Java Memory Model Under The Hood


java membar openjdk internals volatile jmm happens-before

There are many sources where you can get an idea of what JMM is about, but most of them still leave you with lots of unanswered questions. How does that happens-before thing work? Does using volatile result in caches being dropped? Why do we even need a memory model in the first place?

This article is intended to give the readers a level of understanding which allows them to answer all of these questions. It will consist of two large parts; the first of them being a hardware-level outline of what’s happening, and the second is indulging in some digging around OpenJDK sources and experimenting. Thus, even if you’re not exactly into Java, the first part might still be of interest to you.

The Hardware-Related Stuff

The engineers that create hardware are working hard on optimizing their products ever further, enabling you to get more and more performance units out of your code. However, it does come at a price of counter-intuitive execution scenarios that your code may display when it is run. There are countless hardware details obscured from our view by abstractions. And abstractions tend to get leaky.

Processor Caches

A request to the main memory is an expensive operation, which can take hundreds of nanoseconds to execute, even on modern hardware. The execution time of other operations, however, has grown considerably smaller over the years, unlike the main memory access. This problem is commonly named as the Memory Wall, and the obvious workaround for this is introducing caches. To put it simply, the processor has local copies of the contents of the main memory that it frequently uses. You can read further on different cache structures here, while we shall move on to the problem of keeping the cached values up to date.

Although there is apparently no problem when you have only one execution unit (referred to as processor from now on), things get complicated if you have more than one of those.

How does processor A know that processor B has modified some value, if A has it cached?

Or, more generally, how do you ensure cache coherency?

To maintain a consistent view on the state of the world, processors have to communicate with each other. The rules of such communication are called a cache coherency protocol.

Cache Coherency Protocols

There are numerous different protocols, which vary not only from one hardware manufacturer to another, but also constantly develop within a single vendor’s product line. In spite of all this variety, most of the protocols have lots of common aspects. Which is why we will take a closer look at MESI. It does not give the reader a full overview of all the protocols out there, however. There are some (e.g. Directory Based) protocols that are absolutely different. We are not going to look into them.

In MESI, every cache entry can be in one of the following states:

  • Invalid: the cache does not have such entry
  • Exclusive: the entry resides in this cache only, and has not been modified
  • Modified: the processor has modified a value, but has not yet written it back to the main memory or sent to any other processor
  • Shared: more that one processor has the entry in its cache

Transitions between states occur via sending certain messages that are also a part of the protocol. The exact message types are not quite relevant, so they are omitted in this article. There are many other sources which you can use to gain insight into them. Memory Barriers: a Hardware View for Software Hackers is the one that I would recommend.

It is ironical that deep down, messaging is used to change states concurrently. Problem, Actor Model haters?

MESI Optimizations And The Problems They Introduce

Without going into details, we will say that it takes time for messages to be delivered, which introduces more latency into state switching. It is also important to understand that some state transitions require some special handling, which might stall the processor. These things lead to all sorts of scalability and performance problems.

Store Buffers

If you need to write something to a variable that is Shared in the cache, you have to send an Invalidate message to all its other holders, and wait for them to acknowledge it. The processor is going to be stalled for that duration Which is a sad thing, seeing as the time required for that is typically several orders of magnitude higher than executing a simple instruction needs.

In real life, cache entries do not contain just a single variable. The established unit is a cache line, which usually contains more than one variable, and is typically 64 bytes in size.

It can lead to interesting side effects, e.g. cache contention.

To avoid such a waste of time, Store Buffers are used. The processor places the values which it wants to write to its buffer, and goes on executing things. When all the Invalidate Acknowledge messages are received, the data is finally committed.

One can expect a number of hidden dangers here. The easy one is that a processor may try to read some value that it has placed in the store buffer, but which has not yet been committed. The workaround is called Store Forwarding, which causes the loads to return the value in the store buffer, if it is present.

The second pitfall is that there is no guarantee on the order in which the stores will leave the buffer. Consider the following piece of code:

void executedOnCpu0() {
    value = 10;
    finished = true;
void executedOnCpu1() {
    assert value == 10;

Suppose that when the execution starts, CPU 0 has finished in the Exclusive state, while value is not installed in its cache at all (i.e. is Invalid). In such scenario, value will leave the store buffer considerably later than finished will. It is entirely possible that CPU 1 will then load finished as true, while value will not be equal to 10.

