jdk8-mips64-public/hotspot: src/share/vm/runtime/mutex.cpp@63147986a428 (annotated)

src/share/vm/runtime/mutex.cpp@63147986a428 (annotated)

src/share/vm/runtime/mutex.cpp

Wed, 18 Sep 2013 07:02:10 -0700

author: dcubed
date: Wed, 18 Sep 2013 07:02:10 -0700
changeset 5743: 63147986a428
parent 5424: 5e3b6f79d280
child 6198: 55fb97c4c58d
permissions: -rw-r--r--

8019835: Strings interned in different threads equal but does not ==
Summary: Add -XX:+VerifyStringTableAtExit option and code to verify StringTable invariants.
Reviewed-by: rdurbin, sspitsyn, coleenp

 /*
  * Copyright (c) 1998, 2012, Oracle and/or its affiliates. All rights reserved.
  * DO NOT ALTER OR REMOVE COPYRIGHT NOTICES OR THIS FILE HEADER.
  *
  * This code is free software; you can redistribute it and/or modify it
  * under the terms of the GNU General Public License version 2 only, as
  * published by the Free Software Foundation.
  *
  * This code is distributed in the hope that it will be useful, but WITHOUT
  * ANY WARRANTY; without even the implied warranty of MERCHANTABILITY or
  * FITNESS FOR A PARTICULAR PURPOSE.  See the GNU General Public License
  * version 2 for more details (a copy is included in the LICENSE file that
  * accompanied this code).
  *
  * You should have received a copy of the GNU General Public License version
  * 2 along with this work; if not, write to the Free Software Foundation,
  * Inc., 51 Franklin St, Fifth Floor, Boston, MA 02110-1301 USA.
  *
  * Please contact Oracle, 500 Oracle Parkway, Redwood Shores, CA 94065 USA
  * or visit www.oracle.com if you need additional information or have any
  * questions.
  *
  */
 #include "precompiled.hpp"
 #include "runtime/mutex.hpp"
 #include "runtime/osThread.hpp"
 #include "runtime/thread.inline.hpp"
 #include "utilities/events.hpp"
 #ifdef TARGET_OS_FAMILY_linux
 # include "mutex_linux.inline.hpp"
 #endif
 #ifdef TARGET_OS_FAMILY_solaris
 # include "mutex_solaris.inline.hpp"
 #endif
 #ifdef TARGET_OS_FAMILY_windows
 # include "mutex_windows.inline.hpp"
 #endif
 #ifdef TARGET_OS_FAMILY_bsd
 # include "mutex_bsd.inline.hpp"
 #endif
 // o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o
 //
 // Native Monitor-Mutex locking - theory of operations
 //
 // * Native Monitors are completely unrelated to Java-level monitors,
 //   although the "back-end" slow-path implementations share a common lineage.
 //   See objectMonitor:: in synchronizer.cpp.
 //   Native Monitors do *not* support nesting or recursion but otherwise
 //   they're basically Hoare-flavor monitors.
 //
 // * A thread acquires ownership of a Monitor/Mutex by CASing the LockByte
 //   in the _LockWord from zero to non-zero.  Note that the _Owner field
 //   is advisory and is used only to verify that the thread calling unlock()
 //   is indeed the last thread to have acquired the lock.
 //
 // * Contending threads "push" themselves onto the front of the contention
 //   queue -- called the cxq -- with CAS and then spin/park.
 //   The _LockWord contains the LockByte as well as the pointer to the head
 //   of the cxq.  Colocating the LockByte with the cxq precludes certain races.
 //
 // * Using a separately addressable LockByte allows for CAS:MEMBAR or CAS:0
 //   idioms.  We currently use MEMBAR in the uncontended unlock() path, as
 //   MEMBAR often has less latency than CAS.  If warranted, we could switch to
 //   a CAS:0 mode, using timers to close the resultant race, as is done
 //   with Java Monitors in synchronizer.cpp.
 //
 //   See the following for a discussion of the relative cost of atomics (CAS)
 //   MEMBAR, and ways to eliminate such instructions from the common-case paths:
 //   -- http://blogs.sun.com/dave/entry/biased_locking_in_hotspot
 //   -- http://blogs.sun.com/dave/resource/MustangSync.pdf
 //   -- http://blogs.sun.com/dave/resource/synchronization-public2.pdf
 //   -- synchronizer.cpp
 //
 // * Overall goals - desiderata
 //   1. Minimize context switching
 //   2. Minimize lock migration
 //   3. Minimize CPI -- affinity and locality
 //   4. Minimize the execution of high-latency instructions such as CAS or MEMBAR
 //   5. Minimize outer lock hold times
 //   6. Behave gracefully on a loaded system
 //
 // * Thread flow and list residency:
 //
 //   Contention queue --> EntryList --> OnDeck --> Owner --> !Owner
 //   [..resident on monitor list..]
 //   [...........contending..................]
 //
 //   -- The contention queue (cxq) contains recently-arrived threads (RATs).
 //      Threads on the cxq eventually drain into the EntryList.
 //   -- Invariant: a thread appears on at most one list -- cxq, EntryList
 //      or WaitSet -- at any one time.
 //   -- For a given monitor there can be at most one "OnDeck" thread at any
 //      given time but if needbe this particular invariant could be relaxed.
 //
 // * The WaitSet and EntryList linked lists are composed of ParkEvents.
 //   I use ParkEvent instead of threads as ParkEvents are immortal and
 //   type-stable, meaning we can safely unpark() a possibly stale
 //   list element in the unlock()-path.  (That's benign).
 //
 // * Succession policy - providing for progress:
 //
 //   As necessary, the unlock()ing thread identifies, unlinks, and unparks
 //   an "heir presumptive" tentative successor thread from the EntryList.
 //   This becomes the so-called "OnDeck" thread, of which there can be only
 //   one at any given time for a given monitor.  The wakee will recontend
 //   for ownership of monitor.
 //
 //   Succession is provided for by a policy of competitive handoff.
 //   The exiting thread does _not_ grant or pass ownership to the
 //   successor thread.  (This is also referred to as "handoff" succession").
 //   Instead the exiting thread releases ownership and possibly wakes
 //   a successor, so the successor can (re)compete for ownership of the lock.
 //
 //   Competitive handoff provides excellent overall throughput at the expense
 //   of short-term fairness.  If fairness is a concern then one remedy might
 //   be to add an AcquireCounter field to the monitor.  After a thread acquires
 //   the lock it will decrement the AcquireCounter field.  When the count
 //   reaches 0 the thread would reset the AcquireCounter variable, abdicate
 //   the lock directly to some thread on the EntryList, and then move itself to the
 //   tail of the EntryList.
 //
 //   But in practice most threads engage or otherwise participate in resource
 //   bounded producer-consumer relationships, so lock domination is not usually
 //   a practical concern.  Recall too, that in general it's easier to construct
 //   a fair lock from a fast lock, but not vice-versa.
 //
 // * The cxq can have multiple concurrent "pushers" but only one concurrent
 //   detaching thread.  This mechanism is immune from the ABA corruption.
 //   More precisely, the CAS-based "push" onto cxq is ABA-oblivious.
 //   We use OnDeck as a pseudo-lock to enforce the at-most-one detaching
 //   thread constraint.
 //
 // * Taken together, the cxq and the EntryList constitute or form a
 //   single logical queue of threads stalled trying to acquire the lock.
 //   We use two distinct lists to reduce heat on the list ends.
 //   Threads in lock() enqueue onto cxq while threads in unlock() will
 //   dequeue from the EntryList.  (c.f. Michael Scott's "2Q" algorithm).
 //   A key desideratum is to minimize queue & monitor metadata manipulation
 //   that occurs while holding the "outer" monitor lock -- that is, we want to
 //   minimize monitor lock holds times.
