A Tour through the UNIX** C Compiler

D. M. Ritchie
Languages Computing


The Intermediate Language

      Communication between the two phases of the compiler proper is carried out by means of a pair of intermediate files. These files are treated as having identical structure, although the second file contains only the code generated for strings. It is convenient to write strings out separately to reduce the need for multiple location counters in a later assembly phase.

      The intermediate language is not machine-independent; its structure in a number of ways reflects the fact that C was originally a one-pass compiler chopped in two to reduce the maximum memory requirement. In fact, only the latest version of the compiler has a complete intermediate language at all. Until recently, the first phase of the compiler generated assembly code for those constructions it could deal with, and passed expression parse trees, in absolute binary form, to the second phase for code generation. Now, at least, all inter-phase information is passed in a describable form, and there are no absolute pointers involved, so the coupling between the phases is not so strong.

      The areas in which the machine (and system) dependencies are most noticeable are

Storage allocation for automatic variables and arguments has already been performed, and nodes for such variables refer to them by offset from a display pointer. Type conversion (for example, from integer to pointer) has already occurred using the assumption of byte addressing and 2-byte words.
Data representations suitable to the PDP-11 are assumed; in particular, floating point constants are passed as four words in the machine representation.

      As it happens, each intermediate file is represented as a sequence of binary numbers without any explicit demarcations. It consists of a sequence of conceptual lines, each headed by an operator, and possibly containing various operands. The operators are small numbers; to assist in recognizing failure in synchronization, the high-order byte of each operator word is always the octal number 376. Operands are either 16-bit binary numbers or strings of characters representing names. Each name is terminated by a null character. There is no alignment requirement for numerical operands and so there is no padding after a name string.

      The binary representation was chosen to avoid the necessity of converting to and from character form and to minimize the size of the files. It would be very easy to make each operator-operand `line' in the file be a genuine, printable line, with the numbers in octal or decimal; this in fact was the representation originally used.

      The operators fall naturally into two classes: those which represent part of an expression, and all others. Expressions are transmitted in a reverse-Polish notation; as they are being read, a tree is built which is isomorphic to the tree constructed in the first phase. Expressions are passed as a whole, with no non-expression operators intervening. The reader maintains a stack; each leaf of the expression tree (name, constant) is pushed on the stack; each unary operator replaces the top of the stack by a node whose operand is the old top-of-stack; each binary operator replaces the top pair on the stack with a single entry. When the expression is complete there is exactly one item on the stack. Following each expression is a special operator which passes the unique previous expression to the `optimizer' described below and then to the code generator.

      Here is the list of operators not themselves part of expressions.


BDATA flag data ...

WDATA flag data ...





CSPACE name size





SNAME name number

ANAME name number

RNAME name number

SAVE number

SETREG number


SWIT deflab line label value ...

LABEL number

BRANCH number


EXPR line

NAME class type name

NAME class type number

CON type value

FCON type 4-word-value

SFCON type value


CBRANCH label cond

binary-operator type

unary-operator type

Expression Optimization

      Each expression tree, as it is read in, is subjected to a fairly comprehensive analysis. This is performed by the optim routine and a number of subroutines; the major things done are

Modifications and simplifications of the tree so its value may be computed more efficiently and conveniently by the code generator.
Marking each interior node with an estimate of the number of registers required to evaluate it. This register count is needed to guide the code generation algorithm.

      One thing that is definitely not done is discovery or exploitation of common subexpressions, nor is this done anywhere in the compiler.

      The basic organization is simple: a depth-first scan of the tree. Optim does nothing for leaf nodes (except for automatics; see below), and calls unoptim to handle unary operators. For binary operators, it calls itself to process the operands, then treats each operator separately. One important case is commutative and associative operators, which are handled by acommute.

      Here is a brief catalog of the transformations carried out by by optim itself. It is not intended to be complete. Some of the transformations are machine-dependent, although they may well be useful on machines other than the PDP-11.