Such changes in the observable behavior are called reorderings. Note that it does not necessarily mean that your instructions' places have been changed by some malicious (or well-meaning) party.

It just means that some other CPU has observed their results in a different order than what's written in the program.

Invalidate Queues

Executing an invalidation is not a cheap operation as well, and it costs for the processor applying it. Moreover, it is no surprise that Store Buffers are not infinite, so the processors sometimes have to wait for Invalidate Acknowledge to come. These two can make performance degrade considerably. To counter this, Invalidate Queues have been introduced. Their contract is as follows:

  • For all incoming Invalidate requests, Invalidate Acknowledge messages are immediately sent
  • The Invalidate is not in fact applied, but placed to a special queue, to be executed when convenient
  • The processor will not send any messages on the cache entry in question, until it processes the Invalidate

There, too, are cases when such optimization will lead to counter-intuitive results. We return to our code, and assume that CPU 1 has value in the Exclusive state. Here’s a diagram of a possible execution:

# CPU 0: operations CPU 0: value CPU 0: finished CPU 1: operations CPU 1: value CPU 1: finished
0 0 (Shared) false (Exclusive) 0 (Shared) (Invalid)
value = 10;
- store_buffer(value)
← invalidate(value)
0 (Shared)
10 (in store buffer)
false (Exclusive)
while (!finished);
← read(finished)
0 (Shared) (Invalid)
finished = true;
0 (Shared)
10 (in store buffer)
true (Modified)
→ invalidate(value)
← invalidate_ack(value)
- invalidate_queue(value)
0 (Shared)
(in invalidation queue)
→ read(finished)
← read_response(finished)
0 (Shared)
10 (in store buffer)
true (Shared)
→ read_response(finished)
0 (Shared)
(in invalidation queue)
true (Shared)
assert value == 10;
0 (Shared)
(in invalidation queue)
true (Shared)
Assertion fails
- invalidate(value)
(Invalid) true (Shared)

Concurrency is simple and easy, is it not? The problem is in steps (4) — (6). When CPU 1 receives an Invalidate in (4), it queues it without processing. Then CPU 1 gets Read Response in (6), while the corresponding Read has been sent earlier in (2). Despite this, we do not invalidate value, ending up with an assertion that fails. If only operation (N) has executed earlier. But alas, the damn optimization has spoiled everything! On the other hand, it grants us some significant performance boost.

The thing is that hardware engineers cannot know in advance when such an optimization is allowed, and when it is not. Which is why they leave the problem in our capable hands. They also give us a little something, with a note attached to it: “It’s dangerous to go alone! Take this!”

Hardware Memory Model

The Magical Sword that software engineers who are setting out to fight Dragons are given, is not quite a sword. Rather, what the hardware guys have given us are the Rules As Written. They describe which values a processor can observe given the instructions this (or some other) processor has executed. What we could classify as Spells would be the Memory Barriers. For the MESI example of ours, they would be the following:

  • Store Memory Barrier (a.k.a. ST, SMB, smp_wmb) is the instruction that tells the processor to apply all the stores that are already in the store buffer, before it applies any that come after this instruction
  • Load Memory Barrier (a.k.a. LD, RMB, smp_rmb) is the instruction that tells the processor to apply all the invalidates that are already in the invalidate queue, before executing any loads

So, these two Spells can prevent the two situations which we have come across earlier. We should use it:

void executedOnCpu0() {
    value = 10;
    storeMemoryBarrier(); // Mighty Spell!
    finished = true;
void executedOnCpu1() {
    loadMemoryBarrier(); // I am a Wizard!
    assert value == 10;

Yay! We are now safe. Time to write some high-performance and correct concurrent code!

Oh, wait. It doesn’t even compile, says something about missing methods. What a mess.

Write Once @ Run Anywhere

All those cache coherency protocols, memory barriers, dropped caches and whatnot seem to be awfully platform-specific things. Java Developers should not care for those at all. Java Memory Model has no notion of reordering, after all.