 //
 //   The EntryList is ordered by the prevailing queue discipline and
 //   can be organized in any convenient fashion, such as a doubly-linked list or
 //   a circular doubly-linked list.  If we need a priority queue then something akin
 //   to Solaris' sleepq would work nicely.  Viz.,
 //   -- http://agg.eng/ws/on10_nightly/source/usr/src/uts/common/os/sleepq.c.
 //   -- http://cvs.opensolaris.org/source/xref/onnv/onnv-gate/usr/src/uts/common/os/sleepq.c
 //   Queue discipline is enforced at ::unlock() time, when the unlocking thread
 //   drains the cxq into the EntryList, and orders or reorders the threads on the
 //   EntryList accordingly.
 //
 //   Barring "lock barging", this mechanism provides fair cyclic ordering,
 //   somewhat similar to an elevator-scan.
 //
 // * OnDeck
 //   --  For a given monitor there can be at most one OnDeck thread at any given
 //       instant.  The OnDeck thread is contending for the lock, but has been
 //       unlinked from the EntryList and cxq by some previous unlock() operations.
 //       Once a thread has been designated the OnDeck thread it will remain so
 //       until it manages to acquire the lock -- being OnDeck is a stable property.
 //   --  Threads on the EntryList or cxq are _not allowed to attempt lock acquisition.
 //   --  OnDeck also serves as an "inner lock" as follows.  Threads in unlock() will, after
 //       having cleared the LockByte and dropped the outer lock,  attempt to "trylock"
 //       OnDeck by CASing the field from null to non-null.  If successful, that thread
 //       is then responsible for progress and succession and can use CAS to detach and
 //       drain the cxq into the EntryList.  By convention, only this thread, the holder of
 //       the OnDeck inner lock, can manipulate the EntryList or detach and drain the
 //       RATs on the cxq into the EntryList.  This avoids ABA corruption on the cxq as
 //       we allow multiple concurrent "push" operations but restrict detach concurrency
 //       to at most one thread.  Having selected and detached a successor, the thread then
 //       changes the OnDeck to refer to that successor, and then unparks the successor.
 //       That successor will eventually acquire the lock and clear OnDeck.  Beware
 //       that the OnDeck usage as a lock is asymmetric.  A thread in unlock() transiently
 //       "acquires" OnDeck, performs queue manipulations, passes OnDeck to some successor,
 //       and then the successor eventually "drops" OnDeck.  Note that there's never
 //       any sense of contention on the inner lock, however.  Threads never contend
 //       or wait for the inner lock.
 //   --  OnDeck provides for futile wakeup throttling a described in section 3.3 of
 //       See http://www.usenix.org/events/jvm01/full_papers/dice/dice.pdf
 //       In a sense, OnDeck subsumes the ObjectMonitor _Succ and ObjectWaiter
 //       TState fields found in Java-level objectMonitors.  (See synchronizer.cpp).
 //
 // * Waiting threads reside on the WaitSet list -- wait() puts
 //   the caller onto the WaitSet.  Notify() or notifyAll() simply
 //   transfers threads from the WaitSet to either the EntryList or cxq.
 //   Subsequent unlock() operations will eventually unpark the notifyee.
 //   Unparking a notifee in notify() proper is inefficient - if we were to do so
 //   it's likely the notifyee would simply impale itself on the lock held
 //   by the notifier.
 //
 // * The mechanism is obstruction-free in that if the holder of the transient
 //   OnDeck lock in unlock() is preempted or otherwise stalls, other threads
 //   can still acquire and release the outer lock and continue to make progress.
 //   At worst, waking of already blocked contending threads may be delayed,
 //   but nothing worse.  (We only use "trylock" operations on the inner OnDeck
 //   lock).
 //
 // * Note that thread-local storage must be initialized before a thread
 //   uses Native monitors or mutexes.  The native monitor-mutex subsystem
 //   depends on Thread::current().
 //
 // * The monitor synchronization subsystem avoids the use of native
 //   synchronization primitives except for the narrow platform-specific
 //   park-unpark abstraction.  See the comments in os_solaris.cpp regarding
 //   the semantics of park-unpark.  Put another way, this monitor implementation
 //   depends only on atomic operations and park-unpark.  The monitor subsystem
 //   manages all RUNNING->BLOCKED and BLOCKED->READY transitions while the
 //   underlying OS manages the READY<->RUN transitions.
 //
 // * The memory consistency model provide by lock()-unlock() is at least as
 //   strong or stronger than the Java Memory model defined by JSR-133.
 //   That is, we guarantee at least entry consistency, if not stronger.
 //   See http://g.oswego.edu/dl/jmm/cookbook.html.
 //
 // * Thread:: currently contains a set of purpose-specific ParkEvents:
 //   _MutexEvent, _ParkEvent, etc.  A better approach might be to do away with
 //   the purpose-specific ParkEvents and instead implement a general per-thread
 //   stack of available ParkEvents which we could provision on-demand.  The
 //   stack acts as a local cache to avoid excessive calls to ParkEvent::Allocate()
 //   and ::Release().  A thread would simply pop an element from the local stack before it
 //   enqueued or park()ed.  When the contention was over the thread would
 //   push the no-longer-needed ParkEvent back onto its stack.
 //
 // * A slightly reduced form of ILock() and IUnlock() have been partially
 //   model-checked (Murphi) for safety and progress at T=1,2,3 and 4.
 //   It'd be interesting to see if TLA/TLC could be useful as well.
 //
 // * Mutex-Monitor is a low-level "leaf" subsystem.  That is, the monitor
 //   code should never call other code in the JVM that might itself need to
 //   acquire monitors or mutexes.  That's true *except* in the case of the
 //   ThreadBlockInVM state transition wrappers.  The ThreadBlockInVM DTOR handles
 //   mutator reentry (ingress) by checking for a pending safepoint in which case it will
 //   call SafepointSynchronize::block(), which in turn may call Safepoint_lock->lock(), etc.
 //   In that particular case a call to lock() for a given Monitor can end up recursively
 //   calling lock() on another monitor.   While distasteful, this is largely benign
 //   as the calls come from jacket that wraps lock(), and not from deep within lock() itself.
 //
 //   It's unfortunate that native mutexes and thread state transitions were convolved.
 //   They're really separate concerns and should have remained that way.  Melding
 //   them together was facile -- a bit too facile.   The current implementation badly
 //   conflates the two concerns.
 //
 // * TODO-FIXME:
 //
 //   -- Add DTRACE probes for contended acquire, contended acquired, contended unlock
 //      We should also add DTRACE probes in the ParkEvent subsystem for
 //      Park-entry, Park-exit, and Unpark.
 //
 //   -- We have an excess of mutex-like constructs in the JVM, namely:
 //      1. objectMonitors for Java-level synchronization (synchronizer.cpp)
 //      2. low-level muxAcquire and muxRelease
 //      3. low-level spinAcquire and spinRelease
 //      4. native Mutex:: and Monitor::
 //      5. jvm_raw_lock() and _unlock()
 //      6. JVMTI raw monitors -- distinct from (5) despite having a confusingly
 //         similar name.
 //
 // o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o-o
 // CASPTR() uses the canonical argument order that dominates in the literature.
 // Our internal cmpxchg_ptr() uses a bastardized ordering to accommodate Sun .il templates.
 #define CASPTR(a,c,s) intptr_t(Atomic::cmpxchg_ptr ((void *)(s),(void *)(a),(void *)(c)))
 #define UNS(x) (uintptr_t(x))
 #define TRACE(m) { static volatile int ctr = 0 ; int x = ++ctr ; if ((x & (x-1))==0) { ::printf ("%d:%s\n", x, #m); ::fflush(stdout); }}
 // Simplistic low-quality Marsaglia SHIFT-XOR RNG.
