jdk8-mips64-public/hotspot: src/share/vm/runtime/synchronizer.cpp@0fbdb4381b99 (annotated)

src/share/vm/runtime/synchronizer.cpp@0fbdb4381b99 (annotated)

src/share/vm/runtime/synchronizer.cpp

Mon, 09 Mar 2009 13:28:46 -0700

author: xdono
date: Mon, 09 Mar 2009 13:28:46 -0700
changeset 1014: 0fbdb4381b99
parent 948: 2328d1d3f8cf
child 1137: b9fba36710f2
permissions: -rw-r--r--

6814575: Update copyright year
Summary: Update copyright for files that have been modified in 2009, up to 03/09
Reviewed-by: katleman, tbell, ohair

 /*
  * Copyright 1998-2009 Sun Microsystems, Inc.  All Rights Reserved.
  * DO NOT ALTER OR REMOVE COPYRIGHT NOTICES OR THIS FILE HEADER.
  *
  * This code is free software; you can redistribute it and/or modify it
  * under the terms of the GNU General Public License version 2 only, as
  * published by the Free Software Foundation.
  *
  * This code is distributed in the hope that it will be useful, but WITHOUT
  * ANY WARRANTY; without even the implied warranty of MERCHANTABILITY or
  * FITNESS FOR A PARTICULAR PURPOSE.  See the GNU General Public License
  * version 2 for more details (a copy is included in the LICENSE file that
  * accompanied this code).
  *
  * You should have received a copy of the GNU General Public License version
  * 2 along with this work; if not, write to the Free Software Foundation,
  * Inc., 51 Franklin St, Fifth Floor, Boston, MA 02110-1301 USA.
  *
  * Please contact Sun Microsystems, Inc., 4150 Network Circle, Santa Clara,
  * CA 95054 USA or visit www.sun.com if you need additional information or
  * have any questions.
  *
  */
 # include "incls/_precompiled.incl"
 # include "incls/_synchronizer.cpp.incl"
 #if defined(__GNUC__) && !defined(IA64)
   // Need to inhibit inlining for older versions of GCC to avoid build-time failures
   #define ATTR __attribute__((noinline))
 #else
   #define ATTR
 #endif
 // Native markword accessors for synchronization and hashCode().
 //
 // The "core" versions of monitor enter and exit reside in this file.
 // The interpreter and compilers contain specialized transliterated
 // variants of the enter-exit fast-path operations.  See i486.ad fast_lock(),
 // for instance.  If you make changes here, make sure to modify the
 // interpreter, and both C1 and C2 fast-path inline locking code emission.
 //
 // TODO: merge the objectMonitor and synchronizer classes.
 //
 // -----------------------------------------------------------------------------
 #ifdef DTRACE_ENABLED
 // Only bother with this argument setup if dtrace is available
 // TODO-FIXME: probes should not fire when caller is _blocked.  assert() accordingly.
 HS_DTRACE_PROBE_DECL5(hotspot, monitor__wait,
   jlong, uintptr_t, char*, int, long);
 HS_DTRACE_PROBE_DECL4(hotspot, monitor__waited,
   jlong, uintptr_t, char*, int);
 HS_DTRACE_PROBE_DECL4(hotspot, monitor__notify,
   jlong, uintptr_t, char*, int);
 HS_DTRACE_PROBE_DECL4(hotspot, monitor__notifyAll,
   jlong, uintptr_t, char*, int);
 HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__enter,
   jlong, uintptr_t, char*, int);
 HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__entered,
   jlong, uintptr_t, char*, int);
 HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__exit,
   jlong, uintptr_t, char*, int);
 #define DTRACE_MONITOR_PROBE_COMMON(klassOop, thread)                      \
   char* bytes = NULL;                                                      \
   int len = 0;                                                             \
   jlong jtid = SharedRuntime::get_java_tid(thread);                        \
   symbolOop klassname = ((oop)(klassOop))->klass()->klass_part()->name();  \
   if (klassname != NULL) {                                                 \
     bytes = (char*)klassname->bytes();                                     \
     len = klassname->utf8_length();                                        \
   }
 #define DTRACE_MONITOR_WAIT_PROBE(monitor, klassOop, thread, millis)       \
   {                                                                        \
     if (DTraceMonitorProbes) {                                            \
       DTRACE_MONITOR_PROBE_COMMON(klassOop, thread);                       \
       HS_DTRACE_PROBE5(hotspot, monitor__wait, jtid,                       \
                        (monitor), bytes, len, (millis));                   \
     }                                                                      \
   }
 #define DTRACE_MONITOR_PROBE(probe, monitor, klassOop, thread)             \
   {                                                                        \
     if (DTraceMonitorProbes) {                                            \
       DTRACE_MONITOR_PROBE_COMMON(klassOop, thread);                       \
       HS_DTRACE_PROBE4(hotspot, monitor__##probe, jtid,                    \
                        (uintptr_t)(monitor), bytes, len);                  \
     }                                                                      \
   }
 #else //  ndef DTRACE_ENABLED
 #define DTRACE_MONITOR_WAIT_PROBE(klassOop, thread, millis, mon)    {;}
 #define DTRACE_MONITOR_PROBE(probe, klassOop, thread, mon)          {;}
 #endif // ndef DTRACE_ENABLED
 // ObjectWaiter serves as a "proxy" or surrogate thread.
 // TODO-FIXME: Eliminate ObjectWaiter and use the thread-specific
 // ParkEvent instead.  Beware, however, that the JVMTI code
 // knows about ObjectWaiters, so we'll have to reconcile that code.
 // See next_waiter(), first_waiter(), etc.
 class ObjectWaiter : public StackObj {
  public:
   enum TStates { TS_UNDEF, TS_READY, TS_RUN, TS_WAIT, TS_ENTER, TS_CXQ } ;
   enum Sorted  { PREPEND, APPEND, SORTED } ;
   ObjectWaiter * volatile _next;
   ObjectWaiter * volatile _prev;
   Thread*       _thread;
   ParkEvent *   _event;
   volatile int  _notified ;
   volatile TStates TState ;
   Sorted        _Sorted ;           // List placement disposition
   bool          _active ;           // Contention monitoring is enabled
  public:
   ObjectWaiter(Thread* thread) {
     _next     = NULL;
     _prev     = NULL;
     _notified = 0;
     TState    = TS_RUN ;
     _thread   = thread;
     _event    = thread->_ParkEvent ;
     _active   = false;
     assert (_event != NULL, "invariant") ;
   }
   void wait_reenter_begin(ObjectMonitor *mon) {
     JavaThread *jt = (JavaThread *)this->_thread;
     _active = JavaThreadBlockedOnMonitorEnterState::wait_reenter_begin(jt, mon);
   }
   void wait_reenter_end(ObjectMonitor *mon) {
     JavaThread *jt = (JavaThread *)this->_thread;
     JavaThreadBlockedOnMonitorEnterState::wait_reenter_end(jt, _active);
   }
 };
 enum ManifestConstants {
     ClearResponsibleAtSTW   = 0,
     MaximumRecheckInterval  = 1000
 } ;
 #undef TEVENT
 #define TEVENT(nom) {if (SyncVerbose) FEVENT(nom); }
 #define FEVENT(nom) { static volatile int ctr = 0 ; int v = ++ctr ; if ((v & (v-1)) == 0) { ::printf (#nom " : %d \n", v); ::fflush(stdout); }}
 #undef  TEVENT
 #define TEVENT(nom) {;}
 // Performance concern:
 // OrderAccess::storestore() calls release() which STs 0 into the global volatile
 // OrderAccess::Dummy variable.  This store is unnecessary for correctness.
 // Many threads STing into a common location causes considerable cache migration
 // or "sloshing" on large SMP system.  As such, I avoid using OrderAccess::storestore()
 // until it's repaired.  In some cases OrderAccess::fence() -- which incurs local
 // latency on the executing processor -- is a better choice as it scales on SMP
 // systems.  See http://blogs.sun.com/dave/entry/biased_locking_in_hotspot for a
 // discussion of coherency costs.  Note that all our current reference platforms
 // provide strong ST-ST order, so the issue is moot on IA32, x64, and SPARC.
 //
 // As a general policy we use "volatile" to control compiler-based reordering
 // and explicit fences (barriers) to control for architectural reordering performed
 // by the CPU(s) or platform.
 static int  MBFence (int x) { OrderAccess::fence(); return x; }
 struct SharedGlobals {
     // These are highly shared mostly-read variables.
     // To avoid false-sharing they need to be the sole occupants of a $ line.
     double padPrefix [8];
     volatile int stwRandom ;
     volatile int stwCycle ;
     // Hot RW variables -- Sequester to avoid false-sharing
     double padSuffix [16];
     volatile int hcSequence ;
     double padFinal [8] ;
 } ;
 static SharedGlobals GVars ;
 // Tunables ...
 // The knob* variables are effectively final.  Once set they should
 // never be modified hence.  Consider using __read_mostly with GCC.
 static int Knob_LogSpins           = 0 ;       // enable jvmstat tally for spins
 static int Knob_HandOff            = 0 ;
 static int Knob_Verbose            = 0 ;
 static int Knob_ReportSettings     = 0 ;
 static int Knob_SpinLimit          = 5000 ;    // derived by an external tool -
 static int Knob_SpinBase           = 0 ;       // Floor AKA SpinMin
 static int Knob_SpinBackOff        = 0 ;       // spin-loop backoff
 static int Knob_CASPenalty         = -1 ;      // Penalty for failed CAS
 static int Knob_OXPenalty          = -1 ;      // Penalty for observed _owner change
 static int Knob_SpinSetSucc        = 1 ;       // spinners set the _succ field
 static int Knob_SpinEarly          = 1 ;
 static int Knob_SuccEnabled        = 1 ;       // futile wake throttling
 static int Knob_SuccRestrict       = 0 ;       // Limit successors + spinners to at-most-one
 static int Knob_MaxSpinners        = -1 ;      // Should be a function of # CPUs
 static int Knob_Bonus              = 100 ;     // spin success bonus
 static int Knob_BonusB             = 100 ;     // spin success bonus
 static int Knob_Penalty            = 200 ;     // spin failure penalty
 static int Knob_Poverty            = 1000 ;
 static int Knob_SpinAfterFutile    = 1 ;       // Spin after returning from park()
 static int Knob_FixedSpin          = 0 ;
 static int Knob_OState             = 3 ;       // Spinner checks thread state of _owner
 static int Knob_UsePause           = 1 ;
 static int Knob_ExitPolicy         = 0 ;
 static int Knob_PreSpin            = 10 ;      // 20-100 likely better
 static int Knob_ResetEvent         = 0 ;
 static int BackOffMask             = 0 ;
 static int Knob_FastHSSEC          = 0 ;
 static int Knob_MoveNotifyee       = 2 ;       // notify() - disposition of notifyee
 static int Knob_QMode              = 0 ;       // EntryList-cxq policy - queue discipline
 static volatile int InitDone       = 0 ;
 // hashCode() generation :
 //
 // Possibilities:
 // * MD5Digest of {obj,stwRandom}
 // * CRC32 of {obj,stwRandom} or any linear-feedback shift register function.
 // * A DES- or AES-style SBox[] mechanism
 // * One of the Phi-based schemes, such as:
 //   2654435761 = 2^32 * Phi (golden ratio)
 //   HashCodeValue = ((uintptr_t(obj) >> 3) * 2654435761) ^ GVars.stwRandom ;
 // * A variation of Marsaglia's shift-xor RNG scheme.
 // * (obj ^ stwRandom) is appealing, but can result
 //   in undesirable regularity in the hashCode values of adjacent objects
 //   (objects allocated back-to-back, in particular).  This could potentially
 //   result in hashtable collisions and reduced hashtable efficiency.
 //   There are simple ways to "diffuse" the middle address bits over the
 //   generated hashCode values:
 //
 static inline intptr_t get_next_hash(Thread * Self, oop obj) {
   intptr_t value = 0 ;
   if (hashCode == 0) {
      // This form uses an unguarded global Park-Miller RNG,
      // so it's possible for two threads to race and generate the same RNG.
      // On MP system we'll have lots of RW access to a global, so the
      // mechanism induces lots of coherency traffic.
      value = os::random() ;
   } else
   if (hashCode == 1) {
      // This variation has the property of being stable (idempotent)
      // between STW operations.  This can be useful in some of the 1-0
      // synchronization schemes.
      intptr_t addrBits = intptr_t(obj) >> 3 ;
      value = addrBits ^ (addrBits >> 5) ^ GVars.stwRandom ;
   } else
   if (hashCode == 2) {
      value = 1 ;            // for sensitivity testing
   } else
   if (hashCode == 3) {
      value = ++GVars.hcSequence ;
   } else
   if (hashCode == 4) {
      value = intptr_t(obj) ;
   } else {
      // Marsaglia's xor-shift scheme with thread-specific state
      // This is probably the best overall implementation -- we'll
      // likely make this the default in future releases.
      unsigned t = Self->_hashStateX ;
      t ^= (t << 11) ;
      Self->_hashStateX = Self->_hashStateY ;
      Self->_hashStateY = Self->_hashStateZ ;
      Self->_hashStateZ = Self->_hashStateW ;
      unsigned v = Self->_hashStateW ;
      v = (v ^ (v >> 19)) ^ (t ^ (t >> 8)) ;
      Self->_hashStateW = v ;
      value = v ;
   }
   value &= markOopDesc::hash_mask;
   if (value == 0) value = 0xBAD ;
   assert (value != markOopDesc::no_hash, "invariant") ;
   TEVENT (hashCode: GENERATE) ;
   return value;
 }
 void BasicLock::print_on(outputStream* st) const {
   st->print("monitor");
 }
 void BasicLock::move_to(oop obj, BasicLock* dest) {
   // Check to see if we need to inflate the lock. This is only needed
   // if an object is locked using "this" lightweight monitor. In that
   // case, the displaced_header() is unlocked, because the
   // displaced_header() contains the header for the originally unlocked
   // object. However the object could have already been inflated. But it
   // does not matter, the inflation will just a no-op. For other cases,
   // the displaced header will be either 0x0 or 0x3, which are location
   // independent, therefore the BasicLock is free to move.
   //
   // During OSR we may need to relocate a BasicLock (which contains a
   // displaced word) from a location in an interpreter frame to a
   // new location in a compiled frame.  "this" refers to the source
   // basiclock in the interpreter frame.  "dest" refers to the destination
   // basiclock in the new compiled frame.  We *always* inflate in move_to().
   // The always-Inflate policy works properly, but in 1.5.0 it can sometimes
   // cause performance problems in code that makes heavy use of a small # of
   // uncontended locks.   (We'd inflate during OSR, and then sync performance
   // would subsequently plummet because the thread would be forced thru the slow-path).
   // This problem has been made largely moot on IA32 by inlining the inflated fast-path
   // operations in Fast_Lock and Fast_Unlock in i486.ad.
   //
   // Note that there is a way to safely swing the object's markword from
   // one stack location to another.  This avoids inflation.  Obviously,
   // we need to ensure that both locations refer to the current thread's stack.
   // There are some subtle concurrency issues, however, and since the benefit is
   // is small (given the support for inflated fast-path locking in the fast_lock, etc)
   // we'll leave that optimization for another time.
   if (displaced_header()->is_neutral()) {
     ObjectSynchronizer::inflate_helper(obj);
     // WARNING: We can not put check here, because the inflation
     // will not update the displaced header. Once BasicLock is inflated,
     // no one should ever look at its content.
   } else {
     // Typically the displaced header will be 0 (recursive stack lock) or
     // unused_mark.  Naively we'd like to assert that the displaced mark
     // value is either 0, neutral, or 3.  But with the advent of the
     // store-before-CAS avoidance in fast_lock/compiler_lock_object
     // we can find any flavor mark in the displaced mark.
   }
 // [RGV] The next line appears to do nothing!
   intptr_t dh = (intptr_t) displaced_header();
   dest->set_displaced_header(displaced_header());
 }
 // -----------------------------------------------------------------------------
 // standard constructor, allows locking failures
 ObjectLocker::ObjectLocker(Handle obj, Thread* thread, bool doLock) {
   _dolock = doLock;
   _thread = thread;
   debug_only(if (StrictSafepointChecks) _thread->check_for_valid_safepoint_state(false);)
   _obj = obj;
   if (_dolock) {
     TEVENT (ObjectLocker) ;
     ObjectSynchronizer::fast_enter(_obj, &_lock, false, _thread);
   }
 }
 ObjectLocker::~ObjectLocker() {
   if (_dolock) {
     ObjectSynchronizer::fast_exit(_obj(), &_lock, _thread);
   }
 }
 // -----------------------------------------------------------------------------
 PerfCounter * ObjectSynchronizer::_sync_Inflations                  = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_Deflations                  = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_ContendedLockAttempts       = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_FutileWakeups               = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_Parks                       = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_EmptyNotifications          = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_Notifications               = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_PrivateA                    = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_PrivateB                    = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_SlowExit                    = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_SlowEnter                   = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_SlowNotify                  = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_SlowNotifyAll               = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_FailedSpins                 = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_SuccessfulSpins             = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_MonInCirculation            = NULL ;
 PerfCounter * ObjectSynchronizer::_sync_MonScavenged                = NULL ;
 PerfLongVariable * ObjectSynchronizer::_sync_MonExtant              = NULL ;
 // One-shot global initialization for the sync subsystem.
 // We could also defer initialization and initialize on-demand
 // the first time we call inflate().  Initialization would
 // be protected - like so many things - by the MonitorCache_lock.
 void ObjectSynchronizer::Initialize () {
   static int InitializationCompleted = 0 ;
   assert (InitializationCompleted == 0, "invariant") ;
   InitializationCompleted = 1 ;
   if (UsePerfData) {
       EXCEPTION_MARK ;
       #define NEWPERFCOUNTER(n)   {n = PerfDataManager::create_counter(SUN_RT, #n, PerfData::U_Events,CHECK); }
       #define NEWPERFVARIABLE(n)  {n = PerfDataManager::create_variable(SUN_RT, #n, PerfData::U_Events,CHECK); }
       NEWPERFCOUNTER(_sync_Inflations) ;
       NEWPERFCOUNTER(_sync_Deflations) ;
       NEWPERFCOUNTER(_sync_ContendedLockAttempts) ;
       NEWPERFCOUNTER(_sync_FutileWakeups) ;
       NEWPERFCOUNTER(_sync_Parks) ;
       NEWPERFCOUNTER(_sync_EmptyNotifications) ;
       NEWPERFCOUNTER(_sync_Notifications) ;
       NEWPERFCOUNTER(_sync_SlowEnter) ;
       NEWPERFCOUNTER(_sync_SlowExit) ;
       NEWPERFCOUNTER(_sync_SlowNotify) ;
       NEWPERFCOUNTER(_sync_SlowNotifyAll) ;
       NEWPERFCOUNTER(_sync_FailedSpins) ;
       NEWPERFCOUNTER(_sync_SuccessfulSpins) ;
       NEWPERFCOUNTER(_sync_PrivateA) ;
       NEWPERFCOUNTER(_sync_PrivateB) ;
       NEWPERFCOUNTER(_sync_MonInCirculation) ;
       NEWPERFCOUNTER(_sync_MonScavenged) ;
       NEWPERFVARIABLE(_sync_MonExtant) ;
       #undef NEWPERFCOUNTER
   }
 }
 // Compile-time asserts
 // When possible, it's better to catch errors deterministically at
 // compile-time than at runtime.  The down-side to using compile-time
 // asserts is that error message -- often something about negative array
 // indices -- is opaque.
 #define CTASSERT(x) { int tag[1-(2*!(x))]; printf ("Tag @" INTPTR_FORMAT "\n", (intptr_t)tag); }
 void ObjectMonitor::ctAsserts() {
   CTASSERT(offset_of (ObjectMonitor, _header) == 0);
 }
 static int Adjust (volatile int * adr, int dx) {
   int v ;
   for (v = *adr ; Atomic::cmpxchg (v + dx, adr, v) != v; v = *adr) ;
   return v ;
 }
 // Ad-hoc mutual exclusion primitives: SpinLock and Mux
 //
 // We employ SpinLocks _only for low-contention, fixed-length
 // short-duration critical sections where we're concerned
 // about native mutex_t or HotSpot Mutex:: latency.
 // The mux construct provides a spin-then-block mutual exclusion
 // mechanism.
 //
 // Testing has shown that contention on the ListLock guarding gFreeList
 // is common.  If we implement ListLock as a simple SpinLock it's common
 // for the JVM to devolve to yielding with little progress.  This is true
 // despite the fact that the critical sections protected by ListLock are
 // extremely short.
 //
 // TODO-FIXME: ListLock should be of type SpinLock.
 // We should make this a 1st-class type, integrated into the lock
 // hierarchy as leaf-locks.  Critically, the SpinLock structure
 // should have sufficient padding to avoid false-sharing and excessive
 // cache-coherency traffic.
 typedef volatile int SpinLockT ;
 void Thread::SpinAcquire (volatile int * adr, const char * LockName) {
   if (Atomic::cmpxchg (1, adr, 0) == 0) {
      return ;   // normal fast-path return
   }
   // Slow-path : We've encountered contention -- Spin/Yield/Block strategy.
   TEVENT (SpinAcquire - ctx) ;
   int ctr = 0 ;
   int Yields = 0 ;
   for (;;) {
      while (*adr != 0) {
         ++ctr ;
         if ((ctr & 0xFFF) == 0 || !os::is_MP()) {
            if (Yields > 5) {
              // Consider using a simple NakedSleep() instead.
              // Then SpinAcquire could be called by non-JVM threads
              Thread::current()->_ParkEvent->park(1) ;
            } else {
              os::NakedYield() ;
              ++Yields ;
            }
         } else {
            SpinPause() ;
         }
      }
      if (Atomic::cmpxchg (1, adr, 0) == 0) return ;
   }
 }
 void Thread::SpinRelease (volatile int * adr) {
   assert (*adr != 0, "invariant") ;
   OrderAccess::fence() ;      // guarantee at least release consistency.
   // Roach-motel semantics.
   // It's safe if subsequent LDs and STs float "up" into the critical section,
   // but prior LDs and STs within the critical section can't be allowed
   // to reorder or float past the ST that releases the lock.
   *adr = 0 ;
 }
 // muxAcquire and muxRelease:
 //
 // *  muxAcquire and muxRelease support a single-word lock-word construct.
 //    The LSB of the word is set IFF the lock is held.
 //    The remainder of the word points to the head of a singly-linked list
 //    of threads blocked on the lock.
 //
 // *  The current implementation of muxAcquire-muxRelease uses its own
 //    dedicated Thread._MuxEvent instance.  If we're interested in
 //    minimizing the peak number of extant ParkEvent instances then
 //    we could eliminate _MuxEvent and "borrow" _ParkEvent as long
 //    as certain invariants were satisfied.  Specifically, care would need
 //    to be taken with regards to consuming unpark() "permits".
 //    A safe rule of thumb is that a thread would never call muxAcquire()
 //    if it's enqueued (cxq, EntryList, WaitList, etc) and will subsequently
 //    park().  Otherwise the _ParkEvent park() operation in muxAcquire() could
 //    consume an unpark() permit intended for monitorenter, for instance.
 //    One way around this would be to widen the restricted-range semaphore
 //    implemented in park().  Another alternative would be to provide
 //    multiple instances of the PlatformEvent() for each thread.  One
 //    instance would be dedicated to muxAcquire-muxRelease, for instance.
 //
 // *  Usage:
 //    -- Only as leaf locks
 //    -- for short-term locking only as muxAcquire does not perform
 //       thread state transitions.
 //
 // Alternatives:
 // *  We could implement muxAcquire and muxRelease with MCS or CLH locks
 //    but with parking or spin-then-park instead of pure spinning.
 // *  Use Taura-Oyama-Yonenzawa locks.
