Garbage Collection Safepoints in LLVM

Status

This document describes a set of extensions to LLVM to support garbage collection. By now, these mechanisms are well proven with commercial java implementation with a fully relocating collector having shipped using them. There are a couple places where bugs might still linger; these are called out below.

They are still listed as “experimental” to indicate that no forward or backward compatibility guarantees are offered across versions. If your use case is such that you need some form of forward compatibility guarantee, please raise the issue on the llvm-dev mailing list.

LLVM still supports an alternate mechanism for conservative garbage collection support using the gcroot intrinsic. The gcroot mechanism is mostly of historical interest at this point with one exception - its implementation of shadow stacks has been used successfully by a number of language frontends and is still supported.

Overview & Core Concepts

To collect dead objects, garbage collectors must be able to identify any references to objects contained within executing code, and, depending on the collector, potentially update them. The collector does not need this information at all points in code - that would make the problem much harder - but only at well-defined points in the execution known as ‘safepoints’ For most collectors, it is sufficient to track at least one copy of each unique pointer value. However, for a collector which wishes to relocate objects directly reachable from running code, a higher standard is required.

One additional challenge is that the compiler may compute intermediate results (“derived pointers”) which point outside of the allocation or even into the middle of another allocation. The eventual use of this intermediate value must yield an address within the bounds of the allocation, but such “exterior derived pointers” may be visible to the collector. Given this, a garbage collector can not safely rely on the runtime value of an address to indicate the object it is associated with. If the garbage collector wishes to move any object, the compiler must provide a mapping, for each pointer, to an indication of its allocation.

To simplify the interaction between a collector and the compiled code, most garbage collectors are organized in terms of three abstractions: load barriers, store barriers, and safepoints.

1. A load barrier is a bit of code executed immediately after the machine load instruction, but before any use of the value loaded. Depending on the collector, such a barrier may be needed for all loads, merely loads of a particular type (in the original source language), or none at all.

2. Analogously, a store barrier is a code fragment that runs immediately before the machine store instruction, but after the computation of the value stored. The most common use of a store barrier is to update a ‘card table’ in a generational garbage collector.

3. A safepoint is a location at which pointers visible to the compiled code (i.e. currently in registers or on the stack) are allowed to change. After the safepoint completes, the actual pointer value may differ, but the ‘object’ (as seen by the source language) pointed to will not.

Note that the term ‘safepoint’ is somewhat overloaded. It refers to both the location at which the machine state is parsable and the coordination protocol involved in bring application threads to a point at which the collector can safely use that information. The term “statepoint” as used in this document refers exclusively to the former.

This document focuses on the last item - compiler support for safepoints in generated code. We will assume that an outside mechanism has decided where to place safepoints. From our perspective, all safepoints will be function calls. To support relocation of objects directly reachable from values in compiled code, the collector must be able to:

1. identify every copy of a pointer (including copies introduced by the compiler itself) at the safepoint,

2. identify which object each pointer relates to, and

3. potentially update each of those copies.

This document describes the mechanism by which an LLVM based compiler can provide this information to a language runtime/collector, and ensure that all pointers can be read and updated if desired.

Abstract Machine Model

At a high level, LLVM has been extended to support compiling to an abstract machine which extends the actual target with a non-integral pointer type suitable for representing a garbage collected reference to an object. In particular, such non-integral pointer type have no defined mapping to an integer representation. This semantic quirk allows the runtime to pick a integer mapping for each point in the program allowing relocations of objects without visible effects.

This high level abstract machine model is used for most of the optimizer. As a result, transform passes do not need to be extended to look through explicit relocation sequence. Before starting code generation, we switch representations to an explicit form. The exact location chosen for lowering is an implementation detail.

Note that most of the value of the abstract machine model comes for collectors which need to model potentially relocatable objects. For a compiler which supports only a non-relocating collector, you may wish to consider starting with the fully explicit form.

