\part{Compiler Retargeting}

[\#\#\#

In general, it is a danger sign if a generator references a TN that isn't an
operand or temporary, since lifetime analysis hasn't been done for that use.
We are doing weird stuff for the old-cont and return-pc passing locations,
hoping that the conflicts at the called function have the desired effect.
Other stuff?  When a function returns unknown values, we don't reference the
values locations when a single-value return is done.  But nothing is live at a
return point anyway.



Have a way for template conversion to special-case constant arguments?  
How about:
    If an arg restriction is (:satisfies [$<$predicate function$>$]), and the
    corresponding argument is constant, with the constant value satisfying the
    predicate, then (if any other restrictions are satisfied), the template
    will be emitted with the literal value passed as an info argument.  If the
    predicate is omitted, then any constant will do.

    We could sugar this up a bit by allowing (:member $<$object$>$*) for
    (:satisfies (lambda (x) (member x '($<$object$>$*))))

We could allow this to be translated into a Lisp type by adding a new Constant
type specifier.  This could only appear as an argument to a function type.
To satisfy (Constant $<$type$>$), the argument must be a compile-time constant of
the specified type.  Just Constant means any constant (i.e. (Constant *)).
This would be useful for the type constraints on ICR transforms.


Constant TNs: we count on being able to indirect to the leaf, and don't try to
wedge the information into the offset.  We set the FSC to an appropriate
immediate SC.

    Allow "more operands" to VOPs in define-vop.  You can't do much with the
    more operands: define-vop just fills in the cost information according to
    the loading costs for a SC you specify.  You can't restrict more operands,
    and you can't make local preferences.  In the generator, the named variable
    is bound to the TN-ref for the first extra operand.  This should be good
    enough to handle all the variable arg VOPs (primarily function call and
    return).  Usually more operands are used just to get TN lifetimes to work
    out; the generator actually ignores them.

    Variable-arg VOPs can't be used with the VOP macro.  You must use VOP*.
    VOP* doesn't do anything with these extra operand except stick them on the
    ends of the operand lists passed into the template.  VOP* is often useful
    within the convert functions for non-VOP templates, since it can emit a VOP
    using an already prepared TN-Ref list.
    

    It is pretty basic to the whole primitive-type idea that there is only one
    primitive-type for a given lisp type.  This is really the same as saying
    primitive types are disjoint.  A primitive type serves two somewhat
    unrelated purposes:
     -- It is an abstraction of a Lisp type used to select type specific
        operations.  Originally kind of an efficiency hack, but it lets a
        template's type signature be used both for selection and operand
        representation determination.
     -- It represents a set of possible representations for a value (SCs).  The
        primitive type is used to determine the legal SCs for a TN, and is also
        used to determine which type-coercion/move VOP to use.

]

There are basically three levels of target dependence:

 -- Code in the "front end" (before VMR conversion) deals only with Lisp
    semantics, and is totally target independent.

 -- Code after VMR conversion and before code generation depends on the VM,
    but should work with little modification across a wide range of
    "conventional" architectures.

 -- Code generation depends on the machine's instruction set and other
    implementation details, so it will have to be redone for each
    implementation.  Most of the work here is in defining the translation into
    assembly code of all the supported VOPs.



\chapter{Storage bases and classes}
New interface: instead of CURRENT-FRAME-SIZE, have CURRENT-SB-SIZE \verb+<name>+ which
returns the current element size of the named SB.

How can we have primitive types that overlap, i.e. (UNSIGNED-BYTE 32),
(SIGNED-BYTE 32), FIXNUM?
Primitive types are used for two things:
    Representation selection: which SCs can be used to represent this value?
	For this purpose, it isn't necessary that primitive types be disjoint,
	since any primitive type can choose an arbitrary set of
	representations.  For moves between the overlapping representations,
	the move/load operations can just be noops when the locations are the
	same (vanilla MOVE), since any bad moves should be caught out by type
	checking.
    VOP selection:
	Is this operand legal for this VOP?  When ptypes overlap in interesting
	ways, there is a problem with allowing just a simple ptype restriction,
	since we might want to allow multiple ptypes.  This could be handled
	by allowing "union primitive types", or by allowing multiple primitive
	types to be specified (only in the operand restriction.)  The latter
	would be along the lines of other more flexible VOP operand restriction
	mechanisms, (constant, etc.)