 // *  It's possible to construct a 1-0 lock if we encode the lockword as
 //    (List,LockByte).  Acquire will CAS the full lockword while Release
 //    will STB 0 into the LockByte.  The 1-0 scheme admits stranding, so
 //    acquiring threads use timers (ParkTimed) to detect and recover from
 //    the stranding window.  Thread/Node structures must be aligned on 256-byte
 //    boundaries by using placement-new.
 // *  Augment MCS with advisory back-link fields maintained with CAS().
 //    Pictorially:  LockWord -> T1 <-> T2 <-> T3 <-> ... <-> Tn <-> Owner.
 //    The validity of the backlinks must be ratified before we trust the value.
 //    If the backlinks are invalid the exiting thread must back-track through the
 //    the forward links, which are always trustworthy.
 // *  Add a successor indication.  The LockWord is currently encoded as
 //    (List, LOCKBIT:1).  We could also add a SUCCBIT or an explicit _succ variable
 //    to provide the usual futile-wakeup optimization.
 //    See RTStt for details.
 // *  Consider schedctl.sc_nopreempt to cover the critical section.
 //
 typedef volatile intptr_t MutexT ;      // Mux Lock-word
 enum MuxBits { LOCKBIT = 1 } ;
 void Thread::muxAcquire (volatile intptr_t * Lock, const char * LockName) {
   intptr_t w = Atomic::cmpxchg_ptr (LOCKBIT, Lock, 0) ;
   if (w == 0) return ;
   if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
      return ;
   }
   TEVENT (muxAcquire - Contention) ;
   ParkEvent * const Self = Thread::current()->_MuxEvent ;
   assert ((intptr_t(Self) & LOCKBIT) == 0, "invariant") ;
   for (;;) {
      int its = (os::is_MP() ? 100 : 0) + 1 ;
      // Optional spin phase: spin-then-park strategy
      while (--its >= 0) {
        w = *Lock ;
        if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
           return ;
        }
      }
      Self->reset() ;
      Self->OnList = intptr_t(Lock) ;
      // The following fence() isn't _strictly necessary as the subsequent
      // CAS() both serializes execution and ratifies the fetched *Lock value.
      OrderAccess::fence();
      for (;;) {
         w = *Lock ;
         if ((w & LOCKBIT) == 0) {
             if (Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
                 Self->OnList = 0 ;   // hygiene - allows stronger asserts
                 return ;
             }
             continue ;      // Interference -- *Lock changed -- Just retry
         }
         assert (w & LOCKBIT, "invariant") ;
         Self->ListNext = (ParkEvent *) (w & ~LOCKBIT );
         if (Atomic::cmpxchg_ptr (intptr_t(Self)|LOCKBIT, Lock, w) == w) break ;
      }
      while (Self->OnList != 0) {
         Self->park() ;
      }
   }
 }
 void Thread::muxAcquireW (volatile intptr_t * Lock, ParkEvent * ev) {
   intptr_t w = Atomic::cmpxchg_ptr (LOCKBIT, Lock, 0) ;
   if (w == 0) return ;
   if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
     return ;
   }
   TEVENT (muxAcquire - Contention) ;
   ParkEvent * ReleaseAfter = NULL ;
   if (ev == NULL) {
     ev = ReleaseAfter = ParkEvent::Allocate (NULL) ;
   }
   assert ((intptr_t(ev) & LOCKBIT) == 0, "invariant") ;
   for (;;) {
     guarantee (ev->OnList == 0, "invariant") ;
     int its = (os::is_MP() ? 100 : 0) + 1 ;
     // Optional spin phase: spin-then-park strategy
     while (--its >= 0) {
       w = *Lock ;
       if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
         if (ReleaseAfter != NULL) {
           ParkEvent::Release (ReleaseAfter) ;
         }
         return ;
       }
     }
     ev->reset() ;
     ev->OnList = intptr_t(Lock) ;
     // The following fence() isn't _strictly necessary as the subsequent
     // CAS() both serializes execution and ratifies the fetched *Lock value.
     OrderAccess::fence();
     for (;;) {
       w = *Lock ;
       if ((w & LOCKBIT) == 0) {
         if (Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
           ev->OnList = 0 ;
           // We call ::Release while holding the outer lock, thus
           // artificially lengthening the critical section.
           // Consider deferring the ::Release() until the subsequent unlock(),
           // after we've dropped the outer lock.
           if (ReleaseAfter != NULL) {
             ParkEvent::Release (ReleaseAfter) ;
           }
           return ;
         }
         continue ;      // Interference -- *Lock changed -- Just retry
       }
       assert (w & LOCKBIT, "invariant") ;
       ev->ListNext = (ParkEvent *) (w & ~LOCKBIT );
       if (Atomic::cmpxchg_ptr (intptr_t(ev)|LOCKBIT, Lock, w) == w) break ;
     }
     while (ev->OnList != 0) {
       ev->park() ;
     }
   }
 }
 // Release() must extract a successor from the list and then wake that thread.
 // It can "pop" the front of the list or use a detach-modify-reattach (DMR) scheme
 // similar to that used by ParkEvent::Allocate() and ::Release().  DMR-based
 // Release() would :
 // (A) CAS() or swap() null to *Lock, releasing the lock and detaching the list.
 // (B) Extract a successor from the private list "in-hand"
 // (C) attempt to CAS() the residual back into *Lock over null.
 //     If there were any newly arrived threads and the CAS() would fail.
 //     In that case Release() would detach the RATs, re-merge the list in-hand
 //     with the RATs and repeat as needed.  Alternately, Release() might
 //     detach and extract a successor, but then pass the residual list to the wakee.
 //     The wakee would be responsible for reattaching and remerging before it
 //     competed for the lock.
 //
 // Both "pop" and DMR are immune from ABA corruption -- there can be
 // multiple concurrent pushers, but only one popper or detacher.
 // This implementation pops from the head of the list.  This is unfair,
 // but tends to provide excellent throughput as hot threads remain hot.
 // (We wake recently run threads first).
 void Thread::muxRelease (volatile intptr_t * Lock)  {
   for (;;) {
     const intptr_t w = Atomic::cmpxchg_ptr (0, Lock, LOCKBIT) ;
     assert (w & LOCKBIT, "invariant") ;
     if (w == LOCKBIT) return ;
     ParkEvent * List = (ParkEvent *) (w & ~LOCKBIT) ;
     assert (List != NULL, "invariant") ;
     assert (List->OnList == intptr_t(Lock), "invariant") ;
     ParkEvent * nxt = List->ListNext ;
     // The following CAS() releases the lock and pops the head element.
     if (Atomic::cmpxchg_ptr (intptr_t(nxt), Lock, w) != w) {
       continue ;
     }
     List->OnList = 0 ;
     OrderAccess::fence() ;
     List->unpark () ;
     return ;
   }
 }
 // ObjectMonitor Lifecycle
 // -----------------------
 // Inflation unlinks monitors from the global gFreeList and
 // associates them with objects.  Deflation -- which occurs at
 // STW-time -- disassociates idle monitors from objects.  Such
 // scavenged monitors are returned to the gFreeList.
 //
 // The global list is protected by ListLock.  All the critical sections
 // are short and operate in constant-time.
 //
 // ObjectMonitors reside in type-stable memory (TSM) and are immortal.
 //
 // Lifecycle:
 // --   unassigned and on the global free list
 // --   unassigned and on a thread's private omFreeList
 // --   assigned to an object.  The object is inflated and the mark refers
 //      to the objectmonitor.
 //
 // TODO-FIXME:
 //
 // *  We currently protect the gFreeList with a simple lock.
 //    An alternate lock-free scheme would be to pop elements from the gFreeList
 //    with CAS.  This would be safe from ABA corruption as long we only
 //    recycled previously appearing elements onto the list in deflate_idle_monitors()
 //    at STW-time.  Completely new elements could always be pushed onto the gFreeList
 //    with CAS.  Elements that appeared previously on the list could only
 //    be installed at STW-time.
 //
 // *  For efficiency and to help reduce the store-before-CAS penalty
 //    the objectmonitors on gFreeList or local free lists should be ready to install
 //    with the exception of _header and _object.  _object can be set after inflation.
 //    In particular, keep all objectMonitors on a thread's private list in ready-to-install
 //    state with m.Owner set properly.
 //
 // *  We could all diffuse contention by using multiple global (FreeList, Lock)
 //    pairs -- threads could use trylock() and a cyclic-scan strategy to search for
 //    an unlocked free list.
 //
 // *  Add lifecycle tags and assert()s.
 //
 // *  Be more consistent about when we clear an objectmonitor's fields:
 //    A.  After extracting the objectmonitor from a free list.
 //    B.  After adding an objectmonitor to a free list.
 //
 ObjectMonitor * ObjectSynchronizer::gBlockList = NULL ;
 ObjectMonitor * volatile ObjectSynchronizer::gFreeList  = NULL ;
 static volatile intptr_t ListLock = 0 ;      // protects global monitor free-list cache
 #define CHAINMARKER ((oop)-1)
 ObjectMonitor * ATTR ObjectSynchronizer::omAlloc (Thread * Self) {
     // A large MAXPRIVATE value reduces both list lock contention
     // and list coherency traffic, but also tends to increase the
     // number of objectMonitors in circulation as well as the STW
     // scavenge costs.  As usual, we lean toward time in space-time
     // tradeoffs.
     const int MAXPRIVATE = 1024 ;
     for (;;) {
         ObjectMonitor * m ;
         // 1: try to allocate from the thread's local omFreeList.
         // Threads will attempt to allocate first from their local list, then
         // from the global list, and only after those attempts fail will the thread
         // attempt to instantiate new monitors.   Thread-local free lists take
         // heat off the ListLock and improve allocation latency, as well as reducing
         // coherency traffic on the shared global list.
         m = Self->omFreeList ;
         if (m != NULL) {
            Self->omFreeList = m->FreeNext ;
            Self->omFreeCount -- ;
            // CONSIDER: set m->FreeNext = BAD -- diagnostic hygiene
            guarantee (m->object() == NULL, "invariant") ;
            return m ;
         }
         // 2: try to allocate from the global gFreeList
         // CONSIDER: use muxTry() instead of muxAcquire().
         // If the muxTry() fails then drop immediately into case 3.
         // If we're using thread-local free lists then try
         // to reprovision the caller's free list.
         if (gFreeList != NULL) {
             // Reprovision the thread's omFreeList.
             // Use bulk transfers to reduce the allocation rate and heat
             // on various locks.
             Thread::muxAcquire (&ListLock, "omAlloc") ;
             for (int i = Self->omFreeProvision; --i >= 0 && gFreeList != NULL; ) {
                 ObjectMonitor * take = gFreeList ;
                 gFreeList = take->FreeNext ;
                 guarantee (take->object() == NULL, "invariant") ;
                 guarantee (!take->is_busy(), "invariant") ;
                 take->Recycle() ;
                 omRelease (Self, take) ;
             }
             Thread::muxRelease (&ListLock) ;
             Self->omFreeProvision += 1 + (Self->omFreeProvision/2) ;
             if (Self->omFreeProvision > MAXPRIVATE ) Self->omFreeProvision = MAXPRIVATE ;
             TEVENT (omFirst - reprovision) ;
             continue ;
         }
         // 3: allocate a block of new ObjectMonitors
         // Both the local and global free lists are empty -- resort to malloc().
         // In the current implementation objectMonitors are TSM - immortal.
         assert (_BLOCKSIZE > 1, "invariant") ;
         ObjectMonitor * temp = new ObjectMonitor[_BLOCKSIZE];
         // NOTE: (almost) no way to recover if allocation failed.
         // We might be able to induce a STW safepoint and scavenge enough
         // objectMonitors to permit progress.
         if (temp == NULL) {
             vm_exit_out_of_memory (sizeof (ObjectMonitor[_BLOCKSIZE]), "Allocate ObjectMonitors") ;
         }
         // Format the block.
         // initialize the linked list, each monitor points to its next
         // forming the single linked free list, the very first monitor
         // will points to next block, which forms the block list.
         // The trick of using the 1st element in the block as gBlockList
         // linkage should be reconsidered.  A better implementation would
         // look like: class Block { Block * next; int N; ObjectMonitor Body [N] ; }
         for (int i = 1; i < _BLOCKSIZE ; i++) {
            temp[i].FreeNext = &temp[i+1];
         }
         // terminate the last monitor as the end of list
         temp[_BLOCKSIZE - 1].FreeNext = NULL ;
         // Element [0] is reserved for global list linkage
         temp[0].set_object(CHAINMARKER);
         // Consider carving out this thread's current request from the
         // block in hand.  This avoids some lock traffic and redundant
         // list activity.
         // Acquire the ListLock to manipulate BlockList and FreeList.
         // An Oyama-Taura-Yonezawa scheme might be more efficient.
         Thread::muxAcquire (&ListLock, "omAlloc [2]") ;
         // Add the new block to the list of extant blocks (gBlockList).
         // The very first objectMonitor in a block is reserved and dedicated.
         // It serves as blocklist "next" linkage.
         temp[0].FreeNext = gBlockList;
         gBlockList = temp;
         // Add the new string of objectMonitors to the global free list
         temp[_BLOCKSIZE - 1].FreeNext = gFreeList ;
         gFreeList = temp + 1;
         Thread::muxRelease (&ListLock) ;
         TEVENT (Allocate block of monitors) ;
     }
 }
 // Place "m" on the caller's private per-thread omFreeList.
 // In practice there's no need to clamp or limit the number of
 // monitors on a thread's omFreeList as the only time we'll call
 // omRelease is to return a monitor to the free list after a CAS
 // attempt failed.  This doesn't allow unbounded #s of monitors to
 // accumulate on a thread's free list.
 //
 // In the future the usage of omRelease() might change and monitors
 // could migrate between free lists.  In that case to avoid excessive
 // accumulation we could  limit omCount to (omProvision*2), otherwise return
 // the objectMonitor to the global list.  We should drain (return) in reasonable chunks.
 // That is, *not* one-at-a-time.
 void ObjectSynchronizer::omRelease (Thread * Self, ObjectMonitor * m) {
     guarantee (m->object() == NULL, "invariant") ;
     m->FreeNext = Self->omFreeList ;
     Self->omFreeList = m ;
     Self->omFreeCount ++ ;
 }
 // Return the monitors of a moribund thread's local free list to
 // the global free list.  Typically a thread calls omFlush() when
 // it's dying.  We could also consider having the VM thread steal
 // monitors from threads that have not run java code over a few
 // consecutive STW safepoints.  Relatedly, we might decay
 // omFreeProvision at STW safepoints.
 //
 // We currently call omFlush() from the Thread:: dtor _after the thread
 // has been excised from the thread list and is no longer a mutator.
 // That means that omFlush() can run concurrently with a safepoint and
 // the scavenge operator.  Calling omFlush() from JavaThread::exit() might
 // be a better choice as we could safely reason that that the JVM is
 // not at a safepoint at the time of the call, and thus there could
 // be not inopportune interleavings between omFlush() and the scavenge
 // operator.
 void ObjectSynchronizer::omFlush (Thread * Self) {
     ObjectMonitor * List = Self->omFreeList ;  // Null-terminated SLL
     Self->omFreeList = NULL ;
     if (List == NULL) return ;
     ObjectMonitor * Tail = NULL ;
     ObjectMonitor * s ;
     for (s = List ; s != NULL ; s = s->FreeNext) {
         Tail = s ;
         guarantee (s->object() == NULL, "invariant") ;
         guarantee (!s->is_busy(), "invariant") ;
         s->set_owner (NULL) ;   // redundant but good hygiene
         TEVENT (omFlush - Move one) ;
     }
     guarantee (Tail != NULL && List != NULL, "invariant") ;
     Thread::muxAcquire (&ListLock, "omFlush") ;
     Tail->FreeNext = gFreeList ;
     gFreeList = List ;
     Thread::muxRelease (&ListLock) ;
     TEVENT (omFlush) ;
 }
 // Get the next block in the block list.
 static inline ObjectMonitor* next(ObjectMonitor* block) {
   assert(block->object() == CHAINMARKER, "must be a block header");
   block = block->FreeNext ;
   assert(block == NULL || block->object() == CHAINMARKER, "must be a block header");
   return block;
 }
 // Fast path code shared by multiple functions
 ObjectMonitor* ObjectSynchronizer::inflate_helper(oop obj) {
   markOop mark = obj->mark();
   if (mark->has_monitor()) {
     assert(ObjectSynchronizer::verify_objmon_isinpool(mark->monitor()), "monitor is invalid");
     assert(mark->monitor()->header()->is_neutral(), "monitor must record a good object header");
     return mark->monitor();
   }
   return ObjectSynchronizer::inflate(Thread::current(), obj);
 }
 // Note that we could encounter some performance loss through false-sharing as
 // multiple locks occupy the same $ line.  Padding might be appropriate.
 #define NINFLATIONLOCKS 256
 static volatile intptr_t InflationLocks [NINFLATIONLOCKS] ;
 static markOop ReadStableMark (oop obj) {
   markOop mark = obj->mark() ;
   if (!mark->is_being_inflated()) {
     return mark ;       // normal fast-path return
   }
   int its = 0 ;
   for (;;) {
     markOop mark = obj->mark() ;
     if (!mark->is_being_inflated()) {
       return mark ;    // normal fast-path return
     }
     // The object is being inflated by some other thread.
     // The caller of ReadStableMark() must wait for inflation to complete.
     // Avoid live-lock
     // TODO: consider calling SafepointSynchronize::do_call_back() while
     // spinning to see if there's a safepoint pending.  If so, immediately
     // yielding or blocking would be appropriate.  Avoid spinning while
     // there is a safepoint pending.
     // TODO: add inflation contention performance counters.
     // TODO: restrict the aggregate number of spinners.
     ++its ;
     if (its > 10000 || !os::is_MP()) {
        if (its & 1) {
          os::NakedYield() ;
          TEVENT (Inflate: INFLATING - yield) ;
        } else {
          // Note that the following code attenuates the livelock problem but is not
          // a complete remedy.  A more complete solution would require that the inflating
          // thread hold the associated inflation lock.  The following code simply restricts
          // the number of spinners to at most one.  We'll have N-2 threads blocked
          // on the inflationlock, 1 thread holding the inflation lock and using
          // a yield/park strategy, and 1 thread in the midst of inflation.
          // A more refined approach would be to change the encoding of INFLATING
          // to allow encapsulation of a native thread pointer.  Threads waiting for
          // inflation to complete would use CAS to push themselves onto a singly linked
          // list rooted at the markword.  Once enqueued, they'd loop, checking a per-thread flag
          // and calling park().  When inflation was complete the thread that accomplished inflation
          // would detach the list and set the markword to inflated with a single CAS and
          // then for each thread on the list, set the flag and unpark() the thread.
          // This is conceptually similar to muxAcquire-muxRelease, except that muxRelease
          // wakes at most one thread whereas we need to wake the entire list.
          int ix = (intptr_t(obj) >> 5) & (NINFLATIONLOCKS-1) ;
          int YieldThenBlock = 0 ;
          assert (ix >= 0 && ix < NINFLATIONLOCKS, "invariant") ;
          assert ((NINFLATIONLOCKS & (NINFLATIONLOCKS-1)) == 0, "invariant") ;
          Thread::muxAcquire (InflationLocks + ix, "InflationLock") ;
          while (obj->mark() == markOopDesc::INFLATING()) {
            // Beware: NakedYield() is advisory and has almost no effect on some platforms
            // so we periodically call Self->_ParkEvent->park(1).
            // We use a mixed spin/yield/block mechanism.
            if ((YieldThenBlock++) >= 16) {
               Thread::current()->_ParkEvent->park(1) ;
            } else {
               os::NakedYield() ;
            }
          }
          Thread::muxRelease (InflationLocks + ix ) ;
          TEVENT (Inflate: INFLATING - yield/park) ;
        }
     } else {
        SpinPause() ;       // SMP-polite spinning
     }
   }
 }
 ObjectMonitor * ATTR ObjectSynchronizer::inflate (Thread * Self, oop object) {
   // Inflate mutates the heap ...
   // Relaxing assertion for bug 6320749.
   assert (Universe::verify_in_progress() ||
           !SafepointSynchronize::is_at_safepoint(), "invariant") ;
   for (;;) {
       const markOop mark = object->mark() ;
       assert (!mark->has_bias_pattern(), "invariant") ;
       // The mark can be in one of the following states:
       // *  Inflated     - just return
       // *  Stack-locked - coerce it to inflated
       // *  INFLATING    - busy wait for conversion to complete
       // *  Neutral      - aggressively inflate the object.
       // *  BIASED       - Illegal.  We should never see this
       // CASE: inflated
       if (mark->has_monitor()) {
           ObjectMonitor * inf = mark->monitor() ;
           assert (inf->header()->is_neutral(), "invariant");
           assert (inf->object() == object, "invariant") ;
           assert (ObjectSynchronizer::verify_objmon_isinpool(inf), "monitor is invalid");
           return inf ;
       }
       // CASE: inflation in progress - inflating over a stack-lock.
       // Some other thread is converting from stack-locked to inflated.
       // Only that thread can complete inflation -- other threads must wait.
       // The INFLATING value is transient.
       // Currently, we spin/yield/park and poll the markword, waiting for inflation to finish.
       // We could always eliminate polling by parking the thread on some auxiliary list.
       if (mark == markOopDesc::INFLATING()) {
          TEVENT (Inflate: spin while INFLATING) ;
          ReadStableMark(object) ;
          continue ;
       }
       // CASE: stack-locked
       // Could be stack-locked either by this thread or by some other thread.
       //
       // Note that we allocate the objectmonitor speculatively, _before_ attempting
       // to install INFLATING into the mark word.  We originally installed INFLATING,
       // allocated the objectmonitor, and then finally STed the address of the
       // objectmonitor into the mark.  This was correct, but artificially lengthened
       // the interval in which INFLATED appeared in the mark, thus increasing
       // the odds of inflation contention.
       //
       // We now use per-thread private objectmonitor free lists.
       // These list are reprovisioned from the global free list outside the
       // critical INFLATING...ST interval.  A thread can transfer
       // multiple objectmonitors en-mass from the global free list to its local free list.
       // This reduces coherency traffic and lock contention on the global free list.
       // Using such local free lists, it doesn't matter if the omAlloc() call appears
       // before or after the CAS(INFLATING) operation.
       // See the comments in omAlloc().
       if (mark->has_locker()) {
           ObjectMonitor * m = omAlloc (Self) ;
           // Optimistically prepare the objectmonitor - anticipate successful CAS
           // We do this before the CAS in order to minimize the length of time
           // in which INFLATING appears in the mark.
           m->Recycle();
           m->FreeNext      = NULL ;
           m->_Responsible  = NULL ;
           m->OwnerIsThread = 0 ;
           m->_recursions   = 0 ;
           m->_SpinDuration = Knob_SpinLimit ;   // Consider: maintain by type/class
           markOop cmp = (markOop) Atomic::cmpxchg_ptr (markOopDesc::INFLATING(), object->mark_addr(), mark) ;
           if (cmp != mark) {
              omRelease (Self, m) ;
              continue ;       // Interference -- just retry
           }
           // We've successfully installed INFLATING (0) into the mark-word.
           // This is the only case where 0 will appear in a mark-work.
           // Only the singular thread that successfully swings the mark-word
           // to 0 can perform (or more precisely, complete) inflation.
           //
           // Why do we CAS a 0 into the mark-word instead of just CASing the
           // mark-word from the stack-locked value directly to the new inflated state?
           // Consider what happens when a thread unlocks a stack-locked object.
           // It attempts to use CAS to swing the displaced header value from the
           // on-stack basiclock back into the object header.  Recall also that the
           // header value (hashcode, etc) can reside in (a) the object header, or
           // (b) a displaced header associated with the stack-lock, or (c) a displaced
           // header in an objectMonitor.  The inflate() routine must copy the header
           // value from the basiclock on the owner's stack to the objectMonitor, all
           // the while preserving the hashCode stability invariants.  If the owner
           // decides to release the lock while the value is 0, the unlock will fail
           // and control will eventually pass from slow_exit() to inflate.  The owner
           // will then spin, waiting for the 0 value to disappear.   Put another way,
           // the 0 causes the owner to stall if the owner happens to try to
           // drop the lock (restoring the header from the basiclock to the object)
           // while inflation is in-progress.  This protocol avoids races that might
           // would otherwise permit hashCode values to change or "flicker" for an object.
