view hotspot/src/share/vm/runtime/mutex.cpp @ 26683:a02753d5a0b2

8057107: cleanup indent white space issues prior to Contended Locking reorder and cache line bucket Reviewed-by: fparain, sspitsyn, coleenp
author dcubed
date Wed, 10 Sep 2014 11:48:20 -0600
parents 381638db28e6
children d1221849ea3d
line wrap: on
line source
 * Copyright (c) 1998, 2014, Oracle and/or its affiliates. All rights reserved.
 * This code is free software; you can redistribute it and/or modify it
 * under the terms of the GNU General Public License version 2 only, as
 * published by the Free Software Foundation.
 * This code is distributed in the hope that it will be useful, but WITHOUT
 * ANY WARRANTY; without even the implied warranty of MERCHANTABILITY or
 * FITNESS FOR A PARTICULAR PURPOSE.  See the GNU General Public License
 * version 2 for more details (a copy is included in the LICENSE file that
 * accompanied this code).
 * You should have received a copy of the GNU General Public License version
 * 2 along with this work; if not, write to the Free Software Foundation,
 * Inc., 51 Franklin St, Fifth Floor, Boston, MA 02110-1301 USA.
 * Please contact Oracle, 500 Oracle Parkway, Redwood Shores, CA 94065 USA
 * or visit if you need additional information or have any
 * questions.

#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"
#ifdef TARGET_OS_FAMILY_solaris
# include "mutex_solaris.inline.hpp"
#ifdef TARGET_OS_FAMILY_windows
# include "mutex_windows.inline.hpp"
# include "mutex_bsd.inline.hpp"


// 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:
//   --
//   --
//   --
//   -- 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.
//   --
//   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
//       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
// * 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.
//   -- 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;

    if ((flgs & 8) == 0) {

    // Periodically increase Delay -- variable Delay form
    // conceptually: delay *= 1 + 1/Exponent
    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 {

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) {
  } 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()) {
    assert(ILocked(), "invariant");

  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)

  // 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.
  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

  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();

  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.

  // 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)) {

  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.
    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.
    cxq = _LockWord.FullWord;
    if (cxq & _LBIT) 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.
  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

bool Monitor::notify() {
  assert(_owner == Thread::current(), "invariant");
  assert(ILocked(), "invariant");
  if (_WaitSet == NULL) return true;

  // 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 ...
    nfy->Notified = 1;
  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.
  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;

  // 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;

  // 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().

  // 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.

  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

  // 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");
  } 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

// 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) {
  // Clear unhandled oops so we get a crash right away.  Only clear for non-vm
  // or GC threads.
  if (Self->is_Java_thread()) {

  assert(_owner != Self              , "invariant");
  assert(_OnDeck != Self->_MutexEvent, "invariant");

  if (TryFast()) {
    assert(ILocked(), "invariant");
    assert(owner() == NULL, "invariant");

  // 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;

  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);
  } else {
    // Mirabile dictu
  goto Exeunt;

void Monitor::lock() {

// Lock without safepoint check - a degenerate variant of lock().
// Should ONLY be used by safepoint code and other code
// that is guaranteed not to block while running inside the VM. If this is called with
// thread state set to be in VM, the safepoint synchronization code will deadlock!

void Monitor::lock_without_safepoint_check (Thread * Self) {
  assert(_owner != Self, "invariant");
  assert(_owner == NULL, "invariant");

void Monitor::lock_without_safepoint_check() {

// Returns true if thread succeeds in grabbing the lock, otherwise false.

bool Monitor::try_lock() {
  Thread * const Self = Thread::current();
  // 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");
    return true;
  return false;

void Monitor::unlock() {
  assert(_owner  == Thread::current(), "invariant");
  assert(_OnDeck != Thread::current()->_MutexEvent , "invariant");
  if (_snuck) {
    assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak");
    _snuck = 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()) {
    assert(ILocked(), "invariant");
    assert(_owner == NULL, "invariant");
    // This can potentially be called by non-java Threads. Thus, the ThreadLocalStorage
    // might return NULL. Don't call set_owner since it will break on an NULL owner
    // Consider installing a non-null "ANON" distinguished value instead of just NULL.
    _owner = ThreadLocalStorage::thread();

  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);

