JIT Compiler Structure

Introduction

RyuJIT is the code name for the next generation Just in Time Compiler (aka “JIT”) for the AMD64 .NET runtime. Its first implementation is for the AMD64 architecture. It is derived from a code base that is still in use for the other targets of .NET.

The primary design considerations for RyuJIT are to:

• Maintain a high compatibility bar with previous JITs, especially those for x86 (jit32) and x64 (jit64).
• Support and enable good runtime performance through code optimizations, register allocation, and code generation.
• Ensure good throughput via largely linear-order optimizations and transformations, along with limitations on tracked variables for analyses (such as dataflow) that are inherently super-linear.
• Ensure that the JIT architecture is designed to support a range of targets and scenarios.

The first objective was the primary motivation for evolving the existing code base, rather than starting from scratch or departing more drastically from the existing IR and architecture.

Execution Environment and External Interface

RyuJIT provides the just in time compilation service for the .NET runtime. The runtime itself is variously called the EE (execution engine), the VM (virtual machine) or simply the CLR (common language runtime). Depending upon the configuration, the EE and JIT may reside in the same or different executable files. RyuJIT implements the JIT side of the JIT/EE interfaces:

• ICorJitCompiler – this is the interface that the JIT compiler implements. This interface is defined in src/inc/corjit.h and its implementation is in src/jit/ee_il_dll.cpp. The following are the key methods on this interface:
  • compileMethod is the main entry point for the JIT. The EE passes it a ICorJitInfo object, and the “info” containing the IL, the method header, and various other useful tidbits. It returns a pointer to the code, its size, and additional GC, EH and (optionally) debug info.
  • getVersionIdentifier is the mechanism by which the JIT/EE interface is versioned. There is a single GUID (manually generated) which the JIT and EE must agree on.
  • getMaxIntrinsicSIMDVectorLength communicates to the EE the largest SIMD vector length that the JIT can support.
• ICorJitInfo – this is the interface that the EE implements. It has many methods defined on it that allow the JIT to look up metadata tokens, traverse type signatures, compute field and vtable offsets, find method entry points, construct string literals, etc. This bulk of this interface is inherited from ICorDynamicInfo which is defined in src/inc/corinfo.h. The implementation is defined in src/vm/jitinterface.cpp.

Internal Representation (IR)

Compiler object

The Compiler object is the primary data structure of the JIT. While it is not part of the JIT's IR per se, it serves as the root from which the data structures that implement the IR are accessible. For example, the Compiler object points to the head of the function's BasicBlock list with the fgFirstBB field, as well as having additional pointers to the end of the list, and other distinguished locations. ICorJitCompiler::CompileMethod() is invoked for each method, and creates a new Compiler object. Thus, the JIT need not worry about thread synchronization while accessing Compiler state. The EE has the necessary synchronization to ensure there is a single JIT’d copy of a method when two or more threads try to trigger JIT compilation of the same method.

Overview of the IR

RyuJIT represents a function as a doubly-linked list of BasicBlock values. Each BasicBlock has explicit edges to its successors that define the function's non-exceptional control flow. Exceptional control flow is implicit, with protected regions and handlers described in a table of EHblkDsc values. At the beginning of a compilation, each BasicBlock contains nodes in a high-level, statement- and tree-oriented form (HIR); this form persists throughout the JIT's front end. During the first phase of the back end--the rationalization phase--the HIR for each block is lowered to a linearly-ordered, node-oriented form (LIR). The fundamental distinction between HIR and LIR is in ordering semantics, though there are also some restrictions on the types of nodes that may appear in an HIR or LIR block.

Both HIR and LIR blocks are composed of GenTree nodes that define the operations performed by the block. A GenTree node may consume some number of operands and may produce a singly-defined, at-most-singly-used value as a result. These values are referred to interchangably as SDSU temps or tree temps. Defs of SDSU temps are represented by GenTree nodes themselves, and uses are represented by edges from the using node to the defining node. Furthermore, SDSU temps defined in one block may not be used in a different block. In cases where a value must be multiply-defined, multiply-used, or defined in one block and used in another, the IR provides another class of temporary: the local var. Local vars are defined by assignment nodes in HIR or store nodes in LIR, and are used by local var nodes in both forms.

An HIR block is composed of a doubly-linked list of statement nodes (GenTreeStmt), each of which references a single expression tree (gtStmtExpr). The GenTree nodes in this tree execute in "tree order", which is defined as the order produced by a depth-first, left-to-right traversal of the tree, with two notable exceptions:

• Binary nodes marked with the GTF_REVERSE_OPS flag execute their right operand tree (gtOp2) before their left operand tree (gtOp1)
• Dynamically-sized block copy nodes where gtEvalSizeFirst is true execute the gtDynamicSize tree before executing their other operand trees In addition to tree order, HIR also requires that no SDSU temp is defined in one statement and used in another. In situations where the requirements of tree and statement order prove onerous (e.g. when code must execute at a particular point in a function), HIR provides GT_COMMA nodes as an escape valve: these nodes consume and discard the results of their left-hand side while producing a copy of the vaule produced by their right-hand side. This allows the compiler to insert code in the middle of a statement without requiring that the statement be split apart.