           // Critically, while object->mark is 0 mark->displaced_mark_helper() is stable.
           // 0 serves as a "BUSY" inflate-in-progress indicator.
           // fetch the displaced mark from the owner's stack.
           // The owner can't die or unwind past the lock while our INFLATING
           // object is in the mark.  Furthermore the owner can't complete
           // an unlock on the object, either.
           markOop dmw = mark->displaced_mark_helper() ;
           assert (dmw->is_neutral(), "invariant") ;
           // Setup monitor fields to proper values -- prepare the monitor
           m->set_header(dmw) ;
           // Optimization: if the mark->locker stack address is associated
           // with this thread we could simply set m->_owner = Self and
           // m->OwnerIsThread = 1.  Note that a thread can inflate an object
           // that it has stack-locked -- as might happen in wait() -- directly
           // with CAS.  That is, we can avoid the xchg-NULL .... ST idiom.
           m->set_owner (mark->locker());
           m->set_object(object);
           // TODO-FIXME: assert BasicLock->dhw != 0.
           // Must preserve store ordering. The monitor state must
           // be stable at the time of publishing the monitor address.
           guarantee (object->mark() == markOopDesc::INFLATING(), "invariant") ;
           object->release_set_mark(markOopDesc::encode(m));
           // Hopefully the performance counters are allocated on distinct cache lines
           // to avoid false sharing on MP systems ...
           if (_sync_Inflations != NULL) _sync_Inflations->inc() ;
           TEVENT(Inflate: overwrite stacklock) ;
           if (TraceMonitorInflation) {
             if (object->is_instance()) {
               ResourceMark rm;
               tty->print_cr("Inflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s",
                 (intptr_t) object, (intptr_t) object->mark(),
                 Klass::cast(object->klass())->external_name());
             }
           }
           return m ;
       }
       // CASE: neutral
       // TODO-FIXME: for entry we currently inflate and then try to CAS _owner.
       // If we know we're inflating for entry it's better to inflate by swinging a
       // pre-locked objectMonitor pointer into the object header.   A successful
       // CAS inflates the object *and* confers ownership to the inflating thread.
       // In the current implementation we use a 2-step mechanism where we CAS()
       // to inflate and then CAS() again to try to swing _owner from NULL to Self.
       // An inflateTry() method that we could call from fast_enter() and slow_enter()
       // would be useful.
       assert (mark->is_neutral(), "invariant");
       ObjectMonitor * m = omAlloc (Self) ;
       // prepare m for installation - set monitor to initial state
       m->Recycle();
       m->set_header(mark);
       m->set_owner(NULL);
       m->set_object(object);
       m->OwnerIsThread = 1 ;
       m->_recursions   = 0 ;
       m->FreeNext      = NULL ;
       m->_Responsible  = NULL ;
       m->_SpinDuration = Knob_SpinLimit ;       // consider: keep metastats by type/class
       if (Atomic::cmpxchg_ptr (markOopDesc::encode(m), object->mark_addr(), mark) != mark) {
           m->set_object (NULL) ;
           m->set_owner  (NULL) ;
           m->OwnerIsThread = 0 ;
           m->Recycle() ;
           omRelease (Self, m) ;
           m = NULL ;
           continue ;
           // interference - the markword changed - just retry.
           // The state-transitions are one-way, so there's no chance of
           // live-lock -- "Inflated" is an absorbing state.
       }
       // Hopefully the performance counters are allocated on distinct
       // cache lines to avoid false sharing on MP systems ...
       if (_sync_Inflations != NULL) _sync_Inflations->inc() ;
       TEVENT(Inflate: overwrite neutral) ;
       if (TraceMonitorInflation) {
         if (object->is_instance()) {
           ResourceMark rm;
           tty->print_cr("Inflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s",
             (intptr_t) object, (intptr_t) object->mark(),
             Klass::cast(object->klass())->external_name());
         }
       }
       return m ;
   }
 }
 // This the fast monitor enter. The interpreter and compiler use
 // some assembly copies of this code. Make sure update those code
 // if the following function is changed. The implementation is
 // extremely sensitive to race condition. Be careful.
 void ObjectSynchronizer::fast_enter(Handle obj, BasicLock* lock, bool attempt_rebias, TRAPS) {
  if (UseBiasedLocking) {
     if (!SafepointSynchronize::is_at_safepoint()) {
       BiasedLocking::Condition cond = BiasedLocking::revoke_and_rebias(obj, attempt_rebias, THREAD);
       if (cond == BiasedLocking::BIAS_REVOKED_AND_REBIASED) {
         return;
       }
     } else {
       assert(!attempt_rebias, "can not rebias toward VM thread");
       BiasedLocking::revoke_at_safepoint(obj);
     }
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   THREAD->update_highest_lock((address)lock);
   slow_enter (obj, lock, THREAD) ;
 }
 void ObjectSynchronizer::fast_exit(oop object, BasicLock* lock, TRAPS) {
   assert(!object->mark()->has_bias_pattern(), "should not see bias pattern here");
   // if displaced header is null, the previous enter is recursive enter, no-op
   markOop dhw = lock->displaced_header();
   markOop mark ;
   if (dhw == NULL) {
      // Recursive stack-lock.
      // Diagnostics -- Could be: stack-locked, inflating, inflated.
      mark = object->mark() ;
      assert (!mark->is_neutral(), "invariant") ;
      if (mark->has_locker() && mark != markOopDesc::INFLATING()) {
         assert(THREAD->is_lock_owned((address)mark->locker()), "invariant") ;
      }
      if (mark->has_monitor()) {
         ObjectMonitor * m = mark->monitor() ;
         assert(((oop)(m->object()))->mark() == mark, "invariant") ;
         assert(m->is_entered(THREAD), "invariant") ;
      }
      return ;
   }
   mark = object->mark() ;
   // If the object is stack-locked by the current thread, try to
   // swing the displaced header from the box back to the mark.
   if (mark == (markOop) lock) {
      assert (dhw->is_neutral(), "invariant") ;
      if ((markOop) Atomic::cmpxchg_ptr (dhw, object->mark_addr(), mark) == mark) {
         TEVENT (fast_exit: release stacklock) ;
         return;
      }
   }
   ObjectSynchronizer::inflate(THREAD, object)->exit (THREAD) ;
 }
 // This routine is used to handle interpreter/compiler slow case
 // We don't need to use fast path here, because it must have been
 // failed in the interpreter/compiler code.
 void ObjectSynchronizer::slow_enter(Handle obj, BasicLock* lock, TRAPS) {
   markOop mark = obj->mark();
   assert(!mark->has_bias_pattern(), "should not see bias pattern here");
   if (mark->is_neutral()) {
     // Anticipate successful CAS -- the ST of the displaced mark must
     // be visible <= the ST performed by the CAS.
     lock->set_displaced_header(mark);
     if (mark == (markOop) Atomic::cmpxchg_ptr(lock, obj()->mark_addr(), mark)) {
       TEVENT (slow_enter: release stacklock) ;
       return ;
     }
     // Fall through to inflate() ...
   } else
   if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) {
     assert(lock != mark->locker(), "must not re-lock the same lock");
     assert(lock != (BasicLock*)obj->mark(), "don't relock with same BasicLock");
     lock->set_displaced_header(NULL);
     return;
   }
 #if 0
   // The following optimization isn't particularly useful.
   if (mark->has_monitor() && mark->monitor()->is_entered(THREAD)) {
     lock->set_displaced_header (NULL) ;
     return ;
   }
 #endif
   // The object header will never be displaced to this lock,
   // so it does not matter what the value is, except that it
   // must be non-zero to avoid looking like a re-entrant lock,
   // and must not look locked either.
   lock->set_displaced_header(markOopDesc::unused_mark());
   ObjectSynchronizer::inflate(THREAD, obj())->enter(THREAD);
 }
 // This routine is used to handle interpreter/compiler slow case
 // We don't need to use fast path here, because it must have
 // failed in the interpreter/compiler code. Simply use the heavy
 // weight monitor should be ok, unless someone find otherwise.
 void ObjectSynchronizer::slow_exit(oop object, BasicLock* lock, TRAPS) {
   fast_exit (object, lock, THREAD) ;
 }
 // NOTE: must use heavy weight monitor to handle jni monitor enter
 void ObjectSynchronizer::jni_enter(Handle obj, TRAPS) { // possible entry from jni enter
   // the current locking is from JNI instead of Java code
   TEVENT (jni_enter) ;
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   THREAD->set_current_pending_monitor_is_from_java(false);
   ObjectSynchronizer::inflate(THREAD, obj())->enter(THREAD);
   THREAD->set_current_pending_monitor_is_from_java(true);
 }
 // NOTE: must use heavy weight monitor to handle jni monitor enter
 bool ObjectSynchronizer::jni_try_enter(Handle obj, Thread* THREAD) {
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   ObjectMonitor* monitor = ObjectSynchronizer::inflate_helper(obj());
   return monitor->try_enter(THREAD);
 }
 // NOTE: must use heavy weight monitor to handle jni monitor exit
 void ObjectSynchronizer::jni_exit(oop obj, Thread* THREAD) {
   TEVENT (jni_exit) ;
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
   }
   assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj);
   // If this thread has locked the object, exit the monitor.  Note:  can't use
   // monitor->check(CHECK); must exit even if an exception is pending.
   if (monitor->check(THREAD)) {
      monitor->exit(THREAD);
   }
 }
 // complete_exit()/reenter() are used to wait on a nested lock
 // i.e. to give up an outer lock completely and then re-enter
 // Used when holding nested locks - lock acquisition order: lock1 then lock2
 //  1) complete_exit lock1 - saving recursion count
 //  2) wait on lock2
 //  3) when notified on lock2, unlock lock2
 //  4) reenter lock1 with original recursion count
 //  5) lock lock2
 // NOTE: must use heavy weight monitor to handle complete_exit/reenter()
 intptr_t ObjectSynchronizer::complete_exit(Handle obj, TRAPS) {
   TEVENT (complete_exit) ;
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj());
   return monitor->complete_exit(THREAD);
 }
 // NOTE: must use heavy weight monitor to handle complete_exit/reenter()
 void ObjectSynchronizer::reenter(Handle obj, intptr_t recursion, TRAPS) {
   TEVENT (reenter) ;
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj());
   monitor->reenter(recursion, THREAD);
 }
 // This exists only as a workaround of dtrace bug 6254741
 int dtrace_waited_probe(ObjectMonitor* monitor, Handle obj, Thread* thr) {
   DTRACE_MONITOR_PROBE(waited, monitor, obj(), thr);
   return 0;
 }
 // NOTE: must use heavy weight monitor to handle wait()
 void ObjectSynchronizer::wait(Handle obj, jlong millis, TRAPS) {
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   if (millis < 0) {
     TEVENT (wait - throw IAX) ;
     THROW_MSG(vmSymbols::java_lang_IllegalArgumentException(), "timeout value is negative");
   }
   ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj());
   DTRACE_MONITOR_WAIT_PROBE(monitor, obj(), THREAD, millis);
   monitor->wait(millis, true, THREAD);
   /* This dummy call is in place to get around dtrace bug 6254741.  Once
      that's fixed we can uncomment the following line and remove the call */
   // DTRACE_MONITOR_PROBE(waited, monitor, obj(), THREAD);
   dtrace_waited_probe(monitor, obj, THREAD);
 }
 void ObjectSynchronizer::waitUninterruptibly (Handle obj, jlong millis, TRAPS) {
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   if (millis < 0) {
     TEVENT (wait - throw IAX) ;
     THROW_MSG(vmSymbols::java_lang_IllegalArgumentException(), "timeout value is negative");
   }
   ObjectSynchronizer::inflate(THREAD, obj()) -> wait(millis, false, THREAD) ;
 }
 void ObjectSynchronizer::notify(Handle obj, TRAPS) {
  if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   markOop mark = obj->mark();
   if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) {
     return;
   }
   ObjectSynchronizer::inflate(THREAD, obj())->notify(THREAD);
 }
 // NOTE: see comment of notify()
 void ObjectSynchronizer::notifyall(Handle obj, TRAPS) {
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(obj, false, THREAD);
     assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   markOop mark = obj->mark();
   if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) {
     return;
   }
   ObjectSynchronizer::inflate(THREAD, obj())->notifyAll(THREAD);
 }
 intptr_t ObjectSynchronizer::FastHashCode (Thread * Self, oop obj) {
   if (UseBiasedLocking) {
     // NOTE: many places throughout the JVM do not expect a safepoint
     // to be taken here, in particular most operations on perm gen
     // objects. However, we only ever bias Java instances and all of
     // the call sites of identity_hash that might revoke biases have
     // been checked to make sure they can handle a safepoint. The
     // added check of the bias pattern is to avoid useless calls to
     // thread-local storage.
     if (obj->mark()->has_bias_pattern()) {
       // Box and unbox the raw reference just in case we cause a STW safepoint.
       Handle hobj (Self, obj) ;
       // Relaxing assertion for bug 6320749.
       assert (Universe::verify_in_progress() ||
               !SafepointSynchronize::is_at_safepoint(),
              "biases should not be seen by VM thread here");
       BiasedLocking::revoke_and_rebias(hobj, false, JavaThread::current());
       obj = hobj() ;
       assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
     }
   }
   // hashCode() is a heap mutator ...
   // Relaxing assertion for bug 6320749.
   assert (Universe::verify_in_progress() ||
           !SafepointSynchronize::is_at_safepoint(), "invariant") ;
   assert (Universe::verify_in_progress() ||
           Self->is_Java_thread() , "invariant") ;
   assert (Universe::verify_in_progress() ||
          ((JavaThread *)Self)->thread_state() != _thread_blocked, "invariant") ;
   ObjectMonitor* monitor = NULL;
   markOop temp, test;
   intptr_t hash;
   markOop mark = ReadStableMark (obj);
   // object should remain ineligible for biased locking
   assert (!mark->has_bias_pattern(), "invariant") ;
   if (mark->is_neutral()) {
     hash = mark->hash();              // this is a normal header
     if (hash) {                       // if it has hash, just return it
       return hash;
     }
     hash = get_next_hash(Self, obj);  // allocate a new hash code
     temp = mark->copy_set_hash(hash); // merge the hash code into header
     // use (machine word version) atomic operation to install the hash
     test = (markOop) Atomic::cmpxchg_ptr(temp, obj->mark_addr(), mark);
     if (test == mark) {
       return hash;
     }
     // If atomic operation failed, we must inflate the header
     // into heavy weight monitor. We could add more code here
     // for fast path, but it does not worth the complexity.
   } else if (mark->has_monitor()) {
     monitor = mark->monitor();
     temp = monitor->header();
     assert (temp->is_neutral(), "invariant") ;
     hash = temp->hash();
     if (hash) {
       return hash;
     }
     // Skip to the following code to reduce code size
   } else if (Self->is_lock_owned((address)mark->locker())) {
     temp = mark->displaced_mark_helper(); // this is a lightweight monitor owned
     assert (temp->is_neutral(), "invariant") ;
     hash = temp->hash();              // by current thread, check if the displaced
     if (hash) {                       // header contains hash code
       return hash;
     }
     // WARNING:
     //   The displaced header is strictly immutable.
     // It can NOT be changed in ANY cases. So we have
     // to inflate the header into heavyweight monitor
     // even the current thread owns the lock. The reason
     // is the BasicLock (stack slot) will be asynchronously
     // read by other threads during the inflate() function.
     // Any change to stack may not propagate to other threads
     // correctly.
   }
   // Inflate the monitor to set hash code
   monitor = ObjectSynchronizer::inflate(Self, obj);
   // Load displaced header and check it has hash code
   mark = monitor->header();
   assert (mark->is_neutral(), "invariant") ;
   hash = mark->hash();
   if (hash == 0) {
     hash = get_next_hash(Self, obj);
     temp = mark->copy_set_hash(hash); // merge hash code into header
     assert (temp->is_neutral(), "invariant") ;
     test = (markOop) Atomic::cmpxchg_ptr(temp, monitor, mark);
     if (test != mark) {
       // The only update to the header in the monitor (outside GC)
       // is install the hash code. If someone add new usage of
       // displaced header, please update this code
       hash = test->hash();
       assert (test->is_neutral(), "invariant") ;
       assert (hash != 0, "Trivial unexpected object/monitor header usage.");
     }
   }
   // We finally get the hash
   return hash;
 }
 // Deprecated -- use FastHashCode() instead.
 intptr_t ObjectSynchronizer::identity_hash_value_for(Handle obj) {
   return FastHashCode (Thread::current(), obj()) ;
 }
 bool ObjectSynchronizer::current_thread_holds_lock(JavaThread* thread,
                                                    Handle h_obj) {
   if (UseBiasedLocking) {
     BiasedLocking::revoke_and_rebias(h_obj, false, thread);
     assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   assert(thread == JavaThread::current(), "Can only be called on current thread");
   oop obj = h_obj();
   markOop mark = ReadStableMark (obj) ;
   // Uncontended case, header points to stack
   if (mark->has_locker()) {
     return thread->is_lock_owned((address)mark->locker());
   }
   // Contended case, header points to ObjectMonitor (tagged pointer)
   if (mark->has_monitor()) {
     ObjectMonitor* monitor = mark->monitor();
     return monitor->is_entered(thread) != 0 ;
   }
   // Unlocked case, header in place
   assert(mark->is_neutral(), "sanity check");
   return false;
 }
 // Be aware of this method could revoke bias of the lock object.
 // This method querys the ownership of the lock handle specified by 'h_obj'.
 // If the current thread owns the lock, it returns owner_self. If no
 // thread owns the lock, it returns owner_none. Otherwise, it will return
 // ower_other.
 ObjectSynchronizer::LockOwnership ObjectSynchronizer::query_lock_ownership
 (JavaThread *self, Handle h_obj) {
   // The caller must beware this method can revoke bias, and
   // revocation can result in a safepoint.
   assert (!SafepointSynchronize::is_at_safepoint(), "invariant") ;
   assert (self->thread_state() != _thread_blocked , "invariant") ;
   // Possible mark states: neutral, biased, stack-locked, inflated
   if (UseBiasedLocking && h_obj()->mark()->has_bias_pattern()) {
     // CASE: biased
     BiasedLocking::revoke_and_rebias(h_obj, false, self);
     assert(!h_obj->mark()->has_bias_pattern(),
            "biases should be revoked by now");
   }
   assert(self == JavaThread::current(), "Can only be called on current thread");
   oop obj = h_obj();
   markOop mark = ReadStableMark (obj) ;
   // CASE: stack-locked.  Mark points to a BasicLock on the owner's stack.
   if (mark->has_locker()) {
     return self->is_lock_owned((address)mark->locker()) ?
       owner_self : owner_other;
   }
   // CASE: inflated. Mark (tagged pointer) points to an objectMonitor.
   // The Object:ObjectMonitor relationship is stable as long as we're
   // not at a safepoint.
   if (mark->has_monitor()) {
     void * owner = mark->monitor()->_owner ;
     if (owner == NULL) return owner_none ;
     return (owner == self ||
             self->is_lock_owned((address)owner)) ? owner_self : owner_other;
   }
   // CASE: neutral
   assert(mark->is_neutral(), "sanity check");
   return owner_none ;           // it's unlocked
 }
 // FIXME: jvmti should call this
 JavaThread* ObjectSynchronizer::get_lock_owner(Handle h_obj, bool doLock) {
   if (UseBiasedLocking) {
     if (SafepointSynchronize::is_at_safepoint()) {
       BiasedLocking::revoke_at_safepoint(h_obj);
     } else {
       BiasedLocking::revoke_and_rebias(h_obj, false, JavaThread::current());
     }
     assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now");
   }
   oop obj = h_obj();
   address owner = NULL;
   markOop mark = ReadStableMark (obj) ;
   // Uncontended case, header points to stack
   if (mark->has_locker()) {
     owner = (address) mark->locker();
   }
   // Contended case, header points to ObjectMonitor (tagged pointer)
   if (mark->has_monitor()) {
     ObjectMonitor* monitor = mark->monitor();
     assert(monitor != NULL, "monitor should be non-null");
     owner = (address) monitor->owner();
   }
   if (owner != NULL) {
     return Threads::owning_thread_from_monitor_owner(owner, doLock);
   }
   // Unlocked case, header in place
   // Cannot have assertion since this object may have been
   // locked by another thread when reaching here.
   // assert(mark->is_neutral(), "sanity check");
   return NULL;
 }
 // Iterate through monitor cache and attempt to release thread's monitors
 // Gives up on a particular monitor if an exception occurs, but continues
 // the overall iteration, swallowing the exception.
 class ReleaseJavaMonitorsClosure: public MonitorClosure {
 private:
   TRAPS;
 public:
   ReleaseJavaMonitorsClosure(Thread* thread) : THREAD(thread) {}
   void do_monitor(ObjectMonitor* mid) {
     if (mid->owner() == THREAD) {
       (void)mid->complete_exit(CHECK);
     }
   }
 };
 // Release all inflated monitors owned by THREAD.  Lightweight monitors are
 // ignored.  This is meant to be called during JNI thread detach which assumes
 // all remaining monitors are heavyweight.  All exceptions are swallowed.
 // Scanning the extant monitor list can be time consuming.
 // A simple optimization is to add a per-thread flag that indicates a thread
 // called jni_monitorenter() during its lifetime.
 //
 // Instead of No_Savepoint_Verifier it might be cheaper to
 // use an idiom of the form:
 //   auto int tmp = SafepointSynchronize::_safepoint_counter ;
 //   <code that must not run at safepoint>
 //   guarantee (((tmp ^ _safepoint_counter) | (tmp & 1)) == 0) ;
 // Since the tests are extremely cheap we could leave them enabled
 // for normal product builds.
 void ObjectSynchronizer::release_monitors_owned_by_thread(TRAPS) {
   assert(THREAD == JavaThread::current(), "must be current Java thread");
   No_Safepoint_Verifier nsv ;
   ReleaseJavaMonitorsClosure rjmc(THREAD);
   Thread::muxAcquire(&ListLock, "release_monitors_owned_by_thread");
   ObjectSynchronizer::monitors_iterate(&rjmc);
   Thread::muxRelease(&ListLock);
   THREAD->clear_pending_exception();
 }
 // Visitors ...
 void ObjectSynchronizer::monitors_iterate(MonitorClosure* closure) {
   ObjectMonitor* block = gBlockList;
   ObjectMonitor* mid;
   while (block) {
     assert(block->object() == CHAINMARKER, "must be a block header");
     for (int i = _BLOCKSIZE - 1; i > 0; i--) {
       mid = block + i;
       oop object = (oop) mid->object();
       if (object != NULL) {
         closure->do_monitor(mid);
       }
     }
     block = (ObjectMonitor*) block->FreeNext;
   }
 }
 void ObjectSynchronizer::oops_do(OopClosure* f) {
   assert(SafepointSynchronize::is_at_safepoint(), "must be at safepoint");
   for (ObjectMonitor* block = gBlockList; block != NULL; block = next(block)) {
     assert(block->object() == CHAINMARKER, "must be a block header");
     for (int i = 1; i < _BLOCKSIZE; i++) {
       ObjectMonitor* mid = &block[i];
       if (mid->object() != NULL) {
         f->do_oop((oop*)mid->object_addr());
       }
     }
   }
 }
 // Deflate_idle_monitors() is called at all safepoints, immediately
 // after all mutators are stopped, but before any objects have moved.
 // It traverses the list of known monitors, deflating where possible.
 // The scavenged monitor are returned to the monitor free list.
 //
 // Beware that we scavenge at *every* stop-the-world point.
 // Having a large number of monitors in-circulation negatively
 // impacts the performance of some applications (e.g., PointBase).
 // Broadly, we want to minimize the # of monitors in circulation.