 // Bijective except for the trailing mask operation.
 // Useful for spin loops as the compiler can't optimize it away.
 static inline jint MarsagliaXORV (jint x) {
   if (x == 0) x = 1|os::random() ;
   x ^= x << 6;
   x ^= ((unsigned)x) >> 21;
   x ^= x << 7 ;
   return x & 0x7FFFFFFF ;
 }
 static inline jint MarsagliaXOR (jint * const a) {
   jint x = *a ;
   if (x == 0) x = UNS(a)|1 ;
   x ^= x << 6;
   x ^= ((unsigned)x) >> 21;
   x ^= x << 7 ;
   *a = x ;
   return x & 0x7FFFFFFF ;
 }
 static int Stall (int its) {
   static volatile jint rv = 1 ;
   volatile int OnFrame = 0 ;
   jint v = rv ^ UNS(OnFrame) ;
   while (--its >= 0) {
     v = MarsagliaXORV (v) ;
   }
   // Make this impossible for the compiler to optimize away,
   // but (mostly) avoid W coherency sharing on MP systems.
   if (v == 0x12345) rv = v ;
   return v ;
 }
 int Monitor::TryLock () {
   intptr_t v = _LockWord.FullWord ;
   for (;;) {
     if ((v & _LBIT) != 0) return 0 ;
     const intptr_t u = CASPTR (&_LockWord, v, v|_LBIT) ;
     if (v == u) return 1 ;
     v = u ;
   }
 }
 int Monitor::TryFast () {
   // Optimistic fast-path form ...
   // Fast-path attempt for the common uncontended case.
   // Avoid RTS->RTO $ coherence upgrade on typical SMP systems.
   intptr_t v = CASPTR (&_LockWord, 0, _LBIT) ;  // agro ...
   if (v == 0) return 1 ;
   for (;;) {
     if ((v & _LBIT) != 0) return 0 ;
     const intptr_t u = CASPTR (&_LockWord, v, v|_LBIT) ;
     if (v == u) return 1 ;
     v = u ;
   }
 }
 int Monitor::ILocked () {
   const intptr_t w = _LockWord.FullWord & 0xFF ;
   assert (w == 0 || w == _LBIT, "invariant") ;
   return w == _LBIT ;
 }
 // Polite TATAS spinlock with exponential backoff - bounded spin.
 // Ideally we'd use processor cycles, time or vtime to control
 // the loop, but we currently use iterations.
 // All the constants within were derived empirically but work over
 // over the spectrum of J2SE reference platforms.
 // On Niagara-class systems the back-off is unnecessary but
 // is relatively harmless.  (At worst it'll slightly retard
 // acquisition times).  The back-off is critical for older SMP systems
 // where constant fetching of the LockWord would otherwise impair
 // scalability.
 //
 // Clamp spinning at approximately 1/2 of a context-switch round-trip.
 // See synchronizer.cpp for details and rationale.
 int Monitor::TrySpin (Thread * const Self) {
   if (TryLock())    return 1 ;
   if (!os::is_MP()) return 0 ;
   int Probes  = 0 ;
   int Delay   = 0 ;
   int Steps   = 0 ;
   int SpinMax = NativeMonitorSpinLimit ;
   int flgs    = NativeMonitorFlags ;
   for (;;) {
     intptr_t v = _LockWord.FullWord;
     if ((v & _LBIT) == 0) {
       if (CASPTR (&_LockWord, v, v|_LBIT) == v) {
         return 1 ;
       }
       continue ;
     }
     if ((flgs & 8) == 0) {
       SpinPause () ;
     }
     // Periodically increase Delay -- variable Delay form
     // conceptually: delay *= 1 + 1/Exponent
     ++ Probes;
     if (Probes > SpinMax) return 0 ;
     if ((Probes & 0x7) == 0) {
       Delay = ((Delay << 1)|1) & 0x7FF ;
       // CONSIDER: Delay += 1 + (Delay/4); Delay &= 0x7FF ;
     }
     if (flgs & 2) continue ;
     // Consider checking _owner's schedctl state, if OFFPROC abort spin.
     // If the owner is OFFPROC then it's unlike that the lock will be dropped
     // in a timely fashion, which suggests that spinning would not be fruitful
     // or profitable.
     // Stall for "Delay" time units - iterations in the current implementation.
     // Avoid generating coherency traffic while stalled.
     // Possible ways to delay:
     //   PAUSE, SLEEP, MEMBAR #sync, MEMBAR #halt,
     //   wr %g0,%asi, gethrtime, rdstick, rdtick, rdtsc, etc. ...
     // Note that on Niagara-class systems we want to minimize STs in the
     // spin loop.  N1 and brethren write-around the L1$ over the xbar into the L2$.
     // Furthermore, they don't have a W$ like traditional SPARC processors.
     // We currently use a Marsaglia Shift-Xor RNG loop.
     Steps += Delay ;
     if (Self != NULL) {
       jint rv = Self->rng[0] ;
       for (int k = Delay ; --k >= 0; ) {
         rv = MarsagliaXORV (rv) ;
         if ((flgs & 4) == 0 && SafepointSynchronize::do_call_back()) return 0 ;
       }
       Self->rng[0] = rv ;
     } else {
       Stall (Delay) ;
     }
   }
 }
 static int ParkCommon (ParkEvent * ev, jlong timo) {
   // Diagnostic support - periodically unwedge blocked threads
   intx nmt = NativeMonitorTimeout ;
   if (nmt > 0 && (nmt < timo || timo <= 0)) {
      timo = nmt ;
   }
   int err = OS_OK ;
   if (0 == timo) {
     ev->park() ;
   } else {
     err = ev->park(timo) ;
   }
   return err ;
 }
 inline int Monitor::AcquireOrPush (ParkEvent * ESelf) {
   intptr_t v = _LockWord.FullWord ;
   for (;;) {
     if ((v & _LBIT) == 0) {
       const intptr_t u = CASPTR (&_LockWord, v, v|_LBIT) ;
       if (u == v) return 1 ;        // indicate acquired
       v = u ;
     } else {
       // Anticipate success ...
       ESelf->ListNext = (ParkEvent *) (v & ~_LBIT) ;
       const intptr_t u = CASPTR (&_LockWord, v, intptr_t(ESelf)|_LBIT) ;
       if (u == v) return 0 ;        // indicate pushed onto cxq
       v = u ;
     }
     // Interference - LockWord change - just retry
   }
 }
 // ILock and IWait are the lowest level primitive internal blocking
 // synchronization functions.  The callers of IWait and ILock must have
 // performed any needed state transitions beforehand.
 // IWait and ILock may directly call park() without any concern for thread state.
 // Note that ILock and IWait do *not* access _owner.
 // _owner is a higher-level logical concept.
 void Monitor::ILock (Thread * Self) {
   assert (_OnDeck != Self->_MutexEvent, "invariant") ;
   if (TryFast()) {
  Exeunt:
     assert (ILocked(), "invariant") ;
     return ;
   }
   ParkEvent * const ESelf = Self->_MutexEvent ;
   assert (_OnDeck != ESelf, "invariant") ;
   // As an optimization, spinners could conditionally try to set ONDECK to _LBIT
   // Synchronizer.cpp uses a similar optimization.
   if (TrySpin (Self)) goto Exeunt ;
   // Slow-path - the lock is contended.
   // Either Enqueue Self on cxq or acquire the outer lock.
   // LockWord encoding = (cxq,LOCKBYTE)
   ESelf->reset() ;
   OrderAccess::fence() ;
   // Optional optimization ... try barging on the inner lock
   if ((NativeMonitorFlags & 32) && CASPTR (&_OnDeck, NULL, UNS(Self)) == 0) {
     goto OnDeck_LOOP ;
   }
   if (AcquireOrPush (ESelf)) goto Exeunt ;
   // At any given time there is at most one ondeck thread.