As indicated in the discussion of unoptim below, the optimizer can create a node type corresponding to the location addressed by a register plus a constant offset. Since this is precisely the implementation of automatic variables and arguments, where the register is fixed by convention, such variables are changed to the new form to simplify later processing.
Associative and commutative operators are processed by the special routine acommute.
After processing by acommute, the bitwise & operator is turned into a new andn operator; `a & b' becomes `a andn ~b'. This is done because the PDP-11 provides no and operator, but only andn. A similar transformation takes place for `=&'.
Relationals are turned around so the more complicated expression is on the left. (So that `2 > f(x)' becomes `f(x) < 2'). This improves code generation since the algorithm prefers to have the right operand require fewer registers than the left.
An expression minus a constant is turned into the expression plus the negative constant, and the acommute routine is called to take advantage of the properties of addition.
Operators with constant operands are evaluated.
Right shifts (unless by 1) are turned into left shifts with a negated right operand, since the PDP-11 lacks a general right-shift operator.
A number of special cases are simplified, such as division or multiplication by 1, and shifts by 0.

The unoptim routine performs the same sort of processing for unary operators.

`*&x' and `&*x' are simplified to `x'.
If r is a register and c is a constant or the address of a static or external variable, the expressions `*(r+c)' and `*r' are turned into a special kind of name node which expresses the name itself and the offset. This simplifies subsequent processing because such constructions can appear as the the address of a PDP-11 instruction.
When the unary `&' operator is applied to a name node of the special kind just discussed, it is reworked to make the addition explicit again; this is done because the PDP-11 has no `load address' instruction.
Constructions like `*r++' and `*--r' where r is a register are discovered and marked as being implementable using the PDP-11 auto-increment and -decrement modes.
If `!' is applied to a relational, the `!' is discarded and the sense of the relational is reversed.
Special cases involving reflexive use of negation and complementation are discovered.
Operations applying to constants are evaluated.

      The acommute routine, called for associative and commutative operators, discovers clusters of the same operator at the top levels of the current tree, and arranges them in a list: for `a+((b+c)+(d+f))' the list would be`a,b,c,d,e,f'. After each subtree is optimized, the list is sorted in decreasing difficulty of computation; as mentioned above, the code generation algorithm works best when left operands are the difficult ones. The `degree of difficulty' computed is actually finer than the mere number of registers required; a constant is considered simpler than the address of a static or external, which is simpler than reference to a variable. This makes it easy to fold all the constants together, and also to merge together the sum of a constant and the address of a static or external (since in such nodes there is space for an `offset' value). There are also special cases, like multiplication by 1 and addition of 0. A special routine is invoked to handle sums of products. Distrib is based on the fact that it is better to compute `c1*c2*x + c1*y' as `c1*(c2*x + y)' and makes the divisibility tests required to assure the correctness of the transformation. This transformation is rarely possible with code directly written by the user, but it invariably occurs as a result of the implementation of multi-dimensional arrays.

      Finally, acommute reconstructs a tree from the list of expressions which result.

Code Generation

      The grand plan for code-generation is independent of any particular machine; it depends largely on a set of tables. But this fact does not necessarily make it very easy to modify the compiler to produce code for other machines, both because there is a good deal of machine-dependent structure in the tables, and because in any event such tables are non-trivial to prepare.

      The arguments to the basic code generation routine rcexpr are a pointer to a tree representing an expression, the name of a code-generation table, and the number of a register in which the value of the expression should be placed. Rcexpr returns the number of the register in which the value actually ended up; its caller may need to produce a mov instruction if the value really needs to be in the given register. There are four code generation tables.

      Regtab is the basic one, which actually does the job described above: namely, compile code which places the value represented by the expression tree in a register.

      Cctab is used when the value of the expression is not actually needed, but instead the value of the condition codes resulting from evaluation of the expression. This table is used, for example, to evaluate the expression after if. It is clearly silly to calculate the value (0 or 1) of the expression `a==b' in the context `if (a==b) ... '

      The sptab table is used when the value of an expression is to be pushed on the stack, for example when it is an actual argument. For example in the function call `f(a)' it is a bad idea to load a into a register which is then pushed on the stack, when there is a single instruction which does the job.