Warning: There is one currently known semantic hole in the definition of non-integral pointers which has not been addressed upstream. To work around this, you need to disable speculation of loads unless the memory type (non-integral pointer vs anything else) is known to unchanged. That is, it is not safe to speculate a load if doing causes a non-integral pointer value to be loaded as any other type or vice versa. In practice, this restriction is well isolated to isSafeToSpeculate in ValueTracking.cpp.

Explicit Representation

A frontend could directly generate this low level explicit form, but doing so may inhibit optimization. Instead, it is recommended that compilers with relocating collectors target the abstract machine model just described.

The heart of the explicit approach is to construct (or rewrite) the IR in a manner where the possible updates performed by the garbage collector are explicitly visible in the IR. Doing so requires that we:

1. create a new SSA value for each potentially relocated pointer, and ensure that no uses of the original (non relocated) value is reachable after the safepoint,

2. specify the relocation in a way which is opaque to the compiler to ensure that the optimizer can not introduce new uses of an unrelocated value after a statepoint. This prevents the optimizer from performing unsound optimizations.

3. recording a mapping of live pointers (and the allocation they’re associated with) for each statepoint.

At the most abstract level, inserting a safepoint can be thought of as replacing a call instruction with a call to a multiple return value function which both calls the original target of the call, returns its result, and returns updated values for any live pointers to garbage collected objects.

Note that the task of identifying all live pointers to garbage collected values, transforming the IR to expose a pointer giving the base object for every such live pointer, and inserting all the intrinsics correctly is explicitly out of scope for this document. The recommended approach is to use the utility passes described below.

This abstract function call is concretely represented by a sequence of intrinsic calls known collectively as a “statepoint relocation sequence”.

Let’s consider a simple call in LLVM IR:

declare void @foo()
define ptr addrspace(1) @test1(ptr addrspace(1) %obj)
       gc "statepoint-example" {
  call void @foo()
  ret ptr addrspace(1) %obj
}

Depending on our language we may need to allow a safepoint during the execution of foo. If so, we need to let the collector update local values in the current frame. If we don’t, we’ll be accessing a potential invalid reference once we eventually return from the call.

In this example, we need to relocate the SSA value %obj. Since we can’t actually change the value in the SSA value %obj, we need to introduce a new SSA value %obj.relocated which represents the potentially changed value of %obj after the safepoint and update any following uses appropriately. The resulting relocation sequence is:

define ptr addrspace(1) @test(ptr addrspace(1) %obj)
       gc "statepoint-example" {
  %safepoint = call token (i64, i32, ptr, i32, i32, ...) @llvm.experimental.gc.statepoint.p0f_isVoidf(i64 0, i32 0, ptr elementtype(void ()) @foo, i32 0, i32 0, i32 0, i32 0) ["gc-live" (ptr addrspace(1) %obj)]
  %obj.relocated = call ptr addrspace(1) @llvm.experimental.gc.relocate.p1(token %safepoint, i32 0, i32 0)
  ret ptr addrspace(1) %obj.relocated
}

Ideally, this sequence would have been represented as a M argument, N return value function (where M is the number of values being relocated + the original call arguments and N is the original return value + each relocated value), but LLVM does not easily support such a representation.

Instead, the statepoint intrinsic marks the actual site of the safepoint or statepoint. The statepoint returns a token value (which exists only at compile time). To get back the original return value of the call, we use the gc.result intrinsic. To get the relocation of each pointer in turn, we use the gc.relocate intrinsic with the appropriate index. Note that both the gc.relocate and gc.result are tied to the statepoint. The combination forms a “statepoint relocation sequence” and represents the entirety of a parseable call or ‘statepoint’.

When lowered, this example would generate the following x86 assembly:

        .globl        test1
        .align        16, 0x90
        pushq %rax
        callq foo
.Ltmp1:
        movq  (%rsp), %rax  # This load is redundant (oops!)
        popq  %rdx
        retq

Each of the potentially relocated values has been spilled to the stack, and a record of that location has been recorded to the Stack Map section. If the garbage collector needs to update any of these pointers during the call, it knows exactly what to change.