If you do not fully understand this last phrase, you should not continue reading this article. A better idea would be to go and learn some JMM instead. A good start would be this FAQ.

But there are reorderings happening on deeper levels of abstractions. Should be interesting to see how JMM maps to the hardware model. Let’s start with a simple class (github):

public class TestSubject {

    private volatile boolean finished;
    private int value = 0;

    void executedOnCpu0() {
        value = 10;
        finished = true;

    void executedOnCpu1() {
        assert value == 10;


There are many venues we could follow to understand what’s going on: the PrintAssembly fun, checking out the interpreter’s doings, asking someone, mysteriously saying that the caches are being dropped, and many more. I have decided to stick with looking at the C1 (a.k.a. the client compiler) of OpenJDK. While the client compiler is barely used in real applications, it is a good choice for educational purposes.

I have used jdk8 at revision 933:4f8fa4724c14. Things may be different in other versions.

If you have never before digged through the sources of OpenJDK (and even if you have, for that matter), it could be hard to find where the things that interest you lie. An easy way to narrow down the search space is getting the name of the bytecode instruction that interests you, and simply look for it. Alright, let’s do that:

$ javac & javap -c TestSubject
void executedOnCpu0();
     0: aload_0          // Push this to the stack
     1: bipush        10 // Push 10 to the stack
     3: putfield      #2 // Assign 10 to the second field(value) of this
     6: aload_0          // Push this to the stack
     7: iconst_1         // Push 1 to the stack
     8: putfield      #3 // Assign 1 to the third field(finished) of this
    11: return

void executedOnCpu1();
     0: aload_0          // Push this to the stack
     1: getfield      #3 // Load the third field of this(finished) and push it to the stack
     4: ifne          10 // If the top of the stack is not zero, go to label 10
     7: goto          0  // One more iteration of the loop
    10: getstatic     #4 // Get the static system field $assertionsDisabled:Z
    13: ifne          33 // If the assertions are disabled, go to label 33(the end)
    16: aload_0          // Push this to the stack
    17: getfield      #2 // Load the second field of this(value) and push it to the stack
    20: bipush        10 // Push 10 to the stack
    22: if_icmpeq     33 // If the top two elements of the stack are equal, go to label 33(the end)
    25: new           #5 // Create a new java/lang/AssertionError
    28: dup              // Duplicate the top of the stack
    29: invokespecial #6 // Invoke the constructor (the <init> method)
    32: athrow           // Throw what we have at the top of the stack (an AssertionError)
    33: return

You should not try to predict the performance (or even low-level behavior) of your program by looking at the bytecode. When the JIT Compiler is through with it, there will not be much similarities left.

We are only doing this because we need to know who the assassins were working for.

There are two things of interest here:

  1. Assertions are disabled by default, as many people tend to forget. Use -ea to enable them.
  2. The names that we were looking for: getfield and putfield.

Ah, the Field brothers. I knew it was them. It is not long before I put them behind bars for good, now.

Down The Rabbit Hole

As we can see, the instructions used for loading and storing are the same for both volatile and plain fields. So, it is a good idea to find where the compiler learns whether a field is volatile or not. Digging around a little, we end up in share/vm/ci/ciField.hpp. The method of interest is

bool is_volatile    () { return flags().is_volatile(); }

So, what we now are tasked with is finding the methods that handle loading and storing of fields and use investigate all the codepaths conditional on the result of invoking the method above. The Client Compiler processes them on the Low-Level Intermediate Representation (LIR) stage, in the file share/vm/c1/c1_LIRGenerator.cpp.

C1 Intermediate Representation

Let’s start with the stores. The method that we are looking into is void LIRGenerator::do_StoreField(StoreField* x), and resides at lines 1658:1751. The first remarkable action that we see is

if (is_volatile && os::is_MP()) {
  __ membar_release();

Cool, a memory barrier! The two underscores are a macro that expand into gen()->lir()->, and the invoked method is defined in share/vm/c1/c1_LIR.hpp:

void membar_release()                          { append(new LIR_Op0(lir_membar_release)); }

So, what happened is that we have appended one more operation, lir_membar_release, to our representation.

if (is_volatile && !needs_patching) {
  volatile_field_store(value.result(), address, info);

The invoked method has platform-specific implementations. For x86 (cpu/x86/vm/c1\_LIRGenerator\_x86.cpp), it’s fairly simple: for 64-bit fields, we dabble in some Dark Magics to ensure write atomicity. Because the spec says so. This is a bit outdated, and may be reviewed in Java 9. The last thing that we want to see is one more memory barrier at the very end of the method:

if (is_volatile && os::is_MP()) {
  __ membar();
void membar()                                  { append(new LIR_Op0(lir_membar)); }

That’s it for the stores.

