8146801: Allocating short arrays of non-constant size is slow
Reviewed-by: kvn, twisti, vlivanov
/*
* Copyright (c) 1998, 2015, 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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* questions.
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*/
#include "precompiled.hpp"
#include "runtime/atomic.inline.hpp"
#include "runtime/mutex.hpp"
#include "runtime/orderAccess.inline.hpp"
#include "runtime/osThread.hpp"
#include "runtime/thread.inline.hpp"
#include "utilities/events.hpp"
#ifdef TARGET_OS_FAMILY_linux
# include "mutex_linux.inline.hpp"
#endif
#ifdef TARGET_OS_FAMILY_solaris
# include "mutex_solaris.inline.hpp"
#endif
#ifdef TARGET_OS_FAMILY_windows
# include "mutex_windows.inline.hpp"
#endif
#ifdef TARGET_OS_FAMILY_bsd
# include "mutex_bsd.inline.hpp"
#endif
// o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o
//
// Native Monitor-Mutex locking - theory of operations
//
// * Native Monitors are completely unrelated to Java-level monitors,
// although the "back-end" slow-path implementations share a common lineage.
// See objectMonitor:: in synchronizer.cpp.
// Native Monitors do *not* support nesting or recursion but otherwise
// they're basically Hoare-flavor monitors.
//
// * A thread acquires ownership of a Monitor/Mutex by CASing the LockByte
// in the _LockWord from zero to non-zero. Note that the _Owner field
// is advisory and is used only to verify that the thread calling unlock()
// is indeed the last thread to have acquired the lock.
//
// * Contending threads "push" themselves onto the front of the contention
// queue -- called the cxq -- with CAS and then spin/park.
// The _LockWord contains the LockByte as well as the pointer to the head
// of the cxq. Colocating the LockByte with the cxq precludes certain races.
//
// * Using a separately addressable LockByte allows for CAS:MEMBAR or CAS:0
// idioms. We currently use MEMBAR in the uncontended unlock() path, as
// MEMBAR often has less latency than CAS. If warranted, we could switch to
// a CAS:0 mode, using timers to close the resultant race, as is done
// with Java Monitors in synchronizer.cpp.
//
// See the following for a discussion of the relative cost of atomics (CAS)
// MEMBAR, and ways to eliminate such instructions from the common-case paths:
// -- http://blogs.sun.com/dave/entry/biased_locking_in_hotspot
// -- http://blogs.sun.com/dave/resource/MustangSync.pdf
// -- http://blogs.sun.com/dave/resource/synchronization-public2.pdf
// -- synchronizer.cpp
//
// * Overall goals - desiderata
// 1. Minimize context switching
// 2. Minimize lock migration
// 3. Minimize CPI -- affinity and locality
// 4. Minimize the execution of high-latency instructions such as CAS or MEMBAR
// 5. Minimize outer lock hold times
// 6. Behave gracefully on a loaded system
//
// * Thread flow and list residency:
//
// Contention queue --> EntryList --> OnDeck --> Owner --> !Owner
// [..resident on monitor list..]
// [...........contending..................]
//
// -- The contention queue (cxq) contains recently-arrived threads (RATs).
// Threads on the cxq eventually drain into the EntryList.
// -- Invariant: a thread appears on at most one list -- cxq, EntryList
// or WaitSet -- at any one time.
// -- For a given monitor there can be at most one "OnDeck" thread at any
// given time but if needbe this particular invariant could be relaxed.
//
// * The WaitSet and EntryList linked lists are composed of ParkEvents.
// I use ParkEvent instead of threads as ParkEvents are immortal and
// type-stable, meaning we can safely unpark() a possibly stale
// list element in the unlock()-path. (That's benign).
//
// * Succession policy - providing for progress:
//
// As necessary, the unlock()ing thread identifies, unlinks, and unparks
// an "heir presumptive" tentative successor thread from the EntryList.
// This becomes the so-called "OnDeck" thread, of which there can be only
// one at any given time for a given monitor. The wakee will recontend
// for ownership of monitor.
//
// Succession is provided for by a policy of competitive handoff.
// The exiting thread does _not_ grant or pass ownership to the
// successor thread. (This is also referred to as "handoff" succession").
// Instead the exiting thread releases ownership and possibly wakes
// a successor, so the successor can (re)compete for ownership of the lock.
//
// Competitive handoff provides excellent overall throughput at the expense
// of short-term fairness. If fairness is a concern then one remedy might
// be to add an AcquireCounter field to the monitor. After a thread acquires
// the lock it will decrement the AcquireCounter field. When the count
// reaches 0 the thread would reset the AcquireCounter variable, abdicate
// the lock directly to some thread on the EntryList, and then move itself to the
// tail of the EntryList.
//
// But in practice most threads engage or otherwise participate in resource
// bounded producer-consumer relationships, so lock domination is not usually
// a practical concern. Recall too, that in general it's easier to construct
// a fair lock from a fast lock, but not vice-versa.
//
// * The cxq can have multiple concurrent "pushers" but only one concurrent
// detaching thread. This mechanism is immune from the ABA corruption.
// More precisely, the CAS-based "push" onto cxq is ABA-oblivious.
// We use OnDeck as a pseudo-lock to enforce the at-most-one detaching
// thread constraint.
//
// * Taken together, the cxq and the EntryList constitute or form a
// single logical queue of threads stalled trying to acquire the lock.
// We use two distinct lists to reduce heat on the list ends.
// Threads in lock() enqueue onto cxq while threads in unlock() will
// dequeue from the EntryList. (c.f. Michael Scott's "2Q" algorithm).
// A key desideratum is to minimize queue & monitor metadata manipulation
// that occurs while holding the "outer" monitor lock -- that is, we want to
// minimize monitor lock holds times.
//
// The EntryList is ordered by the prevailing queue discipline and
// can be organized in any convenient fashion, such as a doubly-linked list or
// a circular doubly-linked list. If we need a priority queue then something akin
// to Solaris' sleepq would work nicely. Viz.,
// -- http://agg.eng/ws/on10_nightly/source/usr/src/uts/common/os/sleepq.c.
// -- http://cvs.opensolaris.org/source/xref/onnv/onnv-gate/usr/src/uts/common/os/sleepq.c
// Queue discipline is enforced at ::unlock() time, when the unlocking thread
// drains the cxq into the EntryList, and orders or reorders the threads on the
// EntryList accordingly.
//
// Barring "lock barging", this mechanism provides fair cyclic ordering,
// somewhat similar to an elevator-scan.
//
// * OnDeck
// -- For a given monitor there can be at most one OnDeck thread at any given
// instant. The OnDeck thread is contending for the lock, but has been
// unlinked from the EntryList and cxq by some previous unlock() operations.
// Once a thread has been designated the OnDeck thread it will remain so
// until it manages to acquire the lock -- being OnDeck is a stable property.
// -- Threads on the EntryList or cxq are _not allowed to attempt lock acquisition.