Ensure that load/save-operand never need to do representation conversion.

The PRIMITIVE-TYPE more/coerce info would be moved into the SC.  This could
perhaps go along with flushing the TN-COSTS.  We would annotate the TN with
best SC, which implies the representation (boxed or unboxed).  We would still
need to represent the legal SCs for restricted TNs somehow, and also would have to
come up with some other way for pack to keep track of which SCs we have already
tried.

An SC would have a list of "alternate" SCs and a boolean SAVE-P value that
indicates it needs to be saved across calls in some non-SAVE-P SC.  A TN is
initially given its "best" SC.  The SC is annotated with VOPs that are used for
moving between the SC and its alternate SCs (load/save operand, save/restore
register).  It is also annotated with the "move" VOPs used for moving between
this SC and all other SCs it is possible to move between.  We flush the idea
that there is only c-to-t and c-from-t.

But how does this mesh with the idea of putting operand load/save back into the
generator?  Maybe we should instead specify a load/save function?  The
load/save functions would also differ from the move VOPs in that they would
only be called when the TN is in fact in that particular alternate SC, whereas
the move VOPs will be associated with the primary SC, and will be emitted
before it is known whether the TN will be packed in the primary SC or an
alternate.

I guess a packed SC could also have immediate SCs as alternate SCs, and
constant loading functions could be associated with SCs using this mechanism.

So given a TN packed in SC X and an SC restriction for Y and Z, how do we know
which load function to call?  There would be ambiguity if X was an alternate
for both Y and Z and they specified different load functions.  This seems
unlikely to arise in practice, though, so we could just detect the ambiguity
and give an error at define-vop time.  If they are doing something totally
weird, they can always inhibit loading and roll their own.

Note that loading costs can be specified at the same time (same syntax) as
association of loading functions with SCs.  It seems that maybe we will be
rolling DEFINE-SAVE-SCS and DEFINE-MOVE-COSTS into DEFINE-STORAGE-CLASS.

Fortunately, these changes will affect most VOP definitions very little.


A Storage Base represents a physical storage resource such as a register set or
stack frame.  Storage bases for non-global resources such as the stack are
relativized by the environment that the TN is allocated in.  Packing conflict
information is kept in the storage base, but non-packed storage resources such
as closure environments also have storage bases.
Some storage bases:
\begin{verbatim}
    General purpose registers
    Floating point registers
    Boxed (control) stack environment
    Unboxed (number) stack environment
    Closure environment
\end{verbatim}

A storage class is a potentially arbitrary set of the elements in a storage
base.  Although conceptually there may be a hierarchy of storage classes such
as "all registers", "boxed registers", "boxed scratch registers", this doesn't
exist at the implementation level.  Such things can be done by specifying
storage classes whose locations overlap.  A TN shouldn't have lots of
overlapping SC's as legal SC's, since time would be wasted repeatedly
attempting to pack in the same locations.

There will be some SC's whose locations overlap a great deal, since we get Pack
to do our representation analysis by having lots of SC's.  An SC is basically a
way of looking at a storage resource.  Although we could keep a fixnum and an
unboxed representation of the same number in the same register, they correspond
to different SC's since they are different representation choices.

TNs are annotated with the primitive type of the object that they hold:
    T: random boxed object with only one representation.
    Fixnum, Integer, XXX-Float: Object is always of the specified numeric type.
    String-Char: Object is always a string-char.

When a TN is packed, it is annotated with the SC it was packed into.  The code
generator for a VOP must be able to uniquely determine the representation of
its operands from the SC. (debugger also...)

Some SCs:
    Reg: any register (immediate objects)
    Save-Reg: a boxed register near r15 (registers easily saved in a call)
    Boxed-Reg: any boxed register (any boxed object)
    Unboxed-Reg: any unboxed register (any unboxed object)
    Float-Reg, Double-Float-Reg: float in FP register.
    Stack: boxed object on the stack (on cstack)
    Word: any 32bit unboxed object on nstack.
    Double: any 64bit unboxed object on nstack.