   // ondeck implies not resident on cxq and not resident on EntryList
   // Only the OnDeck thread can try to acquire -- contended for -- the lock.
   // CONSIDER: use Self->OnDeck instead of m->OnDeck.
   // Deschedule Self so that others may run.
   while (_OnDeck != ESelf) {
     ParkCommon (ESelf, 0) ;
   }
   // Self is now in the ONDECK position and will remain so until it
   // manages to acquire the lock.
  OnDeck_LOOP:
   for (;;) {
     assert (_OnDeck == ESelf, "invariant") ;
     if (TrySpin (Self)) break ;
     // CONSIDER: if ESelf->TryPark() && TryLock() break ...
     // It's probably wise to spin only if we *actually* blocked
     // CONSIDER: check the lockbyte, if it remains set then
     // preemptively drain the cxq into the EntryList.
     // The best place and time to perform queue operations -- lock metadata --
     // is _before having acquired the outer lock, while waiting for the lock to drop.
     ParkCommon (ESelf, 0) ;
   }
   assert (_OnDeck == ESelf, "invariant") ;
   _OnDeck = NULL ;
   // Note that we current drop the inner lock (clear OnDeck) in the slow-path
   // epilog immediately after having acquired the outer lock.
   // But instead we could consider the following optimizations:
   // A. Shift or defer dropping the inner lock until the subsequent IUnlock() operation.
   //    This might avoid potential reacquisition of the inner lock in IUlock().
   // B. While still holding the inner lock, attempt to opportunistically select
   //    and unlink the next ONDECK thread from the EntryList.
   //    If successful, set ONDECK to refer to that thread, otherwise clear ONDECK.
   //    It's critical that the select-and-unlink operation run in constant-time as
   //    it executes when holding the outer lock and may artificially increase the
   //    effective length of the critical section.
   // Note that (A) and (B) are tantamount to succession by direct handoff for
   // the inner lock.
   goto Exeunt ;
 }
 void Monitor::IUnlock (bool RelaxAssert) {
   assert (ILocked(), "invariant") ;
   // Conceptually we need a MEMBAR #storestore|#loadstore barrier or fence immediately
   // before the store that releases the lock.  Crucially, all the stores and loads in the
   // critical section must be globally visible before the store of 0 into the lock-word
   // that releases the lock becomes globally visible.  That is, memory accesses in the
   // critical section should not be allowed to bypass or overtake the following ST that
   // releases the lock.  As such, to prevent accesses within the critical section
   // from "leaking" out, we need a release fence between the critical section and the
   // store that releases the lock.  In practice that release barrier is elided on
   // platforms with strong memory models such as TSO.
   //
   // Note that the OrderAccess::storeload() fence that appears after unlock store
   // provides for progress conditions and succession and is _not related to exclusion
   // safety or lock release consistency.
   OrderAccess::release_store(&_LockWord.Bytes[_LSBINDEX], 0); // drop outer lock
   OrderAccess::storeload ();
   ParkEvent * const w = _OnDeck ;
   assert (RelaxAssert || w != Thread::current()->_MutexEvent, "invariant") ;
   if (w != NULL) {
     // Either we have a valid ondeck thread or ondeck is transiently "locked"
     // by some exiting thread as it arranges for succession.  The LSBit of
     // OnDeck allows us to discriminate two cases.  If the latter, the
     // responsibility for progress and succession lies with that other thread.
     // For good performance, we also depend on the fact that redundant unpark()
     // operations are cheap.  That is, repeated Unpark()ing of the ONDECK thread
     // is inexpensive.  This approach provides implicit futile wakeup throttling.
     // Note that the referent "w" might be stale with respect to the lock.
     // In that case the following unpark() is harmless and the worst that'll happen
     // is a spurious return from a park() operation.  Critically, if "w" _is stale,
     // then progress is known to have occurred as that means the thread associated
     // with "w" acquired the lock.  In that case this thread need take no further
     // action to guarantee progress.
     if ((UNS(w) & _LBIT) == 0) w->unpark() ;
     return ;
   }
   intptr_t cxq = _LockWord.FullWord ;
   if (((cxq & ~_LBIT)|UNS(_EntryList)) == 0) {
     return ;      // normal fast-path exit - cxq and EntryList both empty
   }
   if (cxq & _LBIT) {
     // Optional optimization ...
     // Some other thread acquired the lock in the window since this
     // thread released it.  Succession is now that thread's responsibility.
     return ;
   }
  Succession:
   // Slow-path exit - this thread must ensure succession and progress.
   // OnDeck serves as lock to protect cxq and EntryList.
   // Only the holder of OnDeck can manipulate EntryList or detach the RATs from cxq.
   // Avoid ABA - allow multiple concurrent producers (enqueue via push-CAS)
   // but only one concurrent consumer (detacher of RATs).
   // Consider protecting this critical section with schedctl on Solaris.
   // Unlike a normal lock, however, the exiting thread "locks" OnDeck,
   // picks a successor and marks that thread as OnDeck.  That successor
   // thread will then clear OnDeck once it eventually acquires the outer lock.
   if (CASPTR (&_OnDeck, NULL, _LBIT) != UNS(NULL)) {
     return ;
   }
   ParkEvent * List = _EntryList ;
   if (List != NULL) {
     // Transfer the head of the EntryList to the OnDeck position.
     // Once OnDeck, a thread stays OnDeck until it acquires the lock.
     // For a given lock there is at most OnDeck thread at any one instant.
    WakeOne:
     assert (List == _EntryList, "invariant") ;
     ParkEvent * const w = List ;
     assert (RelaxAssert || w != Thread::current()->_MutexEvent, "invariant") ;
     _EntryList = w->ListNext ;
     // as a diagnostic measure consider setting w->_ListNext = BAD
     assert (UNS(_OnDeck) == _LBIT, "invariant") ;
     _OnDeck = w ;           // pass OnDeck to w.
                             // w will clear OnDeck once it acquires the outer lock
     // Another optional optimization ...
     // For heavily contended locks it's not uncommon that some other
     // thread acquired the lock while this thread was arranging succession.
     // Try to defer the unpark() operation - Delegate the responsibility
     // for unpark()ing the OnDeck thread to the current or subsequent owners
     // That is, the new owner is responsible for unparking the OnDeck thread.
     OrderAccess::storeload() ;
     cxq = _LockWord.FullWord ;
     if (cxq & _LBIT) return ;
     w->unpark() ;
     return ;
   }
   cxq = _LockWord.FullWord ;
   if ((cxq & ~_LBIT) != 0) {
     // The EntryList is empty but the cxq is populated.
     // drain RATs from cxq into EntryList
     // Detach RATs segment with CAS and then merge into EntryList
     for (;;) {
       // optional optimization - if locked, the owner is responsible for succession
       if (cxq & _LBIT) goto Punt ;
       const intptr_t vfy = CASPTR (&_LockWord, cxq, cxq & _LBIT) ;
       if (vfy == cxq) break ;
       cxq = vfy ;
       // Interference - LockWord changed - Just retry
       // We can see concurrent interference from contending threads
       // pushing themselves onto the cxq or from lock-unlock operations.
       // From the perspective of this thread, EntryList is stable and
       // the cxq is prepend-only -- the head is volatile but the interior
       // of the cxq is stable.  In theory if we encounter interference from threads
       // pushing onto cxq we could simply break off the original cxq suffix and
       // move that segment to the EntryList, avoiding a 2nd or multiple CAS attempts
       // on the high-traffic LockWord variable.   For instance lets say the cxq is "ABCD"
       // when we first fetch cxq above.  Between the fetch -- where we observed "A"
       // -- and CAS -- where we attempt to CAS null over A -- "PQR" arrive,
       // yielding cxq = "PQRABCD".  In this case we could simply set A.ListNext
       // null, leaving cxq = "PQRA" and transfer the "BCD" segment to the EntryList.