      The efftab table is used when an expression is to be evaluated for its side effects, not its value. This occurs mostly for expressions which are statements, which have no value. Thus the code for the statement `a = b' need produce only the approoriate mov instruction, and need not leave the value of b in a register, while in the expression `a + (b = c)' the value of `b = c' will appear in a register.

      All of the tables besides regtab are rather small, and handle only a relatively few special cases. If one of these subsidiary tables does not contain an entry applicable to the given expression tree, rcexpr uses regtab to put the value of the expression into a register and then fixes things up; nothing need be done when the table was efftab, but a tst instruction is produced when the table called for was cctab, and a mov instruction, pushing the register on the stack, when the table was sptab.

      The rcexpr routine itself picks off some special cases, then calls cexpr to do the real work. Cexpr tries to find an entry applicable to the given tree in the given table, and returns -1 if no such entry is found, letting rcexpr try again with a different table. A successful match yields a string containing both literal characters which are written out and pseudo-operations, or macros, which are expanded. Before studying the contents of these strings we will consider how table entries are matched against trees.

      Recall that most non-leaf nodes in an expression tree contain the name of the operator, the type of the value represented, and pointers to the subtrees (operands). They also contain an estimate of the number of registers required to evaluate the expression, placed there by the expression-optimizer routines. The register counts are used to guide the code generation process, which is based on the Sethi-Ullman algorithm.

      The main code generation tables consist of entries each containing an operator number and a pointer to a subtable for the corresponding operator. A subtable consists of a sequence of entries, each with a key describing certain properties of the operands of the operator involved; associated with the key is a code string. Once the subtable corresponding to the operator is found, the subtable is searched linearly until a key is found such that the properties demanded by the key are compatible with the operands of the tree node. A successful match returns the code string; an unsuccessful search, either for the operator in the main table or a compatble key in the subtable, returns a failure indication.

      The tables are all contained in a file which must be processed to obtain an assembly language program. Thus they are written in a special-purpose language. To provided definiteness to the following discussion, here is an example of a subtable entry.

add A2,R
The `%' indicates the key; the information following (up to a blank line) specifies the code string. Very briefly, this entry is in the subtable for `+' of regtab; the key specifies that the left operand is any integer, character, or pointer expression, and the right operand is any word quantity which is directly addressible (e.g. a variable or constant). The code string calls for the generation of the code to compile the left (first) operand into the current register (`F') and then to produce an `add' instruction which adds the second operand (`A2') to the register (`R'). All of the notation will be explained below.

      Only three features of the operands are used in deciding whether a match has occurred. They are:

Is the type of the operand compatible with that demanded?
Is the `degree of difficulty' (in a sense described below) compatible?
The table may demand that the operand have a `*' (indirection operator) as its highest operator.

      As suggested above, the key for a subtable entry is indicated by a `%,' and a comma-separated pair of specifications for the operands. (The second specification is ignored for unary operators). A specification indicates a type requirement by including one of the following letters. If no type letter is present, any integer, character, or pointer operand will satisfy the requirement (not float, double, or long).

A byte (character) operand is required.
A word (integer or pointer) operand is required.
A float or double operand is required.
A double operand is required.
A long (32-bit integer) operand is required.

      Before discussing the `degree of difficulty' specification, the algorithm has to be explained more completely. Rcexpr (and cexpr) are called with a register number in which to place their result. Registers 0, 1, ... are used during evaluation of expressions; the maximum register which can be used in this way depends on the number of register variables, but in any event only registers 0 through 4 are available since r5 is used as a stack frame header and r6 (sp) and r7 (pc) have special hardware properties. The code generation routines assume that when called with register n as argument, they may use n+1, ... (up to the first register variable) as temporaries. Consider the expression `X+Y', where both X and Y are expressions. As a first approximation, there are three ways of compiling code to put this expression in register n.