 // Alternately, we could partition the active monitors into sub-lists
 // of those that need scanning and those that do not.
 // Specifically, we would add a new sub-list of objectmonitors
 // that are in-circulation and potentially active.  deflate_idle_monitors()
 // would scan only that list.  Other monitors could reside on a quiescent
 // list.  Such sequestered monitors wouldn't need to be scanned by
 // deflate_idle_monitors().  omAlloc() would first check the global free list,
 // then the quiescent list, and, failing those, would allocate a new block.
 // Deflate_idle_monitors() would scavenge and move monitors to the
 // quiescent list.
 //
 // Perversely, the heap size -- and thus the STW safepoint rate --
 // typically drives the scavenge rate.  Large heaps can mean infrequent GC,
 // which in turn can mean large(r) numbers of objectmonitors in circulation.
 // This is an unfortunate aspect of this design.
 //
 // Another refinement would be to refrain from calling deflate_idle_monitors()
 // except at stop-the-world points associated with garbage collections.
 //
 // An even better solution would be to deflate on-the-fly, aggressively,
 // at monitorexit-time as is done in EVM's metalock or Relaxed Locks.
 void ObjectSynchronizer::deflate_idle_monitors() {
   assert(SafepointSynchronize::is_at_safepoint(), "must be at safepoint");
   int nInuse = 0 ;              // currently associated with objects
   int nInCirculation = 0 ;      // extant
   int nScavenged = 0 ;          // reclaimed
   ObjectMonitor * FreeHead = NULL ;  // Local SLL of scavenged monitors
   ObjectMonitor * FreeTail = NULL ;
   // Iterate over all extant monitors - Scavenge all idle monitors.
   TEVENT (deflate_idle_monitors) ;
   for (ObjectMonitor* block = gBlockList; block != NULL; block = next(block)) {
     assert(block->object() == CHAINMARKER, "must be a block header");
     nInCirculation += _BLOCKSIZE ;
     for (int i = 1 ; i < _BLOCKSIZE; i++) {
       ObjectMonitor* mid = &block[i];
       oop obj = (oop) mid->object();
       if (obj == NULL) {
         // The monitor is not associated with an object.
         // The monitor should either be a thread-specific private
         // free list or the global free list.
         // obj == NULL IMPLIES mid->is_busy() == 0
         guarantee (!mid->is_busy(), "invariant") ;
         continue ;
       }
       // Normal case ... The monitor is associated with obj.
       guarantee (obj->mark() == markOopDesc::encode(mid), "invariant") ;
       guarantee (mid == obj->mark()->monitor(), "invariant");
       guarantee (mid->header()->is_neutral(), "invariant");
       if (mid->is_busy()) {
          if (ClearResponsibleAtSTW) mid->_Responsible = NULL ;
          nInuse ++ ;
       } else {
          // Deflate the monitor if it is no longer being used
          // It's idle - scavenge and return to the global free list
          // plain old deflation ...
          TEVENT (deflate_idle_monitors - scavenge1) ;
          if (TraceMonitorInflation) {
            if (obj->is_instance()) {
              ResourceMark rm;
                tty->print_cr("Deflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s",
                     (intptr_t) obj, (intptr_t) obj->mark(), Klass::cast(obj->klass())->external_name());
            }
          }
          // Restore the header back to obj
          obj->release_set_mark(mid->header());
          mid->clear();
          assert (mid->object() == NULL, "invariant") ;
          // Move the object to the working free list defined by FreeHead,FreeTail.
          mid->FreeNext = NULL ;
          if (FreeHead == NULL) FreeHead = mid ;
          if (FreeTail != NULL) FreeTail->FreeNext = mid ;
          FreeTail = mid ;
          nScavenged ++ ;
       }
     }
   }
   // Move the scavenged monitors back to the global free list.
   // In theory we don't need the freelist lock as we're at a STW safepoint.
   // omAlloc() and omFree() can only be called while a thread is _not in safepoint state.
   // But it's remotely possible that omFlush() or release_monitors_owned_by_thread()
   // might be called while not at a global STW safepoint.  In the interest of
   // safety we protect the following access with ListLock.
   // An even more conservative and prudent approach would be to guard
   // the main loop in scavenge_idle_monitors() with ListLock.
   if (FreeHead != NULL) {
      guarantee (FreeTail != NULL && nScavenged > 0, "invariant") ;
      assert (FreeTail->FreeNext == NULL, "invariant") ;
      // constant-time list splice - prepend scavenged segment to gFreeList
      Thread::muxAcquire (&ListLock, "scavenge - return") ;
      FreeTail->FreeNext = gFreeList ;
      gFreeList = FreeHead ;
      Thread::muxRelease (&ListLock) ;
   }
   if (_sync_Deflations != NULL) _sync_Deflations->inc(nScavenged) ;
   if (_sync_MonExtant  != NULL) _sync_MonExtant ->set_value(nInCirculation);
   // TODO: Add objectMonitor leak detection.
   // Audit/inventory the objectMonitors -- make sure they're all accounted for.
   GVars.stwRandom = os::random() ;
   GVars.stwCycle ++ ;
 }
 // A macro is used below because there may already be a pending
 // exception which should not abort the execution of the routines
 // which use this (which is why we don't put this into check_slow and
 // call it with a CHECK argument).
 #define CHECK_OWNER()                                                             \
   do {                                                                            \
     if (THREAD != _owner) {                                                       \
       if (THREAD->is_lock_owned((address) _owner)) {                              \
         _owner = THREAD ;  /* Convert from basiclock addr to Thread addr */       \
         _recursions = 0;                                                          \
         OwnerIsThread = 1 ;                                                       \
       } else {                                                                    \
         TEVENT (Throw IMSX) ;                                                     \
         THROW(vmSymbols::java_lang_IllegalMonitorStateException());               \
       }                                                                           \
     }                                                                             \
   } while (false)
 // TODO-FIXME: eliminate ObjectWaiters.  Replace this visitor/enumerator
 // interface with a simple FirstWaitingThread(), NextWaitingThread() interface.
 ObjectWaiter* ObjectMonitor::first_waiter() {
   return _WaitSet;
 }
 ObjectWaiter* ObjectMonitor::next_waiter(ObjectWaiter* o) {
   return o->_next;
 }
 Thread* ObjectMonitor::thread_of_waiter(ObjectWaiter* o) {
   return o->_thread;
 }
 // initialize the monitor, exception the semaphore, all other fields
 // are simple integers or pointers
 ObjectMonitor::ObjectMonitor() {
   _header       = NULL;
   _count        = 0;
   _waiters      = 0,
   _recursions   = 0;
   _object       = NULL;
   _owner        = NULL;
   _WaitSet      = NULL;
   _WaitSetLock  = 0 ;
   _Responsible  = NULL ;
   _succ         = NULL ;
   _cxq          = NULL ;
   FreeNext      = NULL ;
   _EntryList    = NULL ;
   _SpinFreq     = 0 ;
   _SpinClock    = 0 ;
   OwnerIsThread = 0 ;
 }
 ObjectMonitor::~ObjectMonitor() {
    // TODO: Add asserts ...
    // _cxq == 0 _succ == NULL _owner == NULL _waiters == 0
    // _count == 0 _EntryList  == NULL etc
 }
 intptr_t ObjectMonitor::is_busy() const {
   // TODO-FIXME: merge _count and _waiters.
   // TODO-FIXME: assert _owner == null implies _recursions = 0
   // TODO-FIXME: assert _WaitSet != null implies _count > 0
   return _count|_waiters|intptr_t(_owner)|intptr_t(_cxq)|intptr_t(_EntryList ) ;
 }
 void ObjectMonitor::Recycle () {
   // TODO: add stronger asserts ...
   // _cxq == 0 _succ == NULL _owner == NULL _waiters == 0
   // _count == 0 EntryList  == NULL
   // _recursions == 0 _WaitSet == NULL
   // TODO: assert (is_busy()|_recursions) == 0
   _succ          = NULL ;
   _EntryList     = NULL ;
   _cxq           = NULL ;
   _WaitSet       = NULL ;
   _recursions    = 0 ;
   _SpinFreq      = 0 ;
   _SpinClock     = 0 ;
   OwnerIsThread  = 0 ;
 }
 // WaitSet management ...
 inline void ObjectMonitor::AddWaiter(ObjectWaiter* node) {
   assert(node != NULL, "should not dequeue NULL node");
   assert(node->_prev == NULL, "node already in list");
   assert(node->_next == NULL, "node already in list");
   // put node at end of queue (circular doubly linked list)
   if (_WaitSet == NULL) {
     _WaitSet = node;
     node->_prev = node;
     node->_next = node;
   } else {
     ObjectWaiter* head = _WaitSet ;
     ObjectWaiter* tail = head->_prev;
     assert(tail->_next == head, "invariant check");
     tail->_next = node;
     head->_prev = node;
     node->_next = head;
     node->_prev = tail;
   }
 }
 inline ObjectWaiter* ObjectMonitor::DequeueWaiter() {
   // dequeue the very first waiter
   ObjectWaiter* waiter = _WaitSet;
   if (waiter) {
     DequeueSpecificWaiter(waiter);
   }
   return waiter;
 }
 inline void ObjectMonitor::DequeueSpecificWaiter(ObjectWaiter* node) {
   assert(node != NULL, "should not dequeue NULL node");
   assert(node->_prev != NULL, "node already removed from list");
   assert(node->_next != NULL, "node already removed from list");
   // when the waiter has woken up because of interrupt,
   // timeout or other spurious wake-up, dequeue the
   // waiter from waiting list
   ObjectWaiter* next = node->_next;
   if (next == node) {
     assert(node->_prev == node, "invariant check");
     _WaitSet = NULL;
   } else {
     ObjectWaiter* prev = node->_prev;
     assert(prev->_next == node, "invariant check");
     assert(next->_prev == node, "invariant check");
     next->_prev = prev;
     prev->_next = next;
     if (_WaitSet == node) {
       _WaitSet = next;
     }
   }
   node->_next = NULL;
   node->_prev = NULL;
 }
 static char * kvGet (char * kvList, const char * Key) {
     if (kvList == NULL) return NULL ;
     size_t n = strlen (Key) ;
     char * Search ;
     for (Search = kvList ; *Search ; Search += strlen(Search) + 1) {
         if (strncmp (Search, Key, n) == 0) {
             if (Search[n] == '=') return Search + n + 1 ;
             if (Search[n] == 0)   return (char *) "1" ;
         }
     }
     return NULL ;
 }
 static int kvGetInt (char * kvList, const char * Key, int Default) {
     char * v = kvGet (kvList, Key) ;
     int rslt = v ? ::strtol (v, NULL, 0) : Default ;
     if (Knob_ReportSettings && v != NULL) {
         ::printf ("  SyncKnob: %s %d(%d)\n", Key, rslt, Default) ;
         ::fflush (stdout) ;
     }
     return rslt ;
 }
 // By convention we unlink a contending thread from EntryList|cxq immediately
 // after the thread acquires the lock in ::enter().  Equally, we could defer
 // unlinking the thread until ::exit()-time.
 void ObjectMonitor::UnlinkAfterAcquire (Thread * Self, ObjectWaiter * SelfNode)
 {
     assert (_owner == Self, "invariant") ;
     assert (SelfNode->_thread == Self, "invariant") ;
     if (SelfNode->TState == ObjectWaiter::TS_ENTER) {
         // Normal case: remove Self from the DLL EntryList .
         // This is a constant-time operation.
         ObjectWaiter * nxt = SelfNode->_next ;
         ObjectWaiter * prv = SelfNode->_prev ;
         if (nxt != NULL) nxt->_prev = prv ;
         if (prv != NULL) prv->_next = nxt ;
         if (SelfNode == _EntryList ) _EntryList = nxt ;
         assert (nxt == NULL || nxt->TState == ObjectWaiter::TS_ENTER, "invariant") ;
         assert (prv == NULL || prv->TState == ObjectWaiter::TS_ENTER, "invariant") ;
         TEVENT (Unlink from EntryList) ;
     } else {
         guarantee (SelfNode->TState == ObjectWaiter::TS_CXQ, "invariant") ;
         // Inopportune interleaving -- Self is still on the cxq.
         // This usually means the enqueue of self raced an exiting thread.
         // Normally we'll find Self near the front of the cxq, so
         // dequeueing is typically fast.  If needbe we can accelerate
         // this with some MCS/CHL-like bidirectional list hints and advisory
         // back-links so dequeueing from the interior will normally operate
         // in constant-time.
         // Dequeue Self from either the head (with CAS) or from the interior
         // with a linear-time scan and normal non-atomic memory operations.
         // CONSIDER: if Self is on the cxq then simply drain cxq into EntryList
         // and then unlink Self from EntryList.  We have to drain eventually,
         // so it might as well be now.
         ObjectWaiter * v = _cxq ;
         assert (v != NULL, "invariant") ;
         if (v != SelfNode || Atomic::cmpxchg_ptr (SelfNode->_next, &_cxq, v) != v) {
             // The CAS above can fail from interference IFF a "RAT" arrived.
             // In that case Self must be in the interior and can no longer be
             // at the head of cxq.
             if (v == SelfNode) {
                 assert (_cxq != v, "invariant") ;
                 v = _cxq ;          // CAS above failed - start scan at head of list
             }
             ObjectWaiter * p ;
             ObjectWaiter * q = NULL ;
             for (p = v ; p != NULL && p != SelfNode; p = p->_next) {
                 q = p ;
                 assert (p->TState == ObjectWaiter::TS_CXQ, "invariant") ;
             }
             assert (v != SelfNode,  "invariant") ;
             assert (p == SelfNode,  "Node not found on cxq") ;
             assert (p != _cxq,      "invariant") ;
             assert (q != NULL,      "invariant") ;
             assert (q->_next == p,  "invariant") ;
             q->_next = p->_next ;
         }
         TEVENT (Unlink from cxq) ;
     }
     // Diagnostic hygiene ...
     SelfNode->_prev  = (ObjectWaiter *) 0xBAD ;
     SelfNode->_next  = (ObjectWaiter *) 0xBAD ;
     SelfNode->TState = ObjectWaiter::TS_RUN ;
 }
 // Caveat: TryLock() is not necessarily serializing if it returns failure.
 // Callers must compensate as needed.
 int ObjectMonitor::TryLock (Thread * Self) {
    for (;;) {
       void * own = _owner ;
       if (own != NULL) return 0 ;
       if (Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) {
          // Either guarantee _recursions == 0 or set _recursions = 0.
          assert (_recursions == 0, "invariant") ;
          assert (_owner == Self, "invariant") ;
          // CONSIDER: set or assert that OwnerIsThread == 1
          return 1 ;
       }
       // The lock had been free momentarily, but we lost the race to the lock.
       // Interference -- the CAS failed.
       // We can either return -1 or retry.
       // Retry doesn't make as much sense because the lock was just acquired.
       if (true) return -1 ;
    }
 }
 // NotRunnable() -- informed spinning
 //
 // Don't bother spinning if the owner is not eligible to drop the lock.
 // Peek at the owner's schedctl.sc_state and Thread._thread_values and
 // spin only if the owner thread is _thread_in_Java or _thread_in_vm.
 // The thread must be runnable in order to drop the lock in timely fashion.
 // If the _owner is not runnable then spinning will not likely be
 // successful (profitable).
 //
 // Beware -- the thread referenced by _owner could have died
 // so a simply fetch from _owner->_thread_state might trap.
 // Instead, we use SafeFetchXX() to safely LD _owner->_thread_state.
 // Because of the lifecycle issues the schedctl and _thread_state values
 // observed by NotRunnable() might be garbage.  NotRunnable must
 // tolerate this and consider the observed _thread_state value
 // as advisory.
 //
 // Beware too, that _owner is sometimes a BasicLock address and sometimes
 // a thread pointer.  We differentiate the two cases with OwnerIsThread.
 // Alternately, we might tag the type (thread pointer vs basiclock pointer)
 // with the LSB of _owner.  Another option would be to probablistically probe
 // the putative _owner->TypeTag value.
 //
 // Checking _thread_state isn't perfect.  Even if the thread is
 // in_java it might be blocked on a page-fault or have been preempted
 // and sitting on a ready/dispatch queue.  _thread state in conjunction
 // with schedctl.sc_state gives us a good picture of what the
 // thread is doing, however.
 //
 // TODO: check schedctl.sc_state.
 // We'll need to use SafeFetch32() to read from the schedctl block.
 // See RFE #5004247 and http://sac.sfbay.sun.com/Archives/CaseLog/arc/PSARC/2005/351/
 //
 // The return value from NotRunnable() is *advisory* -- the
 // result is based on sampling and is not necessarily coherent.
 // The caller must tolerate false-negative and false-positive errors.
 // Spinning, in general, is probabilistic anyway.
 int ObjectMonitor::NotRunnable (Thread * Self, Thread * ox) {
     // Check either OwnerIsThread or ox->TypeTag == 2BAD.
     if (!OwnerIsThread) return 0 ;
     if (ox == NULL) return 0 ;
     // Avoid transitive spinning ...
     // Say T1 spins or blocks trying to acquire L.  T1._Stalled is set to L.
     // Immediately after T1 acquires L it's possible that T2, also
     // spinning on L, will see L.Owner=T1 and T1._Stalled=L.
     // This occurs transiently after T1 acquired L but before
     // T1 managed to clear T1.Stalled.  T2 does not need to abort
     // its spin in this circumstance.
     intptr_t BlockedOn = SafeFetchN ((intptr_t *) &ox->_Stalled, intptr_t(1)) ;
     if (BlockedOn == 1) return 1 ;
     if (BlockedOn != 0) {
       return BlockedOn != intptr_t(this) && _owner == ox ;
     }
     assert (sizeof(((JavaThread *)ox)->_thread_state == sizeof(int)), "invariant") ;
     int jst = SafeFetch32 ((int *) &((JavaThread *) ox)->_thread_state, -1) ; ;
     // consider also: jst != _thread_in_Java -- but that's overspecific.
     return jst == _thread_blocked || jst == _thread_in_native ;
 }
 // Adaptive spin-then-block - rational spinning
 //
 // Note that we spin "globally" on _owner with a classic SMP-polite TATAS
 // algorithm.  On high order SMP systems it would be better to start with
 // a brief global spin and then revert to spinning locally.  In the spirit of MCS/CLH,
 // a contending thread could enqueue itself on the cxq and then spin locally
 // on a thread-specific variable such as its ParkEvent._Event flag.
 // That's left as an exercise for the reader.  Note that global spinning is
 // not problematic on Niagara, as the L2$ serves the interconnect and has both
 // low latency and massive bandwidth.
 //
 // Broadly, we can fix the spin frequency -- that is, the % of contended lock
 // acquisition attempts where we opt to spin --  at 100% and vary the spin count
 // (duration) or we can fix the count at approximately the duration of
 // a context switch and vary the frequency.   Of course we could also
 // vary both satisfying K == Frequency * Duration, where K is adaptive by monitor.
 // See http://j2se.east/~dice/PERSIST/040824-AdaptiveSpinning.html.
 //
 // This implementation varies the duration "D", where D varies with
 // the success rate of recent spin attempts. (D is capped at approximately
 // length of a round-trip context switch).  The success rate for recent
 // spin attempts is a good predictor of the success rate of future spin
 // attempts.  The mechanism adapts automatically to varying critical
 // section length (lock modality), system load and degree of parallelism.
 // D is maintained per-monitor in _SpinDuration and is initialized
 // optimistically.  Spin frequency is fixed at 100%.
 //
 // Note that _SpinDuration is volatile, but we update it without locks
 // or atomics.  The code is designed so that _SpinDuration stays within
 // a reasonable range even in the presence of races.  The arithmetic
 // operations on _SpinDuration are closed over the domain of legal values,
 // so at worst a race will install and older but still legal value.
 // At the very worst this introduces some apparent non-determinism.
 // We might spin when we shouldn't or vice-versa, but since the spin
 // count are relatively short, even in the worst case, the effect is harmless.
 //
 // Care must be taken that a low "D" value does not become an
 // an absorbing state.  Transient spinning failures -- when spinning
 // is overall profitable -- should not cause the system to converge
 // on low "D" values.  We want spinning to be stable and predictable
 // and fairly responsive to change and at the same time we don't want
 // it to oscillate, become metastable, be "too" non-deterministic,
 // or converge on or enter undesirable stable absorbing states.
 //
 // We implement a feedback-based control system -- using past behavior
 // to predict future behavior.  We face two issues: (a) if the
 // input signal is random then the spin predictor won't provide optimal
 // results, and (b) if the signal frequency is too high then the control
 // system, which has some natural response lag, will "chase" the signal.
 // (b) can arise from multimodal lock hold times.  Transient preemption
 // can also result in apparent bimodal lock hold times.
 // Although sub-optimal, neither condition is particularly harmful, as
 // in the worst-case we'll spin when we shouldn't or vice-versa.
 // The maximum spin duration is rather short so the failure modes aren't bad.
 // To be conservative, I've tuned the gain in system to bias toward
 // _not spinning.  Relatedly, the system can sometimes enter a mode where it
 // "rings" or oscillates between spinning and not spinning.  This happens
 // when spinning is just on the cusp of profitability, however, so the
 // situation is not dire.  The state is benign -- there's no need to add
 // hysteresis control to damp the transition rate between spinning and
 // not spinning.
 //
 // - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - -
 //
 // Spin-then-block strategies ...
 //
 // Thoughts on ways to improve spinning :
 //
 // *  Periodically call {psr_}getloadavg() while spinning, and
 //    permit unbounded spinning if the load average is <
 //    the number of processors.  Beware, however, that getloadavg()
 //    is exceptionally fast on solaris (about 1/10 the cost of a full
 //    spin cycle, but quite expensive on linux.  Beware also, that
 //    multiple JVMs could "ring" or oscillate in a feedback loop.
 //    Sufficient damping would solve that problem.
 //
 // *  We currently use spin loops with iteration counters to approximate
 //    spinning for some interval.  Given the availability of high-precision
 //    time sources such as gethrtime(), %TICK, %STICK, RDTSC, etc., we should
 //    someday reimplement the spin loops to duration-based instead of iteration-based.
 //
 // *  Don't spin if there are more than N = (CPUs/2) threads
 //        currently spinning on the monitor (or globally).
 //    That is, limit the number of concurrent spinners.
 //    We might also limit the # of spinners in the JVM, globally.
 //
 // *  If a spinning thread observes _owner change hands it should
 //    abort the spin (and park immediately) or at least debit
 //    the spin counter by a large "penalty".
 //
 // *  Classically, the spin count is either K*(CPUs-1) or is a
 //        simple constant that approximates the length of a context switch.
 //    We currently use a value -- computed by a special utility -- that
 //    approximates round-trip context switch times.
 //
 // *  Normally schedctl_start()/_stop() is used to advise the kernel
 //    to avoid preempting threads that are running in short, bounded
 //    critical sections.  We could use the schedctl hooks in an inverted
 //    sense -- spinners would set the nopreempt flag, but poll the preempt
 //    pending flag.  If a spinner observed a pending preemption it'd immediately
 //    abort the spin and park.   As such, the schedctl service acts as
 //    a preemption warning mechanism.
 //
 // *  In lieu of spinning, if the system is running below saturation
 //    (that is, loadavg() << #cpus), we can instead suppress futile
 //    wakeup throttling, or even wake more than one successor at exit-time.
 //    The net effect is largely equivalent to spinning.  In both cases,
 //    contending threads go ONPROC and opportunistically attempt to acquire
 //    the lock, decreasing lock handover latency at the expense of wasted
 //    cycles and context switching.
 //
 // *  We might to spin less after we've parked as the thread will
 //    have less $ and TLB affinity with the processor.
 //    Likewise, we might spin less if we come ONPROC on a different
 //    processor or after a long period (>> rechose_interval).
 //
 // *  A table-driven state machine similar to Solaris' dispadmin scheduling
 //    tables might be a better design.  Instead of encoding information in
 //    _SpinDuration, _SpinFreq and _SpinClock we'd just use explicit,
 //    discrete states.   Success or failure during a spin would drive
 //    state transitions, and each state node would contain a spin count.