// -- OnDeck also serves as an "inner lock" as follows. Threads in unlock() will, after
// having cleared the LockByte and dropped the outer lock, attempt to "trylock"
// OnDeck by CASing the field from null to non-null. If successful, that thread
// is then responsible for progress and succession and can use CAS to detach and
// drain the cxq into the EntryList. By convention, only this thread, the holder of
// the OnDeck inner lock, can manipulate the EntryList or detach and drain the
// RATs on the cxq into the EntryList. This avoids ABA corruption on the cxq as
// we allow multiple concurrent "push" operations but restrict detach concurrency
// to at most one thread. Having selected and detached a successor, the thread then
// changes the OnDeck to refer to that successor, and then unparks the successor.
// That successor will eventually acquire the lock and clear OnDeck. Beware
// that the OnDeck usage as a lock is asymmetric. A thread in unlock() transiently
// "acquires" OnDeck, performs queue manipulations, passes OnDeck to some successor,
// and then the successor eventually "drops" OnDeck. Note that there's never
// any sense of contention on the inner lock, however. Threads never contend
// or wait for the inner lock.
// -- OnDeck provides for futile wakeup throttling a described in section 3.3 of
// See http://www.usenix.org/events/jvm01/full_papers/dice/dice.pdf
// In a sense, OnDeck subsumes the ObjectMonitor _Succ and ObjectWaiter
// TState fields found in Java-level objectMonitors. (See synchronizer.cpp).
//
// * Waiting threads reside on the WaitSet list -- wait() puts
// the caller onto the WaitSet. Notify() or notifyAll() simply
// transfers threads from the WaitSet to either the EntryList or cxq.
// Subsequent unlock() operations will eventually unpark the notifyee.
// Unparking a notifee in notify() proper is inefficient - if we were to do so
// it's likely the notifyee would simply impale itself on the lock held
// by the notifier.
//
// * The mechanism is obstruction-free in that if the holder of the transient
// OnDeck lock in unlock() is preempted or otherwise stalls, other threads
// can still acquire and release the outer lock and continue to make progress.
// At worst, waking of already blocked contending threads may be delayed,
// but nothing worse. (We only use "trylock" operations on the inner OnDeck
// lock).
//
// * Note that thread-local storage must be initialized before a thread
// uses Native monitors or mutexes. The native monitor-mutex subsystem
// depends on Thread::current().
//
// * The monitor synchronization subsystem avoids the use of native
// synchronization primitives except for the narrow platform-specific
// park-unpark abstraction. See the comments in os_solaris.cpp regarding
// the semantics of park-unpark. Put another way, this monitor implementation
// depends only on atomic operations and park-unpark. The monitor subsystem
// manages all RUNNING->BLOCKED and BLOCKED->READY transitions while the
// underlying OS manages the READY<->RUN transitions.
//
// * The memory consistency model provide by lock()-unlock() is at least as
// strong or stronger than the Java Memory model defined by JSR-133.
// That is, we guarantee at least entry consistency, if not stronger.
// See http://g.oswego.edu/dl/jmm/cookbook.html.
//
// * Thread:: currently contains a set of purpose-specific ParkEvents:
// _MutexEvent, _ParkEvent, etc. A better approach might be to do away with
// the purpose-specific ParkEvents and instead implement a general per-thread
// stack of available ParkEvents which we could provision on-demand. The
// stack acts as a local cache to avoid excessive calls to ParkEvent::Allocate()
// and ::Release(). A thread would simply pop an element from the local stack before it
// enqueued or park()ed. When the contention was over the thread would
// push the no-longer-needed ParkEvent back onto its stack.
//
// * A slightly reduced form of ILock() and IUnlock() have been partially
// model-checked (Murphi) for safety and progress at T=1,2,3 and 4.
// It'd be interesting to see if TLA/TLC could be useful as well.
//
// * Mutex-Monitor is a low-level "leaf" subsystem. That is, the monitor
// code should never call other code in the JVM that might itself need to
// acquire monitors or mutexes. That's true *except* in the case of the
// ThreadBlockInVM state transition wrappers. The ThreadBlockInVM DTOR handles
// mutator reentry (ingress) by checking for a pending safepoint in which case it will
// call SafepointSynchronize::block(), which in turn may call Safepoint_lock->lock(), etc.
// In that particular case a call to lock() for a given Monitor can end up recursively
// calling lock() on another monitor. While distasteful, this is largely benign
// as the calls come from jacket that wraps lock(), and not from deep within lock() itself.
//
// It's unfortunate that native mutexes and thread state transitions were convolved.
// They're really separate concerns and should have remained that way. Melding
// them together was facile -- a bit too facile. The current implementation badly
// conflates the two concerns.
//
// * TODO-FIXME:
//
// -- Add DTRACE probes for contended acquire, contended acquired, contended unlock
// We should also add DTRACE probes in the ParkEvent subsystem for
// Park-entry, Park-exit, and Unpark.
//
// -- We have an excess of mutex-like constructs in the JVM, namely:
// 1. objectMonitors for Java-level synchronization (synchronizer.cpp)
// 2. low-level muxAcquire and muxRelease
// 3. low-level spinAcquire and spinRelease
// 4. native Mutex:: and Monitor::
// 5. jvm_raw_lock() and _unlock()
// 6. JVMTI raw monitors -- distinct from (5) despite having a confusingly
// similar name.
//
// o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o
// CASPTR() uses the canonical argument order that dominates in the literature.
// Our internal cmpxchg_ptr() uses a bastardized ordering to accommodate Sun .il templates.
#define CASPTR(a, c, s) \
intptr_t(Atomic::cmpxchg_ptr((void *)(s), (void *)(a), (void *)(c)))
#define UNS(x) (uintptr_t(x))
#define TRACE(m) \
{ \
static volatile int ctr = 0; \
int x = ++ctr; \
if ((x & (x - 1)) == 0) { \
::printf("%d:%s\n", x, #m); \
::fflush(stdout); \
} \
}
// Simplistic low-quality Marsaglia SHIFT-XOR RNG.
// Bijective except for the trailing mask operation.
// Useful for spin loops as the compiler can't optimize it away.
static inline jint MarsagliaXORV(jint x) {
if (x == 0) x = 1|os::random();
x ^= x << 6;
x ^= ((unsigned)x) >> 21;
x ^= x << 7;
return x & 0x7FFFFFFF;
}
static int Stall(int its) {
static volatile jint rv = 1;
volatile int OnFrame = 0;
jint v = rv ^ UNS(OnFrame);
while (--its >= 0) {
v = MarsagliaXORV(v);
}
// Make this impossible for the compiler to optimize away,
// but (mostly) avoid W coherency sharing on MP systems.
if (v == 0x12345) rv = v;
return v;
}
int Monitor::TryLock() {
intptr_t v = _LockWord.FullWord;
for (;;) {
if ((v & _LBIT) != 0) return 0;
const intptr_t u = CASPTR(&_LockWord, v, v|_LBIT);
if (v == u) return 1;
v = u;
}
}
int Monitor::TryFast() {
// Optimistic fast-path form ...
