JDK14/Java14源码在线阅读
/*
* Copyright (c) 1998, 2019, Oracle and/or its affiliates. All rights reserved.
* DO NOT ALTER OR REMOVE COPYRIGHT NOTICES OR THIS FILE HEADER.
*
* This code is free software; you can redistribute it and/or modify it
* under the terms of the GNU General Public License version 2 only, as
* published by the Free Software Foundation.
*
* This code is distributed in the hope that it will be useful, but WITHOUT
* ANY WARRANTY; without even the implied warranty of MERCHANTABILITY or
* FITNESS FOR A PARTICULAR PURPOSE. See the GNU General Public License
* version 2 for more details (a copy is included in the LICENSE file that
* accompanied this code).
*
* You should have received a copy of the GNU General Public License version
* 2 along with this work; if not, write to the Free Software Foundation,
* Inc., 51 Franklin St, Fifth Floor, Boston, MA 02110-1301 USA.
*
* Please contact Oracle, 500 Oracle Parkway, Redwood Shores, CA 94065 USA
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*/
#include "precompiled.hpp"
#include "classfile/vmSymbols.hpp"
#include "jfr/jfrEvents.hpp"
#include "jfr/support/jfrThreadId.hpp"
#include "logging/log.hpp"
#include "logging/logStream.hpp"
#include "memory/allocation.inline.hpp"
#include "memory/resourceArea.hpp"
#include "oops/markWord.hpp"
#include "oops/oop.inline.hpp"
#include "runtime/atomic.hpp"
#include "runtime/handles.inline.hpp"
#include "runtime/interfaceSupport.inline.hpp"
#include "runtime/mutexLocker.hpp"
#include "runtime/objectMonitor.hpp"
#include "runtime/objectMonitor.inline.hpp"
#include "runtime/orderAccess.hpp"
#include "runtime/osThread.hpp"
#include "runtime/safepointMechanism.inline.hpp"
#include "runtime/sharedRuntime.hpp"
#include "runtime/stubRoutines.hpp"
#include "runtime/thread.inline.hpp"
#include "services/threadService.hpp"
#include "utilities/dtrace.hpp"
#include "utilities/macros.hpp"
#include "utilities/preserveException.hpp"
#if INCLUDE_JFR
#include "jfr/support/jfrFlush.hpp"
#endif
#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.
#define DTRACE_MONITOR_PROBE_COMMON(obj, thread) \
char* bytes = NULL; \
int len = 0; \
jlong jtid = SharedRuntime::get_java_tid(thread); \
Symbol* klassname = ((oop)obj)->klass()->name(); \
if (klassname != NULL) { \
bytes = (char*)klassname->bytes(); \
len = klassname->utf8_length(); \
}
#define DTRACE_MONITOR_WAIT_PROBE(monitor, obj, thread, millis) \
{ \
if (DTraceMonitorProbes) { \
DTRACE_MONITOR_PROBE_COMMON(obj, thread); \
HOTSPOT_MONITOR_WAIT(jtid, \
(monitor), bytes, len, (millis)); \
} \
}
#define HOTSPOT_MONITOR_contended__enter HOTSPOT_MONITOR_CONTENDED_ENTER
#define HOTSPOT_MONITOR_contended__entered HOTSPOT_MONITOR_CONTENDED_ENTERED
#define HOTSPOT_MONITOR_contended__exit HOTSPOT_MONITOR_CONTENDED_EXIT
#define HOTSPOT_MONITOR_notify HOTSPOT_MONITOR_NOTIFY
#define HOTSPOT_MONITOR_notifyAll HOTSPOT_MONITOR_NOTIFYALL
#define DTRACE_MONITOR_PROBE(probe, monitor, obj, thread) \
{ \
if (DTraceMonitorProbes) { \
DTRACE_MONITOR_PROBE_COMMON(obj, thread); \
HOTSPOT_MONITOR_##probe(jtid, \
(uintptr_t)(monitor), bytes, len); \
} \
}
#else // ndef DTRACE_ENABLED
#define DTRACE_MONITOR_WAIT_PROBE(obj, thread, millis, mon) {;}
#define DTRACE_MONITOR_PROBE(probe, obj, thread, mon) {;}
#endif // ndef DTRACE_ENABLED
// Tunables ...
// The knob* variables are effectively final. Once set they should
// never be modified hence. Consider using __read_mostly with GCC.
int ObjectMonitor::Knob_SpinLimit = 5000; // derived by an external tool -
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_FixedSpin = 0;
static int Knob_PreSpin = 10; // 20-100 likely better
DEBUG_ONLY(static volatile bool InitDone = false;)
// -----------------------------------------------------------------------------
// 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.