       // Note too, that it's safe for this thread to traverse the cxq
       // without taking any special concurrency precautions.
     }
     // We don't currently reorder the cxq segment as we move it onto
     // the EntryList, but it might make sense to reverse the order
     // or perhaps sort by thread priority.  See the comments in
     // synchronizer.cpp objectMonitor::exit().
     assert (_EntryList == NULL, "invariant") ;
     _EntryList = List = (ParkEvent *)(cxq & ~_LBIT) ;
     assert (List != NULL, "invariant") ;
     goto WakeOne ;
   }
   // cxq|EntryList is empty.
   // w == NULL implies that cxq|EntryList == NULL in the past.
   // Possible race - rare inopportune interleaving.
   // A thread could have added itself to cxq since this thread previously checked.
   // Detect and recover by refetching cxq.
  Punt:
   assert (UNS(_OnDeck) == _LBIT, "invariant") ;
   _OnDeck = NULL ;            // Release inner lock.
   OrderAccess::storeload();   // Dekker duality - pivot point
   // Resample LockWord/cxq to recover from possible race.
   // For instance, while this thread T1 held OnDeck, some other thread T2 might
   // acquire the outer lock.  Another thread T3 might try to acquire the outer
   // lock, but encounter contention and enqueue itself on cxq.  T2 then drops the
   // outer lock, but skips succession as this thread T1 still holds OnDeck.
   // T1 is and remains responsible for ensuring succession of T3.
   //
   // Note that we don't need to recheck EntryList, just cxq.
   // If threads moved onto EntryList since we dropped OnDeck
   // that implies some other thread forced succession.
   cxq = _LockWord.FullWord ;
   if ((cxq & ~_LBIT) != 0 && (cxq & _LBIT) == 0) {
     goto Succession ;         // potential race -- re-run succession
   }
   return ;
 }
 bool Monitor::notify() {
   assert (_owner == Thread::current(), "invariant") ;
   assert (ILocked(), "invariant") ;
   if (_WaitSet == NULL) return true ;
   NotifyCount ++ ;
   // Transfer one thread from the WaitSet to the EntryList or cxq.
   // Currently we just unlink the head of the WaitSet and prepend to the cxq.
   // And of course we could just unlink it and unpark it, too, but
   // in that case it'd likely impale itself on the reentry.
   Thread::muxAcquire (_WaitLock, "notify:WaitLock") ;
   ParkEvent * nfy = _WaitSet ;
   if (nfy != NULL) {                  // DCL idiom
     _WaitSet = nfy->ListNext ;
     assert (nfy->Notified == 0, "invariant") ;
     // push nfy onto the cxq
     for (;;) {
       const intptr_t v = _LockWord.FullWord ;
       assert ((v & 0xFF) == _LBIT, "invariant") ;
       nfy->ListNext = (ParkEvent *)(v & ~_LBIT);
       if (CASPTR (&_LockWord, v, UNS(nfy)|_LBIT) == v) break;
       // interference - _LockWord changed -- just retry
     }
     // Note that setting Notified before pushing nfy onto the cxq is
     // also legal and safe, but the safety properties are much more
     // subtle, so for the sake of code stewardship ...
     OrderAccess::fence() ;
     nfy->Notified = 1;
   }
   Thread::muxRelease (_WaitLock) ;
   if (nfy != NULL && (NativeMonitorFlags & 16)) {
     // Experimental code ... light up the wakee in the hope that this thread (the owner)
     // will drop the lock just about the time the wakee comes ONPROC.
     nfy->unpark() ;
   }
   assert (ILocked(), "invariant") ;
   return true ;
 }
 // Currently notifyAll() transfers the waiters one-at-a-time from the waitset
 // to the cxq.  This could be done more efficiently with a single bulk en-mass transfer,
 // but in practice notifyAll() for large #s of threads is rare and not time-critical.
 // Beware too, that we invert the order of the waiters.  Lets say that the
 // waitset is "ABCD" and the cxq is "XYZ".  After a notifyAll() the waitset
 // will be empty and the cxq will be "DCBAXYZ".  This is benign, of course.
 bool Monitor::notify_all() {
   assert (_owner == Thread::current(), "invariant") ;
   assert (ILocked(), "invariant") ;
   while (_WaitSet != NULL) notify() ;
   return true ;
 }
 int Monitor::IWait (Thread * Self, jlong timo) {
   assert (ILocked(), "invariant") ;
   // Phases:
   // 1. Enqueue Self on WaitSet - currently prepend
   // 2. unlock - drop the outer lock
   // 3. wait for either notification or timeout
   // 4. lock - reentry - reacquire the outer lock
   ParkEvent * const ESelf = Self->_MutexEvent ;
   ESelf->Notified = 0 ;
   ESelf->reset() ;
   OrderAccess::fence() ;
   // Add Self to WaitSet
   // Ideally only the holder of the outer lock would manipulate the WaitSet -
   // That is, the outer lock would implicitly protect the WaitSet.
   // But if a thread in wait() encounters a timeout it will need to dequeue itself
   // from the WaitSet _before it becomes the owner of the lock.  We need to dequeue
   // as the ParkEvent -- which serves as a proxy for the thread -- can't reside
   // on both the WaitSet and the EntryList|cxq at the same time..  That is, a thread
   // on the WaitSet can't be allowed to compete for the lock until it has managed to
   // unlink its ParkEvent from WaitSet.  Thus the need for WaitLock.
   // Contention on the WaitLock is minimal.
   //
   // Another viable approach would be add another ParkEvent, "WaitEvent" to the
   // thread class.  The WaitSet would be composed of WaitEvents.  Only the
   // owner of the outer lock would manipulate the WaitSet.  A thread in wait()
   // could then compete for the outer lock, and then, if necessary, unlink itself
   // from the WaitSet only after having acquired the outer lock.  More precisely,
   // there would be no WaitLock.  A thread in in wait() would enqueue its WaitEvent
   // on the WaitSet; release the outer lock; wait for either notification or timeout;
   // reacquire the inner lock; and then, if needed, unlink itself from the WaitSet.
   //
   // Alternatively, a 2nd set of list link fields in the ParkEvent might suffice.
   // One set would be for the WaitSet and one for the EntryList.
   // We could also deconstruct the ParkEvent into a "pure" event and add a
   // new immortal/TSM "ListElement" class that referred to ParkEvents.
   // In that case we could have one ListElement on the WaitSet and another
   // on the EntryList, with both referring to the same pure Event.
   Thread::muxAcquire (_WaitLock, "wait:WaitLock:Add") ;
   ESelf->ListNext = _WaitSet ;
   _WaitSet = ESelf ;
   Thread::muxRelease (_WaitLock) ;
   // Release the outer lock
   // We call IUnlock (RelaxAssert=true) as a thread T1 might
   // enqueue itself on the WaitSet, call IUnlock(), drop the lock,
   // and then stall before it can attempt to wake a successor.
   // Some other thread T2 acquires the lock, and calls notify(), moving
   // T1 from the WaitSet to the cxq.  T2 then drops the lock.  T1 resumes,
   // and then finds *itself* on the cxq.  During the course of a normal
   // IUnlock() call a thread should _never find itself on the EntryList
   // or cxq, but in the case of wait() it's possible.
   // See synchronizer.cpp objectMonitor::wait().