We have a number of non-packed storage classes which serve to represent access
costs associated with values that are not allocated using conflicts
information.  Non-packed TNs appear to already be packed in the appropriate
storage base so that Pack doesn't get confused.  Costs for relevant non-packed
SC's appear in the TN-Ref cost information, but need not ever be summed into
the TN cost vectors, since TNs cannot be packed into them.

There are SCs for non-immediate constants and for each significant kind of
immediate operand in the architecture.  On the RT, 4, 8 and 20 bit integer SCs
are probably worth having.

\begin{verbatim}
Non-packed SCs:
    Constant
    Immediate constant SCs:
        Signed-Byte-<N>, Unsigned-Byte-<N>, for various architecture dependent
	    values of <N>
	String-Char
	XXX-Float
	Magic values: T, NIL, 0.
\end{verbatim}

\chapter{Type system parameterization}

The main aspect of the VM that is likely to vary for good reason is the type
system:

 -- Different systems will have different ways of representing dynamic type
    information.  The primary effect this has on the compiler is causing VMR
    conversion of type tests and checks to be implementation dependent.
    Rewriting this code for each implementation shouldn't be a big problem,
    since the portable semantics of types has already been dealt with.

 -- Different systems will have different specialized number and array types,
    and different VOPs specialized for these types.  It is easy to add this kind
    of knowledge without affecting the rest of the compiler.  All you have to
    do is define the VOPs and translations.

 -- Different systems will offer different specialized storage resources
    such as floating-point registers, and will have additional kinds of
    primitive-types.  The storage class mechanism handles a large part of this,
    but there may be some problem in getting VMR conversion to realize the
    possibly large hidden costs in implicit moves to and from these specialized
    storage resources.  Probably the answer is to have some sort of general
    mechanism for determining the primitive-type for a TN given the Lisp type,
    and then to have some sort of mechanism for automatically using specialized
    Move VOPs when the source or destination has some particular primitive-type.

\#|
How to deal with list/null(symbol)/cons in primitive-type structure?  Since
cons and symbol aren't used for type-specific template selection, it isn't
really all that critical.  Probably Primitive-Type should return the List
primitive type for all of Cons, List and Null (indicating when it is exact).
This would allow type-dispatch for simple sequence functions (such as length)
to be done using the standard template-selection mechanism.  [Not a wired
assumption] 
|\#



\chapter{VOP Definition}

Before the operand TN-refs are passed to the emit function, the following
stuff is done:
 -- The refs in the operand and result lists are linked together in order using
    the Across slot.  This list is properly NIL terminated.
 -- The TN slot in each ref is set, and the ref is linked into that TN's refs
    using the Next slot.
 -- The Write-P slot is set depending on whether the ref is an argument or
    result.
 -- The other slots have the default values.

The template emit function fills in the Vop, Costs, Cost-Function,
SC-Restriction and Preference slots, and links together the Next-Ref chain as
appropriate.


\section{Lifetime model}

\#|
Note in doc that the same TN may not be used as both a more operand and as any
other operand to the same VOP, to simplify more operand LTN number coalescing.
|\#

It seems we need a fairly elaborate model for intra-VOP conflicts in order to
allocate temporaries without introducing spurious conflicts.  Consider the
important case of a VOP such as a miscop that must have operands in certain
registers.  We allocate a wired temporary, create a local preference for the
corresponding operand, and move to (or from) the temporary.  If all temporaries
conflict with all arguments, the result will be correct, but arguments could
never be packed in the actual passing register.  If temporaries didn't conflict
with any arguments, then the temporary for an earlier argument might get packed
in the same location as the operand for a later argument; loading would then
destroy an argument before it was read.

A temporary's intra-VOP lifetime is represented by the times at which its life
starts and ends.  There are various instants during the evaluation that start
and end VOP lifetimes.  Two TNs conflict if the live intervals overlap.
Lifetimes are open intervals: if one TN's lifetime begins at a point where
another's ends, then the TNs don't conflict.

The times within a VOP are the following:

:Load
    This is the beginning of the argument's lives, as far as intra-vop
    conflicts are concerned.  If load-TNs are allocated, then this is the
    beginning of their lives.