 //
 // *  If the processor is operating in a mode intended to conserve power
 //    (such as Intel's SpeedStep) or to reduce thermal output (thermal
 //    step-down mode) then the Java synchronization subsystem should
 //    forgo spinning.
 //
 // *  The minimum spin duration should be approximately the worst-case
 //    store propagation latency on the platform.  That is, the time
 //    it takes a store on CPU A to become visible on CPU B, where A and
 //    B are "distant".
 //
 // *  We might want to factor a thread's priority in the spin policy.
 //    Threads with a higher priority might spin for slightly longer.
 //    Similarly, if we use back-off in the TATAS loop, lower priority
 //    threads might back-off longer.  We don't currently use a
 //    thread's priority when placing it on the entry queue.  We may
 //    want to consider doing so in future releases.
 //
 // *  We might transiently drop a thread's scheduling priority while it spins.
 //    SCHED_BATCH on linux and FX scheduling class at priority=0 on Solaris
 //    would suffice.  We could even consider letting the thread spin indefinitely at
 //    a depressed or "idle" priority.  This brings up fairness issues, however --
 //    in a saturated system a thread would with a reduced priority could languish
 //    for extended periods on the ready queue.
 //
 // *  While spinning try to use the otherwise wasted time to help the VM make
 //    progress:
 //
 //    -- YieldTo() the owner, if the owner is OFFPROC but ready
 //       Done our remaining quantum directly to the ready thread.
 //       This helps "push" the lock owner through the critical section.
 //       It also tends to improve affinity/locality as the lock
 //       "migrates" less frequently between CPUs.
 //    -- Walk our own stack in anticipation of blocking.  Memoize the roots.
 //    -- Perform strand checking for other thread.  Unpark potential strandees.
 //    -- Help GC: trace or mark -- this would need to be a bounded unit of work.
 //       Unfortunately this will pollute our $ and TLBs.  Recall that we
 //       spin to avoid context switching -- context switching has an
 //       immediate cost in latency, a disruptive cost to other strands on a CMT
 //       processor, and an amortized cost because of the D$ and TLB cache
 //       reload transient when the thread comes back ONPROC and repopulates
 //       $s and TLBs.
 //    -- call getloadavg() to see if the system is saturated.  It'd probably
 //       make sense to call getloadavg() half way through the spin.
 //       If the system isn't at full capacity the we'd simply reset
 //       the spin counter to and extend the spin attempt.
 //    -- Doug points out that we should use the same "helping" policy
 //       in thread.yield().
 //
 // *  Try MONITOR-MWAIT on systems that support those instructions.
 //
 // *  The spin statistics that drive spin decisions & frequency are
 //    maintained in the objectmonitor structure so if we deflate and reinflate
 //    we lose spin state.  In practice this is not usually a concern
 //    as the default spin state after inflation is aggressive (optimistic)
 //    and tends toward spinning.  So in the worst case for a lock where
 //    spinning is not profitable we may spin unnecessarily for a brief
 //    period.  But then again, if a lock is contended it'll tend not to deflate
 //    in the first place.
 intptr_t ObjectMonitor::SpinCallbackArgument = 0 ;
 int (*ObjectMonitor::SpinCallbackFunction)(intptr_t, int) = NULL ;
 // Spinning: Fixed frequency (100%), vary duration
 int ObjectMonitor::TrySpin_VaryDuration (Thread * Self) {
     // Dumb, brutal spin.  Good for comparative measurements against adaptive spinning.
     int ctr = Knob_FixedSpin ;
     if (ctr != 0) {
         while (--ctr >= 0) {
             if (TryLock (Self) > 0) return 1 ;
             SpinPause () ;
         }
         return 0 ;
     }
     for (ctr = Knob_PreSpin + 1; --ctr >= 0 ; ) {
       if (TryLock(Self) > 0) {
         // Increase _SpinDuration ...
         // Note that we don't clamp SpinDuration precisely at SpinLimit.
         // Raising _SpurDuration to the poverty line is key.
         int x = _SpinDuration ;
         if (x < Knob_SpinLimit) {
            if (x < Knob_Poverty) x = Knob_Poverty ;
            _SpinDuration = x + Knob_BonusB ;
         }
         return 1 ;
       }
       SpinPause () ;
     }
     // Admission control - verify preconditions for spinning
     //
     // We always spin a little bit, just to prevent _SpinDuration == 0 from
     // becoming an absorbing state.  Put another way, we spin briefly to
     // sample, just in case the system load, parallelism, contention, or lock
     // modality changed.
     //
     // Consider the following alternative:
     // Periodically set _SpinDuration = _SpinLimit and try a long/full
     // spin attempt.  "Periodically" might mean after a tally of
     // the # of failed spin attempts (or iterations) reaches some threshold.
     // This takes us into the realm of 1-out-of-N spinning, where we
     // hold the duration constant but vary the frequency.
     ctr = _SpinDuration  ;
     if (ctr < Knob_SpinBase) ctr = Knob_SpinBase ;
     if (ctr <= 0) return 0 ;
     if (Knob_SuccRestrict && _succ != NULL) return 0 ;
     if (Knob_OState && NotRunnable (Self, (Thread *) _owner)) {
        TEVENT (Spin abort - notrunnable [TOP]);
        return 0 ;
     }
     int MaxSpin = Knob_MaxSpinners ;
     if (MaxSpin >= 0) {
        if (_Spinner > MaxSpin) {
           TEVENT (Spin abort -- too many spinners) ;
           return 0 ;
        }
        // Slighty racy, but benign ...
        Adjust (&_Spinner, 1) ;
     }
     // We're good to spin ... spin ingress.
     // CONSIDER: use Prefetch::write() to avoid RTS->RTO upgrades
     // when preparing to LD...CAS _owner, etc and the CAS is likely
     // to succeed.
     int hits    = 0 ;
     int msk     = 0 ;
     int caspty  = Knob_CASPenalty ;
     int oxpty   = Knob_OXPenalty ;
     int sss     = Knob_SpinSetSucc ;
     if (sss && _succ == NULL ) _succ = Self ;
     Thread * prv = NULL ;
     // There are three ways to exit the following loop:
     // 1.  A successful spin where this thread has acquired the lock.
     // 2.  Spin failure with prejudice
     // 3.  Spin failure without prejudice
     while (--ctr >= 0) {
       // Periodic polling -- Check for pending GC
       // Threads may spin while they're unsafe.
       // We don't want spinning threads to delay the JVM from reaching
       // a stop-the-world safepoint or to steal cycles from GC.
       // If we detect a pending safepoint we abort in order that
       // (a) this thread, if unsafe, doesn't delay the safepoint, and (b)
       // this thread, if safe, doesn't steal cycles from GC.
       // This is in keeping with the "no loitering in runtime" rule.
       // We periodically check to see if there's a safepoint pending.
       if ((ctr & 0xFF) == 0) {
          if (SafepointSynchronize::do_call_back()) {
             TEVENT (Spin: safepoint) ;
             goto Abort ;           // abrupt spin egress
          }
          if (Knob_UsePause & 1) SpinPause () ;
          int (*scb)(intptr_t,int) = SpinCallbackFunction ;
          if (hits > 50 && scb != NULL) {
             int abend = (*scb)(SpinCallbackArgument, 0) ;
          }
       }
       if (Knob_UsePause & 2) SpinPause() ;
       // Exponential back-off ...  Stay off the bus to reduce coherency traffic.
       // This is useful on classic SMP systems, but is of less utility on
       // N1-style CMT platforms.
       //
       // Trade-off: lock acquisition latency vs coherency bandwidth.
       // Lock hold times are typically short.  A histogram
       // of successful spin attempts shows that we usually acquire
       // the lock early in the spin.  That suggests we want to
       // sample _owner frequently in the early phase of the spin,
       // but then back-off and sample less frequently as the spin
       // progresses.  The back-off makes a good citizen on SMP big
       // SMP systems.  Oversampling _owner can consume excessive
       // coherency bandwidth.  Relatedly, if we _oversample _owner we
       // can inadvertently interfere with the the ST m->owner=null.
       // executed by the lock owner.
       if (ctr & msk) continue ;
       ++hits ;
       if ((hits & 0xF) == 0) {
         // The 0xF, above, corresponds to the exponent.
         // Consider: (msk+1)|msk
         msk = ((msk << 2)|3) & BackOffMask ;
       }
       // Probe _owner with TATAS
       // If this thread observes the monitor transition or flicker
       // from locked to unlocked to locked, then the odds that this
       // thread will acquire the lock in this spin attempt go down
       // considerably.  The same argument applies if the CAS fails
       // or if we observe _owner change from one non-null value to
       // another non-null value.   In such cases we might abort
       // the spin without prejudice or apply a "penalty" to the
       // spin count-down variable "ctr", reducing it by 100, say.
       Thread * ox = (Thread *) _owner ;
       if (ox == NULL) {
          ox = (Thread *) Atomic::cmpxchg_ptr (Self, &_owner, NULL) ;
          if (ox == NULL) {
             // The CAS succeeded -- this thread acquired ownership
             // Take care of some bookkeeping to exit spin state.
             if (sss && _succ == Self) {
                _succ = NULL ;
             }
             if (MaxSpin > 0) Adjust (&_Spinner, -1) ;
             // Increase _SpinDuration :
             // The spin was successful (profitable) so we tend toward
             // longer spin attempts in the future.
             // CONSIDER: factor "ctr" into the _SpinDuration adjustment.
             // If we acquired the lock early in the spin cycle it
             // makes sense to increase _SpinDuration proportionally.
             // Note that we don't clamp SpinDuration precisely at SpinLimit.
             int x = _SpinDuration ;
             if (x < Knob_SpinLimit) {
                 if (x < Knob_Poverty) x = Knob_Poverty ;
                 _SpinDuration = x + Knob_Bonus ;
             }
             return 1 ;
          }
          // The CAS failed ... we can take any of the following actions:
          // * penalize: ctr -= Knob_CASPenalty
          // * exit spin with prejudice -- goto Abort;
          // * exit spin without prejudice.
          // * Since CAS is high-latency, retry again immediately.
          prv = ox ;
          TEVENT (Spin: cas failed) ;
          if (caspty == -2) break ;
          if (caspty == -1) goto Abort ;
          ctr -= caspty ;
          continue ;
       }
       // Did lock ownership change hands ?
       if (ox != prv && prv != NULL ) {
           TEVENT (spin: Owner changed)
           if (oxpty == -2) break ;
           if (oxpty == -1) goto Abort ;
           ctr -= oxpty ;
       }
       prv = ox ;
       // Abort the spin if the owner is not executing.
       // The owner must be executing in order to drop the lock.
       // Spinning while the owner is OFFPROC is idiocy.
       // Consider: ctr -= RunnablePenalty ;
       if (Knob_OState && NotRunnable (Self, ox)) {
          TEVENT (Spin abort - notrunnable);
          goto Abort ;
       }
       if (sss && _succ == NULL ) _succ = Self ;
    }
    // Spin failed with prejudice -- reduce _SpinDuration.
    // TODO: Use an AIMD-like policy to adjust _SpinDuration.
    // AIMD is globally stable.
    TEVENT (Spin failure) ;
    {
      int x = _SpinDuration ;
      if (x > 0) {
         // Consider an AIMD scheme like: x -= (x >> 3) + 100
         // This is globally sample and tends to damp the response.
         x -= Knob_Penalty ;
         if (x < 0) x = 0 ;
         _SpinDuration = x ;
      }
    }
  Abort:
    if (MaxSpin >= 0) Adjust (&_Spinner, -1) ;
    if (sss && _succ == Self) {
       _succ = NULL ;
       // Invariant: after setting succ=null a contending thread
       // must recheck-retry _owner before parking.  This usually happens
       // in the normal usage of TrySpin(), but it's safest
       // to make TrySpin() as foolproof as possible.
       OrderAccess::fence() ;
       if (TryLock(Self) > 0) return 1 ;
    }
    return 0 ;
 }
 #define TrySpin TrySpin_VaryDuration
 static void DeferredInitialize () {
   if (InitDone > 0) return ;
   if (Atomic::cmpxchg (-1, &InitDone, 0) != 0) {
       while (InitDone != 1) ;
       return ;
   }
   // One-shot global initialization ...
   // The initialization is idempotent, so we don't need locks.
   // In the future consider doing this via os::init_2().
   // SyncKnobs consist of <Key>=<Value> pairs in the style
   // of environment variables.  Start by converting ':' to NUL.
   if (SyncKnobs == NULL) SyncKnobs = "" ;
   size_t sz = strlen (SyncKnobs) ;
   char * knobs = (char *) malloc (sz + 2) ;
   if (knobs == NULL) {
      vm_exit_out_of_memory (sz + 2, "Parse SyncKnobs") ;
      guarantee (0, "invariant") ;
   }
   strcpy (knobs, SyncKnobs) ;
   knobs[sz+1] = 0 ;
   for (char * p = knobs ; *p ; p++) {
      if (*p == ':') *p = 0 ;
   }
   #define SETKNOB(x) { Knob_##x = kvGetInt (knobs, #x, Knob_##x); }
   SETKNOB(ReportSettings) ;
   SETKNOB(Verbose) ;
   SETKNOB(FixedSpin) ;
   SETKNOB(SpinLimit) ;
   SETKNOB(SpinBase) ;
   SETKNOB(SpinBackOff);
   SETKNOB(CASPenalty) ;
   SETKNOB(OXPenalty) ;
   SETKNOB(LogSpins) ;
   SETKNOB(SpinSetSucc) ;
   SETKNOB(SuccEnabled) ;
   SETKNOB(SuccRestrict) ;
   SETKNOB(Penalty) ;
   SETKNOB(Bonus) ;
   SETKNOB(BonusB) ;
   SETKNOB(Poverty) ;
   SETKNOB(SpinAfterFutile) ;
   SETKNOB(UsePause) ;
   SETKNOB(SpinEarly) ;
   SETKNOB(OState) ;
   SETKNOB(MaxSpinners) ;
   SETKNOB(PreSpin) ;
   SETKNOB(ExitPolicy) ;
   SETKNOB(QMode);
   SETKNOB(ResetEvent) ;
   SETKNOB(MoveNotifyee) ;
   SETKNOB(FastHSSEC) ;
   #undef SETKNOB
   if (os::is_MP()) {
      BackOffMask = (1 << Knob_SpinBackOff) - 1 ;
      if (Knob_ReportSettings) ::printf ("BackOffMask=%X\n", BackOffMask) ;
      // CONSIDER: BackOffMask = ROUNDUP_NEXT_POWER2 (ncpus-1)
   } else {
      Knob_SpinLimit = 0 ;
      Knob_SpinBase  = 0 ;
      Knob_PreSpin   = 0 ;
      Knob_FixedSpin = -1 ;
   }
   if (Knob_LogSpins == 0) {
      ObjectSynchronizer::_sync_FailedSpins = NULL ;
   }
   free (knobs) ;
   OrderAccess::fence() ;
   InitDone = 1 ;
 }
 // Theory of operations -- Monitors lists, thread residency, etc:
 //
 // * A thread acquires ownership of a monitor by successfully
 //   CAS()ing the _owner field from null to non-null.
 //
 // * Invariant: A thread appears on at most one monitor list --
 //   cxq, EntryList or WaitSet -- at any one time.
 //
 // * Contending threads "push" themselves onto the cxq with CAS
 //   and then spin/park.
 //
 // * After a contending thread eventually acquires the lock it must
 //   dequeue itself from either the EntryList or the cxq.
 //
 // * The exiting thread identifies and unparks an "heir presumptive"
 //   tentative successor thread on the EntryList.  Critically, the
 //   exiting thread doesn't unlink the successor thread from the EntryList.
 //   After having been unparked, the wakee will recontend for ownership of
 //   the monitor.   The successor (wakee) will either acquire the lock or
 //   re-park itself.
 //
 //   Succession is provided for by a policy of competitive handoff.
 //   The exiting thread does _not_ grant or pass ownership to the
 //   successor thread.  (This is also referred to as "handoff" succession").
 //   Instead the exiting thread releases ownership and possibly wakes
 //   a successor, so the successor can (re)compete for ownership of the lock.
 //   If the EntryList is empty but the cxq is populated the exiting
 //   thread will drain the cxq into the EntryList.  It does so by
 //   by detaching the cxq (installing null with CAS) and folding
 //   the threads from the cxq into the EntryList.  The EntryList is
 //   doubly linked, while the cxq is singly linked because of the
 //   CAS-based "push" used to enqueue recently arrived threads (RATs).
 //
 // * Concurrency invariants:
 //
 //   -- only the monitor owner may access or mutate the EntryList.
 //      The mutex property of the monitor itself protects the EntryList
 //      from concurrent interference.
 //   -- Only the monitor owner may detach the cxq.
 //
 // * The monitor entry list operations avoid locks, but strictly speaking
 //   they're not lock-free.  Enter is lock-free, exit is not.
 //   See http://j2se.east/~dice/PERSIST/040825-LockFreeQueues.html
 //
 // * The cxq can have multiple concurrent "pushers" but only one concurrent
 //   detaching thread.  This mechanism is immune from the ABA corruption.
 //   More precisely, the CAS-based "push" onto cxq is ABA-oblivious.
 //
 // * Taken together, the cxq and the EntryList constitute or form a
 //   single logical queue of threads stalled trying to acquire the lock.
 //   We use two distinct lists to improve the odds of a constant-time
 //   dequeue operation after acquisition (in the ::enter() epilog) and
 //   to reduce heat on the list ends.  (c.f. Michael Scott's "2Q" algorithm).
 //   A key desideratum is to minimize queue & monitor metadata manipulation
 //   that occurs while holding the monitor lock -- that is, we want to
 //   minimize monitor lock holds times.  Note that even a small amount of
 //   fixed spinning will greatly reduce the # of enqueue-dequeue operations
 //   on EntryList|cxq.  That is, spinning relieves contention on the "inner"
 //   locks and monitor metadata.
 //
 //   Cxq points to the the set of Recently Arrived Threads attempting entry.
 //   Because we push threads onto _cxq with CAS, the RATs must take the form of
 //   a singly-linked LIFO.  We drain _cxq into EntryList  at unlock-time when
 //   the unlocking thread notices that EntryList is null but _cxq is != null.
 //
 //   The EntryList is ordered by the prevailing queue discipline and
 //   can be organized in any convenient fashion, such as a doubly-linked list or
 //   a circular doubly-linked list.  Critically, we want insert and delete operations
 //   to operate in constant-time.  If we need a priority queue then something akin
 //   to Solaris' sleepq would work nicely.  Viz.,
 //   http://agg.eng/ws/on10_nightly/source/usr/src/uts/common/os/sleepq.c.
 //   Queue discipline is enforced at ::exit() time, when the unlocking thread
 //   drains the cxq into the EntryList, and orders or reorders the threads on the
 //   EntryList accordingly.
 //
 //   Barring "lock barging", this mechanism provides fair cyclic ordering,
 //   somewhat similar to an elevator-scan.
 //
 // * The monitor synchronization subsystem avoids the use of native
 //   synchronization primitives except for the narrow platform-specific
 //   park-unpark abstraction.  See the comments in os_solaris.cpp regarding
 //   the semantics of park-unpark.  Put another way, this monitor implementation
 //   depends only on atomic operations and park-unpark.  The monitor subsystem
 //   manages all RUNNING->BLOCKED and BLOCKED->READY transitions while the
 //   underlying OS manages the READY<->RUN transitions.
 //
 // * Waiting threads reside on the WaitSet list -- wait() puts
 //   the caller onto the WaitSet.
 //
 // * notify() or notifyAll() simply transfers threads from the WaitSet to
 //   either the EntryList or cxq.  Subsequent exit() operations will
 //   unpark the notifyee.  Unparking a notifee in notify() is inefficient -
 //   it's likely the notifyee would simply impale itself on the lock held
 //   by the notifier.
 //
 // * An interesting alternative is to encode cxq as (List,LockByte) where
 //   the LockByte is 0 iff the monitor is owned.  _owner is simply an auxiliary
 //   variable, like _recursions, in the scheme.  The threads or Events that form
 //   the list would have to be aligned in 256-byte addresses.  A thread would
 //   try to acquire the lock or enqueue itself with CAS, but exiting threads
 //   could use a 1-0 protocol and simply STB to set the LockByte to 0.
 //   Note that is is *not* word-tearing, but it does presume that full-word
 //   CAS operations are coherent with intermix with STB operations.  That's true
 //   on most common processors.
 //
 // * See also http://blogs.sun.com/dave
 void ATTR ObjectMonitor::EnterI (TRAPS) {
     Thread * Self = THREAD ;
     assert (Self->is_Java_thread(), "invariant") ;
     assert (((JavaThread *) Self)->thread_state() == _thread_blocked   , "invariant") ;
     // Try the lock - TATAS
     if (TryLock (Self) > 0) {
         assert (_succ != Self              , "invariant") ;
         assert (_owner == Self             , "invariant") ;
         assert (_Responsible != Self       , "invariant") ;
         return ;
     }
     DeferredInitialize () ;
     // We try one round of spinning *before* enqueueing Self.
     //
     // If the _owner is ready but OFFPROC we could use a YieldTo()
     // operation to donate the remainder of this thread's quantum
     // to the owner.  This has subtle but beneficial affinity
     // effects.
     if (TrySpin (Self) > 0) {
         assert (_owner == Self        , "invariant") ;
         assert (_succ != Self         , "invariant") ;
         assert (_Responsible != Self  , "invariant") ;
         return ;
     }
     // The Spin failed -- Enqueue and park the thread ...
     assert (_succ  != Self            , "invariant") ;
     assert (_owner != Self            , "invariant") ;
     assert (_Responsible != Self      , "invariant") ;
     // Enqueue "Self" on ObjectMonitor's _cxq.
     //
     // Node acts as a proxy for Self.
     // As an aside, if were to ever rewrite the synchronization code mostly
     // in Java, WaitNodes, ObjectMonitors, and Events would become 1st-class
     // Java objects.  This would avoid awkward lifecycle and liveness issues,
     // as well as eliminate a subset of ABA issues.
     // TODO: eliminate ObjectWaiter and enqueue either Threads or Events.
     //
     ObjectWaiter node(Self) ;
     Self->_ParkEvent->reset() ;
     node._prev   = (ObjectWaiter *) 0xBAD ;
     node.TState  = ObjectWaiter::TS_CXQ ;
     // Push "Self" onto the front of the _cxq.
     // Once on cxq/EntryList, Self stays on-queue until it acquires the lock.
     // Note that spinning tends to reduce the rate at which threads
     // enqueue and dequeue on EntryList|cxq.
     ObjectWaiter * nxt ;
     for (;;) {
         node._next = nxt = _cxq ;
         if (Atomic::cmpxchg_ptr (&node, &_cxq, nxt) == nxt) break ;
         // Interference - the CAS failed because _cxq changed.  Just retry.
         // As an optional optimization we retry the lock.
         if (TryLock (Self) > 0) {
             assert (_succ != Self         , "invariant") ;
             assert (_owner == Self        , "invariant") ;
             assert (_Responsible != Self  , "invariant") ;
             return ;
         }
     }
     // Check for cxq|EntryList edge transition to non-null.  This indicates
     // the onset of contention.  While contention persists exiting threads
     // will use a ST:MEMBAR:LD 1-1 exit protocol.  When contention abates exit
     // operations revert to the faster 1-0 mode.  This enter operation may interleave
     // (race) a concurrent 1-0 exit operation, resulting in stranding, so we
     // arrange for one of the contending thread to use a timed park() operations
     // to detect and recover from the race.  (Stranding is form of progress failure
     // where the monitor is unlocked but all the contending threads remain parked).
     // That is, at least one of the contended threads will periodically poll _owner.
     // One of the contending threads will become the designated "Responsible" thread.