   IUnlock (true) ;
   // Wait for either notification or timeout
   // Beware that in some circumstances we might propagate
   // spurious wakeups back to the caller.
   for (;;) {
     if (ESelf->Notified) break ;
     int err = ParkCommon (ESelf, timo) ;
     if (err == OS_TIMEOUT || (NativeMonitorFlags & 1)) break ;
   }
   // Prepare for reentry - if necessary, remove ESelf from WaitSet
   // ESelf can be:
   // 1. Still on the WaitSet.  This can happen if we exited the loop by timeout.
   // 2. On the cxq or EntryList
   // 3. Not resident on cxq, EntryList or WaitSet, but in the OnDeck position.
   OrderAccess::fence() ;
   int WasOnWaitSet = 0 ;
   if (ESelf->Notified == 0) {
     Thread::muxAcquire (_WaitLock, "wait:WaitLock:remove") ;
     if (ESelf->Notified == 0) {     // DCL idiom
       assert (_OnDeck != ESelf, "invariant") ;   // can't be both OnDeck and on WaitSet
       // ESelf is resident on the WaitSet -- unlink it.
       // A doubly-linked list would be better here so we can unlink in constant-time.
       // We have to unlink before we potentially recontend as ESelf might otherwise
       // end up on the cxq|EntryList -- it can't be on two lists at once.
       ParkEvent * p = _WaitSet ;
       ParkEvent * q = NULL ;            // classic q chases p
       while (p != NULL && p != ESelf) {
         q = p ;
         p = p->ListNext ;
       }
       assert (p == ESelf, "invariant") ;
       if (p == _WaitSet) {      // found at head
         assert (q == NULL, "invariant") ;
         _WaitSet = p->ListNext ;
       } else {                  // found in interior
         assert (q->ListNext == p, "invariant") ;
         q->ListNext = p->ListNext ;
       }
       WasOnWaitSet = 1 ;        // We were *not* notified but instead encountered timeout
     }
     Thread::muxRelease (_WaitLock) ;
   }
   // Reentry phase - reacquire the lock
   if (WasOnWaitSet) {
     // ESelf was previously on the WaitSet but we just unlinked it above
     // because of a timeout.  ESelf is not resident on any list and is not OnDeck
     assert (_OnDeck != ESelf, "invariant") ;
     ILock (Self) ;
   } else {
     // A prior notify() operation moved ESelf from the WaitSet to the cxq.
     // ESelf is now on the cxq, EntryList or at the OnDeck position.
     // The following fragment is extracted from Monitor::ILock()
     for (;;) {
       if (_OnDeck == ESelf && TrySpin(Self)) break ;
       ParkCommon (ESelf, 0) ;
     }
     assert (_OnDeck == ESelf, "invariant") ;
     _OnDeck = NULL ;
   }
   assert (ILocked(), "invariant") ;
   return WasOnWaitSet != 0 ;        // return true IFF timeout
 }
 // ON THE VMTHREAD SNEAKING PAST HELD LOCKS:
 // In particular, there are certain types of global lock that may be held
 // by a Java thread while it is blocked at a safepoint but before it has
 // written the _owner field. These locks may be sneakily acquired by the
 // VM thread during a safepoint to avoid deadlocks. Alternatively, one should
 // identify all such locks, and ensure that Java threads never block at
 // safepoints while holding them (_no_safepoint_check_flag). While it
 // seems as though this could increase the time to reach a safepoint
 // (or at least increase the mean, if not the variance), the latter
 // approach might make for a cleaner, more maintainable JVM design.
 //
 // Sneaking is vile and reprehensible and should be excised at the 1st
 // opportunity.  It's possible that the need for sneaking could be obviated
 // as follows.  Currently, a thread might (a) while TBIVM, call pthread_mutex_lock
 // or ILock() thus acquiring the "physical" lock underlying Monitor/Mutex.
 // (b) stall at the TBIVM exit point as a safepoint is in effect.  Critically,
 // it'll stall at the TBIVM reentry state transition after having acquired the
 // underlying lock, but before having set _owner and having entered the actual
 // critical section.  The lock-sneaking facility leverages that fact and allowed the
 // VM thread to logically acquire locks that had already be physically locked by mutators
 // but where mutators were known blocked by the reentry thread state transition.
 //
 // If we were to modify the Monitor-Mutex so that TBIVM state transitions tightly
 // wrapped calls to park(), then we could likely do away with sneaking.  We'd
 // decouple lock acquisition and parking.  The critical invariant  to eliminating
 // sneaking is to ensure that we never "physically" acquire the lock while TBIVM.
 // An easy way to accomplish this is to wrap the park calls in a narrow TBIVM jacket.
 // One difficulty with this approach is that the TBIVM wrapper could recurse and
 // call lock() deep from within a lock() call, while the MutexEvent was already enqueued.
 // Using a stack (N=2 at minimum) of ParkEvents would take care of that problem.
 //
 // But of course the proper ultimate approach is to avoid schemes that require explicit
 // sneaking or dependence on any any clever invariants or subtle implementation properties
 // of Mutex-Monitor and instead directly address the underlying design flaw.
 void Monitor::lock (Thread * Self) {
 #ifdef CHECK_UNHANDLED_OOPS
   // Clear unhandled oops so we get a crash right away.  Only clear for non-vm
   // or GC threads.
   if (Self->is_Java_thread()) {
     Self->clear_unhandled_oops();
   }
 #endif // CHECK_UNHANDLED_OOPS
   debug_only(check_prelock_state(Self));
   assert (_owner != Self              , "invariant") ;
   assert (_OnDeck != Self->_MutexEvent, "invariant") ;
   if (TryFast()) {
  Exeunt:
     assert (ILocked(), "invariant") ;
     assert (owner() == NULL, "invariant");
     set_owner (Self);
     return ;
   }
   // The lock is contended ...
   bool can_sneak = Self->is_VM_thread() && SafepointSynchronize::is_at_safepoint();
   if (can_sneak && _owner == NULL) {
     // a java thread has locked the lock but has not entered the
     // critical region -- let's just pretend we've locked the lock
     // and go on.  we note this with _snuck so we can also
     // pretend to unlock when the time comes.
     _snuck = true;
     goto Exeunt ;
   }
   // Try a brief spin to avoid passing thru thread state transition ...
   if (TrySpin (Self)) goto Exeunt ;
   check_block_state(Self);
   if (Self->is_Java_thread()) {
     // Horribile dictu - we suffer through a state transition
     assert(rank() > Mutex::special, "Potential deadlock with special or lesser rank mutex");
     ThreadBlockInVM tbivm ((JavaThread *) Self) ;
     ILock (Self) ;
   } else {
     // Mirabile dictu
     ILock (Self) ;
   }
   goto Exeunt ;
 }
 void Monitor::lock() {
   this->lock(Thread::current());
 }
 // Lock without safepoint check - a degenerate variant of lock().
 // Should ONLY be used by safepoint code and other code
 // that is guaranteed not to block while running inside the VM. If this is called with
 // thread state set to be in VM, the safepoint synchronization code will deadlock!
 void Monitor::lock_without_safepoint_check (Thread * Self) {
   assert (_owner != Self, "invariant") ;
   ILock (Self) ;
   assert (_owner == NULL, "invariant");
   set_owner (Self);
 }
 void Monitor::lock_without_safepoint_check () {
   lock_without_safepoint_check (Thread::current()) ;
 }
 // Returns true if thread succeceed [sic] in grabbing the lock, otherwise false.
 bool Monitor::try_lock() {
   Thread * const Self = Thread::current();
   debug_only(check_prelock_state(Self));
   // assert(!thread->is_inside_signal_handler(), "don't lock inside signal handler");
   // Special case, where all Java threads are stopped.
   // The lock may have been acquired but _owner is not yet set.
   // In that case the VM thread can safely grab the lock.