// Fast-path attempt for the common uncontended case.
// Avoid RTS->RTO $ coherence upgrade on typical SMP systems.
intptr_t v = CASPTR(&_LockWord, 0, _LBIT); // agro ...
if (v == 0) return 1;
for (;;) {
if ((v & _LBIT) != 0) return 0;
const intptr_t u = CASPTR(&_LockWord, v, v|_LBIT);
if (v == u) return 1;
v = u;
}
}
int Monitor::ILocked() {
const intptr_t w = _LockWord.FullWord & 0xFF;
assert(w == 0 || w == _LBIT, "invariant");
return w == _LBIT;
}
// Polite TATAS spinlock with exponential backoff - bounded spin.
// Ideally we'd use processor cycles, time or vtime to control
// the loop, but we currently use iterations.
// All the constants within were derived empirically but work over
// over the spectrum of J2SE reference platforms.
// On Niagara-class systems the back-off is unnecessary but
// is relatively harmless. (At worst it'll slightly retard
// acquisition times). The back-off is critical for older SMP systems
// where constant fetching of the LockWord would otherwise impair
// scalability.
//
// Clamp spinning at approximately 1/2 of a context-switch round-trip.
// See synchronizer.cpp for details and rationale.
int Monitor::TrySpin(Thread * const Self) {
if (TryLock()) return 1;
if (!os::is_MP()) return 0;
int Probes = 0;
int Delay = 0;
int Steps = 0;
int SpinMax = NativeMonitorSpinLimit;
int flgs = NativeMonitorFlags;
for (;;) {
intptr_t v = _LockWord.FullWord;
if ((v & _LBIT) == 0) {
if (CASPTR (&_LockWord, v, v|_LBIT) == v) {
return 1;
}
continue;
}
if ((flgs & 8) == 0) {
SpinPause();
}
// Periodically increase Delay -- variable Delay form
// conceptually: delay *= 1 + 1/Exponent
++Probes;
if (Probes > SpinMax) return 0;
if ((Probes & 0x7) == 0) {
Delay = ((Delay << 1)|1) & 0x7FF;
// CONSIDER: Delay += 1 + (Delay/4); Delay &= 0x7FF ;
}
if (flgs & 2) continue;
// Consider checking _owner's schedctl state, if OFFPROC abort spin.
// If the owner is OFFPROC then it's unlike that the lock will be dropped
// in a timely fashion, which suggests that spinning would not be fruitful
// or profitable.
// Stall for "Delay" time units - iterations in the current implementation.
// Avoid generating coherency traffic while stalled.
// Possible ways to delay:
// PAUSE, SLEEP, MEMBAR #sync, MEMBAR #halt,
// wr %g0,%asi, gethrtime, rdstick, rdtick, rdtsc, etc. ...
// Note that on Niagara-class systems we want to minimize STs in the
// spin loop. N1 and brethren write-around the L1$ over the xbar into the L2$.
// Furthermore, they don't have a W$ like traditional SPARC processors.
// We currently use a Marsaglia Shift-Xor RNG loop.
Steps += Delay;
if (Self != NULL) {
jint rv = Self->rng[0];
for (int k = Delay; --k >= 0;) {
rv = MarsagliaXORV(rv);
if ((flgs & 4) == 0 && SafepointSynchronize::do_call_back()) return 0;
}
Self->rng[0] = rv;
} else {
Stall(Delay);
}
}
}
static int ParkCommon(ParkEvent * ev, jlong timo) {
// Diagnostic support - periodically unwedge blocked threads
intx nmt = NativeMonitorTimeout;
if (nmt > 0 && (nmt < timo || timo <= 0)) {
timo = nmt;
}
int err = OS_OK;
if (0 == timo) {
ev->park();
} else {
err = ev->park(timo);
}
return err;
}
inline int Monitor::AcquireOrPush(ParkEvent * ESelf) {
intptr_t v = _LockWord.FullWord;
for (;;) {
if ((v & _LBIT) == 0) {
const intptr_t u = CASPTR(&_LockWord, v, v|_LBIT);
if (u == v) return 1; // indicate acquired
v = u;
} else {
// Anticipate success ...
ESelf->ListNext = (ParkEvent *)(v & ~_LBIT);
const intptr_t u = CASPTR(&_LockWord, v, intptr_t(ESelf)|_LBIT);
if (u == v) return 0; // indicate pushed onto cxq
v = u;
}
// Interference - LockWord change - just retry
}
}
// ILock and IWait are the lowest level primitive internal blocking
// synchronization functions. The callers of IWait and ILock must have
// performed any needed state transitions beforehand.
// IWait and ILock may directly call park() without any concern for thread state.
// Note that ILock and IWait do *not* access _owner.
// _owner is a higher-level logical concept.
void Monitor::ILock(Thread * Self) {
assert(_OnDeck != Self->_MutexEvent, "invariant");
if (TryFast()) {
Exeunt:
assert(ILocked(), "invariant");
return;
}
ParkEvent * const ESelf = Self->_MutexEvent;
assert(_OnDeck != ESelf, "invariant");
// As an optimization, spinners could conditionally try to set ONDECK to _LBIT
// Synchronizer.cpp uses a similar optimization.
if (TrySpin(Self)) goto Exeunt;
// Slow-path - the lock is contended.
// Either Enqueue Self on cxq or acquire the outer lock.
// LockWord encoding = (cxq,LOCKBYTE)
ESelf->reset();
OrderAccess::fence();
// Optional optimization ... try barging on the inner lock
if ((NativeMonitorFlags & 32) && CASPTR (&_OnDeck, NULL, UNS(Self)) == 0) {
goto OnDeck_LOOP;
}
if (AcquireOrPush(ESelf)) goto Exeunt;
// At any given time there is at most one ondeck thread.
// ondeck implies not resident on cxq and not resident on EntryList
// Only the OnDeck thread can try to acquire -- contended for -- the lock.
// CONSIDER: use Self->OnDeck instead of m->OnDeck.
// Deschedule Self so that others may run.
while (_OnDeck != ESelf) {
ParkCommon(ESelf, 0);
}
// Self is now in the ONDECK position and will remain so until it
// manages to acquire the lock.
OnDeck_LOOP:
for (;;) {
assert(_OnDeck == ESelf, "invariant");
if (TrySpin(Self)) break;
// It's probably wise to spin only if we *actually* blocked
// CONSIDER: check the lockbyte, if it remains set then
// preemptively drain the cxq into the EntryList.
// The best place and time to perform queue operations -- lock metadata --
// is _before having acquired the outer lock, while waiting for the lock to drop.
ParkCommon(ESelf, 0);
}
assert(_OnDeck == ESelf, "invariant");
_OnDeck = NULL;
// Note that we current drop the inner lock (clear OnDeck) in the slow-path
// epilogue immediately after having acquired the outer lock.
// But instead we could consider the following optimizations:
// A. Shift or defer dropping the inner lock until the subsequent IUnlock() operation.
// This might avoid potential reacquisition of the inner lock in IUlock().