An LIR block is composed of a doubly-linked list of GenTree nodes, each of which describes a single operation in the method. These nodes execute in the order given by the list; there is no relationship between the order in which a node's operands appear and the order in which the operators that produced those operands execute. The only exception to this rule occurs after the register allocator, which may introduce GT_COPY and GT_RELOAD nodes that execute in "spill order". Spill order is defined as the order in which the register allocator visits a node's operands. For correctness, the code generator must generate code for spills, reloads, and GT_COPY/GT_RELOAD nodes in this order.

In addition to HIR and LIR BasicBlocks, a separate representation--insGroup and instrDesc--is used during the actual instruction encoding.

GenTree Nodes

Each operation is represented as a GenTree node, with an opcode (GT_xxx), zero or more child/operand GenTree nodes, and additional fields as needed to represent the semantics of that node. Every node includes its type, value number, assertions, register assignments, etc. when available.

GenTree nodes are doubly-linked in execution order, but the links are not necessarily valid during all phases of the JIT. In HIR these links are primarily a convenience, as the order produced by a traversal of the links must match the order produced by a "tree order" traversal (see above for details). In LIR these links define the execution order of the nodes.

HIR statement nodes utilize the same GenTree base type as the operation nodes, though they are not truly related.

• The statement nodes are doubly-linked. The first statement node in a block points to the last node in the block via its gtPrev link. Note that the last statement node does not point to the first; that is, the list is not fully circular.
• Each statement node contains two GenTree links – gtStmtExpr points to the top-level node in the statement (i.e. the root of the tree that represents the statement), while gtStmtList points to the first node in execution order (again, this link is not always valid).

Local var descriptors

A LclVarDsc represents a possibly-multiply-defined, possibly-multiply-used temporary. These temporaries may be used to represent user local variables, arguments or JIT-created temps. Each lclVar has a gtLclNum which is the identifier usually associated with the variable in the JIT and its dumps. The LclVarDsc contains the type, use count, weighted use count, frame or register assignment etc. A local var may be "tracked" (lvTracked), in which case it participates in dataflow analysis, and has a secondary name (lvVarIndex) that allows for the use of dense bit vectors.

Example of Post-Import IR

For this snippet of code (extracted from tests/src/JIT/CodeGenBringUpTests/DblRoots.cs):

   r1 = (-b + Math.Sqrt(b*b - 4*a*c))/(2*a);

A stripped-down dump of the GenTree nodes just after they are imported looks like this:

▌ stmtExpr  void  (top level) (IL 0x000...0x026)
│        ┌──▌ lclVar    double V00 arg0
│     ┌──▌ *         double
│     │  └──▌ dconst    double 2.00
│  ┌──▌ /         double
│  │  │  ┌──▌ mathFN    double sqrt
│  │  │  │  │     ┌──▌ lclVar    double V02 arg2
│  │  │  │  │  ┌──▌ *         double
│  │  │  │  │  │  │  ┌──▌ lclVar    double V00 arg0
│  │  │  │  │  │  └──▌ *         double
│  │  │  │  │  │     └──▌ dconst    double 4.00
│  │  │  │  └──▌ -         double
│  │  │  │     │       lclVar    double V01 arg1
│  │  │  │     └──▌ *         double
│  │  │  │        └──▌ lclVar    double V01 arg1
│  │  └──▌ +         double
│  │     └──▌ unary -   double
│  │        └──▌ lclVar    double V01 arg1
└──▌  =         double
   └──▌ indir     double
      └──▌ lclVar    byref  V03 arg3

Types

The JIT is primarily concerned with “primitive” types, i.e. integers, reference types, pointers, and floating point types. It must also be concerned with the format of user-defined value types (i.e. struct types derived from System.ValueType) – specifically, their size and the offset of any GC references they contain, so that they can be correctly initialized and copied. The primitive types are represented in the JIT by the var_types enum, and any additional information required for struct types is obtained from the JIT/EE interface by the use of an opaque CORINFO_CLASS_HANDLE.

Dataflow Information

In order to limit throughput impact, the JIT limits the number of lclVars for which liveness information is computed. These are the tracked lclVars (lvTracked is true), and they are the only candidates for register allocation (i.e. only these lclVars may be assigned registers over their entire lifetime). Defs and uses of untracked lclVars are treated as stores and loads to/from the appropriate stack location, and the corresponding nodes act as normal operators during register allocation.

