMARKER, "must be a block header"); return block; } // Fast path code shared by multiple functions ObjectMonitor* ObjectSynchronizer::inflate_helper(oop obj) { markOop mark = obj->mark(); if (mark->has_monitor()) { assert(ObjectSynchronizer::verify_objmon_isinpool(mark->monitor()), "monitor is invalid"); assert(mark->monitor()->header()->is_neutral(), "monitor must record a good object header"); return mark->monitor(); } return ObjectSynchronizer::inflate(Thread::current(), obj); } // Note that we could encounter some performance loss through false-sharing as // multiple locks occupy the same $ line. Padding might be appropriate. #define NINFLATIONLOCKS 256 static volatile intptr_t InflationLocks [NINFLATIONLOCKS] ; static markOop ReadStableMark (oop obj) { markOop mark = obj->mark() ; if (!mark->is_being_inflated()) { return mark ; // normal fast-path return } int its = 0 ; for (;;) { markOop mark = obj->mark() ; if (!mark->is_being_inflated()) { return mark ; // normal fast-path return } // The object is being inflated by some other thread. // The caller of ReadStableMark() must wait for inflation to complete. // Avoid live-lock // TODO: consider calling SafepointSynchronize::do_call_back() while // spinning to see if there's a safepoint pending. If so, immediately // yielding or blocking would be appropriate. Avoid spinning while // there is a safepoint pending. // TODO: add inflation contention performance counters. // TODO: restrict the aggregate number of spinners. ++its ; if (its > 10000 || !os::is_MP()) { if (its & 1) { os::NakedYield() ; TEVENT (Inflate: INFLATING - yield) ; } else { // Note that the following code attenuates the livelock problem but is not // a complete remedy. A more complete solution would require that the inflating // thread hold the associated inflation lock. The following code simply restricts // the number of spinners to at most one. We'll have N-2 threads blocked // on the inflationlock, 1 thread holding the inflation lock and using // a yield/park strategy, and 1 thread in the midst of inflation. // A more refined approach would be to change the encoding of INFLATING // to allow encapsulation of a native thread pointer. Threads waiting for // inflation to complete would use CAS to push themselves onto a singly linked // list rooted at the markword. Once enqueued, they'd loop, checking a per-thread flag // and calling park(). When inflation was complete the thread that accomplished inflation // would detach the list and set the markword to inflated with a single CAS and // then for each thread on the list, set the flag and unpark() the thread. // This is conceptually similar to muxAcquire-muxRelease, except that muxRelease // wakes at most one thread whereas we need to wake the entire list. int ix = (intptr_t(obj) >> 5) & (NINFLATIONLOCKS-1) ; int YieldThenBlock = 0 ; assert (ix >= 0 && ix < NINFLATIONLOCKS, "invariant") ; assert ((NINFLATIONLOCKS & (NINFLATIONLOCKS-1)) == 0, "invariant") ; Thread::muxAcquire (InflationLocks + ix, "InflationLock") ; while (obj->mark() == markOopDesc::INFLATING()) { // Beware: NakedYield() is advisory and has almost no effect on some platforms // so we periodically call Self->_ParkEvent->park(1). // We use a mixed spin/yield/block mechanism. if ((YieldThenBlock++) >= 16) { Thread::current()->_ParkEvent->park(1) ; } else { os::NakedYield() ; } } Thread::muxRelease (InflationLocks + ix ) ; TEVENT (Inflate: INFLATING - yield/park) ; } } else { SpinPause() ; // SMP-polite spinning } } } ObjectMonitor * ATTR ObjectSynchronizer::inflate (Thread * Self, oop object) { // Inflate mutates the heap ... // Relaxing assertion for bug 6320749. assert (Universe::verify_in_progress() || !SafepointSynchronize::is_at_safepoint(), "invariant") ; for (;;) { const markOop mark = object->mark() ; assert (!mark->has_bias_pattern(), "invariant") ; // The mark can be in one of the following states: // * Inflated - just return // * Stack-locked - coerce it to inflated // * INFLATING - busy wait for conversion to complete // * Neutral - aggressively inflate the object. // * BIASED - Illegal. We should never see this // CASE: inflated if (mark->has_monitor()) { ObjectMonitor * inf = mark->monitor() ; assert (inf->header()->is_neutral(), "invariant"); assert (inf->object() == object, "invariant") ; assert (ObjectSynchronizer::verify_objmon_isinpool(inf), "monitor is invalid"); return inf ; } // CASE: inflation in progress - inflating over a stack-lock. // Some other thread is converting from stack-locked to inflated. // Only that thread can complete inflation -- other threads must wait. // The INFLATING value is transient. // Currently, we spin/yield/park and poll the markword, waiting for inflation to finish. // We could always eliminate polling by parking the thread on some auxiliary list. if (mark == markOopDesc::INFLATING()) { TEVENT (Inflate: spin while INFLATING) ; ReadStableMark(object) ; continue ; } // CASE: stack-locked // Could be stack-locked either by this thread or by some other thread. // // Note that we allocate the objectmonitor speculatively, _before_ attempting // to install INFLATING into the mark word. We originally installed INFLATING, // allocated the objectmonitor, and then finally STed the address of the // objectmonitor into the mark. This was correct, but artificially lengthened // the interval in which INFLATED appeared in the mark, thus increasing // the odds of inflation contention. // // We now use per-thread private objectmonitor free lists. // These list are reprovisioned from the global free list outside the // critical INFLATING...ST interval. A thread can transfer // multiple objectmonitors en-mass from the global free list to its local free list. // This reduces coherency traffic and lock contention on the global free list. // Using such local free lists, it doesn't matter if the omAlloc() call appears // before or after the CAS(INFLATING) operation. // See the comments in omAlloc(). if (mark->has_locker()) { ObjectMonitor * m = omAlloc (Self) ; // Optimistically prepare the objectmonitor - anticipate successful CAS // We do this before the CAS in order to minimize the length of time // in which INFLATING appears in the mark. m->Recycle(); m->FreeNext = NULL ; m->_Responsible = NULL ; m->OwnerIsThread = 0 ; m->_recursions = 0 ; m->_SpinDuration = Knob_SpinLimit ; // Consider: maintain by type/class markOop cmp = (markOop) Atomic::cmpxchg_ptr (markOopDesc::INFLATING(), object->mark_addr(), mark) ; if (cmp != mark) { omRelease (Self, m) ; continue ; // Interference -- just retry } // We've successfully installed INFLATING (0) into the mark-word. // This is the only case where 0 will appear in a mark-work. // Only the singular thread that successfully swings the mark-word // to 0 can perform (or more precisely, complete) inflation. // // Why do we CAS a 0 into the mark-word instead of just CASing the // mark-word from the stack-locked value directly to the new inflated state? // Consider what happens when a thread unlocks a stack-locked object. // It attempts to use CAS to swing the displaced header value from the // on-stack basiclock back into the object header. Recall also that the // header value (hashcode, etc) can reside in (a) the object header, or // (b) a displaced header associated with the stack-lock, or (c) a displaced // header in an objectMonitor. The inflate() routine must copy the header // value from the basiclock on the owner's stack to the objectMonitor, all // the while preserving the hashCode stability invariants. If the owner // decides to release the lock while the value is 0, the unlock will fail // and control will eventually pass from slow_exit() to inflate. The owner // will then spin, waiting for the 0 value to disappear. Put another way, // the 0 causes the owner to stall if the owner happens to try to // drop the lock (restoring the header from the basiclock to the object) // while inflation is in-progress. This protocol avoids races that might // would otherwise permit hashCode values to change or "flicker" for an object. // Critically, while object->mark is 0 mark->displaced_mark_helper() is stable. // 0 serves as a "BUSY" inflate-in-progress indicator. // fetch the displaced mark from the owner's stack. // The owner can't die or unwind past the lock while our INFLATING // object is in the mark. Furthermore the owner can't complete // an unlock on the object, either. markOop dmw = mark->displaced_mark_helper() ; assert (dmw->is_neutral(), "invariant") ; // Setup monitor fields to proper values -- prepare the monitor m->set_header(dmw) ; // Optimization: if the mark->locker stack address is associated // with this thread we could simply set m->_owner = Self and // m->OwnerIsThread = 1. Note that a thread can inflate an object // that it has stack-locked -- as might happen in wait() -- directly // with CAS. That is, we can avoid the xchg-NULL .... ST idiom. m->set_owner (mark->locker()); m->set_object(object); // TODO-FIXME: assert BasicLock->dhw != 0. // Must preserve store ordering. The monitor state must // be stable at the time of publishing the monitor address. guarantee (object->mark() == markOopDesc::INFLATING(), "invariant") ; object->release_set_mark(markOopDesc::encode(m)); // Hopefully the performance counters are allocated on distinct cache lines // to avoid false sharing on MP systems ... if (_sync_Inflations != NULL) _sync_Inflations->inc() ; TEVENT(Inflate: overwrite stacklock) ; if (TraceMonitorInflation) { if (object->is_instance()) { ResourceMark rm; tty->print_cr("Inflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s", (intptr_t) object, (intptr_t) object->mark(), Klass::cast(object->klass())->external_name()); } } return m ; } // CASE: neutral // TODO-FIXME: for entry we currently inflate and then try to CAS _owner. // If we know we're inflating for entry it's better to inflate by swinging a // pre-locked objectMonitor pointer into the object header. A successful // CAS inflates the object *and* confers ownership to the inflating thread. // In the current implementation we use a 2-step mechanism where we CAS() // to inflate and then CAS() again to try to swing _owner from NULL to Self. // An inflateTry() method that we could call from fast_enter() and slow_enter() // would be useful. assert (mark->is_neutral(), "invariant"); ObjectMonitor * m = omAlloc (Self) ; // prepare m for installation - set monitor to initial state m->Recycle(); m->set_header(mark); m->set_owner(NULL); m->set_object(object); m->OwnerIsThread = 1 ; m->_recursions = 0 ; m->FreeNext = NULL ; m->_Responsible = NULL ; m->_SpinDuration = Knob_SpinLimit ; // consider: keep metastats by type/class if (Atomic::cmpxchg_ptr (markOopDesc::encode(m), object->mark_addr(), mark) != mark) { m->set_object (NULL) ; m->set_owner (NULL) ; m->OwnerIsThread = 0 ; m->Recycle() ; omRelease (Self, m) ; m = NULL ; continue ; // interference - the markword changed - just retry. // The state-transitions are one-way, so there's no chance of // live-lock -- "Inflated" is an absorbing state. } // Hopefully the performance counters are allocated on distinct // cache lines to avoid false sharing on MP systems ... if (_sync_Inflations != NULL) _sync_Inflations->inc() ; TEVENT(Inflate: overwrite neutral) ; if (TraceMonitorInflation) { if (object->is_instance()) { ResourceMark rm; tty->print_cr("Inflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s", (intptr_t) object, (intptr_t) object->mark(), Klass::cast(object->klass())->external_name()); } } return m ; } } // This the fast monitor enter. The interpreter and compiler use // some assembly copies of this code. Make sure update those code // if the following function is changed. The implementation is // extremely sensitive to race condition. Be careful. void ObjectSynchronizer::fast_enter(Handle obj, BasicLock* lock, bool attempt_rebias, TRAPS) { if (UseBiasedLocking) { if (!SafepointSynchronize::is_at_safepoint()) { BiasedLocking::Condition cond = BiasedLocking::revoke_and_rebias(obj, attempt_rebias, THREAD); if (cond == BiasedLocking::BIAS_REVOKED_AND_REBIASED) { return; } } else { assert(!attempt_rebias, "can not rebias toward VM thread"); BiasedLocking::revoke_at_safepoint(obj); } assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } THREAD->update_highest_lock((address)lock); slow_enter (obj, lock, THREAD) ; } void ObjectSynchronizer::fast_exit(oop object, BasicLock* lock, TRAPS) { assert(!object->mark()->has_bias_pattern(), "should not see bias pattern here"); // if displaced header is null, the previous enter is recursive enter, no-op markOop dhw = lock->displaced_header(); markOop mark ; if (dhw == NULL) { // Recursive stack-lock. // Diagnostics -- Could be: stack-locked, inflating, inflated. mark = object->mark() ; assert (!mark->is_neutral(), "invariant") ; if (mark->has_locker() && mark != markOopDesc::INFLATING()) { assert(THREAD->is_lock_owned((address)mark->locker()), "invariant") ; } if (mark->has_monitor()) { ObjectMonitor * m = mark->monitor() ; assert(((oop)(m->object()))->mark() == mark, "invariant") ; assert(m->is_entered(THREAD), "invariant") ; } return ; } mark = object->mark() ; // If the object is stack-locked by the current thread, try to // swing the displaced header from the box back to the mark. if (mark == (markOop) lock) { assert (dhw->is_neutral(), "invariant") ; if ((markOop) Atomic::cmpxchg_ptr (dhw, object->mark_addr(), mark) == mark) { TEVENT (fast_exit: release stacklock) ; return; } } ObjectSynchronizer::inflate(THREAD, object)->exit (THREAD) ; } // This routine is used to handle interpreter/compiler slow case // We don't need to use fast path here, because it must have been // failed in the interpreter/compiler code. void ObjectSynchronizer::slow_enter(Handle obj, BasicLock* lock, TRAPS) { markOop mark = obj->mark(); assert(!mark->has_bias_pattern(), "should not see bias pattern here"); if (mark->is_neutral()) { // Anticipate successful CAS -- the ST of the displaced mark must // be visible <= the ST performed by the CAS. lock->set_displaced_header(mark); if (mark == (markOop) Atomic::cmpxchg_ptr(lock, obj()->mark_addr(), mark)) { TEVENT (slow_enter: release stacklock) ; return ; } // Fall through to inflate() ... } else if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) { assert(lock != mark->locker(), "must not re-lock the same lock"); assert(lock != (BasicLock*)obj->mark(), "don't relock with same BasicLock"); lock->set_displaced_header(NULL); return; } #if 0 // The following optimization isn't particularly useful. if (mark->has_monitor() && mark->monitor()->is_entered(THREAD)) { lock->set_displaced_header (NULL) ; return ; } #endif // The object header will never be displaced to this lock, // so it does not matter what the value is, except that it // must be non-zero to avoid looking like a re-entrant lock, // and must not look locked either. lock->set_displaced_header(markOopDesc::unused_mark()); ObjectSynchronizer::inflate(THREAD, obj())->enter(THREAD); } // This routine is used to handle interpreter/compiler slow case // We don't need to use fast path here, because it must have // failed in the interpreter/compiler code. Simply use the heavy // weight monitor should be ok, unless someone find otherwise. void ObjectSynchronizer::slow_exit(oop object, BasicLock* lock, TRAPS) { fast_exit (object, lock, THREAD) ; } // NOTE: must use heavy weight monitor to handle jni monitor enter void ObjectSynchronizer::jni_enter(Handle obj, TRAPS) { // possible entry from jni enter // the current locking is from JNI instead of Java code TEVENT (jni_enter) ; if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } THREAD->set_current_pending_monitor_is_from_java(false); ObjectSynchronizer::inflate(THREAD, obj())->enter(THREAD); THREAD->set_current_pending_monitor_is_from_java(true); } // NOTE: must use heavy weight monitor to handle jni monitor enter bool ObjectSynchronizer::jni_try_enter(Handle obj, Thread* THREAD) { if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } ObjectMonitor* monitor = ObjectSynchronizer::inflate_helper(obj()); return monitor->try_enter(THREAD); } // NOTE: must use heavy weight monitor to handle jni monitor exit void ObjectSynchronizer::jni_exit(oop obj, Thread* THREAD) { TEVENT (jni_exit) ; if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); } assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj); // If this thread has locked the object, exit the monitor. Note: can't use // monitor->check(CHECK); must exit even if an exception is pending. if (monitor->check(THREAD)) { monitor->exit(THREAD); } } // complete_exit()/reenter() are used to wait on a nested lock // i.e. to give up an outer lock completely and then re-enter // Used when holding nested locks - lock acquisition order: lock1 then lock2 // 1) complete_exit lock1 - saving recursion count // 2) wait on lock2 // 3) when notified on lock2, unlock lock2 // 4) reenter lock1 with original recursion count // 5) lock lock2 // NOTE: must use heavy weight monitor to handle complete_exit/reenter() intptr_t ObjectSynchronizer::complete_exit(Handle obj, TRAPS) { TEVENT (complete_exit) ; if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj()); return monitor->complete_exit(THREAD); } // NOTE: must use heavy weight monitor to handle complete_exit/reenter() void ObjectSynchronizer::reenter(Handle obj, intptr_t recursion, TRAPS) { TEVENT (reenter) ; if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj()); monitor->reenter(recursion, THREAD); } // This exists only as a workaround of dtrace bug 6254741 int dtrace_waited_probe(ObjectMonitor* monitor, Handle obj, Thread* thr) { DTRACE_MONITOR_PROBE(waited, monitor, obj(), thr); return 0; } // NOTE: must use heavy weight monitor to handle wait() void ObjectSynchronizer::wait(Handle obj, jlong millis, TRAPS) { if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } if (millis < 0) { TEVENT (wait - throw IAX) ; THROW_MSG(vmSymbols::java_lang_IllegalArgumentException(), "timeout value is negative"); } ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj()); DTRACE_MONITOR_WAIT_PROBE(monitor, obj(), THREAD, millis); monitor->wait(millis, true, THREAD); /* This dummy call is in place to get around dtrace bug 6254741. Once that's fixed we can uncomment the following line and remove the call */ // DTRACE_MONITOR_PROBE(waited, monitor, obj(), THREAD); dtrace_waited_probe(monitor, obj, THREAD); } void ObjectSynchronizer::waitUninterruptibly (Handle obj, jlong millis, TRAPS) { if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } if (millis < 0) { TEVENT (wait - throw IAX) ; THROW_MSG(vmSymbols::java_lang_IllegalArgumentException(), "timeout value is negative"); } ObjectSynchronizer::inflate(THREAD, obj()) -> wait(millis, false, THREAD) ; } void ObjectSynchronizer::notify(Handle obj, TRAPS) { if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } markOop mark = obj->mark(); if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) { return; } ObjectSynchronizer::inflate(THREAD, obj())->notify(THREAD); } // NOTE: see comment of notify() void ObjectSynchronizer::notifyall(Handle obj, TRAPS) { if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(obj, false, THREAD); assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } markOop mark = obj->mark(); if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) { return; } ObjectSynchronizer::inflate(THREAD, obj())->notifyAll(THREAD); } intptr_t ObjectSynchronizer::FastHashCode (Thread * Self, oop obj) { if (UseBiasedLocking) { // NOTE: many places throughout the JVM do not expect a safepoint // to be taken here, in particular most operations on perm gen // objects. However, we only ever bias Java instances and all of // the call sites of identity_hash that might revoke biases have // been checked to make sure they can handle a safepoint. The // added check of the bias pattern is to avoid useless calls to // thread-local storage. if (obj->mark()->has_bias_pattern()) { // Box and unbox the raw reference just in case we cause a STW safepoint. Handle hobj (Self, obj) ; // Relaxing assertion for bug 6320749. assert (Universe::verify_in_progress() || !SafepointSynchronize::is_at_safepoint(), "biases should not be seen by VM thread here"); BiasedLocking::revoke_and_rebias(hobj, false, JavaThread::current()); obj = hobj() ; assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } } // hashCode() is a heap mutator ... // Relaxing assertion for bug 6320749. assert (Universe::verify_in_progress() || !SafepointSynchronize::is_at_safepoint(), "invariant") ; assert (Universe::verify_in_progress() || Self->is_Java_thread() , "invariant") ; assert (Universe::verify_in_progress() || ((JavaThread *)Self)->thread_state() != _thread_blocked, "invariant") ; ObjectMonitor* monitor = NULL; markOop temp, test; intptr_t hash; markOop mark = ReadStableMark (obj); // object should remain ineligible for biased locking assert (!mark->has_bias_pattern(), "invariant") ; if (mark->is_neutral()) { hash = mark->hash(); // this is a normal header if (hash) { // if it has hash, just return it return hash; } hash = get_next_hash(Self, obj); // allocate a new hash code temp = mark->copy_set_hash(hash); // merge the hash code into header // use (machine word version) atomic operation to install the hash test = (markOop) Atomic::cmpxchg_ptr(temp, obj->mark_addr(), mark); if (test == mark) { return hash; } // If atomic operation failed, we must inflate the header // into heavy weight monitor. We could add more code here // for fast path, but it does not worth the complexity. } else if (mark->has_monitor()) { monitor = mark->monitor(); temp = monitor->header(); assert (temp->is_neutral(), "invariant") ; hash = temp->hash(); if (hash) { return hash; } // Skip to the following code to reduce code size } else if (Self->is_lock_owned((address)mark->locker())) { temp = mark->displaced_mark_helper(); // this is a lightweight monitor owned assert (temp->is_neutral(), "invariant") ; hash = temp->hash(); // by current thread, check if the displaced if (hash) { // header contains hash code return hash; } // WARNING: // The displaced header is strictly immutable. // It can NOT be changed in ANY cases. So we have // to inflate the header into heavyweight monitor // even the current thread owns the lock. The reason // is the BasicLock (stack slot) will be asynchronously // read by other threads during the inflate() function. // Any change to stack may not propagate to other threads // correctly. } // Inflate the monitor to set hash code monitor = ObjectSynchronizer::inflate(Self, obj); // Load displaced header and check it has hash code mark = monitor->header(); assert (mark->is_neutral(), "invariant") ; hash = mark->hash(); if (hash == 0) { hash = get_next_hash(Self, obj); temp = mark->copy_set_hash(hash); // merge hash code into header assert (temp->is_neutral(), "invariant") ; test = (markOop) Atomic::cmpxchg_ptr(temp, monitor, mark); if (test != mark) { // The only update to the header in the monitor (outside GC) // is install the hash code. If someone add new usage of // displaced header, please update this code hash = test->hash(); assert (test->is_neutral(), "invariant") ; assert (hash != 0, "Trivial unexpected object/monitor header usage."); } } // We finally get the hash return hash; } // Deprecated -- use FastHashCode() instead. intptr_t ObjectSynchronizer::identity_hash_value_for(Handle obj) { return FastHashCode (Thread::current(), obj()) ; } bool ObjectSynchronizer::current_thread_holds_lock(JavaThread* thread, Handle h_obj) { if (UseBiasedLocking) { BiasedLocking::revoke_and_rebias(h_obj, false, thread); assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } assert(thread == JavaThread::current(), "Can only be called on current thread"); oop obj = h_obj(); markOop mark = ReadStableMark (obj) ; // Uncontended case, header points to stack if (mark->has_locker()) { return thread->is_lock_owned((address)mark->locker()); } // Contended case, header points to ObjectMonitor (tagged pointer) if (mark->has_monitor()) { ObjectMonitor* monitor = mark->monitor(); return monitor->is_entered(thread) != 0 ; } // Unlocked case, header in place assert(mark->is_neutral(), "sanity check"); return false; } // Be aware of this method could revoke bias of the lock object. // This method querys the ownership of the lock handle specified by 'h_obj'. // If the current thread owns the lock, it returns owner_self. If no // thread owns the lock, it returns owner_none. Otherwise, it will return // ower_other. ObjectSynchronizer::LockOwnership ObjectSynchronizer::query_lock_ownership (JavaThread *self, Handle h_obj) { // The caller must beware this method can revoke bias, and // revocation can result in a safepoint. assert (!SafepointSynchronize::is_at_safepoint(), "invariant") ; assert (self->thread_state() != _thread_blocked , "invariant") ; // Possible mark states: neutral, biased, stack-locked, inflated if (UseBiasedLocking && h_obj()->mark()->has_bias_pattern()) { // CASE: biased BiasedLocking::revoke_and_rebias(h_obj, false, self); assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } assert(self == JavaThread::current(), "Can only be called on current thread"); oop obj = h_obj(); markOop mark = ReadStableMark (obj) ; // CASE: stack-locked. Mark points to a BasicLock on the owner's stack. if (mark->has_locker()) { return self->is_lock_owned((address)mark->locker()) ? owner_self : owner_other; } // CASE: inflated. Mark (tagged pointer) points to an objectMonitor. // The Object:ObjectMonitor relationship is stable as long as we're // not at a safepoint. if (mark->has_monitor()) { void * owner = mark->monitor()->_owner ; if (owner == NULL) return owner_none ; return (owner == self || self->is_lock_owned((address)owner)) ? owner_self : owner_other; } // CASE: neutral assert(mark->is_neutral(), "sanity check"); return owner_none ; // it's unlocked } // FIXME: jvmti should call this JavaThread* ObjectSynchronizer::get_lock_owner(Handle h_obj, bool doLock) { if (UseBiasedLocking) { if (SafepointSynchronize::is_at_safepoint()) { BiasedLocking::revoke_at_safepoint(h_obj); } else { BiasedLocking::revoke_and_rebias(h_obj, false, JavaThread::current()); } assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now"); } oop obj = h_obj(); address owner = NULL; markOop mark = ReadStableMark (obj) ; // Uncontended case, header points to stack if (mark->has_locker()) { owner = (address) mark->locker(); } // Contended case, header points to ObjectMonitor (tagged pointer) if (mark->has_monitor()) { ObjectMonitor* monitor = mark->monitor(); assert(monitor != NULL, "monitor should be non-null"); owner = (address) monitor->owner(); } if (owner != NULL) { return Threads::owning_thread_from_monitor_owner(owner, doLock); } // Unlocked case, header in place // Cannot have assertion since this object may have been // locked by another thread when reaching here. // assert(mark->is_neutral(), "sanity check"); return NULL; } // Iterate through monitor cache and attempt to release thread's monitors // Gives up on a particular monitor if an exception occurs, but continues // the overall iteration, swallowing the exception. class ReleaseJavaMonitorsClosure: public MonitorClosure { private: TRAPS; public: ReleaseJavaMonitorsClosure(Thread* thread) : THREAD(thread) {} void do_monitor(ObjectMonitor* mid) { if (mid->owner() == THREAD) { (void)mid->complete_exit(CHECK); } } }; // Release all inflated monitors owned by THREAD. Lightweight monitors are // ignored. This is meant to be called during JNI thread detach which assumes // all remaining monitors are heavyweight. All exceptions are swallowed. // Scanning the extant monitor list can be time consuming. // A simple optimization is to add a per-thread flag that indicates a thread // called jni_monitorenter() during its lifetime. // // Instead of No_Savepoint_Verifier it might be cheaper to // use an idiom of the form: // auto int tmp = SafepointSynchronize::_safepoint_counter ; // <code that must not run at safepoint> // guarantee (((tmp ^ _safepoint_counter) | (tmp & 1)) == 0) ; // Since the tests are extremely cheap we could leave them enabled // for normal product builds. void ObjectSynchronizer::release_monitors_owned_by_thread(TRAPS) { assert(THREAD == JavaThread::current(), "must be current Java thread"); No_Safepoint_Verifier nsv ; ReleaseJavaMonitorsClosure rjmc(THREAD); Thread::muxAcquire(&ListLock, "release_monitors_owned_by_thread"); ObjectSynchronizer::monitors_iterate(&rjmc); Thread::muxRelease(&ListLock); THREAD->clear_pending_exception(); } // Visitors ... void ObjectSynchronizer::monitors_iterate(MonitorClosure* closure) { ObjectMonitor* block = gBlockList; ObjectMonitor* mid; while (block) { assert(block->object() == CHAINMARKER, "must be a block header"); for (int i = _BLOCKSIZE - 1; i > 0; i--) { mid = block + i; oop object = (oop) mid->object(); if (object != NULL) { closure->do_monitor(mid); } } block = (ObjectMonitor*) block->FreeNext; } } void ObjectSynchronizer::oops_do(OopClosure* f) { assert(SafepointSynchronize::is_at_safepoint(), "must be at safepoint"); for (ObjectMonitor* block = gBlockList; block != NULL; block = next(block)) { assert(block->object() == CHAINMARKER, "must be a block header"); for (int i = 1; i < _BLOCKSIZE; i++) { ObjectMonitor* mid = &block[i]; if (mid->object() != NULL) { f->do_oop((oop*)mid->object_addr()); } } } } // Deflate_idle_monitors() is called at all safepoints, immediately // after all mutators are stopped, but before any objects have moved. // It traverses the list of known monitors, deflating where possible. // The scavenged monitor are returned to the monitor free list. // // Beware that we scavenge at *every* stop-the-world point. // Having a large number of monitors in-circulation negatively // impacts the performance of some applications (e.g., PointBase). // Broadly, we want to minimize the # of monitors in circulation. // Alternately, we could partition the active monitors into sub-lists // of those that need scanning and those that do not. // Specifically, we would add a new sub-list of objectmonitors // that are in-circulation and potentially active. deflate_idle_monitors() // would scan only that list. Other monitors could reside on a quiescent // list. Such sequestered monitors wouldn't need to be scanned by // deflate_idle_monitors(). omAlloc() would first check the global free list, // then the quiescent list, and, failing those, would allocate a new block. // Deflate_idle_monitors() would scavenge and move monitors to the // quiescent list. // // Perversely, the heap size -- and thus the STW safepoint rate -- // typically drives the scavenge rate. Large heaps can mean infrequent GC, // which in turn can mean large(r) numbers of objectmonitors in circulation. // This is an unfortunate aspect of this design. // // Another refinement would be to refrain from calling deflate_idle_monitors() // except at stop-the-world points associated with garbage collections. // // An even better solution would be to deflate on-the-fly, aggressively, // at monitorexit-time as is done in EVM's metalock or Relaxed Locks. void ObjectSynchronizer::deflate_idle_monitors() { assert(SafepointSynchronize::is_at_safepoint(), "must be at safepoint"); int nInuse = 0 ; // currently associated with objects int nInCirculation = 0 ; // extant int nScavenged = 0 ; // reclaimed ObjectMonitor * FreeHead = NULL ; // Local SLL of scavenged monitors ObjectMonitor * FreeTail = NULL ; // Iterate over all extant monitors - Scavenge all idle monitors. TEVENT (deflate_idle_monitors) ; for (ObjectMonitor* block = gBlockList; block != NULL; block = next(block)) { assert(block->object() == CHAINMARKER, "must be a block header"); nInCirculation += _BLOCKSIZE ; for (int i = 1 ; i < _BLOCKSIZE; i++) { ObjectMonitor* mid = &block[i]; oop obj = (oop) mid->object(); if (obj == NULL) { // The monitor is not associated with an object. // The monitor should either be a thread-specific private // free list or the global free list. // obj == NULL IMPLIES mid->is_busy() == 0 guarantee (!mid->is_busy(), "invariant") ; continue ; } // Normal case ... The monitor is associated with obj. guarantee (obj->mark() == markOopDesc::encode(mid), "invariant") ; guarantee (mid == obj->mark()->monitor(), "invariant"); guarantee (mid->header()->is_neutral(), "invariant"); if (mid->is_busy()) { if (ClearResponsibleAtSTW) mid->_Responsible = NULL ; nInuse ++ ; } else { // Deflate the monitor if it is no longer being used // It's idle - scavenge and return to the global free list // plain old deflation ... TEVENT (deflate_idle_monitors - scavenge1) ; if (TraceMonitorInflation) { if (obj->is_instance()) { ResourceMark rm; tty->print_cr("Deflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s", (intptr_t) obj, (intptr_t) obj->mark(), Klass::cast(obj->klass())->external_name()); } } // Restore the header back to obj obj->release_set_mark(mid->header()); mid->clear(); assert (mid->object() == NULL, "invariant") ; // Move the object to the working free list defined by FreeHead,FreeTail. mid->FreeNext = NULL ; if (FreeHead == NULL) FreeHead = mid ; if (FreeTail != NULL) FreeTail->FreeNext = mid ; FreeTail = mid ; nScavenged ++ ; } } } // Move the scavenged monitors back to the global free list. // In theory we don't need the freelist lock as we're at a STW safepoint. // omAlloc() and omFree() can only be called while a thread is _not in safepoint state. // But it's remotely possible that omFlush() or release_monitors_owned_by_thread() // might be called while not at a global STW safepoint. In the interest of // safety we protect the following access with ListLock. // An even more conservative