If Y is an addressible cell, (recursively) put X into register n and add Y to it.
If Y is an expression that can be calculated in k registers, where k smaller than the number of registers available, compile X into register n, Y into register n+1, and add register n+1 to n.
Otherwise, compile Y into register n, save the result in a temporary (actually, on the stack) compile X into register n, then add in the temporary.

      The distinction between cases 2 and 3 therefore depends on whether the right operand can be compiled in fewer than k registers, where k is the number of free registers left after registers 0 through n are taken: 0 through n-1 are presumed to contain already computed temporary results; n will, in case 2, contain the value of the left operand while the right is being evaluated.

      These considerations should make clear the specification codes for the degree of difficulty, bearing in mind that a number of special cases are also present:

is satisfied when the operand is zero, so that special code can be produced for expressions like `x = 0'.
is satisfied when the operand is the constant 1, to optimize cases like left and right shift by 1, which can be done efficiently on the PDP-11.
is satisfied when the operand is a positive (16-bit) constant; this takes care of some special cases in long arithmetic.
is satisfied when the operand is addressible; this occurs not only for variables and constants, but also for some more complicated constructions, such as indirection through a simple variable, `*p++' where p is a register variable (because of the PDP-11's auto-increment address mode), and `*(p+c)' where p is a register and c is a constant. Precisely, the requirement is that the operand refers to a cell whose address can be written as a source or destination of a PDP-11 instruction.
is satisfied by an operand whose value can be generated in a register using no more than k registers, where k is the number of registers left (not counting the current register). The `e' stands for `easy.'
is satisfied by any operand. The `n' stands for `anything.'

      These degrees of difficulty are considered to lie in a linear ordering and any operand which satisfies an earlier-mentioned requirement will satisfy a later one. Since the subtables are searched linearly, if a `1' specification is included, almost certainly a `z' must be written first to prevent expressions containing the constant 0 to be compiled as if the 0 were 1.

      Finally, a key specification may contain a `*' which requires the operand to have an indirection as its leading operator. Examples below should clarify the utility of this specification.

      Now let us consider the contents of the code string associated with each subtable entry. Conventionally, lower-case letters in this string represent literal information which is copied directly to the output. Upper-case letters generally introduce specific macro-operations, some of which may be followed by modifying information. The code strings in the tables are written with tabs and new-lines used freely to suggest instructions which will be generated; the table-compiling program compresses tabs (using the 0200 bit of the next character) and throws away some of the new-lines. For example the macro `F' is ordinarily written on a line by itself; but since its expansion will end with a new-line, the new-line after `F' itself is dispensable. This is all to reduce the size of the stored tables.

      The first set of macro-operations is concerned with compiling subtrees. Recall that this is done by the cexpr routine. In the following discussion the `current register' is generally the argument register to cexpr; that is, the place where the result is desired. The `next register' is numbered one higher than the current register. (This explanation isn't fully true because of complications, described below, involving operations which require even-odd register pairs.)

causes a recursive call to the rcexpr routine to compile code which places the value of the first (left) operand of the operator in the current register.
generates code which places the value of the first operand in the next register. It is incorrectly used if there might be no next register; that is, if the degree of difficulty of the first operand is not `easy;' if not, another register might not be available.
generates code which pushes the value of the first operand on the stack, by calling rcexpr specifying sptab as the table.


S, S1, SS
compile the second (right) operand into the current register, the next register, or onto the stack.

To deal with registers, there are

which expands into the name of the current register.
which expands into the name of the next register.
which expands into the the name of the current register plus 1. It was suggested above that this is the same as the next register, except for complications; here is one of them. Long integer variables have 32 bits and require 2 registers; in such cases the next register is the current register plus 2. The code would like to talk about both halves of the long quantity, so R refers to the register with the high-order part and R+ to the low-order part.
This is another complication, involving division and mod. These operations involve a pair of registers of which the odd-numbered contains the left operand. Cexpr arranges that the current register is odd; the R- notation allows the code to refer to the next lower, even-numbered register.