(:Argument $<$n$>$)
    The point at which the N'th argument is read for the last time (by this
    VOP).  If the argument is dead after this VOP, then the argument becomes
    dead at this time, and may be reused as a temporary or result load-TN.

(:Eval $<$n$>$)
    The N'th evaluation step.  There may be any number of evaluation steps, but
    it is unlikely that more than two are needed.

(:Result $<$n$>$) 
    The point at which the N'th result is first written into.  This is the
    point at which that result becomes live.

:Save
    Similar to :Load, but marks the end of time.  This is the point at which result
    load-TNs are stored back to the actual location.

In any of the list-style time specifications, the keyword by itself stands for
the first such time, i.e.
\begin{verbatim}
    :argument  <==>  (:argument 0)
\end{verbatim}

Note that argument/result read/write times don't actually have to be in the
order specified, but they must *appear* to happen in that order as far as
conflict analysis is concerned.  For example, the arguments can be read in any
order as long as no TN is written that has a life beginning at or after
(:Argument $<$n$>$), where N is the number of an argument whose reading was
postponed.

[\#\#\# (???)

We probably also want some syntactic sugar in Define-VOP for automatically
moving operands to/from explicitly allocated temporaries so that this kind of
thing is somewhat easy.  There isn't really any reason to consider the
temporary to be a load-TN, but we want to compute costs as though it was and
want to use the same operand loading routines.

We also might consider allowing the lifetime of an argument/result to be
extended forward/backward.  This would in many cases eliminate the need for
temporaries when operands are read/written out of order.
]


\section{VOP Cost model}

Note that in this model, if an operand has no restrictions, it has no cost.
This makes sense, since the purpose of the cost is to indicate the
relative value of packing in different SCs.  If the operand isn't required to
be in a good SC (i.e. a register), then we might as well leave it in memory.
The SC restriction mechanism can be used even when doing a move into the SC is
too complex to be generated automatically (perhaps requiring temporary
registers), since Define-VOP allows operand loading to be done explicitly.


\section{Efficiency notes}

  In addition to
being used to tell whether a particular unsafe template might get emitted, we
can also use it to give better efficiency notes:
 -- We can say what is wrong with the call types, rather than just saying we
    failed to open-code.
 -- We can tell whether any of the "better" templates could possibly apply,
    i.e. is the inapplicability of a template because of inadequate type
    information or because the type is just plain wrong.  We don't want to
    flame people when a template that couldn't possibly match doesn't match,
    e.g. complaining that we can't use fixnum+ when the arguments are known to
    be floats.


This is how we give better efficiency notes:

The Template-Note is a short noun-like string without capitalization or
punctuation that describes what the template ``does", i.e. we say
"Unable to do ~A, doing ~A instead."

The Cost is moved from the Vop-Info to the Template structure, and is used to
determine the ``goodness" of possibly applicable templates.  [Could flush
Template/Vop-Info distinction]  The cost is used to choose the best applicable
template to emit, and also to determine what better templates we might have
been able to use.

A template is possibly applicable if there is an intersection between all of
the arg/result types and the corresponding arg/result restrictions, i.e. the
template is not clearly impossible: more declarations might allow it to be
emitted.


\chapter{Assembler Retargeting}


\chapter{Writing Assembly Code}

VOP writers expect:
\begin{Lentry}
\item[MOVE]
      You write when you port the assembler.)
\item[EMIT-LABEL]
      Assembler interface like INST.  Takes a label you made and says "stick it
      here."
   \item[GEN-LABEL]
      Returns a new label suitable for use with EMIT-LABEL exactly once and
      for referencing as often as necessary.
   \item[INST]
      Recognizes and dispatches to instructions you defined for assembler.
   \item[ALIGN]
      This takes the number of zero bits you want in the low end of the address
      of the next instruction.
   \item[ASSEMBLE]
   \item[ASSEMBLE-ELSEWHERE]
      Get ready for assembling stuff.  Takes a VOP and arbitrary PROGN-style
      body.  Wrap these around instruction emission code announcing the first
      pass of our assembler.
   \item[CURRENT-NFP-TN]
      This returns a TN for the NFP if the caller uses the number stack, or
      nil.
   \item[SB-ALLOCATED-SIZE]
      This returns the size of some storage base used by the currently
      compiling component.
   \item[...]
\end{Lentry}
;;;
;;; VOP idioms
;;;

\begin{Lentry}
\item[STORE-STACK-TN]
\item[LOAD-STACK-TN]
   These move a value from a register to the control stack, or from the
   control stack to a register.  They take care of checking the TN types,
   modifying offsets according to the address units per word, etc.
\end{Lentry}