   // It strikes me this should appear _after the TryLock() fails, below.
   bool can_sneak = Self->is_VM_thread() && SafepointSynchronize::is_at_safepoint();
   if (can_sneak && _owner == NULL) {
     set_owner(Self); // Do not need to be atomic, since we are at a safepoint
     _snuck = true;
     return true;
   }
   if (TryLock()) {
     // We got the lock
     assert (_owner == NULL, "invariant");
     set_owner (Self);
     return true;
   }
   return false;
 }
 void Monitor::unlock() {
   assert (_owner  == Thread::current(), "invariant") ;
   assert (_OnDeck != Thread::current()->_MutexEvent , "invariant") ;
   set_owner (NULL) ;
   if (_snuck) {
     assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak");
     _snuck = false;
     return ;
   }
   IUnlock (false) ;
 }
 // Yet another degenerate version of Monitor::lock() or lock_without_safepoint_check()
 // jvm_raw_lock() and _unlock() can be called by non-Java threads via JVM_RawMonitorEnter.
 //
 // There's no expectation that JVM_RawMonitors will interoperate properly with the native
 // Mutex-Monitor constructs.  We happen to implement JVM_RawMonitors in terms of
 // native Mutex-Monitors simply as a matter of convenience.  A simple abstraction layer
 // over a pthread_mutex_t would work equally as well, but require more platform-specific
 // code -- a "PlatformMutex".  Alternatively, a simply layer over muxAcquire-muxRelease
 // would work too.
 //
 // Since the caller might be a foreign thread, we don't necessarily have a Thread.MutexEvent
 // instance available.  Instead, we transiently allocate a ParkEvent on-demand if
 // we encounter contention.  That ParkEvent remains associated with the thread
 // until it manages to acquire the lock, at which time we return the ParkEvent
 // to the global ParkEvent free list.  This is correct and suffices for our purposes.
 //
 // Beware that the original jvm_raw_unlock() had a "_snuck" test but that
 // jvm_raw_lock() didn't have the corresponding test.  I suspect that's an
 // oversight, but I've replicated the original suspect logic in the new code ...
 void Monitor::jvm_raw_lock() {
   assert(rank() == native, "invariant");
   if (TryLock()) {
  Exeunt:
     assert (ILocked(), "invariant") ;
     assert (_owner == NULL, "invariant");
     // This can potentially be called by non-java Threads. Thus, the ThreadLocalStorage
     // might return NULL. Don't call set_owner since it will break on an NULL owner
     // Consider installing a non-null "ANON" distinguished value instead of just NULL.
     _owner = ThreadLocalStorage::thread();
     return ;
   }
   if (TrySpin(NULL)) goto Exeunt ;
   // slow-path - apparent contention
   // Allocate a ParkEvent for transient use.
   // The ParkEvent remains associated with this thread until
   // the time the thread manages to acquire the lock.
   ParkEvent * const ESelf = ParkEvent::Allocate(NULL) ;
   ESelf->reset() ;
   OrderAccess::storeload() ;
   // Either Enqueue Self on cxq or acquire the outer lock.
   if (AcquireOrPush (ESelf)) {
     ParkEvent::Release (ESelf) ;      // surrender the ParkEvent
     goto Exeunt ;
   }
   // At any given time there is at most one ondeck thread.
   // ondeck implies not resident on cxq and not resident on EntryList
   // Only the OnDeck thread can try to acquire -- contended for -- the lock.
   // CONSIDER: use Self->OnDeck instead of m->OnDeck.
   for (;;) {
     if (_OnDeck == ESelf && TrySpin(NULL)) break ;
     ParkCommon (ESelf, 0) ;
   }
   assert (_OnDeck == ESelf, "invariant") ;
   _OnDeck = NULL ;
   ParkEvent::Release (ESelf) ;      // surrender the ParkEvent
   goto Exeunt ;
 }
 void Monitor::jvm_raw_unlock() {
   // Nearly the same as Monitor::unlock() ...
   // directly set _owner instead of using set_owner(null)
   _owner = NULL ;
   if (_snuck) {         // ???
     assert(SafepointSynchronize::is_at_safepoint() && Thread::current()->is_VM_thread(), "sneak");
     _snuck = false;
     return ;
   }
   IUnlock(false) ;
 }
 bool Monitor::wait(bool no_safepoint_check, long timeout, bool as_suspend_equivalent) {
   Thread * const Self = Thread::current() ;
   assert (_owner == Self, "invariant") ;
   assert (ILocked(), "invariant") ;
   // as_suspend_equivalent logically implies !no_safepoint_check
   guarantee (!as_suspend_equivalent || !no_safepoint_check, "invariant") ;
   // !no_safepoint_check logically implies java_thread
   guarantee (no_safepoint_check || Self->is_Java_thread(), "invariant") ;
   #ifdef ASSERT
     Monitor * least = get_least_ranked_lock_besides_this(Self->owned_locks());
     assert(least != this, "Specification of get_least_... call above");
     if (least != NULL && least->rank() <= special) {
       tty->print("Attempting to wait on monitor %s/%d while holding"
                  " lock %s/%d -- possible deadlock",
                  name(), rank(), least->name(), least->rank());
       assert(false, "Shouldn't block(wait) while holding a lock of rank special");
     }
   #endif // ASSERT
   int wait_status ;
   // conceptually set the owner to NULL in anticipation of
   // abdicating the lock in wait
   set_owner(NULL);
   if (no_safepoint_check) {
     wait_status = IWait (Self, timeout) ;
   } else {
     assert (Self->is_Java_thread(), "invariant") ;
     JavaThread *jt = (JavaThread *)Self;
     // Enter safepoint region - ornate and Rococo ...
     ThreadBlockInVM tbivm(jt);
     OSThreadWaitState osts(Self->osthread(), false /* not Object.wait() */);
     if (as_suspend_equivalent) {
       jt->set_suspend_equivalent();
       // cleared by handle_special_suspend_equivalent_condition() or
       // java_suspend_self()
     }
     wait_status = IWait (Self, timeout) ;
     // were we externally suspended while we were waiting?
     if (as_suspend_equivalent && jt->handle_special_suspend_equivalent_condition()) {
       // Our event wait has finished and we own the lock, but
       // while we were waiting another thread suspended us. We don't
       // want to hold the lock while suspended because that
       // would surprise the thread that suspended us.
       assert (ILocked(), "invariant") ;
       IUnlock (true) ;
       jt->java_suspend_self();
       ILock (Self) ;
       assert (ILocked(), "invariant") ;
     }
   }
   // Conceptually reestablish ownership of the lock.
   // The "real" lock -- the LockByte -- was reacquired by IWait().