To refer to addressible quantities, there are the notations:

causes generation of the address specified by the first operand. For this to be legal, the operand must be addressible; its key must contain an `a' or a more restrictive specification.
correspondingly generates the address of the second operand providing it has one.

      We now have enough mechanism to show a complete, if suboptimal, table for the + operator on word or byte operands.


inc R

add A2,R

add R1,R

add (sp)+,R
The first two sequences handle some special cases. Actually it turns out that handling a right operand of 0 is unnecessary since the expression-optimizer throws out adds of 0. Adding 1 by using the `increment' instruction is done next, and then the case where the right operand is addressible. It must be a word quantity, since the PDP-11 lacks an `add byte' instruction. Finally the cases where the right operand either can, or cannot, be done in the available registers are treated.

      The next macro-instructions are conveniently introduced by noticing that the above table is suitable for subtraction as well as addition, since no use is made of the commutativity of addition. All that is needed is substitution of `sub' for `add' and `dec' for 'inc.' Considerable saving of space is achieved by factoring out several similar operations.

is replaced by a string from another table indexed by the operator in the node being expanded. This secondary table actually contains two strings per operator.
is replaced by the second string in the side table entry for the current operator.

      Thus, given that the entries for `+' and `-' in the side table (which is called instab) are `add' and `inc,' `sub' and `dec' respectively, the middle of of the above addition table can be written

I' R

I A2,R
and it will be suitable for subtraction, and several other operators, as well.

      Next, there is the question of character and floating-point operations.

generates the letter `b' if the first operand is a character, `f' if it is float or double, and nothing otherwise. It is used in a context like `movB1' which generates a `mov', `movb', or `movf' instruction according to the type of the operand.
is just like B1 but applies to the second operand.
generates `b' if either operand is a character and null otherwise.
generates `f' if the type of the operator node itself is float or double, otherwise null.

      For example, there is an entry in efftab for the `=' operator

Note first that two key specifications can be applied to the same code string. Next, observe that when a word is assigned to a byte or to a word, or a word is assigned to a byte, a single instruction, a mov or movb as appropriate, does the job. However, when a byte is assigned to a word, it must pass through a register to implement the sign-extension rules:
IB1 R,A1

      Next, there is the question of handling indirection properly. Consider the expression `X + *Y', where X and Y are expressions, Assuming that Y is more complicated than just a variable, but on the other hand qualifies as `easy' in the context, the expression would be compiled by placing the value of X in a register, that of *Y in the next register, and adding the registers. It is easy to see that a better job can be done by compiling X, then Y (into the next register), and producing the instruction symbolized by `add (R1),R'. This scheme avoids generating the instruction `mov (R1),R1' required actually to place the value of *Y in a register. A related situation occurs with the expression `X + *(p+6)', which exemplifies a construction frequent in structure and array references. The addition table shown above would produce

[put X in register R]
mov p,R1
add $6,R1
mov (R1),R1
add R1,R
when the best code is
[put X in R]
mov p,R1
add 6(R1),R
As we said above, a key specification for a code table entry may require an operand to have an indirection as its highest operator. To make use of the requirement, the following macros are provided.
the first operand must have the form *X. If in particular it has the form *(Y + c), for some constant c, then code is produced which places the value of Y in the current register. Otherwise, code is produced which loads X into the current register.
resembles F* except that the next register is loaded.
resembles F* except that the second operand is loaded.
resembles S* except that the next register is loaded.
The first operand must have the form `*X'. Push the value of X on the stack.
resembles FS* except that it applies to the second operand.

To capture the constant that may have been skipped over in the above macros, there are

The first operand must have the form *X; if in particular it has the form *(Y + c) for c a constant, then the constant is written out, otherwise a null string.
is the same as #1 except that the second operand is used.