// For a description of 'Methods and apparatus providing non-blocking access
// to a resource,' see U.S. Pat. No. 7844973.
//
// * 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() epilogue) 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 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* ObjectMonitor::operator new (size_t size) throw() {
return AllocateHeap(size, mtInternal);
}
void* ObjectMonitor::operator new[] (size_t size) throw() {
return operator new (size);
}
void ObjectMonitor::operator delete(void* p) {
FreeHeap(p);
}
void ObjectMonitor::operator delete[] (void *p) {
operator delete(p);
}
// -----------------------------------------------------------------------------
// Enter support
void 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 = Atomic::cmpxchg(&_owner, (void*)NULL, Self);
if (cur == NULL) {
assert(_recursions == 0, "invariant");
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;
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 (TrySpin(Self) > 0) {
assert(_owner == Self, "must be Self: owner=" INTPTR_FORMAT, p2i(_owner));
assert(_recursions == 0, "must be 0: recursions=" INTX_FORMAT, _recursions);
assert(((oop)object())->mark() == markWord::encode(this),
"object mark must match encoded this: mark=" INTPTR_FORMAT
", encoded this=" INTPTR_FORMAT, ((oop)object())->mark().value(),
markWord::encode(this).value());
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(_contentions >= 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(&_contentions);
JFR_ONLY(JfrConditionalFlushWithStacktrace<EventJavaMonitorEnter> flush(jt);)
EventJavaMonitorEnter event;
if (event.should_commit()) {
event.set_monitorClass(((oop)this->object())->klass());
event.set_address((uintptr_t)(this->object_addr()));
}
{ // Change java thread status to indicate blocked on monitor enter.
JavaThreadBlockedOnMonitorEnterState jtbmes(jt, this);
Self->set_current_pending_monitor(this);
DTRACE_MONITOR_PROBE(contended__enter, this, object(), jt);
if (JvmtiExport::should_post_monitor_contended_enter()) {
JvmtiExport::post_monitor_contended_enter(jt, this);
// The current thread does not yet own the monitor and does not
// yet appear on any queues that would get it made the successor.
// This means that the JVMTI_EVENT_MONITOR_CONTENDED_ENTER event
// handler cannot accidentally consume an unpark() meant for the
// ParkEvent associated with this ObjectMonitor.
}
OSThreadContendState osts(Self->osthread());
ThreadBlockInVM tbivm(jt);
// 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(false, Self);
jt->java_suspend_self();
}
Self->set_current_pending_monitor(NULL);
// We cleared the pending monitor info since we've just gotten past
// the enter-check-for-suspend dance and we now own the monitor free
// and clear, i.e., it is no longer pending. The ThreadBlockInVM
// destructor can go to a safepoint at the end of this block. If we
// do a thread dump during that safepoint, then this thread will show
// as having "-locked" the monitor, but the OS and java.lang.Thread
// states will still report that the thread is blocked trying to
// acquire it.
}
Atomic::dec(&_contentions);
assert(_contentions >= 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() == markWord::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);
// The current thread already owns the monitor and is not going to
// call park() for the remainder of the monitor enter protocol. So
// it doesn't matter if the JVMTI_EVENT_MONITOR_CONTENDED_ENTERED
// event handler consumed an unpark() issued by the thread that
// just exited the monitor.
}
if (event.should_commit()) {
event.set_previousOwner((uintptr_t)_previous_owner_tid);
event.commit();
}
OM_PERFDATA_OP(ContendedLockAttempts, inc());
}
// Caveat: TryLock() is not necessarily serializing if it returns failure.
// Callers must compensate as needed.
int ObjectMonitor::TryLock(Thread * Self) {
void * own = _owner;
if (own != NULL) return 0;
if (Atomic::replace_if_null(&_owner, Self)) {
assert(_recursions == 0, "invariant");
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.
return -1;
}
// Convert the fields used by is_busy() to a string that can be
// used for diagnostic output.