// B. While still holding the inner lock, attempt to opportunistically select
// and unlink the next ONDECK thread from the EntryList.
// If successful, set ONDECK to refer to that thread, otherwise clear ONDECK.
// It's critical that the select-and-unlink operation run in constant-time as
// it executes when holding the outer lock and may artificially increase the
// effective length of the critical section.
// Note that (A) and (B) are tantamount to succession by direct handoff for
// the inner lock.
goto Exeunt;
}
void Monitor::IUnlock(bool RelaxAssert) {
assert(ILocked(), "invariant");
// Conceptually we need a MEMBAR #storestore|#loadstore barrier or fence immediately
// before the store that releases the lock. Crucially, all the stores and loads in the
// critical section must be globally visible before the store of 0 into the lock-word
// that releases the lock becomes globally visible. That is, memory accesses in the
// critical section should not be allowed to bypass or overtake the following ST that
// releases the lock. As such, to prevent accesses within the critical section
// from "leaking" out, we need a release fence between the critical section and the
// store that releases the lock. In practice that release barrier is elided on
// platforms with strong memory models such as TSO.
//
// Note that the OrderAccess::storeload() fence that appears after unlock store
// provides for progress conditions and succession and is _not related to exclusion
// safety or lock release consistency.
OrderAccess::release_store(&_LockWord.Bytes[_LSBINDEX], 0); // drop outer lock
OrderAccess::storeload();
ParkEvent * const w = _OnDeck;
assert(RelaxAssert || w != Thread::current()->_MutexEvent, "invariant");
if (w != NULL) {
// Either we have a valid ondeck thread or ondeck is transiently "locked"
// by some exiting thread as it arranges for succession. The LSBit of
// OnDeck allows us to discriminate two cases. If the latter, the
// responsibility for progress and succession lies with that other thread.
// For good performance, we also depend on the fact that redundant unpark()
// operations are cheap. That is, repeated Unpark()ing of the ONDECK thread
// is inexpensive. This approach provides implicit futile wakeup throttling.
// Note that the referent "w" might be stale with respect to the lock.
// In that case the following unpark() is harmless and the worst that'll happen
// is a spurious return from a park() operation. Critically, if "w" _is stale,
// then progress is known to have occurred as that means the thread associated
// with "w" acquired the lock. In that case this thread need take no further
// action to guarantee progress.
if ((UNS(w) & _LBIT) == 0) w->unpark();
return;
}
intptr_t cxq = _LockWord.FullWord;
if (((cxq & ~_LBIT)|UNS(_EntryList)) == 0) {
return; // normal fast-path exit - cxq and EntryList both empty
}
if (cxq & _LBIT) {
// Optional optimization ...
// Some other thread acquired the lock in the window since this
// thread released it. Succession is now that thread's responsibility.
return;
}
Succession:
// Slow-path exit - this thread must ensure succession and progress.
// OnDeck serves as lock to protect cxq and EntryList.
// Only the holder of OnDeck can manipulate EntryList or detach the RATs from cxq.
// Avoid ABA - allow multiple concurrent producers (enqueue via push-CAS)
// but only one concurrent consumer (detacher of RATs).
// Consider protecting this critical section with schedctl on Solaris.
// Unlike a normal lock, however, the exiting thread "locks" OnDeck,
// picks a successor and marks that thread as OnDeck. That successor
// thread will then clear OnDeck once it eventually acquires the outer lock.
if (CASPTR (&_OnDeck, NULL, _LBIT) != UNS(NULL)) {
return;
}
ParkEvent * List = _EntryList;
if (List != NULL) {
// Transfer the head of the EntryList to the OnDeck position.
// Once OnDeck, a thread stays OnDeck until it acquires the lock.
// For a given lock there is at most OnDeck thread at any one instant.
WakeOne:
assert(List == _EntryList, "invariant");
ParkEvent * const w = List;
assert(RelaxAssert || w != Thread::current()->_MutexEvent, "invariant");
_EntryList = w->ListNext;
// as a diagnostic measure consider setting w->_ListNext = BAD
assert(UNS(_OnDeck) == _LBIT, "invariant");
_OnDeck = w; // pass OnDeck to w.
// w will clear OnDeck once it acquires the outer lock
// Another optional optimization ...
// For heavily contended locks it's not uncommon that some other
// thread acquired the lock while this thread was arranging succession.
// Try to defer the unpark() operation - Delegate the responsibility
// for unpark()ing the OnDeck thread to the current or subsequent owners
// That is, the new owner is responsible for unparking the OnDeck thread.
OrderAccess::storeload();
cxq = _LockWord.FullWord;
if (cxq & _LBIT) return;
w->unpark();
return;
}
cxq = _LockWord.FullWord;
if ((cxq & ~_LBIT) != 0) {
// The EntryList is empty but the cxq is populated.
// drain RATs from cxq into EntryList
// Detach RATs segment with CAS and then merge into EntryList
for (;;) {
// optional optimization - if locked, the owner is responsible for succession
if (cxq & _LBIT) goto Punt;
const intptr_t vfy = CASPTR(&_LockWord, cxq, cxq & _LBIT);
if (vfy == cxq) break;
cxq = vfy;
// Interference - LockWord changed - Just retry
// We can see concurrent interference from contending threads
// pushing themselves onto the cxq or from lock-unlock operations.
// From the perspective of this thread, EntryList is stable and
// the cxq is prepend-only -- the head is volatile but the interior
// of the cxq is stable. In theory if we encounter interference from threads
// pushing onto cxq we could simply break off the original cxq suffix and
// move that segment to the EntryList, avoiding a 2nd or multiple CAS attempts
// on the high-traffic LockWord variable. For instance lets say the cxq is "ABCD"
// when we first fetch cxq above. Between the fetch -- where we observed "A"
// -- and CAS -- where we attempt to CAS null over A -- "PQR" arrive,
// yielding cxq = "PQRABCD". In this case we could simply set A.ListNext
// null, leaving cxq = "PQRA" and transfer the "BCD" segment to the EntryList.
// Note too, that it's safe for this thread to traverse the cxq
// without taking any special concurrency precautions.
}
// We don't currently reorder the cxq segment as we move it onto
// the EntryList, but it might make sense to reverse the order
// or perhaps sort by thread priority. See the comments in
// synchronizer.cpp objectMonitor::exit().
assert(_EntryList == NULL, "invariant");
_EntryList = List = (ParkEvent *)(cxq & ~_LBIT);
assert(List != NULL, "invariant");
goto WakeOne;
}
// cxq|EntryList is empty.
// w == NULL implies that cxq|EntryList == NULL in the past.
// Possible race - rare inopportune interleaving.
// A thread could have added itself to cxq since this thread previously checked.
// Detect and recover by refetching cxq.
Punt:
assert(UNS(_OnDeck) == _LBIT, "invariant");
_OnDeck = NULL; // Release inner lock.
OrderAccess::storeload(); // Dekker duality - pivot point
// Resample LockWord/cxq to recover from possible race.