and prudent approach would be to guard // the main loop in scavenge_idle_monitors() with ListLock. if (FreeHead != NULL) { guarantee (FreeTail != NULL && nScavenged > 0, "invariant") ; assert (FreeTail->FreeNext == NULL, "invariant") ; // constant-time list splice - prepend scavenged segment to gFreeList Thread::muxAcquire (&ListLock, "scavenge - return") ; FreeTail->FreeNext = gFreeList ; gFreeList = FreeHead ; Thread::muxRelease (&ListLock) ; } if (_sync_Deflations != NULL) _sync_Deflations->inc(nScavenged) ; if (_sync_MonExtant != NULL) _sync_MonExtant ->set_value(nInCirculation); // TODO: Add objectMonitor leak detection. // Audit/inventory the objectMonitors -- make sure they're all accounted for. GVars.stwRandom = os::random() ; GVars.stwCycle ++ ; } // A macro is used below because there may already be a pending // exception which should not abort the execution of the routines // which use this (which is why we don't put this into check_slow and // call it with a CHECK argument). #define CHECK_OWNER() \ do { \ if (THREAD != _owner) { \ if (THREAD->is_lock_owned((address) _owner)) { \ _owner = THREAD ; /* Convert from basiclock addr to Thread addr */ \ _recursions = 0; \ OwnerIsThread = 1 ; \ } else { \ TEVENT (Throw IMSX) ; \ THROW(vmSymbols::java_lang_IllegalMonitorStateException()); \ } \ } \ } while (false) // TODO-FIXME: eliminate ObjectWaiters. Replace this visitor/enumerator // interface with a simple FirstWaitingThread(), NextWaitingThread() interface. ObjectWaiter* ObjectMonitor::first_waiter() { return _WaitSet; } ObjectWaiter* ObjectMonitor::next_waiter(ObjectWaiter* o) { return o->_next; } Thread* ObjectMonitor::thread_of_waiter(ObjectWaiter* o) { return o->_thread; } // initialize the monitor, exception the semaphore, all other fields // are simple integers or pointers ObjectMonitor::ObjectMonitor() { _header = NULL; _count = 0; _waiters = 0, _recursions = 0; _object = NULL; _owner = NULL; _WaitSet = NULL; _WaitSetLock = 0 ; _Responsible = NULL ; _succ = NULL ; _cxq = NULL ; FreeNext = NULL ; _EntryList = NULL ; _SpinFreq = 0 ; _SpinClock = 0 ; OwnerIsThread = 0 ; } ObjectMonitor::~ObjectMonitor() { // TODO: Add asserts ... // _cxq == 0 _succ == NULL _owner == NULL _waiters == 0 // _count == 0 _EntryList == NULL etc } intptr_t ObjectMonitor::is_busy() const { // TODO-FIXME: merge _count and _waiters. // TODO-FIXME: assert _owner == null implies _recursions = 0 // TODO-FIXME: assert _WaitSet != null implies _count > 0 return _count|_waiters|intptr_t(_owner)|intptr_t(_cxq)|intptr_t(_EntryList ) ; } void ObjectMonitor::Recycle () { // TODO: add stronger asserts ... // _cxq == 0 _succ == NULL _owner == NULL _waiters == 0 // _count == 0 EntryList == NULL // _recursions == 0 _WaitSet == NULL // TODO: assert (is_busy()|_recursions) == 0 _succ = NULL ; _EntryList = NULL ; _cxq = NULL ; _WaitSet = NULL ; _recursions = 0 ; _SpinFreq = 0 ; _SpinClock = 0 ; OwnerIsThread = 0 ; } // WaitSet management ... inline void ObjectMonitor::AddWaiter(ObjectWaiter* node) { assert(node != NULL, "should not dequeue NULL node"); assert(node->_prev == NULL, "node already in list"); assert(node->_next == NULL, "node already in list"); // put node at end of queue (circular doubly linked list) if (_WaitSet == NULL) { _WaitSet = node; node->_prev = node; node->_next = node; } else { ObjectWaiter* head = _WaitSet ; ObjectWaiter* tail = head->_prev; assert(tail->_next == head, "invariant check"); tail->_next = node; head->_prev = node; node->_next = head; node->_prev = tail; } } inline ObjectWaiter* ObjectMonitor::DequeueWaiter() { // dequeue the very first waiter ObjectWaiter* waiter = _WaitSet; if (waiter) { DequeueSpecificWaiter(waiter); } return waiter; } inline void ObjectMonitor::DequeueSpecificWaiter(ObjectWaiter* node) { assert(node != NULL, "should not dequeue NULL node"); assert(node->_prev != NULL, "node already removed from list"); assert(node->_next != NULL, "node already removed from list"); // when the waiter has woken up because of interrupt, // timeout or other spurious wake-up, dequeue the // waiter from waiting list ObjectWaiter* next = node->_next; if (next == node) { assert(node->_prev == node, "invariant check"); _WaitSet = NULL; } else { ObjectWaiter* prev = node->_prev; assert(prev->_next == node, "invariant check"); assert(next->_prev == node, "invariant check"); next->_prev = prev; prev->_next = next; if (_WaitSet == node) { _WaitSet = next; } } node->_next = NULL; node->_prev = NULL; } static char * kvGet (char * kvList, const char * Key) { if (kvList == NULL) return NULL ; size_t n = strlen (Key) ; char * Search ; for (Search = kvList ; *Search ; Search += strlen(Search) + 1) { if (strncmp (Search, Key, n) == 0) { if (Search[n] == '=') return Search + n + 1 ; if (Search[n] == 0) return (char *) "1" ; } } return NULL ; } static int kvGetInt (char * kvList, const char * Key, int Default) { char * v = kvGet (kvList, Key) ; int rslt = v ? ::strtol (v, NULL, 0) : Default ; if (Knob_ReportSettings && v != NULL) { ::printf (" SyncKnob: %s %d(%d)\n", Key, rslt, Default) ; ::fflush (stdout) ; } return rslt ; } // By convention we unlink a contending thread from EntryList|cxq immediately // after the thread acquires the lock in ::enter(). Equally, we could defer // unlinking the thread until ::exit()-time. void ObjectMonitor::UnlinkAfterAcquire (Thread * Self, ObjectWaiter * SelfNode) { assert (_owner == Self, "invariant") ; assert (SelfNode->_thread == Self, "invariant") ; if (SelfNode->TState == ObjectWaiter::TS_ENTER) { // Normal case: remove Self from the DLL EntryList . // This is a constant-time operation. ObjectWaiter * nxt = SelfNode->_next ; ObjectWaiter * prv = SelfNode->_prev ; if (nxt != NULL) nxt->_prev = prv ; if (prv != NULL) prv->_next = nxt ; if (SelfNode == _EntryList ) _EntryList = nxt ; assert (nxt == NULL || nxt->TState == ObjectWaiter::TS_ENTER, "invariant") ; assert (prv == NULL || prv->TState == ObjectWaiter::TS_ENTER, "invariant") ; TEVENT (Unlink from EntryList) ; } else { guarantee (SelfNode->TState == ObjectWaiter::TS_CXQ, "invariant") ; // Inopportune interleaving -- Self is still on the cxq. // This usually means the enqueue of self raced an exiting thread. // Normally we'll find Self near the front of the cxq, so // dequeueing is typically fast. If needbe we can accelerate // this with some MCS/CHL-like bidirectional list hints and advisory // back-links so dequeueing from the interior will normally operate // in constant-time. // Dequeue Self from either the head (with CAS) or from the interior // with a linear-time scan and normal non-atomic memory operations. // CONSIDER: if Self is on the cxq then simply drain cxq into EntryList // and then unlink Self from EntryList. We have to drain eventually, // so it might as well be now. ObjectWaiter * v = _cxq ; assert (v != NULL, "invariant") ; if (v != SelfNode || Atomic::cmpxchg_ptr (SelfNode->_next, &_cxq, v) != v) { // The CAS above can fail from interference IFF a "RAT" arrived. // In that case Self must be in the interior and can no longer be // at the head of cxq. if (v == SelfNode) { assert (_cxq != v, "invariant") ; v = _cxq ; // CAS above failed - start scan at head of list } ObjectWaiter * p ; ObjectWaiter * q = NULL ; for (p = v ; p != NULL && p != SelfNode; p = p->_next) { q = p ; assert (p->TState == ObjectWaiter::TS_CXQ, "invariant") ; } assert (v != SelfNode, "invariant") ; assert (p == SelfNode, "Node not found on cxq") ; assert (p != _cxq, "invariant") ; assert (q != NULL, "invariant") ; assert (q->_next == p, "invariant") ; q->_next = p->_next ; } TEVENT (Unlink from cxq) ; } // Diagnostic hygiene ... SelfNode->_prev = (ObjectWaiter *) 0xBAD ; SelfNode->_next = (ObjectWaiter *) 0xBAD ; SelfNode->TState = ObjectWaiter::TS_RUN ; } // Caveat: TryLock() is not necessarily serializing if it returns failure. // Callers must compensate as needed. int ObjectMonitor::TryLock (Thread * Self) { for (;;) { void * own = _owner ; if (own != NULL) return 0 ; if (Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) { // Either guarantee _recursions == 0 or set _recursions = 0. assert (_recursions == 0, "invariant") ; assert (_owner == Self, "invariant") ; // CONSIDER: set or assert that OwnerIsThread == 1 return 1 ; } // The lock had been free momentarily, but we lost the race to the lock. // Interference -- the CAS failed. // We can either return -1 or retry. // Retry doesn't make as much sense because the lock was just acquired. if (true) return -1 ; } } // NotRunnable() -- informed spinning // // Don't bother spinning if the owner is not eligible to drop the lock. // Peek at the owner's schedctl.sc_state and Thread._thread_values and // spin only if the owner thread is _thread_in_Java or _thread_in_vm. // The thread must be runnable in order to drop the lock in timely fashion. // If the _owner is not runnable then spinning will not likely be // successful (profitable). // // Beware -- the thread referenced by _owner could have died // so a simply fetch from _owner->_thread_state might trap. // Instead, we use SafeFetchXX() to safely LD _owner->_thread_state. // Because of the lifecycle issues the schedctl and _thread_state values // observed by NotRunnable() might be garbage. NotRunnable must // tolerate this and consider the observed _thread_state value // as advisory. // // Beware too, that _owner is sometimes a BasicLock address and sometimes // a thread pointer. We differentiate the two cases with OwnerIsThread. // Alternately, we might tag the type (thread pointer vs basiclock pointer) // with the LSB of _owner. Another option would be to probablistically probe // the putative _owner->TypeTag value. // // Checking _thread_state isn't perfect. Even if the thread is // in_java it might be blocked on a page-fault or have been preempted // and sitting on a ready/dispatch queue. _thread state in conjunction // with schedctl.sc_state gives us a good picture of what the // thread is doing, however. // // TODO: check schedctl.sc_state. // We'll need to use SafeFetch32() to read from the schedctl block. // See RFE #5004247 and http://sac.sfbay.sun.com/Archives/CaseLog/arc/PSARC/2005/351/ // // The return value from NotRunnable() is *advisory* -- the // result is based on sampling and is not necessarily coherent. // The caller must tolerate false-negative and false-positive errors. // Spinning, in general, is probabilistic anyway. int ObjectMonitor::NotRunnable (Thread * Self, Thread * ox) { // Check either OwnerIsThread or ox->TypeTag == 2BAD. if (!OwnerIsThread) return 0 ; if (ox == NULL) return 0 ; // Avoid transitive spinning ... // Say T1 spins or blocks trying to acquire L. T1._Stalled is set to L. // Immediately after T1 acquires L it's possible that T2, also // spinning on L, will see L.Owner=T1 and T1._Stalled=L. // This occurs transiently after T1 acquired L but before // T1 managed to clear T1.Stalled. T2 does not need to abort // its spin in this circumstance. intptr_t BlockedOn = SafeFetchN ((intptr_t *) &ox->_Stalled, intptr_t(1)) ; if (BlockedOn == 1) return 1 ; if (BlockedOn != 0) { return BlockedOn != intptr_t(this) && _owner == ox ; } assert (sizeof(((JavaThread *)ox)->_thread_state == sizeof(int)), "invariant") ; int jst = SafeFetch32 ((int *) &((JavaThread *) ox)->_thread_state, -1) ; ; // consider also: jst != _thread_in_Java -- but that's overspecific. return jst == _thread_blocked || jst == _thread_in_native ; } // Adaptive spin-then-block - rational spinning // // Note that we spin "globally" on _owner with a classic SMP-polite TATAS // algorithm. On high order SMP systems it would be better to start with // a brief global spin and then revert to spinning locally. In the spirit of MCS/CLH, // a contending thread could enqueue itself on the cxq and then spin locally // on a thread-specific variable such as its ParkEvent._Event flag. // That's left as an exercise for the reader. Note that global spinning is // not problematic on Niagara, as the L2$ serves the interconnect and has both // low latency and massive bandwidth. // // Broadly, we can fix the spin frequency -- that is, the % of contended lock // acquisition attempts where we opt to spin -- at 100% and vary the spin count // (duration) or we can fix the count at approximately the duration of // a context switch and vary the frequency. Of course we could also // vary both satisfying K == Frequency * Duration, where K is adaptive by monitor. // See http://j2se.east/~dice/PERSIST/040824-AdaptiveSpinning.html. // // This implementation varies the duration "D", where D varies with // the success rate of recent spin attempts. (D is capped at approximately // length of a round-trip context switch). The success rate for recent // spin attempts is a good predictor of the success rate of future spin // attempts. The mechanism adapts automatically to varying critical // section length (lock modality), system load and degree of parallelism. // D is maintained per-monitor in _SpinDuration and is initialized // optimistically. Spin frequency is fixed at 100%. // // Note that _SpinDuration is volatile, but we update it without locks // or atomics. The code is designed so that _SpinDuration stays within // a reasonable range even in the presence of races. The arithmetic // operations on _SpinDuration are closed over the domain of legal values, // so at worst a race will install and older but still legal value. // At the very worst this introduces some apparent non-determinism. // We might spin when we shouldn't or vice-versa, but since the spin // count are relatively short, even in the worst case, the effect is harmless. // // Care must be taken that a low "D" value does not become an // an absorbing state. Transient spinning failures -- when spinning // is overall profitable -- should not cause the system to converge // on low "D" values. We want spinning to be stable and predictable // and fairly responsive to change and at the same time we don't want // it to oscillate, become metastable, be "too" non-deterministic, // or converge on or enter undesirable stable absorbing states. // // We implement a feedback-based control system -- using past behavior // to predict future behavior. We face two issues: (a) if the // input signal is random then the spin predictor won't provide optimal // results, and (b) if the signal frequency is too high then the control // system, which has some natural response lag, will "chase" the signal. // (b) can arise from multimodal lock hold times. Transient preemption // can also result in apparent bimodal lock hold times. // Although sub-optimal, neither condition is particularly harmful, as // in the worst-case we'll spin when we shouldn't or vice-versa. // The maximum spin duration is rather short so the failure modes aren't bad. // To be conservative, I've tuned the gain in system to bias toward // _not spinning. Relatedly, the system can sometimes enter a mode where it // "rings" or oscillates between spinning and not spinning. This happens // when spinning is just on the cusp of profitability, however, so the // situation is not dire. The state is benign -- there's no need to add // hysteresis control to damp the transition rate between spinning and // not spinning. // // - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - // // Spin-then-block strategies ... // // Thoughts on ways to improve spinning : // // * Periodically call {psr_}getloadavg() while spinning, and // permit unbounded spinning if the load average is < // the number of processors. Beware, however, that getloadavg() // is exceptionally fast on solaris (about 1/10 the cost of a full // spin cycle, but quite expensive on linux. Beware also, that // multiple JVMs could "ring" or oscillate in a feedback loop. // Sufficient damping would solve that problem. // // * We currently use spin loops with iteration counters to approximate // spinning for some interval. Given the availability of high-precision // time sources such as gethrtime(), %TICK, %STICK, RDTSC, etc., we should // someday reimplement the spin loops to duration-based instead of iteration-based. // // * Don't spin if there are more than N = (CPUs/2) threads // currently spinning on the monitor (or globally). // That is, limit the number of concurrent spinners. // We might also limit the # of spinners in the JVM, globally. // // * If a spinning thread observes _owner change hands it should // abort the spin (and park immediately) or at least debit // the spin counter by a large "penalty". // // * Classically, the spin count is either K*(CPUs-1) or is a // simple constant that approximates the length of a context switch. // We currently use a value -- computed by a special utility -- that // approximates round-trip context switch times. // // * Normally schedctl_start()/_stop() is used to advise the kernel // to avoid