const char* ObjectMonitor::is_busy_to_string(stringStream* ss) {
ss->print("is_busy: contentions=%d, waiters=%d, owner=" INTPTR_FORMAT
", cxq=" INTPTR_FORMAT ", EntryList=" INTPTR_FORMAT, _contentions,
_waiters, p2i(_owner), p2i(_cxq), p2i(_EntryList));
return ss->base();
}
#define MAX_RECHECK_INTERVAL 1000
void ObjectMonitor::EnterI(TRAPS) {
Thread * const 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;
}
assert(InitDone, "Unexpectedly not initialized");
// 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(&_cxq, nxt, &node) == 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 (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::replace_if_null(&_Responsible, Self);
}
// The lock might 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 ...
int nWakeups = 0;
int recheckInterval = 1;
for (;;) {
if (TryLock(Self) > 0) break;
assert(_owner != Self, "invariant");
// park self
if (_Responsible == Self) {
Self->_ParkEvent->park((jlong) recheckInterval);
// Increase the recheckInterval, but clamp the value.
recheckInterval *= 8;
if (recheckInterval > MAX_RECHECK_INTERVAL) {
recheckInterval = MAX_RECHECK_INTERVAL;
}
} else {
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.
// This PerfData object can be used in parallel with a safepoint.
// See the work around in PerfDataManager::destroy().
OM_PERFDATA_OP(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 (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.
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() == markWord::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;
OrderAccess::fence(); // Dekker pivot-point
// 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 epilogue - 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.
}
// 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.
return;
}
// 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().
void 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() == markWord::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;
// 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();
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.
++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();
// This PerfData object can be used in parallel with a safepoint.
// See the work around in PerfDataManager::destroy().
OM_PERFDATA_OP(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() == markWord::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()
}
// 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");
} else {
assert(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(&_cxq, v, SelfNode->_next) != 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;
}
}
#ifdef ASSERT
// Diagnostic hygiene ...
SelfNode->_prev = (ObjectWaiter *) 0xBAD;
SelfNode->_next = (ObjectWaiter *) 0xBAD;
SelfNode->TState = ObjectWaiter::TS_RUN;
#endif
}
// -----------------------------------------------------------------------------
// Exit support
//
// 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 _contentions 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, e.g., see macroAssembler_x86.cpp: fast_unlock().
// 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 exit 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.
//
// There is also the risk of a futile wake-up. If we drop the lock
// another thread can reacquire the lock immediately, and we can
// then wake a thread unnecessarily. This is benign, and we've
// structured the code so the windows are short and the frequency
// of such futile wakups is low.
void ObjectMonitor::exit(bool not_suspended, TRAPS) {
Thread * const 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;
} else {
// Apparent unbalanced locking ...
// Naively we'd like to throw IllegalMonitorStateException.
// As a practical matter we can neither allocate nor throw an
// exception as ::exit() can be called from leaf routines.
// see x86_32.ad Fast_Unlock() and the I1 and I2 properties.
// Upon deeper reflection, however, in a properly run JVM the only
// way we should encounter this situation is in the presence of
// unbalanced JNI locking. TODO: CheckJNICalls.
// See also: CR4414101
#ifdef ASSERT
LogStreamHandle(Error, monitorinflation) lsh;
lsh.print_cr("ERROR: ObjectMonitor::exit(): thread=" INTPTR_FORMAT
" is exiting an ObjectMonitor it does not own.", p2i(THREAD));
lsh.print_cr("The imbalance is possibly caused by JNI locking.");
print_debug_style_on(&lsh);
#endif
assert(false, "Non-balanced monitor enter/exit!");
return;
}
}
if (_recursions != 0) {
_recursions--; // this is simple recursive enter
return;
}
// Invariant: after setting Responsible=null an thread must execute
// a MEMBAR or other serializing instruction before fetching EntryList|cxq.
_Responsible = NULL;
#if INCLUDE_JFR
// get the owner's thread id for the MonitorEnter event
// if it is enabled and the thread isn't suspended
if (not_suspended && EventJavaMonitorEnter::is_enabled()) {
_previous_owner_tid = JFR_THREAD_ID(Self);
}
#endif
for (;;) {
assert(THREAD == _owner, "invariant");
// release semantics: prior loads and stores from within the critical section
// must not float (reorder) past the following store that drops the lock.
Atomic::release_store(&_owner, (void*)NULL); // drop the lock
OrderAccess::storeload(); // See if we need to wake a successor
if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) {
return;
}
// Other threads are blocked trying to acquire the lock.
// 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::replace_if_null(&_owner, THREAD)) {
return;
}
guarantee(_owner == THREAD, "invariant");
ObjectWaiter * w = NULL;
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.
// Let's 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.
// Let's 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 = Atomic::cmpxchg(&_cxq, w, (ObjectWaiter*)NULL);
if (u == w) break;
w = u;
}
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.