     // The Responsible thread uses a timed park instead of a normal indefinite park
     // operation -- it periodically wakes and checks for and recovers from potential
     // strandings admitted by 1-0 exit operations.   We need at most one Responsible
     // thread per-monitor at any given moment.  Only threads on cxq|EntryList may
     // be responsible for a monitor.
     //
     // Currently, one of the contended threads takes on the added role of "Responsible".
     // A viable alternative would be to use a dedicated "stranding checker" thread
     // that periodically iterated over all the threads (or active monitors) and unparked
     // successors where there was risk of stranding.  This would help eliminate the
     // timer scalability issues we see on some platforms as we'd only have one thread
     // -- the checker -- parked on a timer.
     if ((SyncFlags & 16) == 0 && nxt == NULL && _EntryList == NULL) {
         // Try to assume the role of responsible thread for the monitor.
         // CONSIDER:  ST vs CAS vs { if (Responsible==null) Responsible=Self }
         Atomic::cmpxchg_ptr (Self, &_Responsible, NULL) ;
     }
     // The lock have been released while this thread was occupied queueing
     // itself onto _cxq.  To close the race and avoid "stranding" and
     // progress-liveness failure we must resample-retry _owner before parking.
     // Note the Dekker/Lamport duality: ST cxq; MEMBAR; LD Owner.
     // In this case the ST-MEMBAR is accomplished with CAS().
     //
     // TODO: Defer all thread state transitions until park-time.
     // Since state transitions are heavy and inefficient we'd like
     // to defer the state transitions until absolutely necessary,
     // and in doing so avoid some transitions ...
     TEVENT (Inflated enter - Contention) ;
     int nWakeups = 0 ;
     int RecheckInterval = 1 ;
     for (;;) {
         if (TryLock (Self) > 0) break ;
         assert (_owner != Self, "invariant") ;
         if ((SyncFlags & 2) && _Responsible == NULL) {
            Atomic::cmpxchg_ptr (Self, &_Responsible, NULL) ;
         }
         // park self
         if (_Responsible == Self || (SyncFlags & 1)) {
             TEVENT (Inflated enter - park TIMED) ;
             Self->_ParkEvent->park ((jlong) RecheckInterval) ;
             // Increase the RecheckInterval, but clamp the value.
             RecheckInterval *= 8 ;
             if (RecheckInterval > 1000) RecheckInterval = 1000 ;
         } else {
             TEVENT (Inflated enter - park UNTIMED) ;
             Self->_ParkEvent->park() ;
         }
         if (TryLock(Self) > 0) break ;
         // The lock is still contested.
         // Keep a tally of the # of futile wakeups.
         // Note that the counter is not protected by a lock or updated by atomics.
         // That is by design - we trade "lossy" counters which are exposed to
         // races during updates for a lower probe effect.
         TEVENT (Inflated enter - Futile wakeup) ;
         if (ObjectSynchronizer::_sync_FutileWakeups != NULL) {
            ObjectSynchronizer::_sync_FutileWakeups->inc() ;
         }
         ++ nWakeups ;
         // Assuming this is not a spurious wakeup we'll normally find _succ == Self.
         // We can defer clearing _succ until after the spin completes
         // TrySpin() must tolerate being called with _succ == Self.
         // Try yet another round of adaptive spinning.
         if ((Knob_SpinAfterFutile & 1) && TrySpin (Self) > 0) break ;
         // We can find that we were unpark()ed and redesignated _succ while
         // we were spinning.  That's harmless.  If we iterate and call park(),
         // park() will consume the event and return immediately and we'll
         // just spin again.  This pattern can repeat, leaving _succ to simply
         // spin on a CPU.  Enable Knob_ResetEvent to clear pending unparks().
         // Alternately, we can sample fired() here, and if set, forgo spinning
         // in the next iteration.
         if ((Knob_ResetEvent & 1) && Self->_ParkEvent->fired()) {
            Self->_ParkEvent->reset() ;
            OrderAccess::fence() ;
         }
         if (_succ == Self) _succ = NULL ;
         // Invariant: after clearing _succ a thread *must* retry _owner before parking.
         OrderAccess::fence() ;
     }
     // Egress :
     // Self has acquired the lock -- Unlink Self from the cxq or EntryList.
     // Normally we'll find Self on the EntryList .
     // From the perspective of the lock owner (this thread), the
     // EntryList is stable and cxq is prepend-only.
     // The head of cxq is volatile but the interior is stable.
     // In addition, Self.TState is stable.
     assert (_owner == Self      , "invariant") ;
     assert (object() != NULL    , "invariant") ;
     // I'd like to write:
     //   guarantee (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
     // but as we're at a safepoint that's not safe.
     UnlinkAfterAcquire (Self, &node) ;
     if (_succ == Self) _succ = NULL ;
     assert (_succ != Self, "invariant") ;
     if (_Responsible == Self) {
         _Responsible = NULL ;
         // Dekker pivot-point.
         // Consider OrderAccess::storeload() here
         // We may leave threads on cxq|EntryList without a designated
         // "Responsible" thread.  This is benign.  When this thread subsequently
         // exits the monitor it can "see" such preexisting "old" threads --
         // threads that arrived on the cxq|EntryList before the fence, above --
         // by LDing cxq|EntryList.  Newly arrived threads -- that is, threads
         // that arrive on cxq after the ST:MEMBAR, above -- will set Responsible
         // non-null and elect a new "Responsible" timer thread.
         //
         // This thread executes:
         //    ST Responsible=null; MEMBAR    (in enter epilog - here)
         //    LD cxq|EntryList               (in subsequent exit)
         //
         // Entering threads in the slow/contended path execute:
         //    ST cxq=nonnull; MEMBAR; LD Responsible (in enter prolog)
         //    The (ST cxq; MEMBAR) is accomplished with CAS().
         //
         // The MEMBAR, above, prevents the LD of cxq|EntryList in the subsequent
         // exit operation from floating above the ST Responsible=null.
         //
         // In *practice* however, EnterI() is always followed by some atomic
         // operation such as the decrement of _count in ::enter().  Those atomics
         // obviate the need for the explicit MEMBAR, above.
     }
     // We've acquired ownership with CAS().
     // CAS is serializing -- it has MEMBAR/FENCE-equivalent semantics.
     // But since the CAS() this thread may have also stored into _succ,
     // EntryList, cxq or Responsible.  These meta-data updates must be
     // visible __before this thread subsequently drops the lock.
     // Consider what could occur if we didn't enforce this constraint --
     // STs to monitor meta-data and user-data could reorder with (become
     // visible after) the ST in exit that drops ownership of the lock.
     // Some other thread could then acquire the lock, but observe inconsistent
     // or old monitor meta-data and heap data.  That violates the JMM.
     // To that end, the 1-0 exit() operation must have at least STST|LDST
     // "release" barrier semantics.  Specifically, there must be at least a
     // STST|LDST barrier in exit() before the ST of null into _owner that drops
     // the lock.   The barrier ensures that changes to monitor meta-data and data
     // protected by the lock will be visible before we release the lock, and
     // therefore before some other thread (CPU) has a chance to acquire the lock.
     // See also: http://gee.cs.oswego.edu/dl/jmm/cookbook.html.
     //
     // Critically, any prior STs to _succ or EntryList must be visible before
     // the ST of null into _owner in the *subsequent* (following) corresponding
     // monitorexit.  Recall too, that in 1-0 mode monitorexit does not necessarily
     // execute a serializing instruction.
     if (SyncFlags & 8) {
        OrderAccess::fence() ;
     }
     return ;
 }
 // ExitSuspendEquivalent:
 // A faster alternate to handle_special_suspend_equivalent_condition()
 //
 // handle_special_suspend_equivalent_condition() unconditionally
 // acquires the SR_lock.  On some platforms uncontended MutexLocker()
 // operations have high latency.  Note that in ::enter() we call HSSEC
 // while holding the monitor, so we effectively lengthen the critical sections.
 //
 // There are a number of possible solutions:
 //
 // A.  To ameliorate the problem we might also defer state transitions
 //     to as late as possible -- just prior to parking.
 //     Given that, we'd call HSSEC after having returned from park(),
 //     but before attempting to acquire the monitor.  This is only a
 //     partial solution.  It avoids calling HSSEC while holding the
 //     monitor (good), but it still increases successor reacquisition latency --
 //     the interval between unparking a successor and the time the successor
 //     resumes and retries the lock.  See ReenterI(), which defers state transitions.
 //     If we use this technique we can also avoid EnterI()-exit() loop
 //     in ::enter() where we iteratively drop the lock and then attempt
 //     to reacquire it after suspending.
 //
 // B.  In the future we might fold all the suspend bits into a
 //     composite per-thread suspend flag and then update it with CAS().
 //     Alternately, a Dekker-like mechanism with multiple variables
 //     would suffice:
 //       ST Self->_suspend_equivalent = false
 //       MEMBAR
 //       LD Self_>_suspend_flags
 //
 bool ObjectMonitor::ExitSuspendEquivalent (JavaThread * jSelf) {
    int Mode = Knob_FastHSSEC ;
    if (Mode && !jSelf->is_external_suspend()) {
       assert (jSelf->is_suspend_equivalent(), "invariant") ;
       jSelf->clear_suspend_equivalent() ;
       if (2 == Mode) OrderAccess::storeload() ;
       if (!jSelf->is_external_suspend()) return false ;
       // We raced a suspension -- fall thru into the slow path
       TEVENT (ExitSuspendEquivalent - raced) ;
       jSelf->set_suspend_equivalent() ;
    }
    return jSelf->handle_special_suspend_equivalent_condition() ;
 }
 // ReenterI() is a specialized inline form of the latter half of the
 // contended slow-path from EnterI().  We use ReenterI() only for
 // monitor reentry in wait().
 //
 // In the future we should reconcile EnterI() and ReenterI(), adding
 // Knob_Reset and Knob_SpinAfterFutile support and restructuring the
 // loop accordingly.
 void ATTR ObjectMonitor::ReenterI (Thread * Self, ObjectWaiter * SelfNode) {
     assert (Self != NULL                , "invariant") ;
     assert (SelfNode != NULL            , "invariant") ;
     assert (SelfNode->_thread == Self   , "invariant") ;
     assert (_waiters > 0                , "invariant") ;
     assert (((oop)(object()))->mark() == markOopDesc::encode(this) , "invariant") ;
     assert (((JavaThread *)Self)->thread_state() != _thread_blocked, "invariant") ;
     JavaThread * jt = (JavaThread *) Self ;
     int nWakeups = 0 ;
     for (;;) {
         ObjectWaiter::TStates v = SelfNode->TState ;
         guarantee (v == ObjectWaiter::TS_ENTER || v == ObjectWaiter::TS_CXQ, "invariant") ;
         assert    (_owner != Self, "invariant") ;
         if (TryLock (Self) > 0) break ;
         if (TrySpin (Self) > 0) break ;
         TEVENT (Wait Reentry - parking) ;
         // State transition wrappers around park() ...
         // ReenterI() wisely defers state transitions until
         // it's clear we must park the thread.
         {
            OSThreadContendState osts(Self->osthread());
            ThreadBlockInVM tbivm(jt);
            // cleared by handle_special_suspend_equivalent_condition()
            // or java_suspend_self()
            jt->set_suspend_equivalent();
            if (SyncFlags & 1) {
               Self->_ParkEvent->park ((jlong)1000) ;
            } else {
               Self->_ParkEvent->park () ;
            }
            // were we externally suspended while we were waiting?
            for (;;) {
               if (!ExitSuspendEquivalent (jt)) break ;
               if (_succ == Self) { _succ = NULL; OrderAccess::fence(); }
               jt->java_suspend_self();
               jt->set_suspend_equivalent();
            }
         }
         // Try again, but just so we distinguish between futile wakeups and
         // successful wakeups.  The following test isn't algorithmically
         // necessary, but it helps us maintain sensible statistics.
         if (TryLock(Self) > 0) break ;
         // The lock is still contested.
         // Keep a tally of the # of futile wakeups.
         // Note that the counter is not protected by a lock or updated by atomics.
         // That is by design - we trade "lossy" counters which are exposed to
         // races during updates for a lower probe effect.
         TEVENT (Wait Reentry - futile wakeup) ;
         ++ nWakeups ;
         // Assuming this is not a spurious wakeup we'll normally
         // find that _succ == Self.
         if (_succ == Self) _succ = NULL ;
         // Invariant: after clearing _succ a contending thread
         // *must* retry  _owner before parking.
         OrderAccess::fence() ;
         if (ObjectSynchronizer::_sync_FutileWakeups != NULL) {
           ObjectSynchronizer::_sync_FutileWakeups->inc() ;
         }
     }
     // Self has acquired the lock -- Unlink Self from the cxq or EntryList .
     // Normally we'll find Self on the EntryList.
     // Unlinking from the EntryList is constant-time and atomic-free.
     // From the perspective of the lock owner (this thread), the
     // EntryList is stable and cxq is prepend-only.
     // The head of cxq is volatile but the interior is stable.
     // In addition, Self.TState is stable.
     assert (_owner == Self, "invariant") ;
     assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
     UnlinkAfterAcquire (Self, SelfNode) ;
     if (_succ == Self) _succ = NULL ;
     assert (_succ != Self, "invariant") ;
     SelfNode->TState = ObjectWaiter::TS_RUN ;
     OrderAccess::fence() ;      // see comments at the end of EnterI()
 }
 bool ObjectMonitor::try_enter(Thread* THREAD) {
   if (THREAD != _owner) {
     if (THREAD->is_lock_owned ((address)_owner)) {
        assert(_recursions == 0, "internal state error");
        _owner = THREAD ;
        _recursions = 1 ;
        OwnerIsThread = 1 ;
        return true;
     }
     if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) {
       return false;
     }
     return true;
   } else {
     _recursions++;
     return true;
   }
 }
 void ATTR ObjectMonitor::enter(TRAPS) {
   // The following code is ordered to check the most common cases first
   // and to reduce RTS->RTO cache line upgrades on SPARC and IA32 processors.
   Thread * const Self = THREAD ;
   void * cur ;
   cur = Atomic::cmpxchg_ptr (Self, &_owner, NULL) ;
   if (cur == NULL) {
      // Either ASSERT _recursions == 0 or explicitly set _recursions = 0.
      assert (_recursions == 0   , "invariant") ;
      assert (_owner      == Self, "invariant") ;
      // CONSIDER: set or assert OwnerIsThread == 1
      return ;
   }
   if (cur == Self) {
      // TODO-FIXME: check for integer overflow!  BUGID 6557169.
      _recursions ++ ;
      return ;
   }
   if (Self->is_lock_owned ((address)cur)) {
     assert (_recursions == 0, "internal state error");
     _recursions = 1 ;
     // Commute owner from a thread-specific on-stack BasicLockObject address to
     // a full-fledged "Thread *".
     _owner = Self ;
     OwnerIsThread = 1 ;
     return ;
   }
   // We've encountered genuine contention.
   assert (Self->_Stalled == 0, "invariant") ;
   Self->_Stalled = intptr_t(this) ;
   // Try one round of spinning *before* enqueueing Self
   // and before going through the awkward and expensive state
   // transitions.  The following spin is strictly optional ...
   // Note that if we acquire the monitor from an initial spin
   // we forgo posting JVMTI events and firing DTRACE probes.
   if (Knob_SpinEarly && TrySpin (Self) > 0) {
      assert (_owner == Self      , "invariant") ;
      assert (_recursions == 0    , "invariant") ;
      assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
      Self->_Stalled = 0 ;
      return ;
   }
   assert (_owner != Self          , "invariant") ;
   assert (_succ  != Self          , "invariant") ;
   assert (Self->is_Java_thread()  , "invariant") ;
   JavaThread * jt = (JavaThread *) Self ;
   assert (!SafepointSynchronize::is_at_safepoint(), "invariant") ;
   assert (jt->thread_state() != _thread_blocked   , "invariant") ;
   assert (this->object() != NULL  , "invariant") ;
   assert (_count >= 0, "invariant") ;
   // Prevent deflation at STW-time.  See deflate_idle_monitors() and is_busy().
   // Ensure the object-monitor relationship remains stable while there's contention.
   Atomic::inc_ptr(&_count);
   { // Change java thread status to indicate blocked on monitor enter.
     JavaThreadBlockedOnMonitorEnterState jtbmes(jt, this);
     DTRACE_MONITOR_PROBE(contended__enter, this, object(), jt);
     if (JvmtiExport::should_post_monitor_contended_enter()) {
       JvmtiExport::post_monitor_contended_enter(jt, this);
     }
     OSThreadContendState osts(Self->osthread());
     ThreadBlockInVM tbivm(jt);
     Self->set_current_pending_monitor(this);
     // TODO-FIXME: change the following for(;;) loop to straight-line code.
     for (;;) {
       jt->set_suspend_equivalent();
       // cleared by handle_special_suspend_equivalent_condition()
       // or java_suspend_self()
       EnterI (THREAD) ;
       if (!ExitSuspendEquivalent(jt)) break ;
       //
       // We have acquired the contended monitor, but while we were
       // waiting another thread suspended us. We don't want to enter
       // the monitor while suspended because that would surprise the
       // thread that suspended us.
       //
           _recursions = 0 ;
       _succ = NULL ;
       exit (Self) ;
       jt->java_suspend_self();
     }
     Self->set_current_pending_monitor(NULL);
   }
   Atomic::dec_ptr(&_count);
   assert (_count >= 0, "invariant") ;
   Self->_Stalled = 0 ;
   // Must either set _recursions = 0 or ASSERT _recursions == 0.
   assert (_recursions == 0     , "invariant") ;
   assert (_owner == Self       , "invariant") ;
   assert (_succ  != Self       , "invariant") ;
   assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
   // The thread -- now the owner -- is back in vm mode.
   // Report the glorious news via TI,DTrace and jvmstat.
   // The probe effect is non-trivial.  All the reportage occurs
   // while we hold the monitor, increasing the length of the critical
   // section.  Amdahl's parallel speedup law comes vividly into play.
   //
   // Another option might be to aggregate the events (thread local or
   // per-monitor aggregation) and defer reporting until a more opportune
   // time -- such as next time some thread encounters contention but has
   // yet to acquire the lock.  While spinning that thread could
   // spinning we could increment JVMStat counters, etc.
   DTRACE_MONITOR_PROBE(contended__entered, this, object(), jt);
   if (JvmtiExport::should_post_monitor_contended_entered()) {
     JvmtiExport::post_monitor_contended_entered(jt, this);
   }
   if (ObjectSynchronizer::_sync_ContendedLockAttempts != NULL) {
      ObjectSynchronizer::_sync_ContendedLockAttempts->inc() ;
   }
 }
 void ObjectMonitor::ExitEpilog (Thread * Self, ObjectWaiter * Wakee) {
    assert (_owner == Self, "invariant") ;
    // Exit protocol:
    // 1. ST _succ = wakee
    // 2. membar #loadstore|#storestore;
    // 2. ST _owner = NULL
    // 3. unpark(wakee)
    _succ = Knob_SuccEnabled ? Wakee->_thread : NULL ;
    ParkEvent * Trigger = Wakee->_event ;
    // Hygiene -- once we've set _owner = NULL we can't safely dereference Wakee again.
    // The thread associated with Wakee may have grabbed the lock and "Wakee" may be
    // out-of-scope (non-extant).
    Wakee  = NULL ;
    // Drop the lock
    OrderAccess::release_store_ptr (&_owner, NULL) ;
    OrderAccess::fence() ;                               // ST _owner vs LD in unpark()
    // TODO-FIXME:
    // If there's a safepoint pending the best policy would be to
    // get _this thread to a safepoint and only wake the successor
    // after the safepoint completed.  monitorexit uses a "leaf"
    // state transition, however, so this thread can't become
    // safe at this point in time.  (Its stack isn't walkable).
    // The next best thing is to defer waking the successor by
    // adding to a list of thread to be unparked after at the
    // end of the forthcoming STW).
    if (SafepointSynchronize::do_call_back()) {
       TEVENT (unpark before SAFEPOINT) ;
    }
    // Possible optimizations ...
    //
    // * Consider: set Wakee->UnparkTime = timeNow()
    //   When the thread wakes up it'll compute (timeNow() - Self->UnparkTime()).
    //   By measuring recent ONPROC latency we can approximate the
    //   system load.  In turn, we can feed that information back
    //   into the spinning & succession policies.
    //   (ONPROC latency correlates strongly with load).
    //
    // * Pull affinity:
    //   If the wakee is cold then transiently setting it's affinity
    //   to the current CPU is a good idea.
    //   See http://j2se.east/~dice/PERSIST/050624-PullAffinity.txt
    DTRACE_MONITOR_PROBE(contended__exit, this, object(), Self);
    Trigger->unpark() ;
    // Maintain stats and report events to JVMTI
    if (ObjectSynchronizer::_sync_Parks != NULL) {
       ObjectSynchronizer::_sync_Parks->inc() ;
    }
 }
 // exit()
 // ~~~~~~
 // Note that the collector can't reclaim the objectMonitor or deflate
 // the object out from underneath the thread calling ::exit() as the
 // thread calling ::exit() never transitions to a stable state.
 // This inhibits GC, which in turn inhibits asynchronous (and
 // inopportune) reclamation of "this".
 //
 // We'd like to assert that: (THREAD->thread_state() != _thread_blocked) ;
 // There's one exception to the claim above, however.  EnterI() can call
 // exit() to drop a lock if the acquirer has been externally suspended.
 // In that case exit() is called with _thread_state as _thread_blocked,
 // but the monitor's _count field is > 0, which inhibits reclamation.
 //
 // 1-0 exit
 // ~~~~~~~~
 // ::exit() uses a canonical 1-1 idiom with a MEMBAR although some of
 // the fast-path operators have been optimized so the common ::exit()
 // operation is 1-0.  See i486.ad fast_unlock(), for instance.
 // The code emitted by fast_unlock() elides the usual MEMBAR.  This
 // greatly improves latency -- MEMBAR and CAS having considerable local
 // latency on modern processors -- but at the cost of "stranding".  Absent the
 // MEMBAR, a thread in fast_unlock() can race a thread in the slow
 // ::enter() path, resulting in the entering thread being stranding
 // and a progress-liveness failure.   Stranding is extremely rare.
 // We use timers (timed park operations) & periodic polling to detect
 // and recover from stranding.  Potentially stranded threads periodically
 // wake up and poll the lock.  See the usage of the _Responsible variable.
 //
 // The CAS() in enter provides for safety and exclusion, while the CAS or
 // MEMBAR in exit provides for progress and avoids stranding.  1-0 locking
 // eliminates the CAS/MEMBAR from the exist path, but it admits stranding.
 // We detect and recover from stranding with timers.
 //
 // If a thread transiently strands it'll park until (a) another
 // thread acquires the lock and then drops the lock, at which time the
 // exiting thread will notice and unpark the stranded thread, or, (b)
 // the timer expires.  If the lock is high traffic then the stranding latency
 // will be low due to (a).  If the lock is low traffic then the odds of
 // stranding are lower, although the worst-case stranding latency
 // is longer.  Critically, we don't want to put excessive load in the
 // platform's timer subsystem.  We want to minimize both the timer injection
 // rate (timers created/sec) as well as the number of timers active at
 // any one time.  (more precisely, we want to minimize timer-seconds, which is
 // the integral of the # of active timers at any instant over time).
 // Both impinge on OS scalability.  Given that, at most one thread parked on
 // a monitor will use a timer.
 void ATTR ObjectMonitor::exit(TRAPS) {
    Thread * Self = THREAD ;
    if (THREAD != _owner) {
      if (THREAD->is_lock_owned((address) _owner)) {
        // Transmute _owner from a BasicLock pointer to a Thread address.
        // We don't need to hold _mutex for this transition.
        // Non-null to Non-null is safe as long as all readers can
        // tolerate either flavor.
        assert (_recursions == 0, "invariant") ;
        _owner = THREAD ;
        _recursions = 0 ;
        OwnerIsThread = 1 ;
      } else {
        // NOTE: we need to handle unbalanced monitor enter/exit
        // in native code by throwing an exception.