preempting threads that are running in short, bounded // critical sections. We could use the schedctl hooks in an inverted // sense -- spinners would set the nopreempt flag, but poll the preempt // pending flag. If a spinner observed a pending preemption it'd immediately // abort the spin and park. As such, the schedctl service acts as // a preemption warning mechanism. // // * In lieu of spinning, if the system is running below saturation // (that is, loadavg() << #cpus), we can instead suppress futile // wakeup throttling, or even wake more than one successor at exit-time. // The net effect is largely equivalent to spinning. In both cases, // contending threads go ONPROC and opportunistically attempt to acquire // the lock, decreasing lock handover latency at the expense of wasted // cycles and context switching. // // * We might to spin less after we've parked as the thread will // have less $ and TLB affinity with the processor. // Likewise, we might spin less if we come ONPROC on a different // processor or after a long period (>> rechose_interval). // // * A table-driven state machine similar to Solaris' dispadmin scheduling // tables might be a better design. Instead of encoding information in // _SpinDuration, _SpinFreq and _SpinClock we'd just use explicit, // discrete states. Success or failure during a spin would drive // state transitions, and each state node would contain a spin count. // // * If the processor is operating in a mode intended to conserve power // (such as Intel's SpeedStep) or to reduce thermal output (thermal // step-down mode) then the Java synchronization subsystem should // forgo spinning. // // * The minimum spin duration should be approximately the worst-case // store propagation latency on the platform. That is, the time // it takes a store on CPU A to become visible on CPU B, where A and // B are "distant". // // * We might want to factor a thread's priority in the spin policy. // Threads with a higher priority might spin for slightly longer. // Similarly, if we use back-off in the TATAS loop, lower priority // threads might back-off longer. We don't currently use a // thread's priority when placing it on the entry queue. We may // want to consider doing so in future releases. // // * We might transiently drop a thread's scheduling priority while it spins. // SCHED_BATCH on linux and FX scheduling class at priority=0 on Solaris // would suffice. We could even consider letting the thread spin indefinitely at // a depressed or "idle" priority. This brings up fairness issues, however -- // in a saturated system a thread would with a reduced priority could languish // for extended periods on the ready queue. // // * While spinning try to use the otherwise wasted time to help the VM make // progress: // // -- YieldTo() the owner, if the owner is OFFPROC but ready // Done our remaining quantum directly to the ready thread. // This helps "push" the lock owner through the critical section. // It also tends to improve affinity/locality as the lock // "migrates" less frequently between CPUs. // -- Walk our own stack in anticipation of blocking. Memoize the roots. // -- Perform strand checking for other thread. Unpark potential strandees. // -- Help GC: trace or mark -- this would need to be a bounded unit of work. // Unfortunately this will pollute our $ and TLBs. Recall that we // spin to avoid context switching -- context switching has an // immediate cost in latency, a disruptive cost to other strands on a CMT // processor, and an amortized cost because of the D$ and TLB cache // reload transient when the thread comes back ONPROC and repopulates // $s and TLBs. // -- call getloadavg() to see if the system is saturated. It'd probably // make sense to call getloadavg() half way through the spin. // If the system isn't at full capacity the we'd simply reset // the spin counter to and extend the spin attempt. // -- Doug points out that we should use the same "helping" policy // in thread.yield(). // // * Try MONITOR-MWAIT on systems that support those instructions. // // * The spin statistics that drive spin decisions & frequency are // maintained in the objectmonitor structure so if we deflate and reinflate // we lose spin state. In practice this is not usually a concern // as the default spin state after inflation is aggressive (optimistic) // and tends toward spinning. So in the worst case for a lock where // spinning is not profitable we may spin unnecessarily for a brief // period. But then again, if a lock is contended it'll tend not to deflate // in the first place. intptr_t ObjectMonitor::SpinCallbackArgument = 0 ; int (*ObjectMonitor::SpinCallbackFunction)(intptr_t, int) = NULL ; // Spinning: Fixed frequency (100%), vary duration int ObjectMonitor::TrySpin_VaryDuration (Thread * Self) { // Dumb, brutal spin. Good for comparative measurements against adaptive spinning. int ctr = Knob_FixedSpin ; if (ctr != 0) { while (--ctr >= 0) { if (TryLock (Self) > 0) return 1 ; SpinPause () ; } return 0 ; } for (ctr = Knob_PreSpin + 1; --ctr >= 0 ; ) { if (TryLock(Self) > 0) { // Increase _SpinDuration ... // Note that we don't clamp SpinDuration precisely at SpinLimit. // Raising _SpurDuration to the poverty line is key. int x = _SpinDuration ; if (x < Knob_SpinLimit) { if (x < Knob_Poverty) x = Knob_Poverty ; _SpinDuration = x + Knob_BonusB ; } return 1 ; } SpinPause () ; } // Admission control - verify preconditions for spinning // // We always spin a little bit, just to prevent _SpinDuration == 0 from // becoming an absorbing state. Put another way, we spin briefly to // sample, just in case the system load, parallelism, contention, or lock // modality changed. // // Consider the following alternative: // Periodically set _SpinDuration = _SpinLimit and try a long/full // spin attempt. "Periodically" might mean after a tally of // the # of failed spin attempts (or iterations) reaches some threshold. // This takes us into the realm of 1-out-of-N spinning, where we // hold the duration constant but vary the frequency. ctr = _SpinDuration ; if (ctr < Knob_SpinBase) ctr = Knob_SpinBase ; if (ctr <= 0) return 0 ; if (Knob_SuccRestrict && _succ != NULL) return 0 ; if (Knob_OState && NotRunnable (Self, (Thread *) _owner)) { TEVENT (Spin abort - notrunnable [TOP]); return 0 ; } int MaxSpin = Knob_MaxSpinners ; if (MaxSpin >= 0) { if (_Spinner > MaxSpin) { TEVENT (Spin abort -- too many spinners) ; return 0 ; } // Slighty racy, but benign ... Adjust (&_Spinner, 1) ; } // We're good to spin ... spin ingress. // CONSIDER: use Prefetch::write() to avoid RTS->RTO upgrades // when preparing to LD...CAS _owner, etc and the CAS is likely // to succeed. int hits = 0 ; int msk = 0 ; int caspty = Knob_CASPenalty ; int oxpty = Knob_OXPenalty ; int sss = Knob_SpinSetSucc ; if (sss && _succ == NULL ) _succ = Self ; Thread * prv = NULL ; // There are three ways to exit the following loop: // 1. A successful spin where this thread has acquired the lock. // 2. Spin failure with prejudice // 3. Spin failure without prejudice while (--ctr >= 0) { // Periodic polling -- Check for pending GC // Threads may spin while they're unsafe. // We don't want spinning threads to delay the JVM from reaching // a stop-the-world safepoint or to steal cycles from GC. // If we detect a pending safepoint we abort in order that // (a) this thread, if unsafe, doesn't delay the safepoint, and (b) // this thread, if safe, doesn't steal cycles from GC. // This is in keeping with the "no loitering in runtime" rule. // We periodically check to see if there's a safepoint pending. if ((ctr & 0xFF) == 0) { if (SafepointSynchronize::do_call_back()) { TEVENT (Spin: safepoint) ; goto Abort ; // abrupt spin egress } if (Knob_UsePause & 1) SpinPause () ; int (*scb)(intptr_t,int) = SpinCallbackFunction ; if (hits > 50 && scb != NULL) { int abend = (*scb)(SpinCallbackArgument, 0) ; } } if (Knob_UsePause & 2) SpinPause() ; // Exponential back-off ... Stay off the bus to reduce coherency traffic. // This is useful on classic SMP systems, but is of less utility on // N1-style CMT platforms. // // Trade-off: lock acquisition latency vs coherency bandwidth. // Lock hold times are typically short. A histogram // of successful spin attempts shows that we usually acquire // the lock early in the spin. That suggests we want to // sample _owner frequently in the early phase of the spin, // but then back-off and sample less frequently as the spin // progresses. The back-off makes a good citizen on SMP big // SMP systems. Oversampling _owner can consume excessive // coherency bandwidth. Relatedly, if we _oversample _owner we // can inadvertently interfere with the the ST m->owner=null. // executed by the lock owner. if (ctr & msk) continue ; ++hits ; if ((hits & 0xF) == 0) { // The 0xF, above, corresponds to the exponent. // Consider: (msk+1)|msk msk = ((msk << 2)|3) & BackOffMask ; } // Probe _owner with TATAS // If this thread observes the monitor transition or flicker // from locked to unlocked to locked, then the odds that this // thread will acquire the lock in this spin attempt go down // considerably. The same argument applies if the CAS fails // or if we observe _owner change from one non-null value to // another non-null value. In such cases we might abort // the spin without prejudice or apply a "penalty" to the // spin count-down variable "ctr", reducing it by 100, say. Thread * ox = (Thread *) _owner ; if (ox == NULL) { ox = (Thread *) Atomic::cmpxchg_ptr (Self, &_owner, NULL) ; if (ox == NULL) { // The CAS succeeded -- this thread acquired ownership // Take care of some bookkeeping to exit spin state. if (sss && _succ == Self) { _succ = NULL ; } if (MaxSpin > 0) Adjust (&_Spinner, -1) ; // Increase _SpinDuration : // The spin was successful (profitable) so we tend toward // longer spin attempts in the future. // CONSIDER: factor "ctr" into the _SpinDuration adjustment. // If we acquired the lock early in the spin cycle it // makes sense to increase _SpinDuration proportionally. // Note that we don't clamp SpinDuration precisely at SpinLimit. int x = _SpinDuration ; if (x < Knob_SpinLimit) { if (x < Knob_Poverty) x = Knob_Poverty ; _SpinDuration = x + Knob_Bonus ; } return 1 ; } // The CAS failed ... we can take any of the following actions: // * penalize: ctr -= Knob_CASPenalty // * exit spin with prejudice -- goto Abort; // * exit spin without prejudice. // * Since CAS is high-latency, retry again immediately. prv = ox ; TEVENT (Spin: cas failed) ; if (caspty == -2) break ; if (caspty == -1) goto Abort ; ctr -= caspty ; continue ; } // Did lock ownership change hands ? if (ox != prv && prv != NULL ) { TEVENT (spin: Owner changed) if (oxpty == -2) break ; if (oxpty == -1) goto Abort ; ctr -= oxpty ; } prv = ox ; // Abort the spin if the owner is not executing. // The owner must be executing in order to drop the lock. // Spinning while the owner is OFFPROC is idiocy. // Consider: ctr -= RunnablePenalty ; if (Knob_OState && NotRunnable (Self, ox)) { TEVENT (Spin abort - notrunnable); goto Abort ; } if (sss && _succ == NULL ) _succ = Self ; } // Spin failed with prejudice -- reduce _SpinDuration. // TODO: Use an AIMD-like policy to adjust _SpinDuration. // AIMD is globally stable. TEVENT (Spin failure) ; { int x = _SpinDuration ; if (x > 0) { // Consider an AIMD scheme like: x -= (x >> 3) + 100 // This is globally sample and tends to damp the response. x -= Knob_Penalty ; if (x < 0) x = 0 ; _SpinDuration = x ; } } Abort: if (MaxSpin >= 0) Adjust (&_Spinner, -1) ; if (sss && _succ == Self) { _succ = NULL ; // Invariant: after setting succ=null a contending thread // must recheck-retry _owner before parking. This usually happens // in the normal usage of TrySpin(), but it's safest // to make TrySpin() as foolproof as possible. OrderAccess::fence() ; if (TryLock(Self) > 0) return 1 ; } return 0 ; } #define TrySpin TrySpin_VaryDuration static void DeferredInitialize () { if (InitDone > 0) return ; if (Atomic::cmpxchg (-1, &InitDone, 0) != 0) { while (InitDone != 1) ; return ; } // One-shot global initialization ... // The initialization is idempotent, so we don't need locks. // In the future consider doing this via os::init_2(). // SyncKnobs consist of <Key>=<Value> pairs in the style // of environment variables. Start by converting ':' to NUL. if (SyncKnobs == NULL) SyncKnobs = "" ; size_t sz = strlen (SyncKnobs) ; char * knobs = (char *) malloc (sz + 2) ; if (knobs == NULL) { vm_exit_out_of_memory (sz + 2, "Parse SyncKnobs") ; guarantee (0, "invariant") ; } strcpy (knobs, SyncKnobs) ; knobs[sz+1] = 0 ; for (char * p = knobs ; *p ; p++) { if (*p == ':') *p = 0 ; } #define SETKNOB(x) { Knob_##x = kvGetInt (knobs, #x, Knob_##x); } SETKNOB(ReportSettings) ; SETKNOB(Verbose) ; SETKNOB(FixedSpin) ; SETKNOB(SpinLimit) ; SETKNOB(SpinBase) ; SETKNOB(SpinBackOff); SETKNOB(CASPenalty) ; SETKNOB(OXPenalty) ; SETKNOB(LogSpins) ; SETKNOB(SpinSetSucc) ; SETKNOB(SuccEnabled) ; SETKNOB(SuccRestrict) ; SETKNOB(Penalty) ; SETKNOB(Bonus) ; SETKNOB(BonusB) ; SETKNOB(Poverty) ; SETKNOB(SpinAfterFutile) ; SETKNOB(UsePause) ; SETKNOB(SpinEarly) ; SETKNOB(OState) ; SETKNOB(MaxSpinners) ; SETKNOB(PreSpin) ; SETKNOB(ExitPolicy) ; SETKNOB(QMode); SETKNOB(ResetEvent) ; SETKNOB(MoveNotifyee) ; SETKNOB(FastHSSEC) ; #undef SETKNOB if (os::is_MP()) { BackOffMask = (1 << Knob_SpinBackOff) - 1 ; if (Knob_ReportSettings) ::printf ("BackOffMask=%X\n", BackOffMask) ; // CONSIDER: BackOffMask = ROUNDUP_NEXT_POWER2 (ncpus-1) } else { Knob_SpinLimit = 0 ; Knob_SpinBase = 0 ; Knob_PreSpin = 0 ; Knob_FixedSpin = -1 ; } if (Knob_LogSpins == 0) { ObjectSynchronizer::_sync_FailedSpins = NULL ; } free (knobs) ; OrderAccess::fence() ; InitDone = 1 ; } // Theory of operations -- Monitors lists, thread residency, etc: // // * A thread acquires ownership of a monitor by successfully // CAS()ing the _owner field from null to non-null. // // * Invariant: A thread appears on at most one monitor list -- // cxq, EntryList or WaitSet -- at any one time. // // * Contending threads "push" themselves onto the cxq with CAS // and then spin/park. // // * After a contending thread eventually acquires the lock it must // dequeue itself from either the EntryList or the cxq. // // * The exiting thread identifies and unparks an "heir presumptive" // tentative successor thread on the EntryList. Critically, the // exiting thread doesn't unlink the successor thread from the EntryList. // After having been unparked, the wakee will recontend for ownership of // the monitor. The successor (wakee) will either acquire the lock or // re-park itself. // // 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. // If the EntryList is empty but the cxq is populated the exiting // thread will drain the cxq into the EntryList. It does so by // by detaching the cxq (installing null with CAS) and folding // the threads from the cxq into the EntryList. The EntryList is // doubly linked, while the cxq is singly linked because of the // CAS-based "push" used to enqueue recently arrived threads (RATs). // // * Concurrency invariants: // // -- only the monitor owner may access or mutate the EntryList. // The mutex property of the monitor itself protects the EntryList // from concurrent interference. // -- Only the monitor owner may detach the cxq. // // * The monitor entry list operations avoid locks, but strictly speaking // they're not lock-free. Enter is lock-free, exit is not. // See http://j2se.east/~dice/PERSIST/040825-LockFreeQueues.html // // * 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. // // * 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 improve the odds of a constant-time // dequeue operation after acquisition (in the ::enter() epilog) and // to reduce heat on the list ends. (c.f. Michael Scott's "2Q" algorithm). // A key desideratum is to minimize queue & monitor metadata manipulation // that occurs while holding the monitor lock -- that is, we want to // minimize monitor lock holds times. Note that even a small amount of // fixed spinning will greatly reduce the # of enqueue-dequeue operations // on EntryList|cxq. That is, spinning relieves contention on the "inner" // locks and monitor metadata. // // Cxq points to the the set of Recently Arrived Threads attempting entry. // Because we push threads onto _cxq with CAS, the RATs must take the form of // a singly-linked LIFO. We drain _cxq into EntryList at unlock-time when // the unlocking thread notices that EntryList is null but _cxq is != null. // // 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. Critically, we want insert and delete operations // to operate in constant-time. 