The liveness analysis determines the set of defs, as well as the uses that are upward exposed, for each block. It then propagates the liveness information. The result of the analysis is captured in the following:

• The live-in and live-out sets are captured in the bbLiveIn and bbLiveOut fields of the BasicBlock.
• The GTF_VAR_DEF flag is set on a lclVar node (all of which are of type GenTreeLclVarCommon) that is a definition.
• The GTF_VAR_USEASG flag is set (in addition to the GTF_VAR_DEF flag) for the target of an update (e.g. +=).

SSA

Static single assignment (SSA) form is constructed in a traditional manner [1]. The SSA names are recorded on the lclVar references. While SSA form usually retains a pointer or link to the defining reference, RyuJIT currently retains only the BasicBlock in which the definition of each SSA name resides.

Value Numbering

Value numbering utilizes SSA for lclVar values, but also performs value numbering of expression trees. It takes advantage of type safety by not invalidating the value number for field references with a heap write, unless the write is to the same field. The IR nodes are annotated with the value numbers, which are indexes into a type-specific value number store. Value numbering traverses the trees, performing symbolic evaluation of many operations.

Phases of RyuJIT

The top-level function of interest is Compiler::compCompile. It invokes the following phases in order.

Pre-import

Prior to reading in the IL for the method, the JIT initializes the local variable table, and scans the IL to find branch targets and form BasicBlocks.

Importation

Importation is the phase that creates the IR for the method, reading in one IL instruction at a time, and building up the statements. During this process, it may need to generate IR with multiple, nested expressions. This is the purpose of the non-expression-like IR nodes:

• It may need to evaluate part of the expression into a temp, in which case it will use a comma (GT_COMMA) node to ensure that the temp is evaluated in the proper execution order – i.e. GT_COMMA(GT_ASG(temp, exp), temp) is inserted into the tree where “exp” would go.
• It may need to create conditional expressions, but adding control flow at this point would be quite messy. In this case it generates question mark/colon (?: or GT_QMARK/GT_COLON) trees that may be nested within an expression.

During importation, tail call candidates (either explicitly marked or opportunistically identified) are identified and flagged. They are further validated, and possibly unmarked, during morphing.

Morphing

The fgMorph phase includes a number of transformations:

Inlining

The fgInline phase determines whether each call site is a candidate for inlining. The initial determination is made via a state machine that runs over the candidate method’s IL. It estimates the native code size corresponding to the inline method, and uses a set of heuristics, including the estimated size of the current method) to determine if inlining would be profitable. If so, a separate Compiler object is created, and the importation phase is called to create the tree for the candidate inline method. Inlining may be aborted prior to completion, if any conditions are encountered that indicate that it may be unprofitable (or otherwise incorrect). If inlining is successful, the inlinee compiler’s trees are incorporated into the inliner compiler (the “parent”), with args and returns appropriately transformed.

Struct Promotion

Struct promotion (fgPromoteStructs()) analyzes the local variables and temps, and determines if their fields are candidates for tracking (and possibly enregistering) separately. It first determines whether it is possible to promote, which takes into account whether the layout may have holes or overlapping fields, whether its fields (flattening any contained structs) will fit in registers, etc.

Next, it determines whether it is likely to be profitable, based on the number of fields, and whether the fields are individually referenced.

When a lclVar is promoted, there are now N+1 lclVars for the struct, where N is the number of fields. The original struct lclVar is not considered to be tracked, but its fields may be.

Mark Address-Exposed Locals

This phase traverses the expression trees, propagating the context (e.g. taking the address, indirecting) to determine which lclVars have their address taken, and which therefore will not be register candidates. If a struct lclVar has been promoted, and is then found to be address-taken, it will be considered “dependently promoted”, which is an odd way of saying that the fields will still be separately tracked, but they will not be register candidates.

Morph Blocks

What is often thought of as “morph” involves localized transformations to the trees. In addition to performing simple optimizing transformations, it performs some normalization that is required, such as converting field and array accesses into pointer arithmetic. It can (and must) be called by subsequent phases on newly added or modified trees. During the main Morph phase, the boolean fgGlobalMorph is set on the Compiler argument, which governs which transformations are permissible.

Eliminate Qmarks

This expands most GT_QMARK/GT_COLON trees into blocks, except for the case that is instantiating a condition.