Now we can improve the addition table above. Just before the `%n,e' entry, put

add #2(R1),R
and just before the `%n,n' put
add *(sp)+,R
When using the stacking macros there is no place to use the constant as an index word, so that particular special case doesn't occur.

      The constant mentioned above can actually be more general than a number. Any quantity acceptable to the assembler as an expression will do, in particular the address of a static cell, perhaps with a numeric offset. If x is an external character array, the expression `x[i+5] = 0' will generate the code

mov i,r0
clrb x+5(r0)
via the table entry (in the `=' part of efftab)
I'B1 #1(R)
Some machine operations place restrictions on the registers used. The divide instruction, used to implement the divide and mod operations, requires the dividend to be placed in the odd member of an even-odd pair; other peculiarities of multiplication make it simplest to put the multiplicand in an odd-numbered register. There is no theory which optimally accounts for this kind of requirement. Cexpr handles it by checking for a multiply, divide, or mod operation; in these cases, its argument register number is incremented by one or two so that it is odd, and if the operation was divide or mod, so that it is a member of a free even-odd pair. The routine which determines the number of registers required estimates, conservatively, that at least two registers are required for a multiplication and three for the other peculiar operators. After the expression is compiled, the register where the result actually ended up is returned. (Divide and mod are actually the same operation except for the location of the result).

      These operations are the ones which cause results to end up in unexpected places, and this possibility adds a further level of complexity. The simplest way of handling the problem is always to move the result to the place where the caller expected it, but this will produce unnecessary register moves in many simple cases; `a = b*c' would generate

mov b,r1
mul c,r1
mov r1,r0
mov r0,a
The next thought is used the passed-back information as to where the result landed to change the notion of the current register. While compiling the `=' operation above, which comes from a table entry like
mov R,A1
it is sufficient to redefine the meaning of `R' after processing the `S' which does the multiply. This technique is in fact used; the tables are written in such a way that correct code is produced. The trouble is that the technique cannot be used in general, because it invalidates the count of the number of registers required for an expression. Consider just `a*b + X' where X is some expression. The algorithm assumes that the value of a*b, once computed, requires just one register. If there are three registers available, and X requires two registers to compute, then this expression will match a key specifying `%n,e'. If a*b is computed and left in register 1, then there are, contrary to expectations, no longer two registers available to compute X, but only one, and bad code will be produced. To guard against this possibility, cexpr checks the result returned by recursive calls which implement F, S and their relatives. If the result is not in the expected register, then the number of registers required by the other operand is checked; if it can be done using those registers which remain even after making unavailable the unexpectedly-occupied register, then the notions of the `next register' and possibly the `current register' are redefined. Otherwise a register-copy instruction is produced. A register-copy is also always produced when the current operator is one of those which have odd-even requirements.

      Finally, there are a few loose-end macro operations and facts about the tables. The operators:

is used for long operations. It is written with an address like a machine instruction; it expands into `adc' (add carry) if the operation is an additive operator, `sbc' (subtract carry) if the operation is a subtractive operator, and disappears, along with the rest of the line, otherwise. Its purpose is to allow common treatment of logical operations, which have no carries, and additive and subtractive operations, which generate carries.
generates a `tst' instruction if the first operand of the tree does not set the condition codes correctly. It is used with divide and mod operations, which require a sign-extended 32-bit operand. The code table for the operations contains an `sxt' (sign-extend) instruction to generate the high-order part of the dividend.
is analogous to the `F' and `S' macros, except that it calls for the generation of code for the current tree (not one of its operands) using regtab. It is used in cctab for all the operators which, when executed normally, set the condition codes properly according to the result. It prevents a `tst' instruction from being generated for constructions like `if (a+b) ...' since after calculation of the value of `a+b' a conditional branch can be written immediately.