The relevant parts of the StackMap section for our example are:

# This describes the call site
# Stack Maps: callsite 2882400000
        .quad 2882400000
        .long .Ltmp1-test1
        .short        0
# .. 8 entries skipped ..
# This entry describes the spill slot which is directly addressable
# off RSP with offset 0.  Given the value was spilled with a pushq,
# that makes sense.
# Stack Maps:   Loc 8: Direct RSP     [encoding: .byte 2, .byte 8, .short 7, .int 0]
        .byte 2
        .byte 8
        .short        7
        .long 0

This example was taken from the tests for the RewriteStatepointsForGC utility pass. As such, its full StackMap can be easily examined with the following command.

opt -rewrite-statepoints-for-gc test/Transforms/RewriteStatepointsForGC/basics.ll -S | llc -debug-only=stackmaps

Simplifications for Non-Relocating GCs

Some of the complexity in the previous example is unnecessary for a non-relocating collector. While a non-relocating collector still needs the information about which location contain live references, it doesn’t need to represent explicit relocations. As such, the previously described explicit lowering can be simplified to remove all of the gc.relocate intrinsic calls and leave uses in terms of the original reference value.

Here’s the explicit lowering for the previous example for a non-relocating collector:

define void @manual_frame(ptr %a, ptr %b) gc "statepoint-example" {
  %alloca = alloca ptr
  %allocb = alloca ptr
  store ptr %a, ptr %alloca
  store ptr %b, ptr %allocb
  call token (i64, i32, ptr, i32, i32, ...) @llvm.experimental.gc.statepoint.p0(i64 0, i32 0, ptr elementtype(void ()) @func, i32 0, i32 0, i32 0, i32 0) ["gc-live" (ptr %alloca, ptr %allocb)]
  ret void
}

Recording On Stack Regions

In addition to the explicit relocation form previously described, the statepoint infrastructure also allows the listing of allocas within the gc pointer list. Allocas can be listed with or without additional explicit gc pointer values and relocations.

An alloca in the gc region of the statepoint operand list will cause the address of the stack region to be listed in the stackmap for the statepoint.

This mechanism can be used to describe explicit spill slots if desired. It then becomes the generator’s responsibility to ensure that values are spill/filled to/from the alloca as needed on either side of the safepoint. Note that there is no way to indicate a corresponding base pointer for such an explicitly specified spill slot, so usage is restricted to values for which the associated collector can derive the object base from the pointer itself.

This mechanism can be used to describe on stack objects containing references provided that the collector can map from the location on the stack to a heap map describing the internal layout of the references the collector needs to process.

WARNING: At the moment, this alternate form is not well exercised. It is recommended to use this with caution and expect to have to fix a few bugs. In particular, the RewriteStatepointsForGC utility pass does not do anything for allocas today.

Base & Derived Pointers

A “base pointer” is one which points to the starting address of an allocation (object). A “derived pointer” is one which is offset from a base pointer by some amount. When relocating objects, a garbage collector needs to be able to relocate each derived pointer associated with an allocation to the same offset from the new address.

“Interior derived pointers” remain within the bounds of the allocation they’re associated with. As a result, the base object can be found at runtime provided the bounds of allocations are known to the runtime system.

“Exterior derived pointers” are outside the bounds of the associated object; they may even fall within another allocations address range. As a result, there is no way for a garbage collector to determine which allocation they are associated with at runtime and compiler support is needed.

The gc.relocate intrinsic supports an explicit operand for describing the allocation associated with a derived pointer. This operand is frequently referred to as the base operand, but does not strictly speaking have to be a base pointer, but it does need to lie within the bounds of the associated allocation. Some collectors may require that the operand be an actual base pointer rather than merely an internal derived pointer. Note that during lowering both the base and derived pointer operands are required to be live over the associated call safepoint even if the base is otherwise unused afterwards.

GC Transitions

As a practical consideration, many garbage-collected systems allow code that is collector-aware (“managed code”) to call code that is not collector-aware (“unmanaged code”). It is common that such calls must also be safepoints, since it is desirable to allow the collector to run during the execution of unmanaged code. Furthermore, it is common that coordinating the transition from managed to unmanaged code requires extra code generation at the call site to inform the collector of the transition. In order to support these needs, a statepoint may be marked as a GC transition, and data that is necessary to perform the transition (if any) may be provided as additional arguments to the statepoint.