        // TODO: Throw an IllegalMonitorStateException ?
        TEVENT (Exit - Throw IMSX) ;
        assert(false, "Non-balanced monitor enter/exit!");
        if (false) {
           THROW(vmSymbols::java_lang_IllegalMonitorStateException());
        }
        return;
      }
    }
    if (_recursions != 0) {
      _recursions--;        // this is simple recursive enter
      TEVENT (Inflated exit - recursive) ;
      return ;
    }
    // Invariant: after setting Responsible=null an thread must execute
    // a MEMBAR or other serializing instruction before fetching EntryList|cxq.
    if ((SyncFlags & 4) == 0) {
       _Responsible = NULL ;
    }
    for (;;) {
       assert (THREAD == _owner, "invariant") ;
       // Fast-path monitor exit:
       //
       // Observe the Dekker/Lamport duality:
       // A thread in ::exit() executes:
       //   ST Owner=null; MEMBAR; LD EntryList|cxq.
       // A thread in the contended ::enter() path executes the complementary:
       //   ST EntryList|cxq = nonnull; MEMBAR; LD Owner.
       //
       // Note that there's a benign race in the exit path.  We can drop the
       // lock, another thread can reacquire the lock immediately, and we can
       // then wake a thread unnecessarily (yet another flavor of futile wakeup).
       // This is benign, and we've structured the code so the windows are short
       // and the frequency of such futile wakeups is low.
       //
       // We could eliminate the race by encoding both the "LOCKED" state and
       // the queue head in a single word.  Exit would then use either CAS to
       // clear the LOCKED bit/byte.  This precludes the desirable 1-0 optimization,
       // however.
       //
       // Possible fast-path ::exit() optimization:
       // The current fast-path exit implementation fetches both cxq and EntryList.
       // See also i486.ad fast_unlock().  Testing has shown that two LDs
       // isn't measurably slower than a single LD on any platforms.
       // Still, we could reduce the 2 LDs to one or zero by one of the following:
       //
       // - Use _count instead of cxq|EntryList
       //   We intend to eliminate _count, however, when we switch
       //   to on-the-fly deflation in ::exit() as is used in
       //   Metalocks and RelaxedLocks.
       //
       // - Establish the invariant that cxq == null implies EntryList == null.
       //   set cxq == EMPTY (1) to encode the state where cxq is empty
       //   by EntryList != null.  EMPTY is a distinguished value.
       //   The fast-path exit() would fetch cxq but not EntryList.
       //
       // - Encode succ as follows:
       //   succ = t :  Thread t is the successor -- t is ready or is spinning.
       //               Exiting thread does not need to wake a successor.
       //   succ = 0 :  No successor required -> (EntryList|cxq) == null
       //               Exiting thread does not need to wake a successor
       //   succ = 1 :  Successor required    -> (EntryList|cxq) != null and
       //               logically succ == null.
       //               Exiting thread must wake a successor.
       //
       //   The 1-1 fast-exit path would appear as :
       //     _owner = null ; membar ;
       //     if (_succ == 1 && CAS (&_owner, null, Self) == null) goto SlowPath
       //     goto FastPathDone ;
       //
       //   and the 1-0 fast-exit path would appear as:
       //      if (_succ == 1) goto SlowPath
       //      Owner = null ;
       //      goto FastPathDone
       //
       // - Encode the LSB of _owner as 1 to indicate that exit()
       //   must use the slow-path and make a successor ready.
       //   (_owner & 1) == 0 IFF succ != null || (EntryList|cxq) == null
       //   (_owner & 1) == 0 IFF succ == null && (EntryList|cxq) != null (obviously)
       //   The 1-0 fast exit path would read:
       //      if (_owner != Self) goto SlowPath
       //      _owner = null
       //      goto FastPathDone
       if (Knob_ExitPolicy == 0) {
          // release semantics: prior loads and stores from within the critical section
          // must not float (reorder) past the following store that drops the lock.
          // On SPARC that requires MEMBAR #loadstore|#storestore.
          // But of course in TSO #loadstore|#storestore is not required.
          // I'd like to write one of the following:
          // A.  OrderAccess::release() ; _owner = NULL
          // B.  OrderAccess::loadstore(); OrderAccess::storestore(); _owner = NULL;
          // Unfortunately OrderAccess::release() and OrderAccess::loadstore() both
          // store into a _dummy variable.  That store is not needed, but can result
          // in massive wasteful coherency traffic on classic SMP systems.
          // Instead, I use release_store(), which is implemented as just a simple
          // ST on x64, x86 and SPARC.
          OrderAccess::release_store_ptr (&_owner, NULL) ;   // drop the lock
          OrderAccess::storeload() ;                         // See if we need to wake a successor
          if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) {
             TEVENT (Inflated exit - simple egress) ;
             return ;
          }
          TEVENT (Inflated exit - complex egress) ;
          // Normally the exiting thread is responsible for ensuring succession,
          // but if other successors are ready or other entering threads are spinning
          // then this thread can simply store NULL into _owner and exit without
          // waking a successor.  The existence of spinners or ready successors
          // guarantees proper succession (liveness).  Responsibility passes to the
          // ready or running successors.  The exiting thread delegates the duty.
          // More precisely, if a successor already exists this thread is absolved
          // of the responsibility of waking (unparking) one.
          //
          // The _succ variable is critical to reducing futile wakeup frequency.
          // _succ identifies the "heir presumptive" thread that has been made
          // ready (unparked) but that has not yet run.  We need only one such
          // successor thread to guarantee progress.
          // See http://www.usenix.org/events/jvm01/full_papers/dice/dice.pdf
          // section 3.3 "Futile Wakeup Throttling" for details.
          //
          // Note that spinners in Enter() also set _succ non-null.
          // In the current implementation spinners opportunistically set
          // _succ so that exiting threads might avoid waking a successor.
          // Another less appealing alternative would be for the exiting thread
          // to drop the lock and then spin briefly to see if a spinner managed
          // to acquire the lock.  If so, the exiting thread could exit
          // immediately without waking a successor, otherwise the exiting
          // thread would need to dequeue and wake a successor.
          // (Note that we'd need to make the post-drop spin short, but no
          // shorter than the worst-case round-trip cache-line migration time.
          // The dropped lock needs to become visible to the spinner, and then
          // the acquisition of the lock by the spinner must become visible to
          // the exiting thread).
          //
          // It appears that an heir-presumptive (successor) must be made ready.
          // Only the current lock owner can manipulate the EntryList or
          // drain _cxq, so we need to reacquire the lock.  If we fail
          // to reacquire the lock the responsibility for ensuring succession
          // falls to the new owner.
          //
          if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) {
             return ;
          }
          TEVENT (Exit - Reacquired) ;
       } else {
          if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) {
             OrderAccess::release_store_ptr (&_owner, NULL) ;   // drop the lock
             OrderAccess::storeload() ;
             // Ratify the previously observed values.
             if (_cxq == NULL || _succ != NULL) {
                 TEVENT (Inflated exit - simple egress) ;
                 return ;
             }
             // inopportune interleaving -- the exiting thread (this thread)
             // in the fast-exit path raced an entering thread in the slow-enter
             // path.
             // We have two choices:
             // A.  Try to reacquire the lock.
             //     If the CAS() fails return immediately, otherwise
             //     we either restart/rerun the exit operation, or simply
             //     fall-through into the code below which wakes a successor.
             // B.  If the elements forming the EntryList|cxq are TSM
             //     we could simply unpark() the lead thread and return
             //     without having set _succ.
             if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) {
                TEVENT (Inflated exit - reacquired succeeded) ;
                return ;
             }
             TEVENT (Inflated exit - reacquired failed) ;
          } else {
             TEVENT (Inflated exit - complex egress) ;
          }
       }
       guarantee (_owner == THREAD, "invariant") ;
       // Select an appropriate successor ("heir presumptive") from the EntryList
       // and make it ready.  Generally we just wake the head of EntryList .
       // There's no algorithmic constraint that we use the head - it's just
       // a policy decision.   Note that the thread at head of the EntryList
       // remains at the head until it acquires the lock.  This means we'll
       // repeatedly wake the same thread until it manages to grab the lock.
       // This is generally a good policy - if we're seeing lots of futile wakeups
       // at least we're waking/rewaking a thread that's like to be hot or warm
       // (have residual D$ and TLB affinity).
       //
       // "Wakeup locality" optimization:
       // http://j2se.east/~dice/PERSIST/040825-WakeLocality.txt
       // In the future we'll try to bias the selection mechanism
       // to preferentially pick a thread that recently ran on
       // a processor element that shares cache with the CPU on which
       // the exiting thread is running.   We need access to Solaris'
       // schedctl.sc_cpu to make that work.
       //
       ObjectWaiter * w = NULL ;
       int QMode = Knob_QMode ;
       if (QMode == 2 && _cxq != NULL) {
           // QMode == 2 : cxq has precedence over EntryList.
           // Try to directly wake a successor from the cxq.
           // If successful, the successor will need to unlink itself from cxq.
           w = _cxq ;
           assert (w != NULL, "invariant") ;
           assert (w->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
           ExitEpilog (Self, w) ;
           return ;
       }
       if (QMode == 3 && _cxq != NULL) {
           // Aggressively drain cxq into EntryList at the first opportunity.
           // This policy ensure that recently-run threads live at the head of EntryList.
           // Drain _cxq into EntryList - bulk transfer.
           // First, detach _cxq.
           // The following loop is tantamount to: w = swap (&cxq, NULL)
           w = _cxq ;
           for (;;) {
              assert (w != NULL, "Invariant") ;
              ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ;
              if (u == w) break ;
              w = u ;
           }
           assert (w != NULL              , "invariant") ;
           ObjectWaiter * q = NULL ;
           ObjectWaiter * p ;
           for (p = w ; p != NULL ; p = p->_next) {
               guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
               p->TState = ObjectWaiter::TS_ENTER ;
               p->_prev = q ;
               q = p ;
           }
           // Append the RATs to the EntryList
           // TODO: organize EntryList as a CDLL so we can locate the tail in constant-time.
           ObjectWaiter * Tail ;
           for (Tail = _EntryList ; Tail != NULL && Tail->_next != NULL ; Tail = Tail->_next) ;
           if (Tail == NULL) {
               _EntryList = w ;
           } else {
               Tail->_next = w ;
               w->_prev = Tail ;
           }
           // Fall thru into code that tries to wake a successor from EntryList
       }
       if (QMode == 4 && _cxq != NULL) {
           // Aggressively drain cxq into EntryList at the first opportunity.
           // This policy ensure that recently-run threads live at the head of EntryList.
           // Drain _cxq into EntryList - bulk transfer.
           // First, detach _cxq.
           // The following loop is tantamount to: w = swap (&cxq, NULL)
           w = _cxq ;
           for (;;) {
              assert (w != NULL, "Invariant") ;
              ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ;
              if (u == w) break ;
              w = u ;
           }
           assert (w != NULL              , "invariant") ;
           ObjectWaiter * q = NULL ;
           ObjectWaiter * p ;
           for (p = w ; p != NULL ; p = p->_next) {
               guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
               p->TState = ObjectWaiter::TS_ENTER ;
               p->_prev = q ;
               q = p ;
           }
           // Prepend the RATs to the EntryList
           if (_EntryList != NULL) {
               q->_next = _EntryList ;
               _EntryList->_prev = q ;
           }
           _EntryList = w ;
           // Fall thru into code that tries to wake a successor from EntryList
       }
       w = _EntryList  ;
       if (w != NULL) {
           // I'd like to write: guarantee (w->_thread != Self).
           // But in practice an exiting thread may find itself on the EntryList.
           // Lets say thread T1 calls O.wait().  Wait() enqueues T1 on O's waitset and
           // then calls exit().  Exit release the lock by setting O._owner to NULL.
           // Lets say T1 then stalls.  T2 acquires O and calls O.notify().  The
           // notify() operation moves T1 from O's waitset to O's EntryList. T2 then
           // release the lock "O".  T2 resumes immediately after the ST of null into
           // _owner, above.  T2 notices that the EntryList is populated, so it
           // reacquires the lock and then finds itself on the EntryList.
           // Given all that, we have to tolerate the circumstance where "w" is
           // associated with Self.
           assert (w->TState == ObjectWaiter::TS_ENTER, "invariant") ;
           ExitEpilog (Self, w) ;
           return ;
       }
       // If we find that both _cxq and EntryList are null then just
       // re-run the exit protocol from the top.
       w = _cxq ;
       if (w == NULL) continue ;
       // Drain _cxq into EntryList - bulk transfer.
       // First, detach _cxq.
       // The following loop is tantamount to: w = swap (&cxq, NULL)
       for (;;) {
           assert (w != NULL, "Invariant") ;
           ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ;
           if (u == w) break ;
           w = u ;
       }
       TEVENT (Inflated exit - drain cxq into EntryList) ;
       assert (w != NULL              , "invariant") ;
       assert (_EntryList  == NULL    , "invariant") ;
       // Convert the LIFO SLL anchored by _cxq into a DLL.
       // The list reorganization step operates in O(LENGTH(w)) time.
       // It's critical that this step operate quickly as
       // "Self" still holds the outer-lock, restricting parallelism
       // and effectively lengthening the critical section.
       // Invariant: s chases t chases u.
       // TODO-FIXME: consider changing EntryList from a DLL to a CDLL so
       // we have faster access to the tail.
       if (QMode == 1) {
          // QMode == 1 : drain cxq to EntryList, reversing order
          // We also reverse the order of the list.
          ObjectWaiter * s = NULL ;
          ObjectWaiter * t = w ;
          ObjectWaiter * u = NULL ;
          while (t != NULL) {
              guarantee (t->TState == ObjectWaiter::TS_CXQ, "invariant") ;
              t->TState = ObjectWaiter::TS_ENTER ;
              u = t->_next ;
              t->_prev = u ;
              t->_next = s ;
              s = t;
              t = u ;
          }
          _EntryList  = s ;
          assert (s != NULL, "invariant") ;
       } else {
          // QMode == 0 or QMode == 2
          _EntryList = w ;
          ObjectWaiter * q = NULL ;
          ObjectWaiter * p ;
          for (p = w ; p != NULL ; p = p->_next) {
              guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
              p->TState = ObjectWaiter::TS_ENTER ;
              p->_prev = q ;
              q = p ;
          }
       }
       // In 1-0 mode we need: ST EntryList; MEMBAR #storestore; ST _owner = NULL
       // The MEMBAR is satisfied by the release_store() operation in ExitEpilog().
       // See if we can abdicate to a spinner instead of waking a thread.
       // A primary goal of the implementation is to reduce the
       // context-switch rate.
       if (_succ != NULL) continue;
       w = _EntryList  ;
       if (w != NULL) {
           guarantee (w->TState == ObjectWaiter::TS_ENTER, "invariant") ;
           ExitEpilog (Self, w) ;
           return ;
       }
    }
 }
 // complete_exit exits a lock returning recursion count
 // complete_exit/reenter operate as a wait without waiting
 // complete_exit requires an inflated monitor
 // The _owner field is not always the Thread addr even with an
 // inflated monitor, e.g. the monitor can be inflated by a non-owning
 // thread due to contention.
 intptr_t ObjectMonitor::complete_exit(TRAPS) {
    Thread * const Self = THREAD;
    assert(Self->is_Java_thread(), "Must be Java thread!");
    JavaThread *jt = (JavaThread *)THREAD;
    DeferredInitialize();
    if (THREAD != _owner) {
     if (THREAD->is_lock_owned ((address)_owner)) {
        assert(_recursions == 0, "internal state error");
        _owner = THREAD ;   /* Convert from basiclock addr to Thread addr */
        _recursions = 0 ;
        OwnerIsThread = 1 ;
     }
    }
    guarantee(Self == _owner, "complete_exit not owner");
    intptr_t save = _recursions; // record the old recursion count
    _recursions = 0;        // set the recursion level to be 0
    exit (Self) ;           // exit the monitor
    guarantee (_owner != Self, "invariant");
    return save;
 }
 // reenter() enters a lock and sets recursion count
 // complete_exit/reenter operate as a wait without waiting
 void ObjectMonitor::reenter(intptr_t recursions, TRAPS) {
    Thread * const Self = THREAD;
    assert(Self->is_Java_thread(), "Must be Java thread!");
    JavaThread *jt = (JavaThread *)THREAD;
    guarantee(_owner != Self, "reenter already owner");
    enter (THREAD);       // enter the monitor
    guarantee (_recursions == 0, "reenter recursion");
    _recursions = recursions;
    return;
 }
 // Note: a subset of changes to ObjectMonitor::wait()
 // will need to be replicated in complete_exit above
 void ObjectMonitor::wait(jlong millis, bool interruptible, TRAPS) {
    Thread * const Self = THREAD ;
    assert(Self->is_Java_thread(), "Must be Java thread!");
    JavaThread *jt = (JavaThread *)THREAD;
    DeferredInitialize () ;
    // Throw IMSX or IEX.
    CHECK_OWNER();
    // check for a pending interrupt
    if (interruptible && Thread::is_interrupted(Self, true) && !HAS_PENDING_EXCEPTION) {
      // post monitor waited event.  Note that this is past-tense, we are done waiting.
      if (JvmtiExport::should_post_monitor_waited()) {
         // Note: 'false' parameter is passed here because the
         // wait was not timed out due to thread interrupt.
         JvmtiExport::post_monitor_waited(jt, this, false);
      }
      TEVENT (Wait - Throw IEX) ;
      THROW(vmSymbols::java_lang_InterruptedException());
      return ;
    }
    TEVENT (Wait) ;
    assert (Self->_Stalled == 0, "invariant") ;
    Self->_Stalled = intptr_t(this) ;
    jt->set_current_waiting_monitor(this);
    // create a node to be put into the queue
    // Critically, after we reset() the event but prior to park(), we must check
    // for a pending interrupt.
    ObjectWaiter node(Self);
    node.TState = ObjectWaiter::TS_WAIT ;
    Self->_ParkEvent->reset() ;
    OrderAccess::fence();          // ST into Event; membar ; LD interrupted-flag
    // Enter the waiting queue, which is a circular doubly linked list in this case
    // but it could be a priority queue or any data structure.
    // _WaitSetLock protects the wait queue.  Normally the wait queue is accessed only
    // by the the owner of the monitor *except* in the case where park()
    // returns because of a timeout of interrupt.  Contention is exceptionally rare
    // so we use a simple spin-lock instead of a heavier-weight blocking lock.
    Thread::SpinAcquire (&_WaitSetLock, "WaitSet - add") ;
    AddWaiter (&node) ;
    Thread::SpinRelease (&_WaitSetLock) ;
    if ((SyncFlags & 4) == 0) {
       _Responsible = NULL ;
    }
    intptr_t save = _recursions; // record the old recursion count
    _waiters++;                  // increment the number of waiters
    _recursions = 0;             // set the recursion level to be 1
    exit (Self) ;                    // exit the monitor
    guarantee (_owner != Self, "invariant") ;
    // As soon as the ObjectMonitor's ownership is dropped in the exit()
    // call above, another thread can enter() the ObjectMonitor, do the
    // notify(), and exit() the ObjectMonitor. If the other thread's
    // exit() call chooses this thread as the successor and the unpark()
    // call happens to occur while this thread is posting a
    // MONITOR_CONTENDED_EXIT event, then we run the risk of the event
    // handler using RawMonitors and consuming the unpark().
    //
    // To avoid the problem, we re-post the event. This does no harm
    // even if the original unpark() was not consumed because we are the
    // chosen successor for this monitor.
    if (node._notified != 0 && _succ == Self) {
       node._event->unpark();
    }
    // The thread is on the WaitSet list - now park() it.
    // On MP systems it's conceivable that a brief spin before we park
    // could be profitable.
    //
    // TODO-FIXME: change the following logic to a loop of the form
    //   while (!timeout && !interrupted && _notified == 0) park()
    int ret = OS_OK ;
    int WasNotified = 0 ;
    { // State transition wrappers
      OSThread* osthread = Self->osthread();
      OSThreadWaitState osts(osthread, true);
      {
        ThreadBlockInVM tbivm(jt);
        // Thread is in thread_blocked state and oop access is unsafe.
        jt->set_suspend_equivalent();
        if (interruptible && (Thread::is_interrupted(THREAD, false) || HAS_PENDING_EXCEPTION)) {
            // Intentionally empty
        } else
        if (node._notified == 0) {
          if (millis <= 0) {
             Self->_ParkEvent->park () ;
          } else {
             ret = Self->_ParkEvent->park (millis) ;
          }
        }
        // were we externally suspended while we were waiting?
        if (ExitSuspendEquivalent (jt)) {
           // TODO-FIXME: add -- if succ == Self then succ = null.
           jt->java_suspend_self();
        }
      } // Exit thread safepoint: transition _thread_blocked -> _thread_in_vm
      // Node may be on the WaitSet, the EntryList (or cxq), or in transition
      // from the WaitSet to the EntryList.
      // See if we need to remove Node from the WaitSet.
      // We use double-checked locking to avoid grabbing _WaitSetLock
      // if the thread is not on the wait queue.
      //
      // Note that we don't need a fence before the fetch of TState.
      // In the worst case we'll fetch a old-stale value of TS_WAIT previously
      // written by the is thread. (perhaps the fetch might even be satisfied
      // by a look-aside into the processor's own store buffer, although given
      // the length of the code path between the prior ST and this load that's
      // highly unlikely).  If the following LD fetches a stale TS_WAIT value
      // then we'll acquire the lock and then re-fetch a fresh TState value.
      // That is, we fail toward safety.
      if (node.TState == ObjectWaiter::TS_WAIT) {
          Thread::SpinAcquire (&_WaitSetLock, "WaitSet - unlink") ;
          if (node.TState == ObjectWaiter::TS_WAIT) {
             DequeueSpecificWaiter (&node) ;       // unlink from WaitSet
             assert(node._notified == 0, "invariant");
             node.TState = ObjectWaiter::TS_RUN ;
          }
          Thread::SpinRelease (&_WaitSetLock) ;
      }
      // The thread is now either on off-list (TS_RUN),
      // on the EntryList (TS_ENTER), or on the cxq (TS_CXQ).
      // The Node's TState variable is stable from the perspective of this thread.
      // No other threads will asynchronously modify TState.
      guarantee (node.TState != ObjectWaiter::TS_WAIT, "invariant") ;
      OrderAccess::loadload() ;
      if (_succ == Self) _succ = NULL ;
      WasNotified = node._notified ;
      // Reentry phase -- reacquire the monitor.
      // re-enter contended monitor after object.wait().
      // retain OBJECT_WAIT state until re-enter successfully completes
      // Thread state is thread_in_vm and oop access is again safe,
      // although the raw address of the object may have changed.
      // (Don't cache naked oops over safepoints, of course).
      // post monitor waited event. Note that this is past-tense, we are done waiting.
      if (JvmtiExport::should_post_monitor_waited()) {
        JvmtiExport::post_monitor_waited(jt, this, ret == OS_TIMEOUT);
      }
      OrderAccess::fence() ;
      assert (Self->_Stalled != 0, "invariant") ;
      Self->_Stalled = 0 ;
      assert (_owner != Self, "invariant") ;
      ObjectWaiter::TStates v = node.TState ;
      if (v == ObjectWaiter::TS_RUN) {
          enter (Self) ;
      } else {
          guarantee (v == ObjectWaiter::TS_ENTER || v == ObjectWaiter::TS_CXQ, "invariant") ;
          ReenterI (Self, &node) ;
          node.wait_reenter_end(this);
      }
      // Self has reacquired the lock.
      // Lifecycle - the node representing Self must not appear on any queues.