\chapter{Required VOPS}


Note: the move VOP cannot have any wired temps.  (Move-Argument also?)  This is
so we can move stuff into wired TNs without stepping on our toes.


We create set closure variables using the Value-Cell VOP, which takes a value
and returns a value cell containing the value.  We can basically use this
instead of a Move VOP when initializing the variable.  Value-Cell-Set and
Value-Cell-Ref are used to access the value cell.  We can have a special effect
for value cells so that value cells references can be discovered to be common
subexpressions or loop invariants.




Represent unknown-values continuations as (start, count).  Unknown values
continuations are always outside of the current frame (on stack top).  Within a
function, we always set up and receive values in the standard passing
locations.  If we receive stack values, then we must BLT them down to the start
of our frame, filling in any unsupplied values.  If we generate unknown values
(i.e. PUSH-VALUES), then we set the values up in the standard locations, then
BLT them to stack top.  When doing a tail-return of MVs, we just set them up in
the standard locations and decrement SP: no BLT is necessary.

Unknown argument call (MV-CALL) takes its arguments on stack top (is given a
base pointer).  If not a tail call, then we just set the arg pointer to the
base pointer and call.  If a tail call, we must BLT the arguments down to the
beginning of the current frame.


Implement more args by BLT'ing the more args *on top* of the current frame.
This solves two problems:
\begin{itemize}
\item Any register more arguments can be made uniformly accessibly by copying
    them into memory.  [We can't store the registers in place, since the
    beginning of the frame gets double use for storing the old-cont, return-pc
    and env.]
\item It solves the deallocation problem: the arguments will be deallocated when
    the frame is returned from or a tail full call is done out of it.  So
    keyword args will be properly tail-recursive without any special mechanism
    for squeezing out the more arg once the parsing is done.  Note that a tail
    local call won't blast the more arg, since in local call the callee just
    takes the frame it is given (in this case containing the more arg).
\end{itemize}

More args in local call???  Perhaps we should not attempt local call conversion
in this case.  We already special-case keyword args in local call.  It seems
that the main importance of more args is primarily related to full call: it is
used for defining various kinds of frobs that need to take arbitrary arguments:
\begin{itemize}
\item Keyword arguments
\item Interpreter stubs
\item "Pass through" applications such as dispatch functions
\end{itemize}
Given the marginal importance of more args in local call, it seems unworth
going to any implementation difficulty.  In fact, it seems that it would cause
complications both at the VMR level and also in the VM definition.  This being
the case, we should flush it.


\section{Function Call}



\subsection{Registers and frame format}

These registers are used in function call and return:

A0..A{\it n}
    In full call, the first three arguments.  In unknown values return, the
    first three return values.

CFP
    The current frame pointer.  In full call, this initially points to a
    partial frame large enough to hold the passed stack arguments (zero-length
    if none).

CSP
    The current control stack top pointer. 

OCFP
    In full call, the passing location for the frame to return to.

    In unknown-values return of other than one value, the pointer to returned
    stack values.  In such a return, OCFP is always initialized to point to
    the frame returned from, even when no stack values are returned.  This
    allows OCFP to be used to restore CSP.

LRA
    In full call, the passing location for the return PC.

NARGS
    In full call, the number of arguments passed.  In unknown-values return of
    other than one value, the number of values returned.


\subsection{Full call}

What is our usage of CFP, OCFP and CSP?  

It is an invariant that CSP always points after any useful information so that
at any time an interrupt can come and allocate stuff in the stack.

TR call is also a constraint: we can't deallocate the caller's frame before the
call, since it holds the stack arguments for the call.  