Note that although in many cases statepoints may be inferred to be GC transitions based on the function symbols involved (e.g. a call from a function with GC strategy “foo” to a function with GC strategy “bar”), indirect calls that are also GC transitions must also be supported. This requirement is the driving force behind the decision to require that GC transitions are explicitly marked.

Let’s revisit the sample given above, this time treating the call to @foo as a GC transition. Depending on our target, the transition code may need to access some extra state in order to inform the collector of the transition. Let’s assume a hypothetical GC–somewhat unimaginatively named “hypothetical-gc” –that requires that a TLS variable must be written to before and after a call to unmanaged code. The resulting relocation sequence is:

@flag = thread_local global i32 0, align 4

define i8 addrspace(1)* @test1(i8 addrspace(1) *%obj)
       gc "hypothetical-gc" {

  %0 = call token (i64, i32, void ()*, i32, i32, ...)* @llvm.experimental.gc.statepoint.p0f_isVoidf(i64 0, i32 0, void ()* @foo, i32 0, i32 1, i32* @Flag, i32 0, i8 addrspace(1)* %obj)
  %obj.relocated = call coldcc i8 addrspace(1)* @llvm.experimental.gc.relocate.p1i8(token %0, i32 7, i32 7)
  ret i8 addrspace(1)* %obj.relocated
}

During lowering, this will result in an instruction selection DAG that looks something like:

CALLSEQ_START
...
GC_TRANSITION_START (lowered i32 *@Flag), SRCVALUE i32* Flag
STATEPOINT
GC_TRANSITION_END (lowered i32 *@Flag), SRCVALUE i32 *Flag
...
CALLSEQ_END

In order to generate the necessary transition code, the backend for each target supported by “hypothetical-gc” must be modified to lower GC_TRANSITION_START and GC_TRANSITION_END nodes appropriately when the “hypothetical-gc” strategy is in use for a particular function. Assuming that such lowering has been added for X86, the generated assembly would be:

        .globl        test1
        .align        16, 0x90
        pushq %rax
        movl $1, %fs:Flag@TPOFF
        callq foo
        movl $0, %fs:Flag@TPOFF
.Ltmp1:
        movq  (%rsp), %rax  # This load is redundant (oops!)
        popq  %rdx
        retq

Note that the design as presented above is not fully implemented: in particular, strategy-specific lowering is not present, and all GC transitions are emitted as as single no-op before and after the call instruction. These no-ops are often removed by the backend during dead machine instruction elimination.

Before the abstract machine model is lowered to the explicit statepoint model of relocations by the RewriteStatepointsForGC pass it is possible for any derived pointer to get its base pointer and offset from the base pointer by using the gc.get.pointer.base and the gc.get.pointer.offset intrinsics respectively. These intrinsics are inlined by the RewriteStatepointsForGC pass and must not be used after this pass.

Stack Map Format

Locations for each pointer value which may need read and/or updated by the runtime or collector are provided in a separate section of the generated object file as specified in the PatchPoint documentation. This special section is encoded per the Stack Map format.

The general expectation is that a JIT compiler will parse and discard this format; it is not particularly memory efficient. If you need an alternate format (e.g. for an ahead of time compiler), see discussion under :ref: open work items <OpenWork> below.

Each statepoint generates the following Locations:

• Constant which describes the calling convention of the call target. This constant is a valid calling convention identifier for the version of LLVM used to generate the stackmap. No additional compatibility guarantees are made for this constant over what LLVM provides elsewhere w.r.t. these identifiers.

• Constant which describes the flags passed to the statepoint intrinsic

• Constant which describes number of following deopt Locations (not operands). Will be 0 if no “deopt” bundle is provided.