// For instance, while this thread T1 held OnDeck, some other thread T2 might
// acquire the outer lock. Another thread T3 might try to acquire the outer
// lock, but encounter contention and enqueue itself on cxq. T2 then drops the
// outer lock, but skips succession as this thread T1 still holds OnDeck.
// T1 is and remains responsible for ensuring succession of T3.
//
// Note that we don't need to recheck EntryList, just cxq.
// If threads moved onto EntryList since we dropped OnDeck
// that implies some other thread forced succession.
cxq = _LockWord.FullWord;
if ((cxq & ~_LBIT) != 0 && (cxq & _LBIT) == 0) {
goto Succession; // potential race -- re-run succession
}
return;
}
bool Monitor::notify() {
assert(_owner == Thread::current(), "invariant");
assert(ILocked(), "invariant");
if (_WaitSet == NULL) return true;
NotifyCount++;
// Transfer one thread from the WaitSet to the EntryList or cxq.
// Currently we just unlink the head of the WaitSet and prepend to the cxq.
// And of course we could just unlink it and unpark it, too, but
// in that case it'd likely impale itself on the reentry.
Thread::muxAcquire(_WaitLock, "notify:WaitLock");
ParkEvent * nfy = _WaitSet;
if (nfy != NULL) { // DCL idiom
_WaitSet = nfy->ListNext;
assert(nfy->Notified == 0, "invariant");
// push nfy onto the cxq
for (;;) {
const intptr_t v = _LockWord.FullWord;
assert((v & 0xFF) == _LBIT, "invariant");
nfy->ListNext = (ParkEvent *)(v & ~_LBIT);
if (CASPTR (&_LockWord, v, UNS(nfy)|_LBIT) == v) break;
// interference - _LockWord changed -- just retry
}
// Note that setting Notified before pushing nfy onto the cxq is
// also legal and safe, but the safety properties are much more
// subtle, so for the sake of code stewardship ...
OrderAccess::fence();
nfy->Notified = 1;
}
Thread::muxRelease(_WaitLock);
if (nfy != NULL && (NativeMonitorFlags & 16)) {
// Experimental code ... light up the wakee in the hope that this thread (the owner)
// will drop the lock just about the time the wakee comes ONPROC.
nfy->unpark();
}
assert(ILocked(), "invariant");
return true;
}
// Currently notifyAll() transfers the waiters one-at-a-time from the waitset
// to the cxq. This could be done more efficiently with a single bulk en-mass transfer,
// but in practice notifyAll() for large #s of threads is rare and not time-critical.
// Beware too, that we invert the order of the waiters. Lets say that the
// waitset is "ABCD" and the cxq is "XYZ". After a notifyAll() the waitset
// will be empty and the cxq will be "DCBAXYZ". This is benign, of course.
bool Monitor::notify_all() {
assert(_owner == Thread::current(), "invariant");
assert(ILocked(), "invariant");
while (_WaitSet != NULL) notify();
return true;
}
int Monitor::IWait(Thread * Self, jlong timo) {
assert(ILocked(), "invariant");
// Phases:
// 1. Enqueue Self on WaitSet - currently prepend
// 2. unlock - drop the outer lock
// 3. wait for either notification or timeout
// 4. lock - reentry - reacquire the outer lock
ParkEvent * const ESelf = Self->_MutexEvent;
ESelf->Notified = 0;
ESelf->reset();
OrderAccess::fence();
// Add Self to WaitSet
// Ideally only the holder of the outer lock would manipulate the WaitSet -
// That is, the outer lock would implicitly protect the WaitSet.
// But if a thread in wait() encounters a timeout it will need to dequeue itself
// from the WaitSet _before it becomes the owner of the lock. We need to dequeue
// as the ParkEvent -- which serves as a proxy for the thread -- can't reside
// on both the WaitSet and the EntryList|cxq at the same time.. That is, a thread
// on the WaitSet can't be allowed to compete for the lock until it has managed to
// unlink its ParkEvent from WaitSet. Thus the need for WaitLock.
// Contention on the WaitLock is minimal.
//
// Another viable approach would be add another ParkEvent, "WaitEvent" to the
// thread class. The WaitSet would be composed of WaitEvents. Only the
// owner of the outer lock would manipulate the WaitSet. A thread in wait()
// could then compete for the outer lock, and then, if necessary, unlink itself
// from the WaitSet only after having acquired the outer lock. More precisely,
// there would be no WaitLock. A thread in in wait() would enqueue its WaitEvent
// on the WaitSet; release the outer lock; wait for either notification or timeout;
// reacquire the inner lock; and then, if needed, unlink itself from the WaitSet.
//
// Alternatively, a 2nd set of list link fields in the ParkEvent might suffice.
// One set would be for the WaitSet and one for the EntryList.
// We could also deconstruct the ParkEvent into a "pure" event and add a
// new immortal/TSM "ListElement" class that referred to ParkEvents.
// In that case we could have one ListElement on the WaitSet and another
// on the EntryList, with both referring to the same pure Event.
Thread::muxAcquire(_WaitLock, "wait:WaitLock:Add");
ESelf->ListNext = _WaitSet;
_WaitSet = ESelf;
Thread::muxRelease(_WaitLock);
// Release the outer lock
// We call IUnlock (RelaxAssert=true) as a thread T1 might
// enqueue itself on the WaitSet, call IUnlock(), drop the lock,
// and then stall before it can attempt to wake a successor.
// Some other thread T2 acquires the lock, and calls notify(), moving
// T1 from the WaitSet to the cxq. T2 then drops the lock. T1 resumes,
// and then finds *itself* on the cxq. During the course of a normal
// IUnlock() call a thread should _never find itself on the EntryList
// or cxq, but in the case of wait() it's possible.
// See synchronizer.cpp objectMonitor::wait().
IUnlock(true);
// Wait for either notification or timeout
// Beware that in some circumstances we might propagate
// spurious wakeups back to the caller.
for (;;) {
if (ESelf->Notified) break;
int err = ParkCommon(ESelf, timo);
if (err == OS_TIMEOUT || (NativeMonitorFlags & 1)) break;
}
// Prepare for reentry - if necessary, remove ESelf from WaitSet
// ESelf can be:
// 1. Still on the WaitSet. This can happen if we exited the loop by timeout.
// 2. On the cxq or EntryList
// 3. Not resident on cxq, EntryList or WaitSet, but in the OnDeck position.
OrderAccess::fence();
int WasOnWaitSet = 0;
if (ESelf->Notified == 0) {
Thread::muxAcquire(_WaitLock, "wait:WaitLock:remove");
if (ESelf->Notified == 0) { // DCL idiom
assert(_OnDeck != ESelf, "invariant"); // can't be both OnDeck and on WaitSet
// ESelf is resident on the WaitSet -- unlink it.
// A doubly-linked list would be better here so we can unlink in constant-time.
// We have to unlink before we potentially recontend as ESelf might otherwise
// end up on the cxq|EntryList -- it can't be on two lists at once.