What we do is have the caller set up CFP, and have the callee set CSP to CFP
plus the frame size.  The caller leaves CSP alone: the callee is the one who
does any necessary stack deallocation.

In a TR call, we don't do anything: CFP is left as CFP, and CSP points to the
end of the frame, keeping the stack arguments from being trashed.

In a normal call, CFP is set to CSP, causing the callee's frame to be allocated
after the current frame.


\subsection{Unknown values return}

The unknown values return convention is always used in full call, and is used
in local call when the compiler either can't prove that a fixed number of
values are returned, or decides not to use the fixed values convention to allow
tail-recursive XEP calls.

The unknown-values return convention has variants: single value and variable
values.  We make this distinction to optimize the important case of a returner
who knows exactly one value is being returned.  Note that it is possible to
return a single value using the variable-values convention, but it is less
efficient.

We indicate single-value return by returning at the return-pc+4; variable value
return is indicated by returning at the return PC.

Single-value return makes only the following guarantees:
    A0 holds the value returned.
    CSP has been reset: there is no garbage on the stack.

In variable value return, more information is passed back:
    A0..A2 hold the first three return values.  If fewer than three values are
    returned, then the unused registers are initialized to NIL.

    OCFP points to the frame returned from.  Note that because of our
    tail-recursive implementation of call, the frame receiving the values is
    always immediately under the frame returning the values.  This means that
    we can use OCFP to index the values when we access them, and to restore
    CSP when we want to discard them.

    NARGS holds the number of values returned.

    CSP is always (+ OCFP (* NARGS 4)), i.e. there is room on the stack
    allocated for all returned values, even if they are all actually passed in
    registers.


\subsection{External Entry Points}

Things that need to be done on XEP entry:
 1] Allocate frame
 2] Move more arg above the frame, saving context
 3] Set up env, saving closure pointer if closure
 4] Move arguments from closure to local home
    Move old-cont and return-pc to the save locations
 5] Argument count checking and dispatching

XEP VOPs:
\begin{verbatim}
Allocate-Frame
Copy-More-Arg <nargs-tn> 'fixed {in a3} => <context>, <count>
Setup-Environment
Setup-Closure-Environment => <closure>
Verify-Argument-Count <nargs-tn> 'count {for fixed-arg lambdas}
Argument-Count-Error <nargs-tn> {Drop-thru on hairy arg dispatching}
Use fast-if-=/fixnum and fast-if-</fixnum for dispatching.
\end{verbatim}

Closure vops:
\begin{verbatim}
make-closure <fun entry> <slot count> => <closure>
closure-init <closure> <values> 'slot
\end{verbatim}

Things that need to be done on all function entry:
\begin{itemize}
\item Move arguments to the variable home (consing value cells as necessary)
\item Move environment values to the local home
\item Move old-cont and return-pc to the save locations
\end{itemize}

\section{Calls}

Calling VOP's are a cross product of the following sets (with some members
missing):
   Return values
      multiple (all values)
      fixed (calling with unknown values conventions, wanting a certain
             number.)
      known (only in local call where caller/callee agree on number of
      	     values.)
      tail (doesn't return but does tail call)
   What function
      local
      named (going through symbol, like full but stash fun name for error sys)
      full (have a function)
   Args
      fixed (number of args are known at compile-time)
      variable (MULTIPLE-VALUE-CALL and APPLY)

Note on all jumps for calls and returns that we want to put some instruction
in the jump's delay slot(s).

Register usage at the time of the call:

LEXENV
   This holds the lexical environment to use during the call if it's a closure,
   and it is undefined otherwise.

CNAME
   This holds the symbol for a named call and garbage otherwise.

OCFP
   This holds the frame pointer, which the system restores upon return.  The
   callee saves this if necessary; this is passed as a pseudo-argument.

A0 ... An
   These holds the first n+1 arguments.

NARGS
   This holds the number of arguments, as a fixnum.

LRA
   This holds the lisp-return-address object which indicates where to return.
   For a tail call, this retains its current value.  The callee saves this
   if necessary; this is passed as a pseudo-argument.

CODE
   This holds the function object being called.

CSP
   The caller ignores this.  The callee sets it as necessary based on CFP.