• Variable number of Locations, one for each deopt parameter listed in the “deopt” operand bundle. At the moment, only deopt parameters with a bitwidth of 64 bits or less are supported. Values of a type larger than 64 bits can be specified and reported only if a) the value is constant at the call site, and b) the constant can be represented with less than 64 bits (assuming zero extension to the original bitwidth).

• Variable number of relocation records, each of which consists of exactly two Locations. Relocation records are described in detail below.

Each relocation record provides sufficient information for a collector to relocate one or more derived pointers. Each record consists of a pair of Locations. The second element in the record represents the pointer (or pointers) which need updated. The first element in the record provides a pointer to the base of the object with which the pointer(s) being relocated is associated. This information is required for handling generalized derived pointers since a pointer may be outside the bounds of the original allocation, but still needs to be relocated with the allocation. Additionally:

• It is guaranteed that the base pointer must also appear explicitly as a relocation pair if used after the statepoint.

• There may be fewer relocation records then gc parameters in the IR statepoint. Each unique pair will occur at least once; duplicates are possible.

• The Locations within each record may either be of pointer size or a multiple of pointer size. In the later case, the record must be interpreted as describing a sequence of pointers and their corresponding base pointers. If the Location is of size N x sizeof(pointer), then there will be N records of one pointer each contained within the Location. Both Locations in a pair can be assumed to be of the same size.

Note that the Locations used in each section may describe the same physical location. e.g. A stack slot may appear as a deopt location, a gc base pointer, and a gc derived pointer.

The LiveOut section of the StkMapRecord will be empty for a statepoint record.

Safepoint Semantics & Verification

The fundamental correctness property for the compiled code’s correctness w.r.t. the garbage collector is a dynamic one. It must be the case that there is no dynamic trace such that an operation involving a potentially relocated pointer is observably-after a safepoint which could relocate it. ‘observably-after’ is this usage means that an outside observer could observe this sequence of events in a way which precludes the operation being performed before the safepoint.

To understand why this ‘observable-after’ property is required, consider a null comparison performed on the original copy of a relocated pointer. Assuming that control flow follows the safepoint, there is no way to observe externally whether the null comparison is performed before or after the safepoint. (Remember, the original Value is unmodified by the safepoint.) The compiler is free to make either scheduling choice.

The actual correctness property implemented is slightly stronger than this. We require that there be no static path on which a potentially relocated pointer is ‘observably-after’ it may have been relocated. This is slightly stronger than is strictly necessary (and thus may disallow some otherwise valid programs), but greatly simplifies reasoning about correctness of the compiled code.

By construction, this property will be upheld by the optimizer if correctly established in the source IR. This is a key invariant of the design.

The existing IR Verifier pass has been extended to check most of the local restrictions on the intrinsics mentioned in their respective documentation. The current implementation in LLVM does not check the key relocation invariant, but this is ongoing work on developing such a verifier. Please ask on llvm-dev if you’re interested in experimenting with the current version.

Utility Passes for Safepoint Insertion

RewriteStatepointsForGC

The pass RewriteStatepointsForGC transforms a function’s IR to lower from the abstract machine model described above to the explicit statepoint model of relocations. To do this, it replaces all calls or invokes of functions which might contain a safepoint poll with a gc.statepoint and associated full relocation sequence, including all required gc.relocates.

This pass only applies to GCStrategy instances where the UseRS4GC flag is set. The two builtin GC strategies with this set are the “statepoint-example” and “coreclr” strategies.

As an example, given this code:

define ptr addrspace(1) @test1(ptr addrspace(1) %obj)
       gc "statepoint-example" {
  call void @foo()
  ret ptr addrspace(1) %obj
}

The pass would produce this IR:

define ptr addrspace(1) @test_rs4gc(ptr addrspace(1) %obj) gc "statepoint-example" {
  %statepoint_token = call token (i64, i32, ptr, i32, i32, ...) @llvm.experimental.gc.statepoint.p0(i64 2882400000, i32 0, ptr elementtype(void ()) @foo, i32 0, i32 0, i32 0, i32 0) [ "gc-live"(ptr addrspace(1) %obj) ]
  %obj.relocated = call coldcc ptr addrspace(1) @llvm.experimental.gc.relocate.p1(token %statepoint_token, i32 0, i32 0) ; (%obj, %obj)
  ret ptr addrspace(1) %obj.relocated
}

In the above examples, the addrspace(1) marker on the pointers is the mechanism that the statepoint-example GC strategy uses to distinguish references from non references. This is controlled via GCStrategy::isGCManagedPointer. The statepoint-example and coreclr strategies (the only two default strategies that support statepoints) both use addrspace(1) to determine which pointers are references, however custom strategies don’t have to follow this convention.