ParkEvent * p = _WaitSet;
ParkEvent * q = NULL; // classic q chases p
while (p != NULL && p != ESelf) {
q = p;
p = p->ListNext;
}
assert(p == ESelf, "invariant");
if (p == _WaitSet) { // found at head
assert(q == NULL, "invariant");
_WaitSet = p->ListNext;
} else { // found in interior
assert(q->ListNext == p, "invariant");
q->ListNext = p->ListNext;
}
WasOnWaitSet = 1; // We were *not* notified but instead encountered timeout
}
Thread::muxRelease(_WaitLock);
}
// Reentry phase - reacquire the lock
if (WasOnWaitSet) {
// ESelf was previously on the WaitSet but we just unlinked it above
// because of a timeout. ESelf is not resident on any list and is not OnDeck
assert(_OnDeck != ESelf, "invariant");
ILock(Self);
} else {
// A prior notify() operation moved ESelf from the WaitSet to the cxq.
// ESelf is now on the cxq, EntryList or at the OnDeck position.
// The following fragment is extracted from Monitor::ILock()
for (;;) {
if (_OnDeck == ESelf && TrySpin(Self)) break;
ParkCommon(ESelf, 0);
}
assert(_OnDeck == ESelf, "invariant");
_OnDeck = NULL;
}
assert(ILocked(), "invariant");
return WasOnWaitSet != 0; // return true IFF timeout
}
// ON THE VMTHREAD SNEAKING PAST HELD LOCKS:
// In particular, there are certain types of global lock that may be held
// by a Java thread while it is blocked at a safepoint but before it has
// written the _owner field. These locks may be sneakily acquired by the
// VM thread during a safepoint to avoid deadlocks. Alternatively, one should
// identify all such locks, and ensure that Java threads never block at
// safepoints while holding them (_no_safepoint_check_flag). While it
// seems as though this could increase the time to reach a safepoint
// (or at least increase the mean, if not the variance), the latter
// approach might make for a cleaner, more maintainable JVM design.
//
// Sneaking is vile and reprehensible and should be excised at the 1st
// opportunity. It's possible that the need for sneaking could be obviated
// as follows. Currently, a thread might (a) while TBIVM, call pthread_mutex_lock
// or ILock() thus acquiring the "physical" lock underlying Monitor/Mutex.
// (b) stall at the TBIVM exit point as a safepoint is in effect. Critically,
// it'll stall at the TBIVM reentry state transition after having acquired the
// underlying lock, but before having set _owner and having entered the actual
// critical section. The lock-sneaking facility leverages that fact and allowed the
// VM thread to logically acquire locks that had already be physically locked by mutators
// but where mutators were known blocked by the reentry thread state transition.
//
// If we were to modify the Monitor-Mutex so that TBIVM state transitions tightly
// wrapped calls to park(), then we could likely do away with sneaking. We'd
// decouple lock acquisition and parking. The critical invariant to eliminating
// sneaking is to ensure that we never "physically" acquire the lock while TBIVM.
// An easy way to accomplish this is to wrap the park calls in a narrow TBIVM jacket.
// One difficulty with this approach is that the TBIVM wrapper could recurse and
// call lock() deep from within a lock() call, while the MutexEvent was already enqueued.
// Using a stack (N=2 at minimum) of ParkEvents would take care of that problem.
//
// But of course the proper ultimate approach is to avoid schemes that require explicit
// sneaking or dependence on any any clever invariants or subtle implementation properties
// of Mutex-Monitor and instead directly address the underlying design flaw.
void Monitor::lock(Thread * Self) {
// Ensure that the Monitor requires/allows safepoint checks.
assert(_safepoint_check_required != Monitor::_safepoint_check_never,
"This lock should never have a safepoint check: %s", name());
#ifdef CHECK_UNHANDLED_OOPS
// Clear unhandled oops so we get a crash right away. Only clear for non-vm
// or GC threads.
if (Self->is_Java_thread()) {
Self->clear_unhandled_oops();
}
#endif // CHECK_UNHANDLED_OOPS
debug_only(check_prelock_state(Self));
assert(_owner != Self, "invariant");
assert(_OnDeck != Self->_MutexEvent, "invariant");
if (TryFast()) {
Exeunt:
assert(ILocked(), "invariant");
assert(owner() == NULL, "invariant");
set_owner(Self);
return;
}
// The lock is contended ...
bool can_sneak = Self->is_VM_thread() && SafepointSynchronize::is_at_safepoint();
if (can_sneak && _owner == NULL) {
// a java thread has locked the lock but has not entered the
// critical region -- let's just pretend we've locked the lock
// and go on. we note this with _snuck so we can also
// pretend to unlock when the time comes.
_snuck = true;
goto Exeunt;
}
// Try a brief spin to avoid passing thru thread state transition ...
if (TrySpin(Self)) goto Exeunt;
check_block_state(Self);
if (Self->is_Java_thread()) {
// Horrible dictu - we suffer through a state transition
assert(rank() > Mutex::special, "Potential deadlock with special or lesser rank mutex");
ThreadBlockInVM tbivm((JavaThread *) Self);
ILock(Self);
} else {
// Mirabile dictu
ILock(Self);
}
goto Exeunt;
}
void Monitor::lock() {
this->lock(Thread::current());
}
// Lock without safepoint check - a degenerate variant of lock().
// Should ONLY be used by safepoint code and other code
// that is guaranteed not to block while running inside the VM. If this is called with
// thread state set to be in VM, the safepoint synchronization code will deadlock!
void Monitor::lock_without_safepoint_check(Thread * Self) {
// Ensure that the Monitor does not require or allow safepoint checks.
assert(_safepoint_check_required != Monitor::_safepoint_check_always,
"This lock should always have a safepoint check: %s", name());
assert(_owner != Self, "invariant");
ILock(Self);
assert(_owner == NULL, "invariant");
set_owner(Self);
}
void Monitor::lock_without_safepoint_check() {
lock_without_safepoint_check(Thread::current());
}
// Returns true if thread succeeds in grabbing the lock, otherwise false.
bool Monitor::try_lock() {
Thread * const Self = Thread::current();
debug_only(check_prelock_state(Self));
// assert(!thread->is_inside_signal_handler(), "don't lock inside signal handler");
// Special case, where all Java threads are stopped.
// The lock may have been acquired but _owner is not yet set.
// In that case the VM thread can safely grab the lock.
// It strikes me this should appear _after the TryLock() fails, below.
bool can_sneak = Self->is_VM_thread() && SafepointSynchronize::is_at_safepoint();
if (can_sneak && _owner == NULL) {
set_owner(Self); // Do not need to be atomic, since we are at a safepoint
_snuck = true;
return true;
}
if (TryLock()) {
// We got the lock
assert(_owner == NULL, "invariant");
set_owner(Self);
return true;
}
return false;
}
void Monitor::unlock() {
assert(_owner == Thread::current(), "invariant");
assert(_OnDeck != Thread::current()->_MutexEvent, "invariant");
set_owner(NULL);
if (_snuck) {
assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak");
_snuck = false;
return;
}
IUnlock(false);
}
// Yet another degenerate version of Monitor::lock() or lock_without_safepoint_check()
// jvm_raw_lock() and _unlock() can be called by non-Java threads via JVM_RawMonitorEnter.