The loads are just a bit lower in the source code, and do not contain anything principally new. They have the same Dark Magic stuff for the atomicity of long and double fields, and add a lir_membar_acquire after the load is done.

Note that I have deliberately left out some of the things that are going on, e.g. the GC-related instructions.

Memory Barrier Types And Abstraction Levels

By this time, you must be wondering what the release and acquire memory barriers are, for we have not yet introduced them. This is all because the store and load memory barriers which we have seen before are the operations in the MESI model, while we currently reside a couple of abstraction levels above it (or any other Cache Coherency Protocol). At this level, we have different terminology.

Given that we have two kinds of operations, Load and Store, we have four ordered of pairs of them: LoadLoad, LoadStore, StoreLoad and StoreStore. It is therefore very convenient to have four types of memory barriers with the same names.

If we have a XY memory barrier, it means that all X operations that come before the barrier must complete their execution before any Y operation after the barrier starts.

For instance, all Store operations before a StoreStore barrier must complete earlier than any Store operation that comes after the barrier starts. The JSR-133 Cookbook is a good read on the subject.

Some people get confused and think that memory barriers take a variable as an argument, and then prohibit reorderings of the variable stores or loads across threads.

Memory barriers work within one thread only. By combining them in the right way, you can ascertain that when some thread loads the values stored by another thread, it sees a consistent picture. More generally, all the abstractions that JMM goes on about are granted by the correct combination of memory barrers.

Then there are the Acquire and Release semantics. A write operation that has release semantics requires that all the memory operations that come before it are finished before the operation itself starts its execution. The opposite is true for the read-acquire operations.

One can see that a Release Memory Barrier can be implemented as a LoadStore|StoreStore combination, and the Acquire Memory Barrier is a LoadStore|LoadLoad. The StoreLoad is what we have seen above as lir_membar.

Emitting Assembly Code

Now that we have sorted out the IR and its memory barriers, we can get down to the native level. All the emission happens in the share/vm/c1/c1\_LIRAssembler.cpp file:

case lir_membar_release:

The memory barriers are platform-specific, so for x86 we are looking into the cpu/x86/vm/c1_LIRAssembler_x86.cpp file. Seeing as x86 is an architecture with a rather strict memory model, most of the memory barriers are no-ops.

void LIR_Assembler::membar_acquire() {
  // No x86 machines currently require load fences
  // __ load_fence();

void LIR_Assembler::membar_release() {
  // No x86 machines currently require store fences
  // __ store_fence();

Not all of them, however:

void LIR_Assembler::membar() {
   // QQQ sparc TSO uses this,
   __ membar( Assembler::Membar_mask_bits(Assembler::StoreLoad));

(which we follow into cpu/x86/vm/assembler_x86.hpp)

// Serializes memory and blows flags
void membar(Membar_mask_bits order_constraint) {
  if (os::is_MP()) {
    // We only have to handle StoreLoad
    if (order_constraint & StoreLoad) {
      // All usable chips support "locked" instructions which suffice
      // as barriers, and are much faster than the alternative of
      // using cpuid instruction. We use here a locked add [esp],0.
      // This is conveniently otherwise a no-op except for blowing
      // flags.
      // Any change to this code may need to revisit other places in
      // the code where this idiom is used, in particular the
      // orderAccess code.
      addl(Address(rsp, 0), 0);// Assert the lock# signal here

So, for every volatile write we have to use the relatively expensive StoreLoad barrier in the form of lock addl $0x0,(%rsp). It forces us to execute all the pending stores, and effectively ensures that other threads see the fresh values quickly. And for volatile read, we emit no additional barriers. One should not think that volatile reads are as cheap as regular reads are, however.