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. // Queue discipline is enforced at ::exit() 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. // // * 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. // // * 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 exit() operations will // unpark the notifyee. Unparking a notifee in notify() is inefficient - // it's likely the notifyee would simply impale itself on the lock held // by the notifier. // // * An interesting alternative is to encode cxq as (List,LockByte) where // the LockByte is 0 iff the monitor is owned. _owner is simply an auxiliary // variable, like _recursions, in the scheme. The threads or Events that form // the list would have to be aligned in 256-byte addresses. A thread would // try to acquire the lock or enqueue itself with CAS, but exiting threads // could use a 1-0 protocol and simply STB to set the LockByte to 0. // Note that is is *not* word-tearing, but it does presume that full-word // CAS operations are coherent with intermix with STB operations. That's true // on most common processors. // // * See also http://blogs.sun.com/dave void ATTR ObjectMonitor::EnterI (TRAPS) { Thread * Self = THREAD ; assert (Self->is_Java_thread(), "invariant") ; assert (((JavaThread *) Self)->thread_state() == _thread_blocked , "invariant") ; // Try the lock - TATAS if (TryLock (Self) > 0) { assert (_succ != Self , "invariant") ; assert (_owner == Self , "invariant") ; assert (_Responsible != Self , "invariant") ; return ; } DeferredInitialize () ; // We try one round of spinning *before* enqueueing Self. // // If the _owner is ready but OFFPROC we could use a YieldTo() // operation to donate the remainder of this thread's quantum // to the owner. This has subtle but beneficial affinity // effects. if (TrySpin (Self) > 0) { assert (_owner == Self , "invariant") ; assert (_succ != Self , "invariant") ; assert (_Responsible != Self , "invariant") ; return ; } // The Spin failed -- Enqueue and park the thread ... assert (_succ != Self , "invariant") ; assert (_owner != Self , "invariant") ; assert (_Responsible != Self , "invariant") ; // Enqueue "Self" on ObjectMonitor's _cxq. // // Node acts as a proxy for Self. // As an aside, if were to ever rewrite the synchronization code mostly // in Java, WaitNodes, ObjectMonitors, and Events would become 1st-class // Java objects. This would avoid awkward lifecycle and liveness issues, // as well as eliminate a subset of ABA issues. // TODO: eliminate ObjectWaiter and enqueue either Threads or Events. // ObjectWaiter node(Self) ; Self->_ParkEvent->reset() ; node._prev = (ObjectWaiter *) 0xBAD ; node.TState = ObjectWaiter::TS_CXQ ; // Push "Self" onto the front of the _cxq. // Once on cxq/EntryList, Self stays on-queue until it acquires the lock. // Note that spinning tends to reduce the rate at which threads // enqueue and dequeue on EntryList|cxq. ObjectWaiter * nxt ; for (;;) { node._next = nxt = _cxq ; if (Atomic::cmpxchg_ptr (&node, &_cxq, nxt) == nxt) break ; // Interference - the CAS failed because _cxq changed. Just retry. // As an optional optimization we retry the lock. if (TryLock (Self) > 0) { assert (_succ != Self , "invariant") ; assert (_owner == Self , "invariant") ; assert (_Responsible != Self , "invariant") ; return ; } } // Check for cxq|EntryList edge transition to non-null. This indicates // the onset of contention. While contention persists exiting threads // will use a ST:MEMBAR:LD 1-1 exit protocol. When contention abates exit // operations revert to the faster 1-0 mode. This enter operation may interleave // (race) a concurrent 1-0 exit operation, resulting in stranding, so we // arrange for one of the contending thread to use a timed park() operations // to detect and recover from the race. (Stranding is form of progress failure // where the monitor is unlocked but all the contending threads remain parked). // That is, at least one of the contended threads will periodically poll _owner. // One of the contending threads will become the designated "Responsible" thread. // The Responsible thread uses a timed park instead of a normal indefinite park // operation -- it periodically wakes and checks for and recovers from potential // strandings admitted by 1-0 exit operations. We need at most one Responsible // thread per-monitor at any given moment. Only threads on cxq|EntryList may // be responsible for a monitor. // // Currently, one of the contended threads takes on the added role of "Responsible". // A viable alternative would be to use a dedicated "stranding checker" thread // that periodically iterated over all the threads (or active monitors) and unparked // successors where there was risk of stranding. This would help eliminate the // timer scalability issues we see on some platforms as we'd only have one thread // -- the checker -- parked on a timer. if ((SyncFlags & 16) == 0 && nxt == NULL && _EntryList == NULL) { // Try to assume the role of responsible thread for the monitor. // CONSIDER: ST vs CAS vs { if (Responsible==null) Responsible=Self } Atomic::cmpxchg_ptr (Self, &_Responsible, NULL) ; } // The lock have been released while this thread was occupied queueing // itself onto _cxq. To close the race and avoid "stranding" and // progress-liveness failure we must resample-retry _owner before parking. // Note the Dekker/Lamport duality: ST cxq; MEMBAR; LD Owner. // In this case the ST-MEMBAR is accomplished with CAS(). // // TODO: Defer all thread state transitions until park-time. // Since state transitions are heavy and inefficient we'd like // to defer the state transitions until absolutely necessary, // and in doing so avoid some transitions ... TEVENT (Inflated enter - Contention) ; int nWakeups = 0 ; int RecheckInterval = 1 ; for (;;) { if (TryLock (Self) > 0) break ; assert (_owner != Self, "invariant") ; if ((SyncFlags & 2) && _Responsible == NULL) { Atomic::cmpxchg_ptr (Self, &_Responsible, NULL) ; } // park self if (_Responsible == Self || (SyncFlags & 1)) { TEVENT (Inflated enter - park TIMED) ; Self->_ParkEvent->park ((jlong) RecheckInterval) ; // Increase the RecheckInterval, but clamp the value. RecheckInterval *= 8 ; if (RecheckInterval > 1000) RecheckInterval = 1000 ; } else { TEVENT (Inflated enter - park UNTIMED) ; Self->_ParkEvent->park() ; } if (TryLock(Self) > 0) break ; // The lock is still contested. // Keep a tally of the # of futile wakeups. // Note that the counter is not protected by a lock or updated by atomics. // That is by design - we trade "lossy" counters which are exposed to // races during updates for a lower probe effect. TEVENT (Inflated enter - Futile wakeup) ; if (ObjectSynchronizer::_sync_FutileWakeups != NULL) { ObjectSynchronizer::_sync_FutileWakeups->inc() ; } ++ nWakeups ; // Assuming this is not a spurious wakeup we'll normally find _succ == Self. // We can defer clearing _succ until after the spin completes // TrySpin() must tolerate being called with _succ == Self. // Try yet another round of adaptive spinning. if ((Knob_SpinAfterFutile & 1) && TrySpin (Self) > 0) break ; // We can find that we were unpark()ed and redesignated _succ while // we were spinning. That's harmless. If we iterate and call park(), // park() will consume the event and return immediately and we'll // just spin again. This pattern can repeat, leaving _succ to simply // spin on a CPU. Enable Knob_ResetEvent to clear pending unparks(). // Alternately, we can sample fired() here, and if set, forgo spinning // in the next iteration. if ((Knob_ResetEvent & 1) && Self->_ParkEvent->fired()) { Self->_ParkEvent->reset() ; OrderAccess::fence() ; } if (_succ == Self) _succ = NULL ; // Invariant: after clearing _succ a thread *must* retry _owner before parking. OrderAccess::fence() ; } // Egress : // Self has acquired the lock -- Unlink Self from the cxq or EntryList. // Normally we'll find Self on the EntryList . // From the perspective of the lock owner (this thread), the // EntryList is stable and cxq is prepend-only. // The head of cxq is volatile but the interior is stable. // In addition, Self.TState is stable. assert (_owner == Self , "invariant") ; assert (object() != NULL , "invariant") ; // I'd like to write: // guarantee (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ; // but as we're at a safepoint that's not safe. UnlinkAfterAcquire (Self, &node) ; if (_succ == Self) _succ = NULL ; assert (_succ != Self, "invariant") ; if (_Responsible == Self) { _Responsible = NULL ; // Dekker pivot-point. // Consider OrderAccess::storeload() here // We may leave threads on cxq|EntryList without a designated // "Responsible" thread. This is benign. When this thread subsequently // exits the monitor it can "see" such preexisting "old" threads -- // threads that arrived on the cxq|EntryList before the fence, above -- // by LDing cxq|EntryList. Newly arrived threads -- that is, threads // that arrive on cxq after the ST:MEMBAR, above -- will set Responsible // non-null and elect a new "Responsible" timer thread. // // This thread executes: // ST Responsible=null; MEMBAR (in enter epilog - here) // LD cxq|EntryList (in subsequent exit) // // Entering threads in the slow/contended path execute: // ST cxq=nonnull; MEMBAR; LD Responsible (in enter prolog) // The (ST cxq; MEMBAR) is accomplished with CAS(). // // The MEMBAR, above, prevents the LD of cxq|EntryList in the subsequent // exit operation from floating above the ST Responsible=null. // // In *practice* however, EnterI() is always followed by some atomic // operation such as the decrement of _count in ::enter(). Those atomics // obviate the need for the explicit MEMBAR, above. } // We've acquired ownership with CAS(). // CAS is serializing -- it has MEMBAR/FENCE-equivalent semantics. // But since the CAS() this thread may have also stored into _succ, // EntryList, cxq or Responsible. These meta-data updates must be // visible __before this thread subsequently drops the lock. // Consider what could occur if we didn't enforce this constraint -- // STs to monitor meta-data and user-data could reorder with (become // visible after) the ST in exit that drops ownership of the lock. // Some other thread could then acquire the lock, but observe inconsistent // or old monitor meta-data and heap data. That violates the JMM. // To that end, the 1-0 exit() operation must have at least STST|LDST // "release" barrier semantics. Specifically, there must be at least a // STST|LDST barrier in exit() before the ST of null into _owner that drops // the lock. The barrier ensures that changes to monitor meta-data and data // protected by the lock will be visible before we release the lock, and // therefore before some other thread (CPU) has a chance to acquire the lock. // See also: http://gee.cs.oswego.edu/dl/jmm/cookbook.html. // // Critically, any prior STs to _succ or EntryList must be visible before // the ST of null into _owner in the *subsequent* (following) corresponding // monitorexit. Recall too, that in 1-0 mode monitorexit does not necessarily // execute a serializing instruction. if (SyncFlags & 8) { OrderAccess::fence() ; } return ; } // ExitSuspendEquivalent: // A faster alternate to handle_special_suspend_equivalent_condition() // // handle_special_suspend_equivalent_condition() unconditionally // acquires the SR_lock. On some platforms uncontended MutexLocker() // operations have high latency. Note that in ::enter() we call HSSEC // while holding the monitor, so we effectively lengthen the critical sections. // // There are a number of possible solutions: // // A. To ameliorate the problem we might also defer state transitions // to as late as possible -- just prior to parking. // Given that, we'd call HSSEC after having returned from park(), // but before attempting to acquire the monitor. This is only a // partial solution. It avoids calling HSSEC while holding the // monitor (good), but it still increases successor reacquisition latency -- // the interval between unparking a successor and the time the successor // resumes and retries the lock. See ReenterI(), which defers state transitions. // If we use this technique we can also avoid EnterI()-exit() loop // in ::enter() where we iteratively drop the lock and then attempt // to reacquire it after suspending. // // B. In the future we might fold all the suspend bits into a // composite per-thread suspend flag and then update it with CAS(). // Alternately, a Dekker-like mechanism with multiple variables // would suffice: // ST Self->_suspend_equivalent = false // MEMBAR // LD Self_>_suspend_flags // bool ObjectMonitor::ExitSuspendEquivalent (JavaThread * jSelf) { int Mode = Knob_FastHSSEC ; if (Mode && !jSelf->is_external_suspend()) { assert (jSelf->is_suspend_equivalent(), "invariant") ; jSelf->clear_suspend_equivalent() ; if (2 == Mode) OrderAccess::storeload() ; if (!jSelf->is_external_suspend()) return false ; // We raced a suspension -- fall thru into the slow path TEVENT (ExitSuspendEquivalent - raced) ; jSelf->set_suspend_equivalent() ; } return jSelf->handle_special_suspend_equivalent_condition() ; } // ReenterI() is a specialized inline form of the latter half of the // contended slow-path from EnterI(). We use ReenterI() only for // monitor reentry in wait(). // // In the future we should reconcile EnterI() and ReenterI(), adding // Knob_Reset and Knob_SpinAfterFutile support and restructuring the // loop accordingly. void ATTR ObjectMonitor::ReenterI (Thread * Self, ObjectWaiter * SelfNode) { assert (Self != NULL , "invariant") ; assert (SelfNode != NULL , "invariant") ; assert (SelfNode->_thread == Self , "invariant") ; assert (_waiters > 0 , "invariant") ; assert (((oop)(object()))->mark() == markOopDesc::encode(this) , "invariant") ; assert (((JavaThread *)Self)->thread_state() != _thread_blocked, "invariant") ; JavaThread * jt = (JavaThread *) Self ; int nWakeups = 0 ; for (;;) { ObjectWaiter::TStates v = SelfNode->TState ; guarantee (v == ObjectWaiter::TS_ENTER || v == ObjectWaiter::TS_CXQ, "invariant") ; assert (_owner != Self, "invariant") ; if (TryLock (Self) > 0) break ; if (TrySpin (Self) > 0) break ; TEVENT (Wait Reentry - parking) ; // State transition wrappers around park() ... // ReenterI() wisely defers state transitions until // it's clear we must park the thread. { OSThreadContendState osts(Self->osthread()); ThreadBlockInVM tbivm(jt); // cleared by handle_special_suspend_equivalent_condition() // or java_suspend_self() jt->set_suspend_equivalent(); if (SyncFlags & 1) { Self->_ParkEvent->park ((jlong)1000) ; } else { Self->_ParkEvent->park () ; } // were we externally suspended while we were waiting? for (;;) { if (!ExitSuspendEquivalent (jt)) break ; if (_succ == Self) { _succ = NULL; OrderAccess::fence(); } jt->java_suspend_self(); jt->set_suspend_equivalent(); } } // Try again, but just so we distinguish between futile wakeups and // successful wakeups. The following test isn't algorithmically // necessary, but it helps us maintain sensible statistics. if (TryLock(Self) > 0) break ; // The lock is still contested. // Keep a tally of the # of futile wakeups. // Note that the counter is not protected by a lock or updated by atomics. // That is by design - we trade "lossy" counters which are exposed to // races during updates for a lower probe effect. TEVENT (Wait Reentry - futile wakeup) ; ++ nWakeups ; // Assuming this is not a spurious wakeup we'll normally // find that _succ == Self. if (_succ == Self) _succ = NULL ; // Invariant: after clearing _succ a contending thread // *must* retry _owner before parking. OrderAccess::fence() ; if (ObjectSynchronizer::_sync_FutileWakeups != NULL) { ObjectSynchronizer::_sync_FutileWakeups->inc() ; } } // Self has acquired the lock -- Unlink Self from the cxq or EntryList . // Normally we'll find Self on the EntryList. // Unlinking from the EntryList is constant-time and atomic-free. // From the perspective of the lock owner (this thread), the // EntryList is stable and cxq is prepend-only. // The head of cxq is volatile but the interior is stable. // In addition, Self.TState is stable. assert (_owner == Self, "invariant") ; assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ; UnlinkAfterAcquire (Self, SelfNode) ; if (_succ == Self) _succ = NULL ; assert (_succ != Self, "invariant") ; SelfNode->TState = ObjectWaiter::TS_RUN ; OrderAccess::fence() ; // see comments at the end of EnterI() } bool ObjectMonitor::try_enter(Thread* THREAD) { if (THREAD != _owner) { if (THREAD->is_lock_owned ((address)_owner)) { assert(_recursions == 0, "internal state error"); _owner = THREAD ; _recursions = 1 ; OwnerIsThread = 1 ; return true; } if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) { return false; } return true; } else { _recursions++; return true; } } void ATTR ObjectMonitor::enter(TRAPS) { // The following code is ordered to check the most common cases first // and to reduce RTS->RTO cache line upgrades on SPARC and IA32 processors. Thread * const Self = THREAD ; void * cur ; cur = Atomic::cmpxchg_ptr (Self, &_owner, NULL) ; if (cur == NULL) { // Either ASSERT _recursions == 0 or explicitly set _recursions = 0. assert (_recursions == 0 , "invariant") ; assert (_owner == Self, "invariant") ; // CONSIDER: set or assert OwnerIsThread == 1 return ; } if (cur == Self) { // TODO-FIXME: check for integer overflow! BUGID 6557169. _recursions ++ ; return ; } if (Self->is_lock_owned ((address)cur)) { assert (_recursions == 0, "internal state error"); _recursions = 1 ; // Commute owner from a thread-specific on-stack BasicLockObject address to // a full-fledged "Thread *". _owner = Self ; OwnerIsThread = 1 ; return ; } // We've encountered genuine contention. assert (Self->_Stalled == 0, "invariant") ; Self->_Stalled = intptr_t(this) ; // Try one round of spinning *before* enqueueing Self // and before going through the awkward and expensive state // transitions. The following spin is strictly optional ... // Note that if we acquire the monitor from an initial spin // we forgo posting JVMTI events and firing DTRACE probes. if (Knob_SpinEarly && TrySpin (Self) > 0) { assert (_owner == Self , "invariant") ; assert (_recursions == 0 , "invariant") ; assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ; Self->_Stalled = 0 ; return ; } assert (_owner != Self , "invariant") ; assert (_succ != Self , "invariant") ; assert (Self->is_Java_thread() , "invariant") ; JavaThread * jt = (JavaThread *) Self ; assert (!SafepointSynchronize::is_at_safepoint(), "invariant") ; assert (jt->thread_state() != _thread_blocked , "invariant") ; assert (this->object() != NULL , "invariant") ; assert (_count >= 0, "invariant") ; // Prevent deflation at STW-time. See deflate_idle_monitors() and is_busy(). // Ensure the object-monitor relationship remains stable while there's contention. Atomic::inc_ptr(&_count); { // Change java thread status to indicate blocked on monitor enter. JavaThreadBlockedOnMonitorEnterState jtbmes(jt, this); DTRACE_MONITOR_PROBE(contended__enter, this, object(), jt); if (JvmtiExport::should_post_monitor_contended_enter()) { JvmtiExport::post_monitor_contended_enter(jt, this); } OSThreadContendState osts(Self->osthread()); ThreadBlockInVM tbivm(jt); Self->set_current_pending_monitor(this); // TODO-FIXME: change the following for(;;) loop to straight-line code. for (;;) { jt->set_suspend_equivalent(); // cleared by handle_special_suspend_equivalent_condition() // or java_suspend_self() EnterI (THREAD) ; if (!ExitSuspendEquivalent(jt)) break ; // // We have acquired the contended monitor, but while we were // waiting another thread suspended us. We don't want to enter // the monitor while suspended because that would surprise the // thread that suspended us. // _recursions = 0 ; _succ = NULL ; exit (Self) ; jt->java_suspend_self(); } Self->set_current_pending_monitor(NULL); } Atomic::dec_ptr(&_count); assert (_count >= 0, "invariant") ; Self->_Stalled = 0 ; // Must either set _recursions = 0 or ASSERT _recursions == 0. assert (_recursions == 0 , "invariant") ; assert (_owner == Self , "invariant") ; assert (_succ != Self , "invariant") ; assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ; // The thread -- now the owner -- is back in vm mode. // Report the glorious news via TI,DTrace and jvmstat. // The probe effect is non-trivial. All the reportage occurs // while we hold the monitor, increasing the length of the critical // section. Amdahl's parallel speedup law comes vividly into play. // // Another option might be to aggregate the events (thread local or // per-monitor aggregation) and defer reporting until a more opportune // time -- such as next time some thread encounters contention but has // yet to acquire the lock. While spinning that thread could // spinning we could increment JVMStat counters, etc. DTRACE_MONITOR_PROBE(contended__entered, this, object(), jt); if (JvmtiExport::should_post_monitor_contended_entered()) { JvmtiExport::post_monitor_contended_entered(jt, this); } if (ObjectSynchronizer::_sync_ContendedLockAttempts != NULL) { ObjectSynchronizer::_sync_ContendedLockAttempts->inc() ; } } void ObjectMonitor::ExitEpilog (Thread * Self, ObjectWaiter * Wakee) { assert (_owner == Self, "invariant") ; // Exit protocol: // 1. ST _succ = wakee // 2. membar #loadstore|#storestore; // 2. ST _owner = NULL // 3. unpark(wakee) _succ = Knob_SuccEnabled ? Wakee->_thread : NULL ; ParkEvent * Trigger = Wakee->_event ; // Hygiene -- once we've set _owner = NULL we can't safely dereference Wakee again. // The thread associated with Wakee may have grabbed the lock and "Wakee" may be // out-of-scope (non-extant). Wakee = NULL ; // Drop the lock OrderAccess::release_store_ptr (&_owner, NULL) ; OrderAccess::fence() ; // ST _owner vs LD in unpark() // TODO-FIXME: // If there's a safepoint pending the best policy would be to // get _this thread to a safepoint and only wake the successor // after the safepoint completed. monitorexit uses a "leaf" // state transition, however, so this thread can't become // safe at this point in time. (Its stack isn't walkable). // The next best thing is to defer waking the successor by // adding to a list of thread to be unparked after at the // end of the forthcoming STW). if (SafepointSynchronize::do_call_back()) { TEVENT (unpark before SAFEPOINT) ; } // Possible optimizations ... // // * Consider: set Wakee->UnparkTime = timeNow() // When the thread wakes up it'll compute (timeNow() - Self->UnparkTime()). // By measuring recent ONPROC latency we can approximate the // system load. In turn, we can feed that information back // into the spinning & succession policies. // (ONPROC latency correlates strongly with load). // // * Pull affinity: // If the wakee is cold then transiently setting it's affinity // to the current CPU is a good idea. // See http://j2se.east/~dice/PERSIST/050624-PullAffinity.txt DTRACE_MONITOR_PROBE(contended__exit, this, object(), Self); Trigger->unpark() ; // Maintain stats and report events to JVMTI if (ObjectSynchronizer::_sync_Parks != NULL) { ObjectSynchronizer::_sync_Parks->inc() ; } } // exit() // ~~~~~~ // Note that the collector can't reclaim the objectMonitor or deflate // the object out from underneath the thread calling ::exit() as the // thread calling ::exit() never transitions to a stable state. // This inhibits GC, which in turn inhibits asynchronous (and // inopportune) reclamation of "this". // // We'd like to assert that: (THREAD->thread_state() != _thread_blocked) ; // There's one exception to the claim above, however. EnterI() can call // exit() to drop a lock if the acquirer has been externally suspended. // In that case exit() is called with _thread_state as _thread_blocked, // but the monitor's _count field is > 0, which inhibits reclamation. // // 1-0 exit // ~~~~~~~~ // ::exit() uses a canonical 1-1 idiom with a MEMBAR although some of // the fast-path operators have been optimized so the common ::exit() // operation is 1-0. See i486.ad fast_unlock(), for instance. // The code emitted by fast_unlock() elides the usual MEMBAR. This // greatly improves latency -- MEMBAR and CAS having considerable local // latency on modern processors -- but at the cost of "stranding". Absent the // MEMBAR, a thread in fast_unlock() can race a thread in the slow // ::enter() path, resulting in the entering thread being stranding // and a progress-liveness failure. Stranding is extremely rare. // We use timers (timed park operations) & periodic polling to detect // and recover from stranding. Potentially stranded threads periodically // wake up and poll the lock. See the usage of the _Responsible variable. // // The CAS() in enter provides for safety and exclusion, while the CAS or // MEMBAR in exit provides for progress and avoids stranding. 1-0 locking // eliminates the CAS/MEMBAR from the exist path, but it admits stranding. // We detect and recover from stranding with timers. // // If a thread transiently strands it'll park until (a) another // thread acquires the lock and then drops the lock, at which time the // exiting thread will notice and unpark the stranded thread, or, (b) // the timer expires. If the lock is high traffic then the stranding latency // will be low due to (a). If the lock is low traffic then the odds of // stranding are lower, although the worst-case stranding latency // is longer. Critically, we don't want to put excessive load in the // platform's timer subsystem. We want to minimize both the timer injection // rate (timers created/sec) as well as the number of timers active at // any one time. (more precisely, we want to minimize timer-seconds, which is // the integral of the # of active timers at any instant over time). // Both impinge on OS scalability. Given that, at most one thread parked on // a monitor will use a timer. void ATTR ObjectMonitor::exit(TRAPS) { Thread * Self = THREAD ; if (THREAD != _owner) { if (THREAD->is_lock_owned((address) _owner)) { // Transmute _owner from a BasicLock pointer to a Thread address. // We don't need to hold _mutex for this transition. // Non-null to Non-null is safe as long as all readers can // tolerate either flavor. assert (_recursions == 0, "invariant") ; _owner = THREAD ; _recursions = 0 ; OwnerIsThread = 1 ; } else { // NOTE: we need to handle unbalanced monitor enter/exit // in native code by throwing an exception. // TODO: Throw an IllegalMonitorStateException ? TEVENT (Exit - Throw IMSX) ; assert(false, "Non-balanced monitor enter/exit!"); if (false) { THROW(vmSymbols::java_lang_IllegalMonitorStateException()); } return; } } if (_recursions != 0) { _recursions--; // this is simple recursive enter TEVENT (Inflated exit - recursive) ; return ; } // Invariant: after setting Responsible=null an thread must execute // a MEMBAR or other serializing instruction before fetching EntryList|cxq. if ((SyncFlags & 4) == 0) { _Responsible = NULL ; } for (;;) { assert (THREAD == _owner, "invariant") ; // Fast-path monitor exit: // // Observe the Dekker/Lamport duality: // A thread in ::exit() executes: // ST Owner=null; MEMBAR; LD EntryList|cxq. // A thread in the contended ::enter() path executes the complementary: // ST EntryList|cxq = nonnull; MEMBAR; LD Owner. // // Note that there's a benign race in the exit path. We can drop the // lock, another thread can reacquire the lock immediately, and we can // then wake a thread unnecessarily (yet another flavor of futile wakeup). // This is benign, and we've structured the code so the windows are short // and the frequency of such futile wakeups is low. // // We could eliminate the race by encoding both the "LOCKED" state and // the queue head in a single word. Exit would then use either CAS to // clear the LOCKED bit/byte. This precludes the desirable 1-0 optimization, // however. // // Possible fast-path ::exit() optimization: // The current fast-path exit implementation fetches both cxq and EntryList. // See also i486.ad fast_unlock(). Testing has shown that two LDs // isn't measurably slower than a single LD on any platforms. // Still, we could reduce the 2 LDs to one or zero by one of the following: // // - Use _count instead of cxq|EntryList // We intend to eliminate _count, however, when we switch // to on-the-fly deflation in ::exit() as is used in // Metalocks and RelaxedLocks. // // - Establish the invariant that cxq == null implies EntryList == null. // set cxq == EMPTY (1) to encode the state where cxq is empty // by EntryList != null. EMPTY is a distinguished value. // The fast-path exit() would fetch cxq but not EntryList. // // - Encode succ as follows: // succ = t : Thread t is the successor -- t is ready or is spinning. // Exiting thread does not need to wake a successor. // succ = 0 : No successor required -> (EntryList|cxq) == null // Exiting thread does not need to wake a successor // succ = 1 : Successor required -> (EntryList|cxq) != null and // logically succ == null. // Exiting thread must wake a successor. // // The 1-1 fast-exit path would appear as : // _owner = null ; membar ; // if (_succ == 1 && CAS (&_owner, null, Self) == null) goto SlowPath // goto FastPathDone ; // // and the 1-0 fast-exit path would appear as: // if (_succ == 1) goto SlowPath // Owner = null ; // goto FastPathDone // // - Encode the LSB of _owner as 1 to indicate that exit() // must use the slow-path and make a successor ready. // (_owner & 1) == 0 IFF succ != null || (EntryList|cxq) == null // (_owner & 1) == 0 IFF succ == null && (EntryList|cxq) != null (obviously) // The 1-0 fast exit path would read: // if (_owner != Self) goto SlowPath // _owner = null // goto FastPathDone if (Knob_ExitPolicy == 0) { // release semantics: prior loads and stores from within the critical section // must not float (reorder) past the following store that drops the lock. // On SPARC that requires MEMBAR #loadstore|#storestore. // But of course in TSO #loadstore|#storestore is not required. // I'd like to write one of the following: // A. OrderAccess::release() ; _owner = NULL // B. OrderAccess::loadstore(); OrderAccess::storestore(); _owner = NULL; // Unfortunately OrderAccess::release() and OrderAccess::loadstore() both // store into a _dummy variable. That store is not needed, but can result // in massive wasteful coherency traffic on classic SMP systems. // Instead, I use release_store(), which is implemented as just a simple // ST on x64, x86 and SPARC. OrderAccess::release_store_ptr (&_owner, NULL) ; // drop the lock OrderAccess::storeload() ; // See if we need to wake a successor if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) { TEVENT (Inflated exit - simple egress) ; return ; } TEVENT (Inflated exit - complex egress) ; // Normally the exiting thread is responsible for ensuring succession, // but if other successors are ready or other entering threads are spinning // then this thread can simply store NULL into _owner and exit without // waking a successor. The existence of spinners or ready successors // guarantees proper succession (liveness). Responsibility passes to the // ready or running successors. The exiting thread delegates the duty. // More precisely, if a successor already exists this thread is absolved // of the responsibility of waking (unparking) one. // // The _succ variable is critical to reducing futile wakeup frequency. // _succ identifies the "heir presumptive" thread that has been made // ready (unparked) but that has not yet run. We need only one such // successor thread to guarantee progress. // See http://www.usenix.org/events/jvm01/full_papers/dice/dice.pdf // section 3.3 "Futile Wakeup Throttling" for details. // // Note that spinners in Enter() also set _succ non-null. // In the current implementation spinners opportunistically set // _succ so that exiting threads might avoid waking a successor. // Another less appealing alternative would be for the exiting thread // to drop the lock and then spin briefly to see if a spinner managed // to acquire the lock. If so, the exiting thread could exit // immediately without waking a successor, otherwise the exiting // thread would need to dequeue and wake a successor. // (Note that we'd need to make the post-drop spin short, but no // shorter than the worst-case round-trip cache-line migration time. // The dropped lock needs to become visible to the spinner, and then // the acquisition of the lock by the spinner must become visible to // the exiting thread). // // It appears that an heir-presumptive (successor) must be made ready. // Only the current lock owner can manipulate the EntryList or // drain _cxq, so we need to reacquire the lock. If we fail // to reacquire the lock the responsibility for ensuring succession // falls to the new owner. // if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) { return ; } TEVENT (Exit - Reacquired) ; } else { if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) { OrderAccess::release_store_ptr (&_owner, NULL) ; // drop the lock OrderAccess::storeload() ; // Ratify the previously observed values. if (_cxq == NULL || _succ != NULL) { TEVENT (Inflated exit - simple egress) ; return ; } // inopportune interleaving -- the exiting thread (this thread) // in the fast-exit path raced an entering thread in the slow-enter // path. // We have two choices: // A. Try to reacquire the lock. // If the CAS() fails return immediately, otherwise // we either restart/rerun the exit operation, or simply // fall-through into the code below which wakes a successor. // B. If the elements forming the EntryList|cxq are TSM // we could simply unpark() the lead thread and return // without having set _succ. if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) { TEVENT (Inflated exit - reacquired succeeded) ; return ; } TEVENT (Inflated exit - reacquired failed) ; } else { TEVENT (Inflated exit - complex egress) ; } } guarantee (_owner == THREAD, "invariant") ; // Select an appropriate