CFP
   This holds the callee's frame pointer.  Caller sets this to the new frame
   pointer, which it remembered when it started computing arguments; this is
   CSP if there were no stack arguments.  For a tail call CFP retains its
   current value.

NSP
   The system uses this within a single function.  A function using NSP must
   allocate and deallocate before returning or making a tail call.

Register usage at the time of the return for single value return, which
goes with the unknown-values convention the caller used.

A0
   This holds the value.

CODE
   This holds the lisp-return-address at which the system continues executing.

CSP
   This holds the CFP.  That is, the stack is guaranteed to be clean, and there
   is no code at the return site to adjust the CSP.

CFP
   This holds the OCFP.

Additional register usage for multiple value return:

NARGS
   This holds the number of values returned.

A0 ... An
   These holds the first n+1 values, or NIL if there are less than n+1 values.

CSP
   Returner stores CSP to hold its CFP + NARGS * \verb+<address units per word>+

OCFP
   Returner stores this as its CFP, so the returnee has a handle on either
   the start of the returned values on the stack.


ALLOCATE FULL CALL FRAME.

If the number of call arguments (passed to the VOP as an info argument)
indicates that there are stack arguments, then it makes some callee frame for
arguments:
\begin{verbatim}
   VOP-result <- CSP
   CSP <- CSP + value of VOP info arg times address units per word.
\end{verbatim}

In a call sequence, move some arguments to the right places.

There's a variety of MOVE-ARGUMENT VOP's.

FULL CALL VOP'S
(variations determined by whether it's named, it's a tail call, there
is a variable arg count, etc.)
\begin{verbatim}
  if variable number of arguments
    NARGS <- (CSP - value of VOP argument) shift right by address units per word.
    A0...An <- values off of VOP argument (just fill them all)
  else
    NARGS <- value of VOP info argument (always a constant)

  if tail call
    OCFP <- value from VOP argument
    LRA <- value from VOP argument
    CFP stays the same since we reuse the frame
    NSP <- NFP
  else
    OCFP <- CFP
    LRA <- compute LRA by adding an assemble-time determined constant to
    	   CODE.
    CFP <- new frame pointer (remembered when starting to compute args)
           This is CSP if no stack args.
    when (current-nfp-tn VOP-self-pointer)
      stack-temp <- NFP

  if named
    CNAME <- function symbol name
    the-fun <- function object out of symbol

  LEXENV <- the-fun (from previous line or VOP argument)
  CODE <- function-entry (the first word after the-fun)
  LIP <- calc first instruction addr (CODE + constant-offset)
  jump and run off temp

  <emit Lisp return address data-block>
  <default and move return values OR receive return values>
  when (current-nfp-tn VOP-self-pointer)
    NFP <- stack-temp
\end{verbatim}
Callee:

\begin{verbatim}
XEP-ALLOCATE-FRAME
  emit function header (maybe initializes offset back to component start,
  			but other pointers are set up at load-time.  Pads
			to dual-word boundary.)
  CSP <- CFP + compile-time determined constant (frame size)
  if the function uses the number stack
    NFP <- NSP
    NSP <- NSP + compile-time determined constant (number stack frame size)
\end{verbatim}

\begin{verbatim}
SETUP-ENVIRONMENT
(either use this or the next one)

CODE <- CODE - assembler-time determined offset from function-entry back to
	       the code data-block address.
\end{verbatim}

\begin{verbatim}
SETUP-CLOSURE-ENVIRONMENT
(either use this or the previous one)
After this the CLOSURE-REF VOP can reference closure variables.

VOP-result <- LEXENV
CODE <- CODE - assembler-time determined offset from function-entry back to
	       the code data-block address.
\end{verbatim}

Return VOP's
RETURN and RETURN-MULTIPLE are for the unknown-values return convention.
For some previous caller this is either it wants n values (and it doesn't
know how many are coming), or it wants all the values returned (and it 
doesn't know how many are coming).