   assert (ILocked(), "invariant") ;
   assert (_owner == NULL, "invariant") ;
   set_owner (Self) ;
   return wait_status != 0 ;          // return true IFF timeout
 }
 Monitor::~Monitor() {
   assert ((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "") ;
 }
 void Monitor::ClearMonitor (Monitor * m, const char *name) {
   m->_owner             = NULL ;
   m->_snuck             = false ;
   if (name == NULL) {
     strcpy(m->_name, "UNKNOWN") ;
   } else {
     strncpy(m->_name, name, MONITOR_NAME_LEN - 1);
     m->_name[MONITOR_NAME_LEN - 1] = '\0';
   }
   m->_LockWord.FullWord = 0 ;
   m->_EntryList         = NULL ;
   m->_OnDeck            = NULL ;
   m->_WaitSet           = NULL ;
   m->_WaitLock[0]       = 0 ;
 }
 Monitor::Monitor() { ClearMonitor(this); }
 Monitor::Monitor (int Rank, const char * name, bool allow_vm_block) {
   ClearMonitor (this, name) ;
 #ifdef ASSERT
   _allow_vm_block  = allow_vm_block;
   _rank            = Rank ;
 #endif
 }
 Mutex::~Mutex() {
   assert ((UNS(_owner)|UNS(_LockWord.FullWord)|UNS(_EntryList)|UNS(_WaitSet)|UNS(_OnDeck)) == 0, "") ;
 }
 Mutex::Mutex (int Rank, const char * name, bool allow_vm_block) {
   ClearMonitor ((Monitor *) this, name) ;
 #ifdef ASSERT
  _allow_vm_block   = allow_vm_block;
  _rank             = Rank ;
 #endif
 }
 bool Monitor::owned_by_self() const {
   bool ret = _owner == Thread::current();
   assert (!ret || _LockWord.Bytes[_LSBINDEX] != 0, "invariant") ;
   return ret;
 }
 void Monitor::print_on_error(outputStream* st) const {
   st->print("[" PTR_FORMAT, this);
   st->print("] %s", _name);
   st->print(" - owner thread: " PTR_FORMAT, _owner);
 }
 // ----------------------------------------------------------------------------------
 // Non-product code
 #ifndef PRODUCT
 void Monitor::print_on(outputStream* st) const {
   st->print_cr("Mutex: [0x%lx/0x%lx] %s - owner: 0x%lx", this, _LockWord.FullWord, _name, _owner);
 }
 #endif
 #ifndef PRODUCT
 #ifdef ASSERT
 Monitor * Monitor::get_least_ranked_lock(Monitor * locks) {
   Monitor *res, *tmp;
   for (res = tmp = locks; tmp != NULL; tmp = tmp->next()) {
     if (tmp->rank() < res->rank()) {
       res = tmp;
     }
   }
   if (!SafepointSynchronize::is_at_safepoint()) {
     // In this case, we expect the held locks to be
     // in increasing rank order (modulo any native ranks)
     for (tmp = locks; tmp != NULL; tmp = tmp->next()) {
       if (tmp->next() != NULL) {
         assert(tmp->rank() == Mutex::native ||
                tmp->rank() <= tmp->next()->rank(), "mutex rank anomaly?");
       }
     }
   }
   return res;
 }
 Monitor* Monitor::get_least_ranked_lock_besides_this(Monitor* locks) {
   Monitor *res, *tmp;
   for (res = NULL, tmp = locks; tmp != NULL; tmp = tmp->next()) {
     if (tmp != this && (res == NULL || tmp->rank() < res->rank())) {
       res = tmp;
     }
   }
   if (!SafepointSynchronize::is_at_safepoint()) {
     // In this case, we expect the held locks to be
     // in increasing rank order (modulo any native ranks)
     for (tmp = locks; tmp != NULL; tmp = tmp->next()) {
       if (tmp->next() != NULL) {
         assert(tmp->rank() == Mutex::native ||
                tmp->rank() <= tmp->next()->rank(), "mutex rank anomaly?");
       }
     }
   }
   return res;
 }
 bool Monitor::contains(Monitor* locks, Monitor * lock) {
   for (; locks != NULL; locks = locks->next()) {
     if (locks == lock)
       return true;
   }
   return false;
 }
 #endif
 // Called immediately after lock acquisition or release as a diagnostic
 // to track the lock-set of the thread and test for rank violations that
 // might indicate exposure to deadlock.
 // Rather like an EventListener for _owner (:>).
 void Monitor::set_owner_implementation(Thread *new_owner) {
   // This function is solely responsible for maintaining
   // and checking the invariant that threads and locks
   // are in a 1/N relation, with some some locks unowned.
   // It uses the Mutex::_owner, Mutex::_next, and
   // Thread::_owned_locks fields, and no other function
   // changes those fields.
   // It is illegal to set the mutex from one non-NULL
   // owner to another--it must be owned by NULL as an
   // intermediate state.
   if (new_owner != NULL) {
     // the thread is acquiring this lock
     assert(new_owner == Thread::current(), "Should I be doing this?");
     assert(_owner == NULL, "setting the owner thread of an already owned mutex");
     _owner = new_owner; // set the owner
     // link "this" into the owned locks list
     #ifdef ASSERT  // Thread::_owned_locks is under the same ifdef
       Monitor* locks = get_least_ranked_lock(new_owner->owned_locks());
                     // Mutex::set_owner_implementation is a friend of Thread
       assert(this->rank() >= 0, "bad lock rank");
       // Deadlock avoidance rules require us to acquire Mutexes only in
       // a global total order. For example m1 is the lowest ranked mutex
       // that the thread holds and m2 is the mutex the thread is trying
       // to acquire, then  deadlock avoidance rules require that the rank
       // of m2 be less  than the rank of m1.
       // The rank Mutex::native  is an exception in that it is not subject
       // to the verification rules.
       // Here are some further notes relating to mutex acquisition anomalies:
       // . under Solaris, the interrupt lock gets acquired when doing
       //   profiling, so any lock could be held.
       // . it is also ok to acquire Safepoint_lock at the very end while we
       //   already hold Terminator_lock - may happen because of periodic safepoints
       if (this->rank() != Mutex::native &&
           this->rank() != Mutex::suspend_resume &&
           locks != NULL && locks->rank() <= this->rank() &&
           !SafepointSynchronize::is_at_safepoint() &&
           this != Interrupt_lock && this != ProfileVM_lock &&
           !(this == Safepoint_lock && contains(locks, Terminator_lock) &&
             SafepointSynchronize::is_synchronizing())) {
         new_owner->print_owned_locks();
         fatal(err_msg("acquiring lock %s/%d out of order with lock %s/%d -- "
                       "possible deadlock", this->name(), this->rank(),
                       locks->name(), locks->rank()));
       }
       this->_next = new_owner->_owned_locks;
       new_owner->_owned_locks = this;
     #endif
   } else {
     // the thread is releasing this lock
     Thread* old_owner = _owner;
     debug_only(_last_owner = old_owner);
     assert(old_owner != NULL, "removing the owner thread of an unowned mutex");
     assert(old_owner == Thread::current(), "removing the owner thread of an unowned mutex");
     _owner = NULL; // set the owner
     #ifdef ASSERT
       Monitor *locks = old_owner->owned_locks();
       // remove "this" from the owned locks list
       Monitor *prev = NULL;
       bool found = false;
       for (; locks != NULL; prev = locks, locks = locks->next()) {
         if (locks == this) {
           found = true;
           break;
         }
       }
       assert(found, "Removing a lock not owned");
       if (prev == NULL) {
         old_owner->_owned_locks = _next;
       } else {
         prev->_next = _next;
       }
       _next = NULL;
     #endif
   }
 }
 // Factored out common sanity checks for locking mutex'es. Used by lock() and try_lock()
 void Monitor::check_prelock_state(Thread *thread) {
   assert((!thread->is_Java_thread() || ((JavaThread *)thread)->thread_state() == _thread_in_vm)
          || rank() == Mutex::special, "wrong thread state for using locks");
   if (StrictSafepointChecks) {
     if (thread->is_VM_thread() && !allow_vm_block()) {
       fatal(err_msg("VM thread using lock %s (not allowed to block on)",
                     name()));
     }
     debug_only(if (rank() != Mutex::special) \
       thread->check_for_valid_safepoint_state(false);)
   }
   if (thread->is_Watcher_thread()) {
     assert(!WatcherThread::watcher_thread()->has_crash_protection(),
         "locking not allowed when crash protection is set");
   }
 }
 void Monitor::check_block_state(Thread *thread) {
   if (!_allow_vm_block && thread->is_VM_thread()) {
     warning("VM thread blocked on lock");
     print();
     BREAKPOINT;
   }
   assert(_owner != thread, "deadlock: blocking on monitor owned by current thread");
 }
 #endif // PRODUCT

Mercurial > jdk8-mips64-public > hotspot / annotate

src/share/vm/runtime/mutex.cpp@63147986a428 (annotated)

src/share/vm/runtime/mutex.cpp