_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;
}
}
}
// 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) {
return jSelf->handle_special_suspend_equivalent_condition();
}
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 = Wakee->_thread;
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
Atomic::release_store(&_owner, (void*)NULL);
OrderAccess::fence(); // ST _owner vs LD in unpark()
DTRACE_MONITOR_PROBE(contended__exit, this, object(), Self);
Trigger->unpark();
// Maintain stats and report events to JVMTI
OM_PERFDATA_OP(Parks, inc());
}
// -----------------------------------------------------------------------------
// Class Loader deadlock handling.
//
// 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.
intx ObjectMonitor::complete_exit(TRAPS) {
Thread * const Self = THREAD;
assert(Self->is_Java_thread(), "Must be Java thread!");
JavaThread *jt = (JavaThread *)THREAD;
assert(InitDone, "Unexpectedly not initialized");
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;
}
}
guarantee(Self == _owner, "complete_exit not owner");
intx save = _recursions; // record the old recursion count
_recursions = 0; // set the recursion level to be 0
exit(true, 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(intx 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;
}
// Checks that the current THREAD owns this monitor and causes an
// immediate return if it doesn't. We don't use the CHECK macro
// because we want the IMSE to be the only exception that is thrown
// from the call site when false is returned. Any other pending
// exception is ignored.
#define CHECK_OWNER() \
do { \
if (!check_owner(THREAD)) { \
assert(HAS_PENDING_EXCEPTION, "expected a pending IMSE here."); \
return; \
} \
} while (false)
// Returns true if the specified thread owns the ObjectMonitor.
// Otherwise returns false and throws IllegalMonitorStateException
// (IMSE). If there is a pending exception and the specified thread
// is not the owner, that exception will be replaced by the IMSE.
bool ObjectMonitor::check_owner(Thread* THREAD) {
if (_owner == THREAD) {
return true;
}
if (THREAD->is_lock_owned((address)_owner)) {
_owner = THREAD; // convert from BasicLock addr to Thread addr
_recursions = 0;
return true;
}
THROW_MSG_(vmSymbols::java_lang_IllegalMonitorStateException(),
"current thread is not owner", false);
}
static void post_monitor_wait_event(EventJavaMonitorWait* event,
ObjectMonitor* monitor,
jlong notifier_tid,
jlong timeout,
bool timedout) {
assert(event != NULL, "invariant");
assert(monitor != NULL, "invariant");
event->set_monitorClass(((oop)monitor->object())->klass());
event->set_timeout(timeout);
event->set_address((uintptr_t)monitor->object_addr());
event->set_notifier(notifier_tid);
event->set_timedOut(timedout);
event->commit();
}
// -----------------------------------------------------------------------------
// Wait/Notify/NotifyAll
//
// Note: a subset of changes to ObjectMonitor::wait()
// will need to be replicated in complete_exit
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;
assert(InitDone, "Unexpectedly not initialized");
CHECK_OWNER(); // Throws IMSE if not owner.
EventJavaMonitorWait event;
// check for a pending interrupt
if (interruptible && jt->is_interrupted(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);
// In this short circuit of the monitor wait protocol, the
// current thread never drops ownership of the monitor and
// never gets added to the wait queue so the current thread
// cannot be made the successor. This means that the
// JVMTI_EVENT_MONITOR_WAITED event handler cannot accidentally
// consume an unpark() meant for the ParkEvent associated with
// this ObjectMonitor.
}
if (event.should_commit()) {
post_monitor_wait_event(&event, this, 0, millis, false);
}
THROW(vmSymbols::java_lang_InterruptedException());
return;
}
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);
_Responsible = NULL;
intx save = _recursions; // record the old recursion count
_waiters++; // increment the number of waiters
_recursions = 0; // set the recursion level to be 1
exit(true, Self); // exit the monitor
guarantee(_owner != Self, "invariant");
// 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;
// Need to check interrupt state whilst still _thread_in_vm
bool interrupted = interruptible && jt->is_interrupted(false);
{ // 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 (interrupted || 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);
if (node._notified != 0 && _succ == Self) {
// In this part of the monitor wait-notify-reenter protocol it
// is possible (and normal) for another thread to do a fastpath
// monitor enter-exit while this thread is still trying to get
// to the reenter portion of the protocol.
//
// The ObjectMonitor was notified and the current thread is
// the successor which also means that an unpark() has already
// been done. The JVMTI_EVENT_MONITOR_WAITED event handler can
// consume the unpark() that was done when the successor was
// set because the same ParkEvent is shared between Java
// monitors and JVM/TI RawMonitors (for now).