This pass can be used an utility function by a language frontend that doesn’t want to manually reason about liveness, base pointers, or relocation when constructing IR. As currently implemented, RewriteStatepointsForGC must be run after SSA construction (i.e. mem2ref).

RewriteStatepointsForGC will ensure that appropriate base pointers are listed for every relocation created. It will do so by duplicating code as needed to propagate the base pointer associated with each pointer being relocated to the appropriate safepoints. The implementation assumes that the following IR constructs produce base pointers: loads from the heap, addresses of global variables, function arguments, function return values. Constant pointers (such as null) are also assumed to be base pointers. In practice, this constraint can be relaxed to producing interior derived pointers provided the target collector can find the associated allocation from an arbitrary interior derived pointer.

By default RewriteStatepointsForGC passes in 0xABCDEF00 as the statepoint ID and 0 as the number of patchable bytes to the newly constructed gc.statepoint. These values can be configured on a per-callsite basis using the attributes "statepoint-id" and "statepoint-num-patch-bytes". If a call site is marked with a "statepoint-id" function attribute and its value is a positive integer (represented as a string), then that value is used as the ID of the newly constructed gc.statepoint. If a call site is marked with a "statepoint-num-patch-bytes" function attribute and its value is a positive integer, then that value is used as the ‘num patch bytes’ parameter of the newly constructed gc.statepoint. The "statepoint-id" and "statepoint-num-patch-bytes" attributes are not propagated to the gc.statepoint call or invoke if they could be successfully parsed.

In practice, RewriteStatepointsForGC should be run much later in the pass pipeline, after most optimization is already done. This helps to improve the quality of the generated code when compiled with garbage collection support.

RewriteStatepointsForGC intrinsic lowering

As a part of lowering to the explicit model of relocations RewriteStatepointsForGC performs GC specific lowering for the following intrinsics:

• gc.get.pointer.base

• gc.get.pointer.offset

• llvm.memcpy.element.unordered.atomic.*

• llvm.memmove.element.unordered.atomic.*

There are two possible lowerings for the memcpy and memmove operations: GC leaf lowering and GC parseable lowering. If a call is explicitly marked with “gc-leaf-function” attribute the call is lowered to a GC leaf call to ‘__llvm_memcpy_element_unordered_atomic_*’ or ‘__llvm_memmove_element_unordered_atomic_*’ symbol. Such a call can not take a safepoint. Otherwise, the call is made GC parseable by wrapping the call into a statepoint. This makes it possible to take a safepoint during copy operation. Note that a GC parseable copy operation is not required to take a safepoint. For example, a short copy operation may be performed without taking a safepoint.

GC parseable calls to ‘llvm.memcpy.element.unordered.atomic.*’, ‘llvm.memmove.element.unordered.atomic.*’ intrinsics are lowered to calls to ‘__llvm_memcpy_element_unordered_atomic_safepoint_*’, ‘__llvm_memmove_element_unordered_atomic_safepoint_*’ symbols respectively. This way the runtime can provide implementations of copy operations with and without safepoints.

GC parseable lowering also involves adjusting the arguments for the call. Memcpy and memmove intrinsics take derived pointers as source and destination arguments. If a copy operation takes a safepoint it might need to relocate the underlying source and destination objects. This requires the corresponding base pointers to be available in the copy operation. In order to make the base pointers available RewriteStatepointsForGC replaces derived pointers with base pointer and offset pairs. For example:

declare void @__llvm_memcpy_element_unordered_atomic_safepoint_1(
  i8 addrspace(1)*  %dest_base, i64 %dest_offset,
  i8 addrspace(1)*  %src_base, i64 %src_offset,
  i64 %length)

PlaceSafepoints

The pass PlaceSafepoints inserts safepoint polls sufficient to ensure running code checks for a safepoint request on a timely manner. This pass is expected to be run before RewriteStatepointsForGC and thus does not produce full relocation sequences.