      // Node is about to go out-of-scope, but even if it were immortal we wouldn't
      // want residual elements associated with this thread left on any lists.
      guarantee (node.TState == ObjectWaiter::TS_RUN, "invariant") ;
      assert    (_owner == Self, "invariant") ;
      assert    (_succ != Self , "invariant") ;
    } // OSThreadWaitState()
    jt->set_current_waiting_monitor(NULL);
    guarantee (_recursions == 0, "invariant") ;
    _recursions = save;     // restore the old recursion count
    _waiters--;             // decrement the number of waiters
    // Verify a few postconditions
    assert (_owner == Self       , "invariant") ;
    assert (_succ  != Self       , "invariant") ;
    assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
    if (SyncFlags & 32) {
       OrderAccess::fence() ;
    }
    // check if the notification happened
    if (!WasNotified) {
      // no, it could be timeout or Thread.interrupt() or both
      // check for interrupt event, otherwise it is timeout
      if (interruptible && Thread::is_interrupted(Self, true) && !HAS_PENDING_EXCEPTION) {
        TEVENT (Wait - throw IEX from epilog) ;
        THROW(vmSymbols::java_lang_InterruptedException());
      }
    }
    // NOTE: Spurious wake up will be consider as timeout.
    // Monitor notify has precedence over thread interrupt.
 }
 // Consider:
 // If the lock is cool (cxq == null && succ == null) and we're on an MP system
 // then instead of transferring a thread from the WaitSet to the EntryList
 // we might just dequeue a thread from the WaitSet and directly unpark() it.
 void ObjectMonitor::notify(TRAPS) {
   CHECK_OWNER();
   if (_WaitSet == NULL) {
      TEVENT (Empty-Notify) ;
      return ;
   }
   DTRACE_MONITOR_PROBE(notify, this, object(), THREAD);
   int Policy = Knob_MoveNotifyee ;
   Thread::SpinAcquire (&_WaitSetLock, "WaitSet - notify") ;
   ObjectWaiter * iterator = DequeueWaiter() ;
   if (iterator != NULL) {
      TEVENT (Notify1 - Transfer) ;
      guarantee (iterator->TState == ObjectWaiter::TS_WAIT, "invariant") ;
      guarantee (iterator->_notified == 0, "invariant") ;
      // Disposition - what might we do with iterator ?
      // a.  add it directly to the EntryList - either tail or head.
      // b.  push it onto the front of the _cxq.
      // For now we use (a).
      if (Policy != 4) {
         iterator->TState = ObjectWaiter::TS_ENTER ;
      }
      iterator->_notified = 1 ;
      ObjectWaiter * List = _EntryList ;
      if (List != NULL) {
         assert (List->_prev == NULL, "invariant") ;
         assert (List->TState == ObjectWaiter::TS_ENTER, "invariant") ;
         assert (List != iterator, "invariant") ;
      }
      if (Policy == 0) {       // prepend to EntryList
          if (List == NULL) {
              iterator->_next = iterator->_prev = NULL ;
              _EntryList = iterator ;
          } else {
              List->_prev = iterator ;
              iterator->_next = List ;
              iterator->_prev = NULL ;
              _EntryList = iterator ;
         }
      } else
      if (Policy == 1) {      // append to EntryList
          if (List == NULL) {
              iterator->_next = iterator->_prev = NULL ;
              _EntryList = iterator ;
          } else {
             // CONSIDER:  finding the tail currently requires a linear-time walk of
             // the EntryList.  We can make tail access constant-time by converting to
             // a CDLL instead of using our current DLL.
             ObjectWaiter * Tail ;
             for (Tail = List ; Tail->_next != NULL ; Tail = Tail->_next) ;
             assert (Tail != NULL && Tail->_next == NULL, "invariant") ;
             Tail->_next = iterator ;
             iterator->_prev = Tail ;
             iterator->_next = NULL ;
         }
      } else
      if (Policy == 2) {      // prepend to cxq
          // prepend to cxq
          if (List == NULL) {
              iterator->_next = iterator->_prev = NULL ;
              _EntryList = iterator ;
          } else {
             iterator->TState = ObjectWaiter::TS_CXQ ;
             for (;;) {
                 ObjectWaiter * Front = _cxq ;
                 iterator->_next = Front ;
                 if (Atomic::cmpxchg_ptr (iterator, &_cxq, Front) == Front) {
                     break ;
                 }
             }
          }
      } else
      if (Policy == 3) {      // append to cxq
         iterator->TState = ObjectWaiter::TS_CXQ ;
         for (;;) {
             ObjectWaiter * Tail ;
             Tail = _cxq ;
             if (Tail == NULL) {
                 iterator->_next = NULL ;
                 if (Atomic::cmpxchg_ptr (iterator, &_cxq, NULL) == NULL) {
                    break ;
                 }
             } else {
                 while (Tail->_next != NULL) Tail = Tail->_next ;
                 Tail->_next = iterator ;
                 iterator->_prev = Tail ;
                 iterator->_next = NULL ;
                 break ;
             }
         }
      } else {
         ParkEvent * ev = iterator->_event ;
         iterator->TState = ObjectWaiter::TS_RUN ;
         OrderAccess::fence() ;
         ev->unpark() ;
      }
      if (Policy < 4) {
        iterator->wait_reenter_begin(this);
      }
      // _WaitSetLock protects the wait queue, not the EntryList.  We could
      // move the add-to-EntryList operation, above, outside the critical section
      // protected by _WaitSetLock.  In practice that's not useful.  With the
      // exception of  wait() timeouts and interrupts the monitor owner
      // is the only thread that grabs _WaitSetLock.  There's almost no contention
      // on _WaitSetLock so it's not profitable to reduce the length of the
      // critical section.
   }
   Thread::SpinRelease (&_WaitSetLock) ;
   if (iterator != NULL && ObjectSynchronizer::_sync_Notifications != NULL) {
      ObjectSynchronizer::_sync_Notifications->inc() ;
   }
 }
 void ObjectMonitor::notifyAll(TRAPS) {
   CHECK_OWNER();
   ObjectWaiter* iterator;
   if (_WaitSet == NULL) {
       TEVENT (Empty-NotifyAll) ;
       return ;
   }
   DTRACE_MONITOR_PROBE(notifyAll, this, object(), THREAD);
   int Policy = Knob_MoveNotifyee ;
   int Tally = 0 ;
   Thread::SpinAcquire (&_WaitSetLock, "WaitSet - notifyall") ;
   for (;;) {
      iterator = DequeueWaiter () ;
      if (iterator == NULL) break ;
      TEVENT (NotifyAll - Transfer1) ;
      ++Tally ;
      // Disposition - what might we do with iterator ?
      // a.  add it directly to the EntryList - either tail or head.
      // b.  push it onto the front of the _cxq.
      // For now we use (a).
      //
      // TODO-FIXME: currently notifyAll() transfers the waiters one-at-a-time from the waitset
      // to the EntryList.  This could be done more efficiently with a single bulk transfer,
      // but in practice it's not time-critical.  Beware too, that in prepend-mode we invert the
      // order of the waiters.  Lets say that the waitset is "ABCD" and the EntryList is "XYZ".
      // After a notifyAll() in prepend mode the waitset will be empty and the EntryList will
      // be "DCBAXYZ".
      guarantee (iterator->TState == ObjectWaiter::TS_WAIT, "invariant") ;
      guarantee (iterator->_notified == 0, "invariant") ;
      iterator->_notified = 1 ;
      if (Policy != 4) {
         iterator->TState = ObjectWaiter::TS_ENTER ;
      }
      ObjectWaiter * List = _EntryList ;
      if (List != NULL) {
         assert (List->_prev == NULL, "invariant") ;
         assert (List->TState == ObjectWaiter::TS_ENTER, "invariant") ;
         assert (List != iterator, "invariant") ;
      }
      if (Policy == 0) {       // prepend to EntryList
          if (List == NULL) {
              iterator->_next = iterator->_prev = NULL ;
              _EntryList = iterator ;
          } else {
              List->_prev = iterator ;
              iterator->_next = List ;
              iterator->_prev = NULL ;
              _EntryList = iterator ;
         }
      } else
      if (Policy == 1) {      // append to EntryList
          if (List == NULL) {
              iterator->_next = iterator->_prev = NULL ;
              _EntryList = iterator ;
          } else {
             // CONSIDER:  finding the tail currently requires a linear-time walk of
             // the EntryList.  We can make tail access constant-time by converting to
             // a CDLL instead of using our current DLL.
             ObjectWaiter * Tail ;
             for (Tail = List ; Tail->_next != NULL ; Tail = Tail->_next) ;
             assert (Tail != NULL && Tail->_next == NULL, "invariant") ;
             Tail->_next = iterator ;
             iterator->_prev = Tail ;
             iterator->_next = NULL ;
         }
      } else
      if (Policy == 2) {      // prepend to cxq
          // prepend to cxq
          iterator->TState = ObjectWaiter::TS_CXQ ;
          for (;;) {
              ObjectWaiter * Front = _cxq ;
              iterator->_next = Front ;
              if (Atomic::cmpxchg_ptr (iterator, &_cxq, Front) == Front) {
                  break ;
              }
          }
      } else
      if (Policy == 3) {      // append to cxq
         iterator->TState = ObjectWaiter::TS_CXQ ;
         for (;;) {
             ObjectWaiter * Tail ;
             Tail = _cxq ;
             if (Tail == NULL) {
                 iterator->_next = NULL ;
                 if (Atomic::cmpxchg_ptr (iterator, &_cxq, NULL) == NULL) {
                    break ;
                 }
             } else {
                 while (Tail->_next != NULL) Tail = Tail->_next ;
                 Tail->_next = iterator ;
                 iterator->_prev = Tail ;
                 iterator->_next = NULL ;
                 break ;
             }
         }
      } else {
         ParkEvent * ev = iterator->_event ;
         iterator->TState = ObjectWaiter::TS_RUN ;
         OrderAccess::fence() ;
         ev->unpark() ;
      }
      if (Policy < 4) {
        iterator->wait_reenter_begin(this);
      }
      // _WaitSetLock protects the wait queue, not the EntryList.  We could
      // move the add-to-EntryList operation, above, outside the critical section
      // protected by _WaitSetLock.  In practice that's not useful.  With the
      // exception of  wait() timeouts and interrupts the monitor owner
      // is the only thread that grabs _WaitSetLock.  There's almost no contention
      // on _WaitSetLock so it's not profitable to reduce the length of the
      // critical section.
   }
   Thread::SpinRelease (&_WaitSetLock) ;
   if (Tally != 0 && ObjectSynchronizer::_sync_Notifications != NULL) {
      ObjectSynchronizer::_sync_Notifications->inc(Tally) ;
   }
 }
 // check_slow() is a misnomer.  It's called to simply to throw an IMSX exception.
 // TODO-FIXME: remove check_slow() -- it's likely dead.
 void ObjectMonitor::check_slow(TRAPS) {
   TEVENT (check_slow - throw IMSX) ;
   assert(THREAD != _owner && !THREAD->is_lock_owned((address) _owner), "must not be owner");
   THROW_MSG(vmSymbols::java_lang_IllegalMonitorStateException(), "current thread not owner");
 }
 // -------------------------------------------------------------------------
 // The raw monitor subsystem is entirely distinct from normal
 // java-synchronization or jni-synchronization.  raw monitors are not
 // associated with objects.  They can be implemented in any manner
 // that makes sense.  The original implementors decided to piggy-back
 // the raw-monitor implementation on the existing Java objectMonitor mechanism.
 // This flaw needs to fixed.  We should reimplement raw monitors as sui-generis.
 // Specifically, we should not implement raw monitors via java monitors.
 // Time permitting, we should disentangle and deconvolve the two implementations
 // and move the resulting raw monitor implementation over to the JVMTI directories.
 // Ideally, the raw monitor implementation would be built on top of
 // park-unpark and nothing else.
 //
 // raw monitors are used mainly by JVMTI
 // The raw monitor implementation borrows the ObjectMonitor structure,
 // but the operators are degenerate and extremely simple.
 //
 // Mixed use of a single objectMonitor instance -- as both a raw monitor
 // and a normal java monitor -- is not permissible.
 //
 // Note that we use the single RawMonitor_lock to protect queue operations for
 // _all_ raw monitors.  This is a scalability impediment, but since raw monitor usage
 // is deprecated and rare, this is not of concern.  The RawMonitor_lock can not
 // be held indefinitely.  The critical sections must be short and bounded.
 //
 // -------------------------------------------------------------------------
 int ObjectMonitor::SimpleEnter (Thread * Self) {
   for (;;) {
     if (Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) {
        return OS_OK ;
     }
     ObjectWaiter Node (Self) ;
     Self->_ParkEvent->reset() ;     // strictly optional
     Node.TState = ObjectWaiter::TS_ENTER ;
     RawMonitor_lock->lock_without_safepoint_check() ;
     Node._next  = _EntryList ;
     _EntryList  = &Node ;
     OrderAccess::fence() ;
     if (_owner == NULL && Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) {
         _EntryList = Node._next ;
         RawMonitor_lock->unlock() ;
         return OS_OK ;
     }
     RawMonitor_lock->unlock() ;
     while (Node.TState == ObjectWaiter::TS_ENTER) {
        Self->_ParkEvent->park() ;
     }
   }
 }
 int ObjectMonitor::SimpleExit (Thread * Self) {
   guarantee (_owner == Self, "invariant") ;
   OrderAccess::release_store_ptr (&_owner, NULL) ;
   OrderAccess::fence() ;
   if (_EntryList == NULL) return OS_OK ;
   ObjectWaiter * w ;
   RawMonitor_lock->lock_without_safepoint_check() ;
   w = _EntryList ;
   if (w != NULL) {
       _EntryList = w->_next ;
   }
   RawMonitor_lock->unlock() ;
   if (w != NULL) {
       guarantee (w ->TState == ObjectWaiter::TS_ENTER, "invariant") ;
       ParkEvent * ev = w->_event ;
       w->TState = ObjectWaiter::TS_RUN ;
       OrderAccess::fence() ;
       ev->unpark() ;
   }
   return OS_OK ;
 }
 int ObjectMonitor::SimpleWait (Thread * Self, jlong millis) {
   guarantee (_owner == Self  , "invariant") ;
   guarantee (_recursions == 0, "invariant") ;
   ObjectWaiter Node (Self) ;
   Node._notified = 0 ;
   Node.TState    = ObjectWaiter::TS_WAIT ;
   RawMonitor_lock->lock_without_safepoint_check() ;
   Node._next     = _WaitSet ;
   _WaitSet       = &Node ;
   RawMonitor_lock->unlock() ;
   SimpleExit (Self) ;
   guarantee (_owner != Self, "invariant") ;
   int ret = OS_OK ;
   if (millis <= 0) {
     Self->_ParkEvent->park();
   } else {
     ret = Self->_ParkEvent->park(millis);
   }
   // If thread still resides on the waitset then unlink it.
   // Double-checked locking -- the usage is safe in this context
   // as we TState is volatile and the lock-unlock operators are
   // serializing (barrier-equivalent).
   if (Node.TState == ObjectWaiter::TS_WAIT) {
     RawMonitor_lock->lock_without_safepoint_check() ;
     if (Node.TState == ObjectWaiter::TS_WAIT) {
       // Simple O(n) unlink, but performance isn't critical here.
       ObjectWaiter * p ;
       ObjectWaiter * q = NULL ;
       for (p = _WaitSet ; p != &Node; p = p->_next) {
          q = p ;
       }
       guarantee (p == &Node, "invariant") ;
       if (q == NULL) {
         guarantee (p == _WaitSet, "invariant") ;
         _WaitSet = p->_next ;
       } else {
         guarantee (p == q->_next, "invariant") ;
         q->_next = p->_next ;
       }
       Node.TState = ObjectWaiter::TS_RUN ;
     }
     RawMonitor_lock->unlock() ;
   }
   guarantee (Node.TState == ObjectWaiter::TS_RUN, "invariant") ;
   SimpleEnter (Self) ;
   guarantee (_owner == Self, "invariant") ;
   guarantee (_recursions == 0, "invariant") ;
   return ret ;
 }
 int ObjectMonitor::SimpleNotify (Thread * Self, bool All) {
   guarantee (_owner == Self, "invariant") ;
   if (_WaitSet == NULL) return OS_OK ;
   // We have two options:
   // A. Transfer the threads from the WaitSet to the EntryList
   // B. Remove the thread from the WaitSet and unpark() it.
   //
   // We use (B), which is crude and results in lots of futile
   // context switching.  In particular (B) induces lots of contention.
   ParkEvent * ev = NULL ;       // consider using a small auto array ...
   RawMonitor_lock->lock_without_safepoint_check() ;
   for (;;) {
       ObjectWaiter * w = _WaitSet ;
       if (w == NULL) break ;
       _WaitSet = w->_next ;
       if (ev != NULL) { ev->unpark(); ev = NULL; }
       ev = w->_event ;
       OrderAccess::loadstore() ;
       w->TState = ObjectWaiter::TS_RUN ;
       OrderAccess::storeload();
       if (!All) break ;
   }
   RawMonitor_lock->unlock() ;
   if (ev != NULL) ev->unpark();
   return OS_OK ;
 }
 // Any JavaThread will enter here with state _thread_blocked
 int ObjectMonitor::raw_enter(TRAPS) {
   TEVENT (raw_enter) ;
   void * Contended ;
   // don't enter raw monitor if thread is being externally suspended, it will
   // surprise the suspender if a "suspended" thread can still enter monitor
   JavaThread * jt = (JavaThread *)THREAD;
   if (THREAD->is_Java_thread()) {
     jt->SR_lock()->lock_without_safepoint_check();
     while (jt->is_external_suspend()) {
       jt->SR_lock()->unlock();
       jt->java_suspend_self();
       jt->SR_lock()->lock_without_safepoint_check();
     }
     // guarded by SR_lock to avoid racing with new external suspend requests.
     Contended = Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) ;
     jt->SR_lock()->unlock();
   } else {
     Contended = Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) ;
   }
   if (Contended == THREAD) {
      _recursions ++ ;
      return OM_OK ;
   }
   if (Contended == NULL) {
      guarantee (_owner == THREAD, "invariant") ;
      guarantee (_recursions == 0, "invariant") ;
      return OM_OK ;
   }
   THREAD->set_current_pending_monitor(this);
   if (!THREAD->is_Java_thread()) {
      // No other non-Java threads besides VM thread would acquire
      // a raw monitor.
      assert(THREAD->is_VM_thread(), "must be VM thread");
      SimpleEnter (THREAD) ;
    } else {
      guarantee (jt->thread_state() == _thread_blocked, "invariant") ;
      for (;;) {
        jt->set_suspend_equivalent();
        // cleared by handle_special_suspend_equivalent_condition() or
        // java_suspend_self()
        SimpleEnter (THREAD) ;
        // were we externally suspended while we were waiting?
        if (!jt->handle_special_suspend_equivalent_condition()) break ;
        // This thread was externally suspended
        //
        // This logic isn't needed for JVMTI raw monitors,
        // but doesn't hurt just in case the suspend rules change. This
            // logic is needed for the ObjectMonitor.wait() reentry phase.
            // We have reentered the contended monitor, but while we were
            // waiting another thread suspended us. We don't want to reenter
            // the monitor while suspended because that would surprise the
            // thread that suspended us.
            //
            // Drop the lock -
        SimpleExit (THREAD) ;
            jt->java_suspend_self();
          }
      assert(_owner == THREAD, "Fatal error with monitor owner!");
      assert(_recursions == 0, "Fatal error with monitor recursions!");
   }
   THREAD->set_current_pending_monitor(NULL);
   guarantee (_recursions == 0, "invariant") ;
   return OM_OK;
 }
 // Used mainly for JVMTI raw monitor implementation
 // Also used for ObjectMonitor::wait().
 int ObjectMonitor::raw_exit(TRAPS) {
   TEVENT (raw_exit) ;
   if (THREAD != _owner) {
     return OM_ILLEGAL_MONITOR_STATE;
   }
   if (_recursions > 0) {
     --_recursions ;
     return OM_OK ;
   }
   void * List = _EntryList ;
   SimpleExit (THREAD) ;
   return OM_OK;
 }
 // Used for JVMTI raw monitor implementation.
 // All JavaThreads will enter here with state _thread_blocked
 int ObjectMonitor::raw_wait(jlong millis, bool interruptible, TRAPS) {
   TEVENT (raw_wait) ;
   if (THREAD != _owner) {
     return OM_ILLEGAL_MONITOR_STATE;
   }
   // To avoid spurious wakeups we reset the parkevent -- This is strictly optional.
   // The caller must be able to tolerate spurious returns from raw_wait().
   THREAD->_ParkEvent->reset() ;
   OrderAccess::fence() ;
   // check interrupt event
   if (interruptible && Thread::is_interrupted(THREAD, true)) {
     return OM_INTERRUPTED;
   }
   intptr_t save = _recursions ;
   _recursions = 0 ;
   _waiters ++ ;
   if (THREAD->is_Java_thread()) {
     guarantee (((JavaThread *) THREAD)->thread_state() == _thread_blocked, "invariant") ;
     ((JavaThread *)THREAD)->set_suspend_equivalent();
   }
   int rv = SimpleWait (THREAD, millis) ;
   _recursions = save ;
   _waiters -- ;
   guarantee (THREAD == _owner, "invariant") ;
   if (THREAD->is_Java_thread()) {
      JavaThread * jSelf = (JavaThread *) THREAD ;
      for (;;) {
         if (!jSelf->handle_special_suspend_equivalent_condition()) break ;
         SimpleExit (THREAD) ;
         jSelf->java_suspend_self();
         SimpleEnter (THREAD) ;
         jSelf->set_suspend_equivalent() ;
      }
   }
   guarantee (THREAD == _owner, "invariant") ;
   if (interruptible && Thread::is_interrupted(THREAD, true)) {
     return OM_INTERRUPTED;
   }
   return OM_OK ;
 }
 int ObjectMonitor::raw_notify(TRAPS) {
   TEVENT (raw_notify) ;
   if (THREAD != _owner) {
     return OM_ILLEGAL_MONITOR_STATE;
   }
   SimpleNotify (THREAD, false) ;
   return OM_OK;
 }
 int ObjectMonitor::raw_notifyAll(TRAPS) {
   TEVENT (raw_notifyAll) ;
   if (THREAD != _owner) {
     return OM_ILLEGAL_MONITOR_STATE;
   }
   SimpleNotify (THREAD, true) ;
   return OM_OK;
 }
 #ifndef PRODUCT
 void ObjectMonitor::verify() {
 }
 void ObjectMonitor::print() {
 }
 #endif
 //------------------------------------------------------------------------------
 // Non-product code
 #ifndef PRODUCT
 void ObjectSynchronizer::trace_locking(Handle locking_obj, bool is_compiled,
                                        bool is_method, bool is_locking) {
   // Don't know what to do here
 }
 // Verify all monitors in the monitor cache, the verification is weak.
 void ObjectSynchronizer::verify() {
   ObjectMonitor* block = gBlockList;
   ObjectMonitor* mid;
   while (block) {
     assert(block->object() == CHAINMARKER, "must be a block header");
     for (int i = 1; i < _BLOCKSIZE; i++) {
       mid = block + i;
       oop object = (oop) mid->object();
       if (object != NULL) {
         mid->verify();
       }
     }
     block = (ObjectMonitor*) block->FreeNext;
   }
 }
 // Check if monitor belongs to the monitor cache
 // The list is grow-only so it's *relatively* safe to traverse
 // the list of extant blocks without taking a lock.
 int ObjectSynchronizer::verify_objmon_isinpool(ObjectMonitor *monitor) {
   ObjectMonitor* block = gBlockList;
   while (block) {
     assert(block->object() == CHAINMARKER, "must be a block header");
     if (monitor > &block[0] && monitor < &block[_BLOCKSIZE]) {
       address mon = (address) monitor;
       address blk = (address) block;
       size_t diff = mon - blk;
       assert((diff % sizeof(ObjectMonitor)) == 0, "check");
       return 1;
     }
     block = (ObjectMonitor*) block->FreeNext;
   }
   return 0;
 }
 #endif

Mercurial > jdk8-mips64-public > hotspot / annotate

src/share/vm/runtime/synchronizer.cpp@0fbdb4381b99 (annotated)

src/share/vm/runtime/synchronizer.cpp