  // 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;

bool Monitor::wait(bool no_safepoint_check, long timeout, bool as_suspend_equivalent) {
  Thread * const Self = Thread::current();
  assert(_owner == Self, "invariant");
  assert(ILocked(), "invariant");

  // as_suspend_equivalent logically implies !no_safepoint_check
  guarantee(!as_suspend_equivalent || !no_safepoint_check, "invariant");
  // !no_safepoint_check logically implies java_thread
  guarantee(no_safepoint_check || Self->is_Java_thread(), "invariant");

  #ifdef ASSERT
  Monitor * least = get_least_ranked_lock_besides_this(Self->owned_locks());
  assert(least != this, "Specification of get_least_... call above");
  if (least != NULL && least->rank() <= special) {
    tty->print("Attempting to wait on monitor %s/%d while holding"
               " lock %s/%d -- possible deadlock",
               name(), rank(), least->name(), least->rank());
    assert(false, "Shouldn't block(wait) while holding a lock of rank special");
  #endif // ASSERT

  int wait_status;
  // conceptually set the owner to NULL in anticipation of
  // abdicating the lock in wait
  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) {
      // 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");
      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");
  return wait_status != 0;          // return true IFF timeout

Monitor::~Monitor() {
  assert((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "");

void Monitor::ClearMonitor (Monitor * m, const char *name) {
  m->_owner             = NULL;
  m->_snuck             = false;
  if (name == NULL) {
    strcpy(m->_name, "UNKNOWN");
  } else {
    strncpy(m->_name, name, MONITOR_NAME_LEN - 1);
    m->_name[MONITOR_NAME_LEN - 1] = '\0';
  m->_LockWord.FullWord = 0;
  m->_EntryList         = NULL;
  m->_OnDeck            = NULL;
  m->_WaitSet           = NULL;
  m->_WaitLock[0]       = 0;

Monitor::Monitor() { ClearMonitor(this); }

Monitor::Monitor (int Rank, const char * name, bool allow_vm_block) {
  ClearMonitor(this, name);
#ifdef ASSERT
  _allow_vm_block  = allow_vm_block;
  _rank            = Rank;

Mutex::~Mutex() {
  assert((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "");

Mutex::Mutex (int Rank, const char * name, bool allow_vm_block) {
  ClearMonitor((Monitor *) this, name);
#ifdef ASSERT
  _allow_vm_block   = allow_vm_block;
  _rank             = Rank;

bool Monitor::owned_by_self() const {
  bool ret = _owner == Thread::current();
  assert(!ret || _LockWord.Bytes[_LSBINDEX] != 0, "invariant");
  return ret;

void Monitor::print_on_error(outputStream* st) const {
  st->print("[" PTR_FORMAT, this);
  st->print("] %s", _name);
  st->print(" - owner thread: " PTR_FORMAT, _owner);

// ----------------------------------------------------------------------------------
// Non-product code

#ifndef PRODUCT
void Monitor::print_on(outputStream* st) const {
  st->print_cr("Mutex: [0x%lx/0x%lx] %s - owner: 0x%lx", this, _LockWord.FullWord, _name, _owner);

#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;

// 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())) {
      fatal(err_msg("acquiring lock %s/%d out of order with lock %s/%d -- "
                    "possible deadlock", this->name(), this->rank(),
                    locks->name(), locks->rank()));

    this->_next = new_owner->_owned_locks;
    new_owner->_owned_locks = this;

  } 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;
    assert(found, "Removing a lock not owned");
    if (prev == NULL) {
      old_owner->_owned_locks = _next;
    } else {
      prev->_next = _next;
    _next = NULL;

// Factored out common sanity checks for locking mutex'es. Used by lock() and try_lock()
void Monitor::check_prelock_state(Thread *thread) {
  assert((!thread->is_Java_thread() || ((JavaThread *)thread)->thread_state() == _thread_in_vm)
         || rank() == Mutex::special, "wrong thread state for using locks");
  if (StrictSafepointChecks) {
    if (thread->is_VM_thread() && !allow_vm_block()) {
      fatal(err_msg("VM thread using lock %s (not allowed to block on)",
    debug_only(if (rank() != Mutex::special) \
  if (thread->is_Watcher_thread()) {
           "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");
  assert(_owner != thread, "deadlock: blocking on monitor owned by current thread");

#endif // PRODUCT