successor ("heir presumptive") from the EntryList // and make it ready. Generally we just wake the head of EntryList . // There's no algorithmic constraint that we use the head - it's just // a policy decision. Note that the thread at head of the EntryList // remains at the head until it acquires the lock. This means we'll // repeatedly wake the same thread until it manages to grab the lock. // This is generally a good policy - if we're seeing lots of futile wakeups // at least we're waking/rewaking a thread that's like to be hot or warm // (have residual D$ and TLB affinity). // // "Wakeup locality" optimization: // http://j2se.east/~dice/PERSIST/040825-WakeLocality.txt // In the future we'll try to bias the selection mechanism // to preferentially pick a thread that recently ran on // a processor element that shares cache with the CPU on which // the exiting thread is running. We need access to Solaris' // schedctl.sc_cpu to make that work. // ObjectWaiter * w = NULL ; int QMode = Knob_QMode ; if (QMode == 2 && _cxq != NULL) { // QMode == 2 : cxq has precedence over EntryList. // Try to directly wake a successor from the cxq. // If successful, the successor will need to unlink itself from cxq. w = _cxq ; assert (w != NULL, "invariant") ; assert (w->TState == ObjectWaiter::TS_CXQ, "Invariant") ; ExitEpilog (Self, w) ; return ; } if (QMode == 3 && _cxq != NULL) { // Aggressively drain cxq into EntryList at the first opportunity. // This policy ensure that recently-run threads live at the head of EntryList. // Drain _cxq into EntryList - bulk transfer. // First, detach _cxq. // The following loop is tantamount to: w = swap (&cxq, NULL) w = _cxq ; for (;;) { assert (w != NULL, "Invariant") ; ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ; if (u == w) break ; w = u ; } assert (w != NULL , "invariant") ; ObjectWaiter * q = NULL ; ObjectWaiter * p ; for (p = w ; p != NULL ; p = p->_next) { guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ; p->TState = ObjectWaiter::TS_ENTER ; p->_prev = q ; q = p ; } // Append the RATs to the EntryList // TODO: organize EntryList as a CDLL so we can locate the tail in constant-time. ObjectWaiter * Tail ; for (Tail = _EntryList ; Tail != NULL && Tail->_next != NULL ; Tail = Tail->_next) ; if (Tail == NULL) { _EntryList = w ; } else { Tail->_next = w ; w->_prev = Tail ; } // Fall thru into code that tries to wake a successor from EntryList } if (QMode == 4 && _cxq != NULL) { // Aggressively drain cxq into EntryList at the first opportunity. // This policy ensure that recently-run threads live at the head of EntryList. // Drain _cxq into EntryList - bulk transfer. // First, detach _cxq. // The following loop is tantamount to: w = swap (&cxq, NULL) w = _cxq ; for (;;) { assert (w != NULL, "Invariant") ; ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ; if (u == w) break ; w = u ; } assert (w != NULL , "invariant") ; ObjectWaiter * q = NULL ; ObjectWaiter * p ; for (p = w ; p != NULL ; p = p->_next) { guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ; p->TState = ObjectWaiter::TS_ENTER ; p->_prev = q ; q = p ; } // Prepend the RATs to the EntryList if (_EntryList != NULL) { q->_next = _EntryList ; _EntryList->_prev = q ; } _EntryList = w ; // Fall thru into code that tries to wake a successor from EntryList } w = _EntryList ; if (w != NULL) { // I'd like to write: guarantee (w->_thread != Self). // But in practice an exiting thread may find itself on the EntryList. // Lets say thread T1 calls O.wait(). Wait() enqueues T1 on O's waitset and // then calls exit(). Exit release the lock by setting O._owner to NULL. // Lets say T1 then stalls. T2 acquires O and calls O.notify(). The // notify() operation moves T1 from O's waitset to O's EntryList. T2 then // release the lock "O". T2 resumes immediately after the ST of null into // _owner, above. T2 notices that the EntryList is populated, so it // reacquires the lock and then finds itself on the EntryList. // Given all that, we have to tolerate the circumstance where "w" is // associated with Self. assert (w->TState == ObjectWaiter::TS_ENTER, "invariant") ; ExitEpilog (Self, w) ; return ; } // If we find that both _cxq and EntryList are null then just // re-run the exit protocol from the top. w = _cxq ; if (w == NULL) continue ; // Drain _cxq into EntryList - bulk transfer. // First, detach _cxq. // The following loop is tantamount to: w = swap (&cxq, NULL) for (;;) { assert (w != NULL, "Invariant") ; ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ; if (u == w) break ; w = u ; } TEVENT (Inflated exit - drain cxq into EntryList) ; assert (w != NULL , "invariant") ; assert (_EntryList == NULL , "invariant") ; // Convert the LIFO SLL anchored by _cxq into a DLL. // The list reorganization step operates in O(LENGTH(w)) time. // It's critical that this step operate quickly as // "Self" still holds the outer-lock, restricting parallelism // and effectively lengthening the critical section. // Invariant: s chases t chases u. // TODO-FIXME: consider changing EntryList from a DLL to a CDLL so // we have faster access to the tail. if (QMode == 1) { // QMode == 1 : drain cxq to EntryList, reversing order // We also reverse the order of the list. ObjectWaiter * s = NULL ; ObjectWaiter * t = w ; ObjectWaiter * u = NULL ; while (t != NULL) { guarantee (t->TState == ObjectWaiter::TS_CXQ, "invariant") ; t->TState = ObjectWaiter::TS_ENTER ; u = t->_next ; t->_prev = u ; t->_next = s ; s = t; t = u ; } _EntryList = s ; assert (s != NULL, "invariant") ; } else { // QMode == 0 or QMode == 2 _EntryList = w ; ObjectWaiter * q = NULL ; ObjectWaiter * p ; for (p = w ; p != NULL ; p = p->_next) { guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ; p->TState = ObjectWaiter::TS_ENTER ; p->_prev = q ; q = p ; } } // In 1-0 mode we need: ST EntryList; MEMBAR #storestore; ST _owner = NULL // The MEMBAR is satisfied by the release_store() operation in ExitEpilog(). // See if we can abdicate to a spinner instead of waking a thread. // A primary goal of the implementation is to reduce the // context-switch rate. if (_succ != NULL) continue; w = _EntryList ; if (w != NULL) { guarantee (w->TState == ObjectWaiter::TS_ENTER, "invariant") ; ExitEpilog (Self, w) ; return ; } } } // complete_exit exits a lock returning recursion count // complete_exit/reenter operate as a wait without waiting // complete_exit requires an inflated monitor // The _owner field is not always the Thread addr even with an // inflated monitor, e.g. the monitor can be inflated by a non-owning // thread due to contention. intptr_t ObjectMonitor::complete_exit(TRAPS) { Thread * const Self = THREAD; assert(Self->is_Java_thread(), "Must be Java thread!"); JavaThread *jt = (JavaThread *)THREAD; DeferredInitialize(); if (THREAD != _owner) { if (THREAD->is_lock_owned ((address)_owner)) { assert(_recursions == 0, "internal state error"); _owner = THREAD ; /* Convert from basiclock addr to Thread addr */ _recursions = 0 ; OwnerIsThread = 1 ; } } guarantee(Self == _owner, "complete_exit not owner"); intptr_t save = _recursions; // record the old recursion count _recursions = 0; // set the recursion level to be 0 exit (Self) ; // exit the monitor guarantee (_owner != Self, "invariant"); return save; } // reenter() enters a lock and sets recursion count // complete_exit/reenter operate as a wait without waiting void ObjectMonitor::reenter(intptr_t recursions, TRAPS) { Thread * const Self = THREAD; assert(Self->is_Java_thread(), "Must be Java thread!"); JavaThread *jt = (JavaThread *)THREAD; guarantee(_owner != Self, "reenter already owner"); enter (THREAD); // enter the monitor guarantee (_recursions == 0, "reenter recursion"); _recursions = recursions; return; } // Note: a subset of changes to ObjectMonitor::wait() // will need to be replicated in complete_exit above void ObjectMonitor::wait(jlong millis, bool interruptible, TRAPS) { Thread * const Self = THREAD ; assert(Self->is_Java_thread(), "Must be Java thread!"); JavaThread *jt = (JavaThread *)THREAD; DeferredInitialize () ; // Throw IMSX or IEX. CHECK_OWNER(); // check for a pending interrupt if (interruptible && Thread::is_interrupted(Self, true) && !HAS_PENDING_EXCEPTION) { // post monitor waited event. Note that this is past-tense, we are done waiting. if (JvmtiExport::should_post_monitor_waited()) { // Note: 'false' parameter is passed here because the // wait was not timed out due to thread interrupt. JvmtiExport::post_monitor_waited(jt, this, false); } TEVENT (Wait - Throw IEX) ; THROW(vmSymbols::java_lang_InterruptedException()); return ; } TEVENT (Wait) ; assert (Self->_Stalled == 0, "invariant") ; Self->_Stalled = intptr_t(this) ; jt->set_current_waiting_monitor(this); // create a node to be put into the queue // Critically, after we reset() the event but prior to park(), we must check // for a pending interrupt. ObjectWaiter node(Self); node.TState = ObjectWaiter::TS_WAIT ; Self->_ParkEvent->reset() ; OrderAccess::fence(); // ST into Event; membar ; LD interrupted-flag // Enter the waiting queue, which is a circular doubly linked list in this case // but it could be a priority queue or any data structure. // _WaitSetLock protects the wait queue. Normally the wait queue is accessed only // by the the owner of the monitor *except* in the case where park() // returns because of a timeout of interrupt. Contention is exceptionally rare // so we use a simple spin-lock instead of a heavier-weight blocking lock. Thread::SpinAcquire (&_WaitSetLock, "WaitSet - add") ; AddWaiter (&node) ; Thread::SpinRelease (&_WaitSetLock) ; if ((SyncFlags & 4) == 0) { _Responsible = NULL ; } intptr_t save = _recursions; // record the old recursion count _waiters++; // increment the number of waiters _recursions = 0; // set the recursion level to be 1 exit (Self) ; // exit the monitor guarantee (_owner != Self, "invariant") ; // As soon as the ObjectMonitor's ownership is dropped in the exit() // call above, another thread can enter() the ObjectMonitor, do the // notify(), and exit() the ObjectMonitor. If the other thread's // exit() call chooses this thread as the successor and the unpark() // call happens to occur while this thread is posting a // MONITOR_CONTENDED_EXIT event, then we run the risk of the event // handler using RawMonitors and consuming the unpark(). // // To avoid the problem, we re-post the event. This does no harm // even if the original unpark() was not consumed because we are the // chosen successor for this monitor. if (node._notified != 0 && _succ == Self) { node._event->unpark(); } // The thread is on the WaitSet list - now park() it. // On MP systems it's conceivable that a brief spin before we park // could be profitable. // // TODO-FIXME: change the following logic to a loop of the form // while (!timeout && !interrupted && _notified == 0) park() int ret = OS_OK ; int WasNotified = 0 ; { // State transition wrappers OSThread* osthread = Self->osthread(); OSThreadWaitState osts(osthread, true); { ThreadBlockInVM tbivm(jt); // Thread is in thread_blocked state and oop access is unsafe. jt->set_suspend_equivalent(); if (interruptible && (Thread::is_interrupted(THREAD, false) || HAS_PENDING_EXCEPTION)) { // Intentionally empty } else if (node._notified == 0) { if (millis <= 0) { Self->_ParkEvent->park () ; } else { ret = Self->_ParkEvent->park (millis) ; } } // were we externally suspended while we were waiting? if (ExitSuspendEquivalent (jt)) { // TODO-FIXME: add -- if succ == Self then succ = null. jt->java_suspend_self(); } } // Exit thread safepoint: transition _thread_blocked -> _thread_in_vm // Node may be on the WaitSet, the EntryList (or cxq), or in transition // from the WaitSet to the EntryList. // See if we need to remove Node from the WaitSet. // We use double-checked locking to avoid grabbing _WaitSetLock // if the thread is not on the wait queue. // // Note that we don't need a fence before the fetch of TState. // In the worst case we'll fetch a old-stale value of TS_WAIT previously // written by the is thread. (perhaps the fetch might even be satisfied // by a look-aside into the processor's own store buffer, although given // the length of the code path between the prior ST and this load that's // highly unlikely). If the following LD fetches a stale TS_WAIT value // then we'll acquire the lock and then re-fetch a fresh TState value. // That is, we fail toward safety. if (node.TState == ObjectWaiter::TS_WAIT) { Thread::SpinAcquire (&_WaitSetLock, "WaitSet - unlink") ; if (node.TState == ObjectWaiter::TS_WAIT) { DequeueSpecificWaiter (&node) ; // unlink from WaitSet assert(node._notified == 0, "invariant"); node.TState = ObjectWaiter::TS_RUN ; } Thread::SpinRelease (&_WaitSetLock) ; } // The thread is now either on off-list (TS_RUN), // on the EntryList (TS_ENTER), or on the cxq (TS_CXQ). // The Node's TState variable is stable from the perspective of this thread. // No other threads will asynchronously modify TState. guarantee (node.TState != ObjectWaiter::TS_WAIT, "invariant") ; OrderAccess::loadload() ; if (_succ == Self) _succ = NULL ; WasNotified = node._notified ; // Reentry phase -- reacquire the monitor. // re-enter contended monitor after object.wait(). // retain OBJECT_WAIT state until re-enter successfully completes // Thread state is thread_in_vm and oop access is again safe, // although the raw address of the object may have changed. // (Don't cache naked oops over safepoints, of course). // post monitor waited event. Note that this is past-tense, we are done waiting. if (JvmtiExport::should_post_monitor_waited()) { JvmtiExport::post_monitor_waited(jt, this, ret == OS_TIMEOUT); } OrderAccess::fence() ; assert (Self->_Stalled != 0, "invariant") ; Self->_Stalled = 0 ; assert (_owner != Self, "invariant") ; ObjectWaiter::TStates v = node.TState ; if (v == ObjectWaiter::TS_RUN) { enter (Self) ; } else { guarantee (v == ObjectWaiter::TS_ENTER || v == ObjectWaiter::TS_CXQ, "invariant") ; ReenterI (Self, &node) ; node.wait_reenter_end(this); } // Self has reacquired the lock. // Lifecycle - the node representing Self must not appear on any queues. // Node is about to go out-of-scope, but even if it were immortal we wouldn't // want residual elements associated with this thread left on any lists. guarantee (node.TState == ObjectWaiter::TS_RUN, "invariant") ; assert (_owner == Self, "invariant") ; assert (_succ != Self , "invariant") ; } // OSThreadWaitState() jt->set_current_waiting_monitor(NULL); guarantee (_recursions == 0, "invariant") ; _recursions = save; // restore the old recursion count _waiters--; // decrement the number of waiters // Verify a few postconditions assert (_owner == Self , "invariant") ; assert (_succ != Self , "invariant") ; assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ; if (SyncFlags & 32) { OrderAccess::fence() ; } // check if the notification happened if (!WasNotified) { // no, it could be timeout or Thread.interrupt() or both // check for interrupt event, otherwise it is timeout if (interruptible && Thread::is_interrupted(Self, true) && !HAS_PENDING_EXCEPTION) { TEVENT (Wait - throw IEX from epilog) ; THROW(vmSymbols::java_lang_InterruptedException()); } } // NOTE: Spurious wake up will be consider as timeout. // Monitor notify has precedence over thread interrupt. } // Consider: // If the lock is cool (cxq == null && succ == null) and we're on an MP system // then instead of transferring a thread from the WaitSet to the EntryList // we might just dequeue a thread from the WaitSet and directly unpark() it. void ObjectMonitor::notify(TRAPS) { CHECK_OWNER(); if (_WaitSet == NULL) { TEVENT (Empty-Notify) ; return ; } DTRACE_MONITOR_PROBE(notify, this, object(), THREAD); int Policy = Knob_MoveNotifyee ; Thread::SpinAcquire (&_WaitSetLock, "WaitSet - notify") ; ObjectWaiter * iterator = DequeueWaiter() ; if (iterator != NULL) { TEVENT (Notify1 - Transfer) ; guarantee (iterator->TState == ObjectWaiter::TS_WAIT, "invariant") ; guarantee (iterator->_notified == 0, "invariant") ; // Disposition - what might we do with iterator ? // a. add it directly to the EntryList - either tail or head. // b. push it onto the front of the _cxq. // For now we use (a). if (Policy != 4) { iterator->TState = ObjectWaiter::TS_ENTER ; } iterator->_notified = 1 ; ObjectWaiter * List = _EntryList ; if (List != NULL) { assert (List->_prev == NULL, "invariant") ; assert (List->TState == ObjectWaiter::TS_ENTER, "invariant") ; assert (List != iterator, "invariant") ; } if (Policy == 0) { // prepend to EntryList if (List == NULL) { iterator->_next = iterator->_prev = NULL ; _EntryList = iterator ; } else { List->_prev = iterator ; iterator->_next = List ; iterator->_prev = NULL ; _EntryList = iterator ; } } else if (Policy == 1) { // append to EntryList if (List == NULL) { iterator->_next = iterator->_prev = NULL ; _EntryList = iterator ; } else { // CONSIDER: finding the tail currently requires a linear-time walk of // the EntryList. We can make tail access constant-time by converting to // a CDLL instead of using our current DLL. ObjectWaiter * Tail ; for (Tail = List ; Tail->_next != NULL ; Tail = Tail->_next) ; assert (Tail != NULL && Tail->_next == NULL, "invariant") ; Tail->_next = iterator ; iterator->_prev = Tail ; iterator->_next = NULL ; } } else if (Policy == 2) { // prepend to cxq // prepend to cxq if (List == NULL) { iterator->_next = iterator->_prev = NULL ; _EntryList = iterator ; } else { iterator->TState = ObjectWaiter::TS_CXQ ; for (;;) { ObjectWaiter * Front = _cxq ; iterator->_next = Front ; if (Atomic::cmpxchg_ptr (iterator, &_cxq, Front) == Front) { break ; } } } } else if (Policy == 3) { // append to cxq iterator->TState = ObjectWaiter::TS_CXQ ; for (;;) { ObjectWaiter * Tail ; Tail = _cxq ; if (Tail == NULL) { iterator->_next = NULL ; if (Atomic::cmpxchg_ptr (iterator, &_cxq, NULL) == NULL) { break ; } } else { while (Tail->_next != NULL) Tail = Tail->_next ; Tail->_next = iterator ; iterator->_prev = Tail ; iterator->_next = NULL ; break ; } } } else { ParkEvent * ev = iterator->_event ; iterator->TState = ObjectWaiter::TS_RUN ; OrderAccess::fence() ; ev->unpark() ; } if (Policy < 4) { iterator->wait_reenter_begin(this); } // _WaitSetLock protects the wait queue, not the EntryList. We could // move the add-to-EntryList operation, above, outside the critical section // protected by _WaitSetLock. In practice that's not useful. With the // exception of wait() timeouts and interrupts the monitor owner // is the only thread that grabs _WaitSetLock. There's almost no contention // on _WaitSetLock so it's not profitable to reduce the length of the // critical section. } Thread::SpinRelease (&_WaitSetLock) ; if (iterator != NULL && ObjectSynchronizer::_sync_Notifications != NULL) { ObjectSynchronizer::_sync_Notifications->inc() ; } } void ObjectMonitor::notifyAll(TRAPS) { CHECK_OWNER(); ObjectWaiter* iterator; if (_WaitSet == NULL) { TEVENT (Empty-NotifyAll) ; return ; } DTRACE_MONITOR_PROBE(notifyAll, this, object(), THREAD); int Policy = Knob_MoveNotifyee ; int Tally = 0 ; Thread::SpinAcquire (&_WaitSetLock, "WaitSet - notifyall") ; for (;;) { iterator = DequeueWaiter () ; if (iterator == NULL) break ; TEVENT (NotifyAll - Transfer1) ; ++Tally ; // Disposition - what might we do with iterator ? // a. add it directly to the EntryList - either tail or head. // b. push it onto the front of the _cxq. // For now we use (a). // // TODO-FIXME: currently notifyAll() transfers the waiters one-at-a-time from the waitset // to the EntryList. This could be done more efficiently with a single bulk transfer, // but in practice it's not time-critical. Beware too, that in prepend-mode we invert the // order of the waiters. Lets say that the waitset is "ABCD" and the EntryList is "XYZ". // After a notifyAll() in prepend mode the waitset will be empty and the EntryList will // be "DCBAXYZ". guarantee (iterator->TState == ObjectWaiter::TS_WAIT, "invariant") ; guarantee (iterator->_notified == 0, "invariant") ; iterator->_notified = 1 ; if (Policy != 4) { iterator->TState = ObjectWaiter::TS_ENTER ; } ObjectWaiter * List = _EntryList ; if (List != NULL) { assert (List->_prev == NULL, "invariant") ; assert (List->TState == ObjectWaiter::TS_ENTER, "invariant") ; assert (List != iterator, "invariant") ; } if (Policy == 0) { // prepend to EntryList if (List == NULL) { iterator->_next = iterator->_prev = NULL ; _EntryList = iterator ; } else { List->_prev = iterator ; iterator->_next = List ; iterator->_prev = NULL ; _EntryList = iterator ; } } else if (Policy == 1) { // append to EntryList if (List == NULL) { iterator->_next = iterator->_prev = NULL ; _EntryList = iterator ; } else { // CONSIDER: finding the tail currently requires a linear-time walk of // the EntryList. We can make tail access constant-time by converting to // a CDLL instead of using our current DLL. ObjectWaiter * Tail ; for (Tail = List ; Tail->_next != NULL ; Tail = Tail->_next) ; assert (Tail != NULL && Tail->_next == NULL, "invariant") ; Tail->_next = iterator ; iterator->_prev = Tail ; iterator->_next = NULL ; } } else if (Policy == 2) { // prepend to cxq // prepend to cxq iterator->TState = ObjectWaiter::TS_CXQ ; for (;;) { ObjectWaiter * Front = _cxq ; iterator->_next = Front ; if (Atomic::cmpxchg_ptr (iterator, &_cxq, Front) == Front) { break ; } } } else if (Policy == 3) { // append to cxq iterator->TState = ObjectWaiter::TS_CXQ ; for (;;) { ObjectWaiter * Tail ; Tail = _cxq ; if (Tail == NULL) { iterator->_next = NULL ; if (Atomic::cmpxchg_ptr (iterator, &_cxq, NULL) == NULL) { break ; } } else { while (Tail->_next != NULL) Tail = Tail->_next ; Tail->_next = iterator ; iterator->_prev = Tail ; iterator->_next = NULL ; break ; } } } else { ParkEvent * ev = iterator->_event ; iterator->TState = ObjectWaiter::TS_RUN ; OrderAccess::fence() ; ev->unpark() ; } if (Policy < 4) { iterator->wait_reenter_begin(this); } // _WaitSetLock protects the wait queue, not the EntryList. We could // move the add-to-EntryList operation, above, outside the critical section // protected by _WaitSetLock. In practice that's not useful. With the // exception of wait() timeouts and interrupts the monitor owner // is the only thread that grabs _WaitSetLock. There's almost no contention // on _WaitSetLock so it's not profitable to reduce the length of the // critical section. } Thread::SpinRelease (&_WaitSetLock) ; if (Tally != 0 && ObjectSynchronizer::_sync_Notifications != NULL) { ObjectSynchronizer::_sync_Notifications->inc(Tally) ; } } // check_slow() is a misnomer. It's called to simply to throw an IMSX exception. // TODO-FIXME: remove check_slow() -- it's likely dead. void ObjectMonitor::check_slow(TRAPS) { TEVENT (check_slow - throw IMSX) ; assert(THREAD != _owner && !THREAD->is_lock_owned((address) _owner), "must not be owner"); THROW_MSG(vmSymbols::java_lang_IllegalMonitorStateException(), "current thread not owner"); } // ------------------------------------------------------------------------- // The raw monitor subsystem is entirely distinct from normal // java-synchronization or jni-synchronization. raw monitors are not // associated with objects. They can be implemented in any manner // that makes sense. The original implementors decided to piggy-back // the raw-monitor implementation on the existing Java objectMonitor mechanism. // This flaw needs to fixed. We should reimplement raw monitors as sui-generis. // Specifically, we should not implement raw monitors via java monitors. // Time permitting, we should disentangle and deconvolve the two implementations // and move the resulting raw monitor implementation over to the JVMTI directories. // Ideally, the raw monitor implementation would be built on top of // park-unpark and nothing else. // // raw monitors are used mainly by JVMTI // The raw monitor implementation borrows the ObjectMonitor structure, // but the operators are degenerate and extremely simple. // // Mixed use of a single objectMonitor instance -- as both a raw monitor // and a normal java monitor -- is not permissible. // // Note that we use the single RawMonitor_lock to protect queue operations for // _all_ raw monitors. This is a scalability impediment, but since raw monitor usage // is deprecated and rare, this is not of concern. The RawMonitor_lock can not // be held indefinitely. The critical sections must be short and bounded. // // ------------------------------------------------------------------------- int ObjectMonitor::SimpleEnter (Thread * Self) { for (;;) { if (Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) { return OS_OK ; } ObjectWaiter Node (Self) ; Self->_ParkEvent->reset() ; // strictly optional Node.TState = ObjectWaiter::TS_ENTER ; RawMonitor_lock->lock_without_safepoint_check() ; Node._next = _EntryList ; _EntryList = &Node ; OrderAccess::fence() ; if (_owner == NULL && Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) { _EntryList = Node._next ; RawMonitor_lock->unlock() ; return OS_OK ; } RawMonitor_lock->unlock() ; while (Node.TState == ObjectWaiter::TS_ENTER) { Self->_ParkEvent->park() ; } } } int ObjectMonitor::SimpleExit (Thread * Self) { guarantee (_owner == Self, "invariant") ; OrderAccess::release_store_ptr (&_owner, NULL) ; OrderAccess::fence() ; if (_EntryList == NULL) return OS_OK ; ObjectWaiter * w ; RawMonitor_lock->lock_without_safepoint_check() ; w = _EntryList ; if (w != NULL) { _EntryList = w->_next ; } RawMonitor_lock->unlock() ; if (w != NULL) { guarantee (w ->TState == ObjectWaiter::TS_ENTER, "invariant") ; ParkEvent * ev = w->_event ; w->TState = ObjectWaiter::TS_RUN ; OrderAccess::fence() ; ev->unpark() ; } return OS_OK ; } int ObjectMonitor::SimpleWait (Thread * Self, jlong millis) { guarantee (_owner == Self , "invariant") ; guarantee (_recursions == 0, "invariant") ; ObjectWaiter Node (Self) ; Node._notified = 0 ; Node.TState = ObjectWaiter::TS_WAIT ; RawMonitor_lock->lock_without_safepoint_check() ; Node._next = _WaitSet ; _WaitSet = &Node ; RawMonitor_lock->unlock() ; SimpleExit (Self) ; guarantee (_owner != Self, "invariant") ; int ret = OS_OK ; if (millis <= 0) { Self->_ParkEvent->park(); } else { ret = Self->_ParkEvent->park(millis); } // If thread still resides on the waitset then unlink it. // Double-checked locking -- the usage is safe in this context // as we TState is volatile and the lock-unlock operators are // serializing (barrier-equivalent). if (Node.TState == ObjectWaiter::TS_WAIT) { RawMonitor_lock->lock_without_safepoint_check() ; if (Node.TState == ObjectWaiter::TS_WAIT) { // Simple O(n) unlink, but performance isn't critical here. ObjectWaiter * p ; ObjectWaiter * q = NULL ; for (p = _WaitSet ; p != &Node; p = p->_next) { q = p ; } guarantee (p == &Node, "invariant") ; if (q == NULL) { guarantee (p == _WaitSet, "invariant") ; _WaitSet = p->_next ; } else { guarantee (p == q->_next, "invariant") ; q->_next = p->_next ; } Node.TState = ObjectWaiter::TS_RUN ; } RawMonitor_lock->unlock() ; } guarantee (Node.TState == ObjectWaiter::TS_RUN, "invariant") ; SimpleEnter (Self) ; guarantee (_owner == Self, "invariant") ; guarantee (_recursions == 0, "invariant") ; return ret ; } int ObjectMonitor::SimpleNotify (Thread * Self, bool All) { guarantee (_owner == Self, "invariant") ; if (_WaitSet == NULL) return OS_OK ; // We have two options: // A. Transfer the threads from the WaitSet to the EntryList // B. Remove the thread from the WaitSet and unpark() it. // // We use (B), which is crude and results in lots of futile // context switching. In particular (B) induces lots of contention. ParkEvent * ev = NULL ; // consider using a small auto array ... RawMonitor_lock->lock_without_safepoint_check() ; for (;;) { ObjectWaiter * w = _WaitSet ; if (w == NULL) break ; _WaitSet = w->_next ; if (ev != NULL) { ev->unpark(); ev = NULL; } ev = w->_event ; OrderAccess::loadstore() ; w->TState = ObjectWaiter::TS_RUN ; OrderAccess::storeload(); if (!All) break ; } RawMonitor_lock->unlock() ; if (ev != NULL) ev->unpark(); return OS_OK ; } // Any JavaThread will enter here with state _thread_blocked int ObjectMonitor::raw_enter(TRAPS) { TEVENT (raw_enter) ; void * Contended ; // don't enter raw monitor if thread is being externally suspended, it will // surprise the suspender if a "suspended" thread can still enter monitor JavaThread * jt = (JavaThread *)THREAD; if (THREAD->is_Java_thread()) { jt->SR_lock()->lock_without_safepoint_check(); while (jt->is_external_suspend()) { jt->SR_lock()->unlock(); jt->java_suspend_self(); jt->SR_lock()->lock_without_safepoint_check(); } // guarded by SR_lock to avoid racing with new external suspend requests. Contended = Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) ; jt->SR_lock()->unlock(); } else { Contended = Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) ; } if (Contended == THREAD) { _recursions ++ ; return OM_OK ; } if (Contended == NULL) { guarantee (_owner == THREAD, "invariant") ; guarantee (_recursions == 0, "invariant") ; return OM_OK ; } THREAD->set_current_pending_monitor(this); if (!THREAD->is_Java_thread()) { // No other non-Java threads besides VM thread would acquire // a raw monitor. assert(THREAD->is_VM_thread(), "must be VM thread"); SimpleEnter (THREAD) ; } else { guarantee (jt->thread_state() == _thread_blocked, "invariant") ; for (;;) { jt->set_suspend_equivalent(); // cleared by handle_special_suspend_equivalent_condition() or // java_suspend_self() SimpleEnter (THREAD) ; // were we externally suspended while we were waiting? if (!jt->handle_special_suspend_equivalent_condition()) break ; // This thread was externally suspended // // This logic isn't needed for JVMTI raw monitors, // but doesn't hurt just in case the suspend rules change. This // logic is needed for the ObjectMonitor.wait() reentry phase. // We have reentered the contended monitor, but while we were // waiting another thread suspended us. We don't want to reenter // the monitor while suspended because that would surprise the // thread that suspended us. // // Drop the lock - SimpleExit (THREAD) ; jt->java_suspend_self(); } assert(_owner == THREAD, "Fatal error with monitor owner!"); assert(_recursions == 0, "Fatal error with monitor recursions!"); } THREAD->set_current_pending_monitor(NULL); guarantee (_recursions == 0, "invariant") ; return OM_OK; } // Used mainly for JVMTI raw monitor implementation // Also used for ObjectMonitor::wait(). int ObjectMonitor::raw_exit(TRAPS) { TEVENT (raw_exit) ; if (THREAD != _owner) { return OM_ILLEGAL_MONITOR_STATE; } if (_recursions > 0) { --_recursions ; return OM_OK ; } void * List = _EntryList ; SimpleExit (THREAD) ; return OM_OK; } // Used for JVMTI raw monitor implementation. // All JavaThreads will enter here with state _thread_blocked int ObjectMonitor::raw_wait(jlong millis, bool interruptible, TRAPS) { TEVENT (raw_wait) ; if (THREAD != _owner) { return OM_ILLEGAL_MONITOR_STATE; } // To avoid spurious wakeups we reset the parkevent -- This is strictly optional. // The caller must be able to tolerate spurious returns from raw_wait(). THREAD->_ParkEvent->reset() ; OrderAccess::fence() ; // check interrupt event if (interruptible && Thread::is_interrupted(THREAD, true)) { return OM_INTERRUPTED; } intptr_t save = _recursions ; _recursions = 0 ; _waiters ++ ; if (THREAD->is_Java_thread()) { guarantee (((JavaThread *) THREAD)->thread_state() == _thread_blocked, "invariant") ; ((JavaThread *)THREAD)->set_suspend_equivalent(); } int rv = SimpleWait (THREAD, millis) ; _recursions = save ; _waiters -- ; guarantee (THREAD == _owner, "invariant") ; if (THREAD->is_Java_thread()) { JavaThread * jSelf = (JavaThread *) THREAD ; for (;;) { if (!jSelf->handle_special_suspend_equivalent_condition()) break ; SimpleExit (THREAD) ; jSelf->java_suspend_self(); SimpleEnter (THREAD) ; jSelf->set_suspend_equivalent() ; } } guarantee (THREAD == _owner, "invariant") ; if (interruptible && Thread::is_interrupted(THREAD, true)) { return OM_INTERRUPTED; } return OM_OK ; } int ObjectMonitor::raw_notify(TRAPS) { TEVENT (raw_notify) ; if (THREAD != _owner) { return OM_ILLEGAL_MONITOR_STATE; } SimpleNotify (THREAD, false) ; return OM_OK; } int ObjectMonitor::raw_notifyAll(TRAPS) { TEVENT (raw_notifyAll) ; if (THREAD != _owner) { return OM_ILLEGAL_MONITOR_STATE; } SimpleNotify (THREAD, true) ; return OM_OK; } #ifndef PRODUCT void ObjectMonitor::verify() { } void ObjectMonitor::print() { } #endif //------------------------------------------------------------------------------ // Non-product code #ifndef PRODUCT void ObjectSynchronizer::trace_locking(Handle locking_obj, bool is_compiled, bool is_method, bool is_locking) { // Don't know what to do here } // Verify all monitors in the monitor cache, the verification is weak. void ObjectSynchronizer::verify() { ObjectMonitor* block = gBlockList; ObjectMonitor* mid; while (block) { assert(block->object() == CHAINMARKER, "must be a block header"); for (int i = 1; i < _BLOCKSIZE; i++) { mid = block + i; oop object = (oop) mid->object(); if (object != NULL) { mid->verify(); } } block = (ObjectMonitor*) block->FreeNext; } } // Check if monitor belongs to the monitor cache // The list is grow-only so it's *relatively* safe to traverse // the list of extant blocks without taking a lock. int ObjectSynchronizer::verify_objmon_isinpool(ObjectMonitor *monitor) { ObjectMonitor* block = gBlockList; while (block) { assert(block->object() == CHAINMARKER, "must be a block header"); if (monitor > &block[0] && monitor < &block[_BLOCKSIZE]) { address mon = (address) monitor; address blk = (address) block; size_t diff = mon - blk; assert((diff % sizeof(ObjectMonitor)) == 0, "check"); return 1; } block = (ObjectMonitor*) block->FreeNext; } return 0; } #endif