Flowgraph Analysis

At this point, a number of analyses and transformations are done on the flowgraph:

• Computing the predecessors of each block
• Computing edge weights, if profile information is available
• Computing reachability and dominators
• Identifying and normalizing loops (transforming while loops to “do while”)
• Cloning and unrolling of loops

Normalize IR for Optimization

At this point, a number of properties are computed on the IR, and must remain valid for the remaining phases. We will call this “normalization”

• lvaMarkLocalVars – set the reference counts (raw and weighted) for lclVars, sort them, and determine which will be tracked (currently up to 128). Note that after this point any transformation that adds or removes lclVar references must update the reference counts.
• optOptimizeBools – this optimizes Boolean expressions, and may change the flowgraph (why is it not done prior to reachability and dominators?)
• Link the trees in evaluation order (setting gtNext and gtPrev fields): and fgFindOperOrder() and fgSetBlockOrder().

SSA and Value Numbering Optimizations

The next set of optimizations are built on top of SSA and value numbering. First, the SSA representation is built (during which dataflow analysis, aka liveness, is computed on the lclVars), then value numbering is done using SSA.

Loop Invariant Code Hoisting

This phase traverses all the loop nests, in outer-to-inner order (thus hoisting expressions outside the largest loop in which they are invariant). It traverses all of the statements in the blocks in the loop that are always executed. If the statement is:

• A valid CSE candidate
• Has no side-effects
• Does not raise an exception OR occurs in the loop prior to any side-effects
• Has a valid value number, and it is a lclVar defined outside the loop, or its children (the value numbers from which it was computed) are invariant.

Copy Propagation

This phase walks each block in the graph (in dominator-first order, maintaining context between dominator and child) keeping track of every live definition. When it encounters a variable that shares the VN with a live definition, it is replaced with the variable in the live definition.

The JIT currently requires that the IR be maintained in conventional SSA form, as there is no “out of SSA” translation (see the comments on optVnCopyProp() for more information).

Common Subexpression Elimination (CSE)

Utilizes value numbers to identify redundant computations, which are then evaluated to a new temp lclVar, and then reused.

Assertion Propagation

Utilizes value numbers to propagate and transform based on properties such as non-nullness.

Range analysis

Optimize array index range checks based on value numbers and assertions.

Rationalization

As the JIT has evolved, changes have been made to improve the ability to reason over the tree in both “tree order” and “linear order”. These changes have been termed the “rationalization” of the IR. In the spirit of reuse and evolution, some of the changes have been made only in the later (“backend”) components of the JIT. The corresponding transformations are made to the IR by a “Rationalizer” component. It is expected that over time some of these changes will migrate to an earlier place in the JIT phase order:

• Elimination of assignment nodes (GT_ASG). The assignment node was problematic because the semantics of its destination (left hand side of the assignment) could not be determined without context. For example, a GT_LCL_VAR on the left-hand side of an assignment is a definition of the local variable, but on the right-hand side it is a use. Furthermore, since the execution order requires that the children be executed before the parent, it is unnatural that the left-hand side of the assignment appears in execution order before the assignment operator.
  • During rationalization, all assignments are replaced by stores, which either represent their destination on the store node itself (e.g. GT_LCL_VAR), or by the use of a child address node (e.g. GT_STORE_IND).
• Elimination of address nodes (GT_ADDR). These are problematic because of the need for parent context to analyze the child.
• Elimination of “comma” nodes (GT_COMMA). These nodes are introduced for convenience during importation, during which a single tree is constructed at a time, and not incorporated into the statement list until it is completed. When it is necessary, for example, to store a partially-constructed tree into a temporary variable, a GT_COMMA node is used to link it into the tree. However, in later phases, these comma nodes are an impediment to analysis, and thus are eliminated.
  • In some cases, it is not possible to fully extract the tree into a separate statement, due to execution order dependencies. In these cases, an “embedded” statement is created. While these are conceptually very similar to the GT_COMMA nodes, they do not masquerade as expressions.
• Elimination of “QMark” (GT_QMARK/GT_COLON) nodes is actually done at the end of morphing, long before the current rationalization phase. The presence of these nodes made analyses (especially dataflow) overly complex.
• Elimination of statements. Without statements, the execution order of a basic block's contents is fully defined by the gtNext/gtPrev links between GenTree nodes.

For our earlier example (Example of Post-Import IR), here is what the simplified dump looks like just prior to Rationalization (the $ annotations are value numbers). Note that some common subexpressions have been computed into new temporary lclVars, and that computation has been inserted as a GT_COMMA (comma) node in the IR:

▌  stmtExpr  void  (top level) (IL 0x000...0x026)
│        ┌──▌  lclVar    double V07 cse1          $185
│     ┌──▌  comma     double                      $185
│     │  │     ┌──▌  dconst    double 2.00        $143
│     │  │  ┌──▌  \*         double                $185
│     │  │  │  └──▌  lclVar   double V00 arg0 u:2 $80
│     │  └──▌  =         double                   $VN.Void
│     │     └──▌  lclVar    double V07 cse1       $185
│  ┌──▌  /         double                         $186
│  │  │  ┌──▌  unary -   double                   $84
│  │  │  │  └──▌  lclVar    double V01 arg1   u:2 $81
│  │  └──▌  +         double                      $184
│  │     │  ┌──▌  lclVar    double V06 cse0       $83
│  │     └──▌  comma     double                   $83
│  │        │  ┌──▌  mathFN    double sqrt        $83
│  │        │  │  │     ┌──▌  lclVar double V02 arg2 u:2 $82
│  │        │  │  │  ┌──▌  \*         double              $182
│  │        │  │  │  │  │  ┌──▌  dconst    double 4.00   $141
│  │        │  │  │  │  └──▌  \*         double           $181
│  │        │  │  │  │     └──▌  lclVar double V00 arg0 u:2 $80
│  │        │  │  └──▌  -         double                    $183
│  │        │  │     │  ┌──▌  lclVar    double V01 arg1 u:2 $81
│  │        │  │     └──▌  \*         double                 $180
│  │        │  │        └──▌  lclVar    double V01 arg1 u:2 $81
│  │        └──▌  =         double                          $VN.Void
│  │           └──▌  lclVar    double V06 cse0              $83
└──▌  =         double                                      $VN.Void
   └──▌  indir     double $186
      └──▌  lclVar    byref  V03 arg3        u:2 (last use) $c0

After rationalization, the nodes are presented in execution order, and the GT_COMMA (comma), GT_ASG (=), and GT_STMT nodes have been eliminated:

t0   = lclVar     double V01 arg1
t1   = lclVar     double V01 arg1
       ┌──▌  t0   double
       ├──▌  t1   double
t2   = \*          double
t3   = lclVar     double V00 arg0
t4   = dconst     double 4.00
       ┌──▌  t3   double
       ├──▌  t4   double
t5   = \*          double
t6   = lclVar     double V02 arg2
       ┌──▌  t5   double
       ├──▌  t6   double
t7   = \*          double
       ┌──▌  t2   double
       ├──▌  t7   double
t8   = -          double
       ┌──▌  t8   double
t9   = mathFN     double sqrt
       ┌──▌  t9   double
       st.lclVar  double V06
t10  = lclVar     double V06
t11  = lclVar     double V01 arg1
       ┌──▌  t11  double
t12  = unary -    double
       ┌──▌  t10  double
       ├──▌  t12  double
t13  = +          double
t14  = lclVar     double V00 arg0
t15  = dconst     double 2.00
       ┌──▌  t14  double
       ├──▌  t15  double
t16  = \*          double
       ┌──▌  t16  double
       st.lclVar  double V07
t18  = lclVar     double V07
       ┌──▌  t17  double
       ├──▌  t18  double
t19  = /          double
t20  = lclVar     byref  V03 arg3
       ┌──▌  t20  double
       storeIndir double

Lowering

Lowering is responsible for transforming the IR in such a way that the control flow, and any register requirements, are fully exposed.

It accomplishes this in two passes.

The first pass is an execution-order traversal that performs context-dependent transformations such as expanding switch statements (using a switch table or a series of conditional branches), constructing addressing modes, etc. For example, this:

t0   = lclVar     ref    V00 arg0
t1   = lclVar     int    V03 loc1
       ┌──▌  t1   int
t2   = cast       long <- int
t3   = const      long   2
       ┌──▌  t2   long
       ├──▌  t3   long
t4   = <<         long
       ┌──▌  t0   ref
       ├──▌  t4   long
t5   = +          byref
t6   = const      long   16
       ┌──▌  t5   ref
       ├──▌  t6   long
t7   = +          byref
       ┌──▌  t7   ref
t8   = indir      int

Is transformed into this, in which the addressing mode is explicit:

t0   = lclVar     ref    V00 arg0
t1   = lclVar     int    V03 loc1
       ┌──▌  t1   int
t2   = cast       long <- int
       ┌──▌  t0   ref
       ├──▌  t2   long
t7   = lea(b+(i*4)+16) byref
       ┌──▌  t7   ref
t8   = indir      int

The next pass annotates the nodes with register requirements. This is done in an execution order traversal in order to ensure that each node's operands are visited prior to the node itself. This pass may also do some transformations that do not require the parent context, such as determining the code generation strategy for block assignments (e.g. GT_COPYBLK) which may become helper calls, unrolled loops, or an instruction like rep stos.