      All of the discussion above is in terms of operators with operands. Leaves of the expression tree (variables and constants), however, are peculiar in that they have no operands. In order to regularize the matching process, cexpr examines its operand to determine if it is a leaf; if so, it creates a special `load' operator whose operand is the leaf, and substitutes it for the argument tree; this allows the table entry for the created operator to use the `A1' notation to load the leaf into a register.

      Purely to save space in the tables, pieces of subtables can be labelled and referred to later. It turns out, for example, that rather large portions of the the efftab table for the `=' and `=+' operators are identical. Thus `=' has an entry

while part of the `=+' table is
% [move3]
Labels are written as `%[ ... : ]', before the key specifications; references are written with `% [ ... ]' after the key. Peculiarities in the implementation make it necessary that labels appear before references to them.

      The example illustrates the utility of allowing separate keys to point to the same code string. The assignment code works properly if either the right operand is a word, or the left operand is a byte; but since there is no `add byte' instruction the addition code has to be restricted to word operands.

Delaying and reordering

      Intertwined with the code generation routines are two other, interrelated processes. The first, implemented by a routine called delay, is based on the observation that naive code generation for the expression `a = b++' would produce

mov b,r0
inc b
mov r0,a
The point is that the table for postfix ++ has to preserve the value of b before incrementing it; the general way to do this is to preserve its value in a register. A cleverer scheme would generate
mov b,a
inc b
Delay is called for each expression input to rcexpr, and it searches for postfix ++ and -- operators. If one is found applied to a variable, the tree is patched to bypass the operator and compiled as it stands; then the increment or decrement itself is done. The effect is as if `a = b; b++' had been written. In this example, of course, the user himself could have done the same job, but more complicated examples are easily constructed, for example `switch (x++)'. An essential restriction is that the condition codes not be required. It would be incorrect to compile `if (a++) ...' as
tst a
inc a
beq ...
because the `inc' destroys the required setting of the condition codes.

      Reordering is a similar sort of optimization. Many cases which it detects are useful mainly with register variables. If r is a register variable, the expression `r = x+y' is best compiled as

mov x,r
add y,r
but the codes tables would produce
mov x,r0
add y,r0
mov r0,r
which is in fact preferred if r is not a register. (If r is not a register, the two sequences are the same size, but the second is slightly faster.) The scheme is to compile the expression as if it had been written `r = x; r =+ y'. The reorder routine is called with a pointer to each tree that rcexpr is about to compile; if it has the right characteristics, the `r = x' tree is constructed and passed recursively to rcexpr; then the original tree is modified to read `r =+ y' and the calling instance of rcexpr compiles that instead. Of course the whole business is itself recursive so that more extended forms of the same phenomenon are handled, like `r = x + y | z'.

      Care does have to be taken to avoid `optimizing' an expression like `r = x + r' into `r = x; r =+ r'. It is required that the right operand of the expression on the right of the `=' be a ', distinct from the register variable.

      The second case that reorder handles is expressions of the form `r = X' used as a subexpression. Again, the code out of the tables for `x = r = y' would be

mov y,r0
mov r0,r
mov r0,x
whereas if r were a register it would be better to produce
mov y,r
mov r,x
When reorder discovers that a register variable is being assigned to in a subexpression, it calls rcexpr recursively to compile the subexpression, then fiddles the tree passed to it so that the register variable itself appears as the operand instead of the whole subexpression. Here care has to be taken to avoid an infinite regress, with rcexpr and reorder calling each other forever to handle assignments to registers.

      A third set of cases treated by reorder comes up when any name, not necessarily a register, occurs as a left operand of an assignment operator other than `=' or as an operand of prefix `++' or `--'. Unless condition-code tests are involved, when a subexpression like `(a =+ b)' is seen, the assignment is performed and the argument tree modified so that a is its operand; effectively `x + (y =+ z)' is compiled as `y =+ z; x + y'. Similarly, prefix increment and decrement are pulled out and performed first, then the remainder of the expression.

      Throughout code generation, the expression optimizer is called whenever delay or reorder change the expression tree. This allows some special cases to be found that otherwise would not be seen.

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