As an example, given input IR of the following:

define void @test() gc "statepoint-example" {
  call void @foo()
  ret void
}

declare void @do_safepoint()
define void @gc.safepoint_poll() {
  call void @do_safepoint()
  ret void
}

This pass would produce the following IR:

define void @test() gc "statepoint-example" {
  call void @do_safepoint()
  call void @foo()
  ret void
}

In this case, we’ve added an (unconditional) entry safepoint poll. Note that despite appearances, the entry poll is not necessarily redundant. We’d have to know that foo and test were not mutually recursive for the poll to be redundant. In practice, you’d probably want to your poll definition to contain a conditional branch of some form.

At the moment, PlaceSafepoints can insert safepoint polls at method entry and loop backedges locations. Extending this to work with return polls would be straight forward if desired.

PlaceSafepoints includes a number of optimizations to avoid placing safepoint polls at particular sites unless needed to ensure timely execution of a poll under normal conditions. PlaceSafepoints does not attempt to ensure timely execution of a poll under worst case conditions such as heavy system paging.

The implementation of a safepoint poll action is specified by looking up a function of the name gc.safepoint_poll in the containing Module. The body of this function is inserted at each poll site desired. While calls or invokes inside this method are transformed to a gc.statepoints, recursive poll insertion is not performed.

This pass is useful for any language frontend which only has to support garbage collection semantics at safepoints. If you need other abstract frame information at safepoints (e.g. for deoptimization or introspection), you can insert safepoint polls in the frontend. If you have the later case, please ask on llvm-dev for suggestions. There’s been a good amount of work done on making such a scheme work well in practice which is not yet documented here.

Supported Architectures

Support for statepoint generation requires some code for each backend. Today, only Aarch64 and X86_64 are supported.

Limitations and Half Baked Ideas

Mixing References and Raw Pointers

Support for languages which allow unmanaged pointers to garbage collected objects (i.e. pass a pointer to an object to a C routine) in the abstract machine model. At the moment, the best idea on how to approach this involves an intrinsic or opaque function which hides the connection between the reference value and the raw pointer. The problem is that having a ptrtoint or inttoptr cast (which is common for such use cases) breaks the rules used for inferring base pointers for arbitrary references when lowering out of the abstract model to the explicit physical model. Note that a frontend which lowers directly to the physical model doesn’t have any problems here.

Objects on the Stack

As noted above, the explicit lowering supports objects allocated on the stack provided the collector can find a heap map given the stack address.

The missing pieces are a) integration with rewriting (RS4GC) from the abstract machine model and b) support for optionally decomposing on stack objects so as not to require heap maps for them. The later is required for ease of integration with some collectors.

Lowering Quality and Representation Overhead

The current statepoint lowering is known to be somewhat poor. In the very long term, we’d like to integrate statepoints with the register allocator; in the near term this is unlikely to happen. We’ve found the quality of lowering to be relatively unimportant as hot-statepoints are almost always inliner bugs.

Concerns have been raised that the statepoint representation results in a large amount of IR being produced for some examples and that this contributes to higher than expected memory usage and compile times. There’s no immediate plans to make changes due to this, but alternate models may be explored in the future.

Relocations Along Exceptional Edges

Relocations along exceptional paths are currently broken in ToT. In particular, there is current no way to represent a rethrow on a path which also has relocations. See this llvm-dev discussion for more detail.

Bugs and Enhancements

Currently known bugs and enhancements under consideration can be tracked by performing a bugzilla search for [Statepoint] in the summary field. When filing new bugs, please use this tag so that interested parties see the newly filed bug. As with most LLVM features, design discussions take place on the Discourse forums and patches should be sent to llvm-commits for review.