The register requirements of a node are expressed by its TreeNodeInfo (gtLsraInfo) value. For example, for the copyBlk node in this snippet:

t0   = const(h)   long   0xCA4000 static
t1   = &lclVar    byref  V04 loc4
t2   = const      int    34
       ┌──▌  t0   long
       ├──▌  t1   byref
       ├──▌  t2   int
       copyBlk    void

The TreeNodeInfo would be as follows:

+<TreeNodeInfo @ 15 0=1 1i 1f
      src=[allInt]
      int=[rax rcx rdx rbx rbp rsi rdi r8-r15 mm0-mm5]
      dst=[allInt] I>

The “@ 15” is the location number of the node. The “0=1” indicates that there are zero destination registers (because this defines only memory), and 1 source register (the address of lclVar V04). The “1i” indicates that it requires 1 internal integer register (for copying the remainder after copying 16-byte sized chunks), the “1f” indicates that it requires 1 internal floating point register (for copying the two 16-byte chunks). The src, int and dst fields are encoded masks that indicate the register constraints for the source, internal and destination registers, respectively.

Register allocation

The RyuJIT register allocator uses a Linear Scan algorithm, with an approach similar to [2]. In brief, it operates on two main data structures:

• Intervals (representing live ranges of variables or tree expressions) and RegRecords (representing physical registers), both of which derive from Referenceable.
• RefPositions, which represent uses or defs (or variants thereof, such as ExposedUses) of either Intervals or physical registers.

Pre-conditions:

• The NodeInfo is initialized for each tree node to indicate:
  • Number of registers consumed and produced by the node.
  • Number and type (int versus float) of internal registers required.

Allocation proceeds in 4 phases:

• Determine the order in which the BasicBlocks will be allocated, and which predecessor of each block will be used to determine the starting location for variables live-in to the BasicBlock.
• Construct Intervals for each tracked lclVar, then walk the BasicBlocks in the determined order building RefPositions for each register use, def, or kill.
• Allocate the registers by traversing the RefPositions.
• Write back the register assignments, and perform any necessary moves at block boundaries where the allocations don’t match.

Post-conditions:

• The gtRegNum property of all GenTree nodes that require a register has been set to a valid register number.
• The gtRsvdRegs field (a set/mask of registers) has the requested number of registers specified for internal use.
• All spilled values (lclVar or expression) are marked with GTF_SPILL at their definition. For lclVars, they are also marked with GTF_SPILLED at any use at which the value must be reloaded.
• For all lclVars that are register candidates:
  • lvRegNum = initial register location (or REG_STK)
  • lvRegister flag set if it always lives in the same register
  • lvSpilled flag is set if it is ever spilled
• The maximum number of simultaneously-live spill locations of each type (used for spilling expression trees) has been communicated via calls to compiler->tmpPreAllocateTemps(type).

Code Generation

The process of code generation is relatively straightforward, as Lowering has done some of the work already. Code generation proceeds roughly as follows:

• Determine the frame layout – allocating space on the frame for any lclVars that are not fully enregistered, as well as any spill temps required for spilling non-lclVar expressions.
• For each BasicBlock, in layout order, and each GenTree node in the block, in execution order:
  • If the node is “contained” (i.e. its operation is subsumed by a parent node), do nothing.
  • Otherwise, “consume” all the register operands of the node.
    • This updates the liveness information (i.e. marking a lclVar as dead if this is the last use), and performs any needed copies.
    • This must be done in "spill order" so that any spill/restore code inserted by the register allocator to resolve register conflicts is generated in the correct order. "
  • Track the live variables in registers, as well as the live stack variables that contain GC refs.
  • Produce the instrDesc(s) for the operation, with the current live GC references.
  • Update the scope information (debug info) at block boundaries.
• Generate the prolog and epilog code.
• Write the final instruction bytes. It does this by invoking the emitter, which holds all the instrDescs.

Phase-dependent Properties and Invariants of the IR

There are several properties of the IR that are valid only during (or after) specific phases of the JIT. This section describes the phase transitions, and how the IR properties are affected.

Phase Transitions

• Flowgraph analysis
  • Sets the predecessors of each block, which must be kept valid after this phase.
  • Computes reachability and dominators. These may be invalidated by changes to the flowgraph.
  • Computes edge weights, if profile information is available.
  • Identifies and normalizes loops. These may be invalidated, but must be marked as such.
• Normalization
  • The lclVar reference counts are set by lvaMarkLocalVars().
  • Statement ordering is determined by fgSetBlockOrder(). Execution order is a depth-first preorder traversal of the nodes, with the operands usually executed in order. The exceptions are:
    • Binary operators, which can have the GTF_REVERSE_OPS flag set to indicate that the RHS (gtOp2) should be evaluated before the LHS (gtOp1).
    • Dynamically-sized block copy nodes, which can have gtEvalSizeFirst set to true to indicate that their gtDynamicSize tree should be evaluated before executing their other operands.
• Rationalization
  • All GT_ASG trees are transformed into GT_STORE variants (e.g. GT_STORE_LCL_VAR).
  • All GT_ADDR nodes are eliminated (e.g. with GT_LCL_VAR_ADDR).
  • All GT_COMMA and GT_STMT nodes are removed and their constituent nodes linked into execution order.
• Lowering
  • GenTree nodes are split or transformed as needed to expose all of their register requirements and any necessary flowgraph changes (e.g., for switch statements).