It should be clear that while the barriers may emit no assembly code, they are still there in the IR. If they were ignored by the components that can modify the code (say, the compiler), there would be bugs like this one.

Sanity Checks

While making up theories by looking at the sources of OpenJDK is all nice and good, all the real scientists go out there and test their theories. Let us not get too out of the loop and try it as well.

Java Concurrency Stress Fun

The first thing we want to check is that things will actually get bad if we remove volatile from our code. The problem with demonstrating such a “reordering” is that the prior probability of it happening is fairly low. And on some architectures, the HMM prohibits it altogether. So, we have to rely on the compiler, and also try it a lot of times.

The good thing is that we have no need to invent the wheel, as there’s the [jcstress]( tool that executes the code lots of times and keeps an aggregated track of the outcomes. It also very conveniently does all the dirty work for us, including the dirty work we did not even suspect we had to do.

Moreover, jcstress already has the very test that we need:

static class State {
    int x;
    int y; // acq/rel var

public void actor1(State s, IntResult2 r) {
    s.x = 1;
    s.x = 2;
    s.y = 1;
    s.x = 3;

public void actor2(State s, IntResult2 r) {
    r.r1 = s.y;
    r.r2 = s.x;

We have one thread executing stores, and another thread doing reads, and then reporting the observed states. The framework aggregates the results for us, and then matches it against some certain rules. We are interested in two possible observations made by the second thread: [1, 0] and [1, 1]. In these two cases, it has loaded y == 1, but has either failed to see any writes to x, or the loaded version was not the most recent one at the time y was written. According to our theory, such events should happen without the volatile modifier. Let’s see:

$ java -jar tests-all/target/jcstress.jar -v -t ".*UnfencedAcquireReleaseTest.*"

Observed state  Occurrence      Expectation       Interpretation
 [0, 0]          32725135        ACCEPTABLE       Before observing releasing write to, any value is OK for $x.
 [0, 1]             15           ACCEPTABLE       Before observing releasing write to, any value is OK for $x.
 [0, 2]             36           ACCEPTABLE       Before observing releasing write to, any value is OK for $x.
 [0, 3]           10902          ACCEPTABLE       Before observing releasing write to, any value is OK for $x.
 [1, 0]           65960    ACCEPTABLE_INTERESTING Can read the default or old value for $x after $y is observed.
 [1, 3]          50929785        ACCEPTABLE       Can see a released value of $x if $y is observed.
 [1, 2]             7            ACCEPTABLE       Can see a released value of $x if $y is observed.

So, in 65960 cases out 83731840 (≈ 0.07%) the second thread has observed y == 1 && x == 0, which confirms that the reorderings can indeed happen.

PrintAssembly Fun

The second thing we want to check is if we have correctly predicted the generated assembly code. So, we add lots of invocations of the required code, disable inlining for easier result interpretation, enable assertions and run in the client VM:

$ java -client -ea -XX:+UnlockDiagnosticVMOptions -XX:+PrintAssembly -XX:MaxInlineSize=0 TestSubject
  # {method} 'executedOnCpu0' '()V' in 'TestSubject'
  0x00007f6d1d07405c: movl   $0xa,0xc(%rsi)
  0x00007f6d1d074063: movb   $0x1,0x10(%rsi)
  0x00007f6d1d074067: lock addl $0x0,(%rsp)     ;*putfield finished
                                                ; - TestSubject::executedOnCpu0@8 (line 15)
  # {method} 'executedOnCpu1' '()V' in 'TestSubject'
  0x00007f6d1d061126: movzbl 0x10(%rbx),%r11d   ;*getfield finished
                                                ; - TestSubject::executedOnCpu1@1 (line 19)
  0x00007f6d1d06112b: test   %r11d,%r11d

Yay, just as planned! Means that it’s time to go wrap up.

Let me remind you the questions that you should be able to answer by now:

  • How does that happens-before thing work?
  • Does using volatile result in caches being dropped?
  • Why do we even need a memory model in the first place?

Do you think you can answer these? Welcome to the comments in any case!

P.S. This is the translation of my earlier blog entry in Russian, which has been reviewed by Aleksey Shipilёv, so thank him.
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