//
// We redo the unpark() to ensure forward progress, i.e., we
// don't want all pending threads hanging (parked) with none
// entering the unlocked monitor.
node._event->unpark();
}
}
if (event.should_commit()) {
post_monitor_wait_event(&event, this, node._notifier_tid, millis, 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() == markWord::encode(this), "invariant");
// 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 && jt->is_interrupted(true) && !HAS_PENDING_EXCEPTION) {
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::INotify(Thread * Self) {
Thread::SpinAcquire(&_WaitSetLock, "WaitSet - notify");
ObjectWaiter * iterator = DequeueWaiter();
if (iterator != NULL) {
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 (policy == 1)
// or head (policy == 0).
// b. push it onto the front of the _cxq (policy == 2).
// For now we use (b).
iterator->TState = ObjectWaiter::TS_ENTER;
iterator->_notified = 1;
iterator->_notifier_tid = JFR_THREAD_ID(Self);
ObjectWaiter * list = _EntryList;
if (list != NULL) {
assert(list->_prev == NULL, "invariant");
assert(list->TState == ObjectWaiter::TS_ENTER, "invariant");
assert(list != iterator, "invariant");
}
// 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(&_cxq, front, iterator) == front) {
break;
}
}
}
// _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.
iterator->wait_reenter_begin(this);
}
Thread::SpinRelease(&_WaitSetLock);
}
// Consider: a not-uncommon synchronization bug is to use notify() when
// notifyAll() is more appropriate, potentially resulting in stranded
// threads; this is one example of a lost wakeup. A useful diagnostic
// option is to force all notify() operations to behave as notifyAll().
//
// Note: We can also detect many such problems with a "minimum wait".
// When the "minimum wait" is set to a small non-zero timeout value
// and the program does not hang whereas it did absent "minimum wait",
// that suggests a lost wakeup bug.
void ObjectMonitor::notify(TRAPS) {
CHECK_OWNER(); // Throws IMSE if not owner.
if (_WaitSet == NULL) {
return;
}
DTRACE_MONITOR_PROBE(notify, this, object(), THREAD);
INotify(THREAD);
OM_PERFDATA_OP(Notifications, inc(1));
}
// The current implementation of 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. Let's 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".
void ObjectMonitor::notifyAll(TRAPS) {
CHECK_OWNER(); // Throws IMSE if not owner.
if (_WaitSet == NULL) {
return;
}
DTRACE_MONITOR_PROBE(notifyAll, this, object(), THREAD);
int tally = 0;
while (_WaitSet != NULL) {
tally++;
INotify(THREAD);
}
OM_PERFDATA_OP(Notifications, inc(tally));
}
// -----------------------------------------------------------------------------
// Adaptive Spinning Support
//
// 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 cache 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.
// For a description of 'Adaptive spin-then-block mutual exclusion in
// multi-threaded processing,' see U.S. Pat. No. 8046758.
//
// 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.
// Spinning: Fixed frequency (100%), vary duration
int ObjectMonitor::TrySpin(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 <= 0) return 0;
if (NotRunnable(Self, (Thread *) _owner)) {
return 0;
}
// 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.
if (_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 (SafepointMechanism::should_block(Self)) {
goto Abort; // abrupt spin egress
}
SpinPause();
}
// 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(&_owner, (void*)NULL, Self);
if (ox == NULL) {
// The CAS succeeded -- this thread acquired ownership
// Take care of some bookkeeping to exit spin state.
if (_succ == Self) {
_succ = NULL;
}
// 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 -= CASPenalty
// * exit spin with prejudice -- goto Abort;
// * exit spin without prejudice.
// * Since CAS is high-latency, retry again immediately.
prv = ox;
goto Abort;
}
// Did lock ownership change hands ?
if (ox != prv && prv != NULL) {
goto Abort;
}
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 (NotRunnable(Self, ox)) {
goto Abort;
}
if (_succ == NULL) {
_succ = Self;
}
}
// Spin failed with prejudice -- reduce _SpinDuration.
// TODO: Use an AIMD-like policy to adjust _SpinDuration.
// AIMD is globally stable.
{
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 (_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;
}
// NotRunnable() -- informed spinning
//
// Don't bother spinning if the owner is not eligible to drop the lock.
// 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 _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.
// Alternately, we might tag the type (thread pointer vs basiclock pointer)
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