GenTree phase-dependent properties

Ordering:

• For GenTreeStmt nodes, the gtNext and gtPrev fields must always be consistent. The last statement in the BasicBlock must have gtNext equal to null. By convention, the gtPrev of the first statement in the BasicBlock must be the last statement of the BasicBlock.
  • In all phases, gtStmtExpr points to the top-level node of the expression.
• For non-statement nodes, the gtNext and gtPrev fields are either null, prior to ordering, or they are consistent (i.e. A->gtPrev->gtNext = A, and A->gtNext->gtPrev == A, if they are non-null).
• After normalization the gtStmtList of the containing statement points to the first node to be executed.
• Prior to normalization, the gtNext and gtPrev pointers on the expression (non-statement) GenTree nodes are invalid. The expression nodes are only traversed via the links from parent to child (e.g. node->gtGetOp1(), or node->gtOp.gtOp1). The gtNext/gtPrev links are set by fgSetBlockOrder().
  • After normalization, and prior to rationalization, the parent/child links remain the primary traversal mechanism. The evaluation order of any nested expression-statements (usually assignments) is enforced by the GT_COMMA in which they are contained.
• After rationalization, all GT_COMMA nodes are eliminated, statements are flattened, and the primary traversal mechanism becomes the gtNext/gtPrev links which define the execution order.
• In tree ordering:
  • The gtPrev of the first node (gtStmtList) is always null.
  • The gtNext of the last node (gtStmtExpr) is always null.

TreeNodeInfo:

• The TreeNodeInfo (gtLsraInfo) is set during the Lowering phase, and communicates the register requirements of the node, including the number and types of registers used as sources, destinations and internal registers. Currently only a single destination per node is supported.

LclVar phase-dependent properties

Prior to normalization, the reference counts (lvRefCnt and lvRefCntWtd) are not valid. After normalization they must be updated when lclVar references are added or removed.

Supporting technologies and components

Instruction encoding

Instruction encoding is performed by the emitter (emit.h), using the insGroup/instrDesc representation. The code generator calls methods on the emitter to construct instrDescs. The encodings information is captured in the following:

• The “instruction” enumeration itemizes the different instructions available on each target, and is used as an index into the various encoding tables (e.g. instInfo[], emitInsModeFmtTab[]) generated from the instrs{tgt}.h (e.g., instrsxarch.h).
• The skeleton encodings are contained in the tables, and then there are methods on the emitter that handle the special encoding constraints for the various instructions, addressing modes, register types, etc.

GC Info

Reporting of live GC references is done in two ways:

• For stack locations that are not tracked (these could be spill locations or lclVars – local variables or temps – that are not register candidates), they are initialized to null in the prolog, and reported as live for the entire method.
• For lclVars with tracked lifetimes, or for expression involving GC references, we report the range over which the reference is live. This is done by the emitter, which adds this information to the instruction group, and which terminates instruction groups when the GC info changes.

The tracking of GC reference lifetimes is done via the GCInfo class in the JIT. It is declared in src/jit/jitgcinfo.h (to differentiate it from src/inc/gcinfo.h), and implemented in src/jit/gcinfo.cpp.

In a JitDump, the generated GC info can be seen following the “In gcInfoBlockHdrSave()” line.

Debugger info

Debug info consists primarily of two types of information in the JIT:

• Mapping of IL offsets to native code offsets. This is accomplished via:
  • the gtStmtILoffsx on the statement nodes (GenTreeStmt)
  • the gtLclILoffs on lclVar references (GenTreeLclVar)
  • The IL offsets are captured during CodeGen by calling CodeGen::genIPmappingAdd(), and then written to debug tables by CodeGen::genIPmappingGen().
• Mapping of user locals to location (register or stack). This is accomplished via:
  • Struct siVarLoc (in compiler.h) captures the location
  • VarScopeDsc (compiler.h) captures the live range of a local variable in a given location.

Exception handling

Exception handling information is captured in an EHblkDsc for each exception handling region. Each region includes the first and last blocks of the try and handler regions, exception type, enclosing region, among other things. Look at jiteh.h and jiteh.cpp, especially, for details. Look at Compiler::fgVerifyHandlerTab() to see how the exception table constraints are verified.

Reading a JitDump

One of the best ways of learning about the JIT compiler is examining a compilation dump in detail. The dump shows you all the really important details of the basic data structures without all the implementation detail of the code. Debugging a JIT bug almost always begins with a JitDump. Only after the problem is isolated by the dump does it make sense to start debugging the JIT code itself.