//
// There's no expectation that JVM_RawMonitors will interoperate properly with the native
// Mutex-Monitor constructs. We happen to implement JVM_RawMonitors in terms of
// native Mutex-Monitors simply as a matter of convenience. A simple abstraction layer
// over a pthread_mutex_t would work equally as well, but require more platform-specific
// code -- a "PlatformMutex". Alternatively, a simply layer over muxAcquire-muxRelease
// would work too.
//
// Since the caller might be a foreign thread, we don't necessarily have a Thread.MutexEvent
// instance available. Instead, we transiently allocate a ParkEvent on-demand if
// we encounter contention. That ParkEvent remains associated with the thread
// until it manages to acquire the lock, at which time we return the ParkEvent
// to the global ParkEvent free list. This is correct and suffices for our purposes.
//
// Beware that the original jvm_raw_unlock() had a "_snuck" test but that
// jvm_raw_lock() didn't have the corresponding test. I suspect that's an
// oversight, but I've replicated the original suspect logic in the new code ...
void Monitor::jvm_raw_lock() {
assert(rank() == native, "invariant");
if (TryLock()) {
Exeunt:
assert(ILocked(), "invariant");
assert(_owner == NULL, "invariant");
// This can potentially be called by non-java Threads. Thus, the Thread::current_or_null()
// might return NULL. Don't call set_owner since it will break on an NULL owner
// Consider installing a non-null "ANON" distinguished value instead of just NULL.
_owner = Thread::current_or_null();
return;
}
if (TrySpin(NULL)) goto Exeunt;
// slow-path - apparent contention
// Allocate a ParkEvent for transient use.
// The ParkEvent remains associated with this thread until
// the time the thread manages to acquire the lock.
ParkEvent * const ESelf = ParkEvent::Allocate(NULL);
ESelf->reset();
OrderAccess::storeload();
// Either Enqueue Self on cxq or acquire the outer lock.
if (AcquireOrPush (ESelf)) {
ParkEvent::Release(ESelf); // surrender the ParkEvent
goto Exeunt;
}
// At any given time there is at most one ondeck thread.
// ondeck implies not resident on cxq and not resident on EntryList
// Only the OnDeck thread can try to acquire -- contended for -- the lock.
// CONSIDER: use Self->OnDeck instead of m->OnDeck.
for (;;) {
if (_OnDeck == ESelf && TrySpin(NULL)) break;
ParkCommon(ESelf, 0);
}
assert(_OnDeck == ESelf, "invariant");
_OnDeck = NULL;
ParkEvent::Release(ESelf); // surrender the ParkEvent
goto Exeunt;
}
void Monitor::jvm_raw_unlock() {
// Nearly the same as Monitor::unlock() ...
// directly set _owner instead of using set_owner(null)
_owner = NULL;
if (_snuck) { // ???
assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak");
_snuck = false;
return;
}
IUnlock(false);
}
bool Monitor::wait(bool no_safepoint_check, long timeout,
bool as_suspend_equivalent) {
// Make sure safepoint checking is used properly.
assert(!(_safepoint_check_required == Monitor::_safepoint_check_never && no_safepoint_check == false),
"This lock should never have a safepoint check: %s", name());
assert(!(_safepoint_check_required == Monitor::_safepoint_check_always && no_safepoint_check == true),
"This lock should always have a safepoint check: %s", name());
Thread * const Self = Thread::current();
assert(_owner == Self, "invariant");
assert(ILocked(), "invariant");
// as_suspend_equivalent logically implies !no_safepoint_check
guarantee(!as_suspend_equivalent || !no_safepoint_check, "invariant");
// !no_safepoint_check logically implies java_thread
guarantee(no_safepoint_check || Self->is_Java_thread(), "invariant");
#ifdef ASSERT
Monitor * least = get_least_ranked_lock_besides_this(Self->owned_locks());
assert(least != this, "Specification of get_least_... call above");
if (least != NULL && least->rank() <= special) {
tty->print("Attempting to wait on monitor %s/%d while holding"
" lock %s/%d -- possible deadlock",
name(), rank(), least->name(), least->rank());
assert(false, "Shouldn't block(wait) while holding a lock of rank special");
}
#endif // ASSERT
int wait_status;
// conceptually set the owner to NULL in anticipation of
// abdicating the lock in wait
set_owner(NULL);
if (no_safepoint_check) {
wait_status = IWait(Self, timeout);
} else {
assert(Self->is_Java_thread(), "invariant");
JavaThread *jt = (JavaThread *)Self;
// Enter safepoint region - ornate and Rococo ...
ThreadBlockInVM tbivm(jt);
OSThreadWaitState osts(Self->osthread(), false /* not Object.wait() */);
if (as_suspend_equivalent) {
jt->set_suspend_equivalent();
// cleared by handle_special_suspend_equivalent_condition() or
// java_suspend_self()
}
wait_status = IWait(Self, timeout);
// were we externally suspended while we were waiting?
if (as_suspend_equivalent && jt->handle_special_suspend_equivalent_condition()) {
// Our event wait has finished and we own the lock, but
// while we were waiting another thread suspended us. We don't
// want to hold the lock while suspended because that
// would surprise the thread that suspended us.
assert(ILocked(), "invariant");
IUnlock(true);
jt->java_suspend_self();
ILock(Self);
assert(ILocked(), "invariant");
}
}
// Conceptually reestablish ownership of the lock.
// The "real" lock -- the LockByte -- was reacquired by IWait().