Dumps are also useful because they give you good places to place breakpoints. If you want to see what is happening at some point in the dump, simply search for the dump text in the source code. This gives you a great place to put a conditional breakpoint.

There is not a strong convention about what or how the information is dumped, but generally you can find phase-specific information by searching for the phase name. Some useful points follow.

How to create a JitDump

You can enable dumps by setting the COMPlus_JitDump environment variable to a space-separated list of the method(s) you want to dump. For example:

:: Print out lots of useful info when
:: compiling methods named Main/GetEnumerator
set "COMPlus_JitDump=Main GetEnumerator"

See Setting configuration variables for more details on this.

Full instructions for dumping the compilation of some managed code can be found here: viewing-jit-dumps.md

Reading expression trees

It takes some time to learn to “read” the expression trees, which are printed with the children indented from the parent, and, for binary operators, with the first operand below the parent and the second operand above.

Here is an example dump

[000027] ------------             ▌  stmtExpr  void  (top level) (IL 0x010...  ???)
[000026] --C-G-------             └──▌  return    double
[000024] --C-G-------                └──▌  call      double BringUpTest.DblSqrt
[000021] ------------                   │     ┌──▌  lclVar    double V02 arg2
[000022] ------------                   │  ┌──▌  -         double
[000020] ------------                   │  │  └──▌  lclVar    double V03 loc0
[000023] ------------ arg0              └──▌  *         double
[000017] ------------                      │     ┌──▌  lclVar    double V01 arg1
[000018] ------------                      │  ┌──▌  -         double
[000016] ------------                      │  │  └──▌  lclVar    double V03 loc0
[000019] ------------                      └──▌  *         double
[000013] ------------                         │     ┌──▌  lclVar    double V00 arg0
[000014] ------------                         │  ┌──▌  -         double
[000012] ------------                         │  │  └──▌  lclVar    double V03 loc0
[000015] ------------                         └──▌  *         double
[000011] ------------                            └──▌  lclVar    double V03 loc0

The tree nodes are indented to represent the parent-child relationship. Binary operators print first the right hand side, then the operator node itself, then the left hand side. This scheme makes sense if you look at the dump “sideways” (lean your head to the left). Oriented this way, the left hand side operator is actually on the left side, and the right hand operator is on the right side so you can almost visualize the tree if you look at it sideways. The indentation level is also there as a backup.

Tree nodes are identified by their gtTreeID. This field only exists in DEBUG builds, but is quite useful for debugging, since all tree nodes are created from the routine gtNewNode (in src/jit/gentree.cpp). If you find a bad tree and wish to understand how it got corrupted, you can place a conditional breakpoint at the end of gtNewNode to see when it is created, and then a data breakpoint on the field that you believe is corrupted.

The trees are connected by line characters (either in ASCII, by default, or in slightly more readable Unicode when COMPlus_JitDumpAscii=0 is specified), to make it a bit easier to read.

N037 (  0,  0) [000391] ----------L- arg0 SETUP  │  ┌──▌  argPlace  ref    REG NA $1c1
N041 (  2,  8) [000389] ------------             │  │     ┌──▌  const(h) long 0xB410A098 REG rcx $240
N043 (  4, 10) [000390] ----G-------             │  │  ┌──▌  indir     ref    REG rcx $1c1
N045 (  4, 10) [000488] ----G------- arg0 in rcx │  ├──▌  putarg_reg ref    REG rcx
N049 ( 18, 16) [000269] --C-G-------             └──▌  call void System.Diagnostics.TraceInternal.Fail $VN.Void

Variable naming

The dump uses the index into the local variable table as its name. The arguments to the function come first, then the local variables, then any compiler generated temps. Thus in a function with 2 parameters (remember “this” is also a parameter), and one local variable, the first argument would be variable 0, the second argument variable 1, and the local variable would be variable 2. As described earlier, tracked variables are given a tracked variable index which identifies the bit for that variable in the dataflow bit vectors. This can lead to confusion as to whether the variable number is its index into the local variable table, or its tracked index. In the dumps when we refer to a variable by its local variable table index we use the ‘V’ prefix, and when we print the tracked index we prefix it by a ‘T’.

References

[1] P. Briggs, K. D. Cooper, T. J. Harvey, and L. T. Simpson, "Practical improvements to the construction and destruction of static single assignment form," Software --- Practice and Experience, vol. 28, no. 8, pp. 859---881, Jul. 1998.

[2] Wimmer, C. and Mössenböck, D. "Optimized Interval Splitting in a Linear Scan Register Allocator," ACM VEE 2005, pp. 132-141. http://portal.acm.org/citation.cfm?id=1064998&dl=ACM&coll=ACM&CFID=105967773&CFTOKEN=80545349