assert(ILocked(), "invariant");
assert(_owner == NULL, "invariant");
set_owner(Self);
return wait_status != 0; // return true IFF timeout
}
Monitor::~Monitor() {
assert((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "");
}
void Monitor::ClearMonitor(Monitor * m, const char *name) {
m->_owner = NULL;
m->_snuck = false;
if (name == NULL) {
strcpy(m->_name, "UNKNOWN");
} else {
strncpy(m->_name, name, MONITOR_NAME_LEN - 1);
m->_name[MONITOR_NAME_LEN - 1] = '\0';
}
m->_LockWord.FullWord = 0;
m->_EntryList = NULL;
m->_OnDeck = NULL;
m->_WaitSet = NULL;
m->_WaitLock[0] = 0;
}
Monitor::Monitor() { ClearMonitor(this); }
Monitor::Monitor(int Rank, const char * name, bool allow_vm_block,
SafepointCheckRequired safepoint_check_required) {
ClearMonitor(this, name);
#ifdef ASSERT
_allow_vm_block = allow_vm_block;
_rank = Rank;
NOT_PRODUCT(_safepoint_check_required = safepoint_check_required;)
#endif
}
Mutex::~Mutex() {
assert((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "");
}
Mutex::Mutex(int Rank, const char * name, bool allow_vm_block,
SafepointCheckRequired safepoint_check_required) {
ClearMonitor((Monitor *) this, name);
#ifdef ASSERT
_allow_vm_block = allow_vm_block;
_rank = Rank;
NOT_PRODUCT(_safepoint_check_required = safepoint_check_required;)
#endif
}
bool Monitor::owned_by_self() const {
bool ret = _owner == Thread::current();
assert(!ret || _LockWord.Bytes[_LSBINDEX] != 0, "invariant");
return ret;
}
void Monitor::print_on_error(outputStream* st) const {
st->print("[" PTR_FORMAT, p2i(this));
st->print("] %s", _name);
st->print(" - owner thread: " PTR_FORMAT, p2i(_owner));
}
// ----------------------------------------------------------------------------------
// Non-product code
#ifndef PRODUCT
void Monitor::print_on(outputStream* st) const {
st->print_cr("Mutex: [" PTR_FORMAT "/" PTR_FORMAT "] %s - owner: " PTR_FORMAT,
p2i(this), _LockWord.FullWord, _name, p2i(_owner));
}
#endif
#ifndef PRODUCT
#ifdef ASSERT
Monitor * Monitor::get_least_ranked_lock(Monitor * locks) {
Monitor *res, *tmp;
for (res = tmp = locks; tmp != NULL; tmp = tmp->next()) {
if (tmp->rank() < res->rank()) {
res = tmp;
}
}
if (!SafepointSynchronize::is_at_safepoint()) {
// In this case, we expect the held locks to be
// in increasing rank order (modulo any native ranks)
for (tmp = locks; tmp != NULL; tmp = tmp->next()) {
if (tmp->next() != NULL) {
assert(tmp->rank() == Mutex::native ||
tmp->rank() <= tmp->next()->rank(), "mutex rank anomaly?");
}
}
}
return res;
}
Monitor* Monitor::get_least_ranked_lock_besides_this(Monitor* locks) {
Monitor *res, *tmp;
for (res = NULL, tmp = locks; tmp != NULL; tmp = tmp->next()) {
if (tmp != this && (res == NULL || tmp->rank() < res->rank())) {
res = tmp;
}
}
if (!SafepointSynchronize::is_at_safepoint()) {
// In this case, we expect the held locks to be
// in increasing rank order (modulo any native ranks)
for (tmp = locks; tmp != NULL; tmp = tmp->next()) {
if (tmp->next() != NULL) {
assert(tmp->rank() == Mutex::native ||
tmp->rank() <= tmp->next()->rank(), "mutex rank anomaly?");
}
}
}
return res;
}
bool Monitor::contains(Monitor* locks, Monitor * lock) {
for (; locks != NULL; locks = locks->next()) {
if (locks == lock) {
return true;
}
}
return false;
}
#endif
// Called immediately after lock acquisition or release as a diagnostic
// to track the lock-set of the thread and test for rank violations that
// might indicate exposure to deadlock.
// Rather like an EventListener for _owner (:>).
void Monitor::set_owner_implementation(Thread *new_owner) {
// This function is solely responsible for maintaining
// and checking the invariant that threads and locks
// are in a 1/N relation, with some some locks unowned.
// It uses the Mutex::_owner, Mutex::_next, and
// Thread::_owned_locks fields, and no other function
// changes those fields.
// It is illegal to set the mutex from one non-NULL
// owner to another--it must be owned by NULL as an
// intermediate state.
if (new_owner != NULL) {
// the thread is acquiring this lock
assert(new_owner == Thread::current(), "Should I be doing this?");
assert(_owner == NULL, "setting the owner thread of an already owned mutex");
_owner = new_owner; // set the owner
// link "this" into the owned locks list
#ifdef ASSERT // Thread::_owned_locks is under the same ifdef
Monitor* locks = get_least_ranked_lock(new_owner->owned_locks());
// Mutex::set_owner_implementation is a friend of Thread
assert(this->rank() >= 0, "bad lock rank");
// Deadlock avoidance rules require us to acquire Mutexes only in
// a global total order. For example m1 is the lowest ranked mutex
// that the thread holds and m2 is the mutex the thread is trying
// to acquire, then deadlock avoidance rules require that the rank
// of m2 be less than the rank of m1.
// The rank Mutex::native is an exception in that it is not subject
// to the verification rules.
// Here are some further notes relating to mutex acquisition anomalies:
// . under Solaris, the interrupt lock gets acquired when doing
// profiling, so any lock could be held.
// . it is also ok to acquire Safepoint_lock at the very end while we
// already hold Terminator_lock - may happen because of periodic safepoints
if (this->rank() != Mutex::native &&
this->rank() != Mutex::suspend_resume &&
locks != NULL && locks->rank() <= this->rank() &&
!SafepointSynchronize::is_at_safepoint() &&
this != Interrupt_lock && this != ProfileVM_lock &&
!(this == Safepoint_lock && contains(locks, Terminator_lock) &&
SafepointSynchronize::is_synchronizing())) {
new_owner->print_owned_locks();
fatal("acquiring lock %s/%d out of order with lock %s/%d -- "
"possible deadlock", this->name(), this->rank(),
locks->name(), locks->rank());
}
this->_next = new_owner->_owned_locks;
new_owner->_owned_locks = this;
#endif
} else {
// the thread is releasing this lock
Thread* old_owner = _owner;
debug_only(_last_owner = old_owner);
assert(old_owner != NULL, "removing the owner thread of an unowned mutex");
assert(old_owner == Thread::current(), "removing the owner thread of an unowned mutex");
_owner = NULL; // set the owner
#ifdef ASSERT
Monitor *locks = old_owner->owned_locks();
// remove "this" from the owned locks list
Monitor *prev = NULL;
bool found = false;
for (; locks != NULL; prev = locks, locks = locks->next()) {
if (locks == this) {
found = true;
break;
}
}
assert(found, "Removing a lock not owned");
if (prev == NULL) {
old_owner->_owned_locks = _next;
} else {
prev->_next = _next;
}
_next = NULL;
#endif
}
}
// Factored out common sanity checks for locking mutex'es. Used by lock() and try_lock()
void Monitor::check_prelock_state(Thread *thread) {
assert((!thread->is_Java_thread() || ((JavaThread *)thread)->thread_state() == _thread_in_vm)
|| rank() == Mutex::special, "wrong thread state for using locks");
if (StrictSafepointChecks) {
if (thread->is_VM_thread() && !allow_vm_block()) {
fatal("VM thread using lock %s (not allowed to block on)", name());
}
debug_only(if (rank() != Mutex::special) \
thread->check_for_valid_safepoint_state(false);)
}
if (thread->is_Watcher_thread()) {
assert(!WatcherThread::watcher_thread()->has_crash_protection(),
"locking not allowed when crash protection is set");
}
}
void Monitor::check_block_state(Thread *thread) {
if (!_allow_vm_block && thread->is_VM_thread()) {
warning("VM thread blocked on lock");
print();
BREAKPOINT;
}
assert(_owner != thread, "deadlock: blocking on monitor owned by current thread");
}
#endif // PRODUCT