RETURN
(known fixed number of values, used with the unknown-values convention
 in the caller.)
When compiler invokes VOP, all values are already where they should be;
just get back to caller.

\begin{verbatim}
when (current-nfp-tn VOP-self-pointer)
  ;; The number stack grows down in memory.
  NSP <- NFP + number stack frame size for calls within the currently
                  compiling component
	       times address units per word
CODE <- value of VOP argument with LRA
if VOP info arg is 1 (number of values we know we're returning)
  CSP <- CFP
  LIP <- calc target addr
          (CODE + skip over LRA header word + skip over address units per branch)
	  (The branch is in the caller to skip down to the MV code.)
else
  NARGS <- value of VOP info arg
  nil out unused arg regs
  OCFP <- CFP  (This indicates the start of return values on the stack,
  		but you leave space for those in registers for convenience.)
  CSP <- CFP + NARGS * address-units-per-word
  LIP <- calc target addr (CODE + skip over LRA header word)
CFP <- value of VOP argument with OCFP
jump and run off LIP
\end{verbatim}

RETURN-MULTIPLE
(unknown number of values, used with the unknown-values convention in
 the caller.)
When compiler invokes VOP, it gets TN's representing a pointer to the
values on the stack and how many values were computed.

\begin{verbatim}
when (current-nfp-tn VOP-self-pointer)
  ;; The number stack grows down in memory.
  NSP <- NFP + number stack frame size for calls within the currently
                  compiling component
	       times address units per word
NARGS <- value of VOP argument
copy the args to the beginning of the current (returner's) frame.
   Actually some go into the argument registers.  When putting the rest at
   the beginning of the frame, leave room for those in the argument registers.
CSP <- CFP + NARGS * address-units-per-word
nil out unused arg regs
OCFP <- CFP  (This indicates the start of return values on the stack,
	      but you leave space for those in registers for convenience.)
CFP <- value of VOP argument with OCFP
CODE <- value of VOP argument with LRA
LIP <- calc target addr (CODE + skip over LRA header word)
jump and run off LIP
\end{verbatim}

Returnee
The call VOP's call DEFAULT-UNKNOWN-VALUES or RECEIVE-UNKNOWN-VALUES after
spitting out transfer control to get stuff from the returner.

DEFAULT-UNKNOWN-VALUES
(We know what we want and we got something.)
If returnee wants one value, it never does anything to deal with a shortage
of return values.  However, if start at PC, then it has to adjust the stack
pointer to dump extra values (move OCFP into CSP).  If it starts at PC+N,
then it just goes along with the "want one value, got it" case.
If the returnee wants multiple values, and there's a shortage of return
values, there are two cases to handle.  One, if the returnee wants fewer
values than there are return registers, and we start at PC+N, then it fills
in return registers \verb|A1..A<desired values necessary>|; if we start at PC,
then the returnee is fine since the returning conventions have filled in
the unused return registers with nil, but the returnee must adjust the
stack pointer to dump possible stack return values (move OCFP to CSP).
Two, if the returnee wants more values than the number of return registers,
and it starts at PC+N (got one value), then it sets up returnee state as if
an unknown number of values came back:
\begin{verbatim}
   A0 has the one value
   A1..An get nil
   NARGS gets 1
   OCFP gets CSP, so general code described below can move OCFP into CSP
If we start at PC, then branch down to the general "got k values, wanted n"
code which takes care of the following issues:
   If k < n, fill in stack return values of nil for shortage of return
      values and move OCFP into CSP
   If k >= n, move OCFP into CSP
This also restores CODE from LRA by subtracting an assemble-time constant.
\end{verbatim}

RECEIVE-UKNOWN-VALUES
(I want whatever I get.)
We want these at the end of our frame.  When the returnee starts at
PC, it moves the return value registers to OCFP..OCFP[An] ignoring where
the end of the stack is and whether all the return value registers had
values.  The returner left room on the stack before the stack return values
for the register return values.  When the returnee starts at PC+N, bump CSP
by 1 and copy A0 there.
This also restores CODE from LRA by subtracting an assemble-time constant.


Local call

There are three flavors:
   1] KNOWN-CALL-LOCAL
      Uses known call convention where caller and callee agree where all
      the values are, and there's a fixed number of return values.
   2] CALL-LOCAL
      Uses the unknown-values convention, but we expect a particular
      number of values in return.
   3] MULTIPLE-CALL-LOCAL
      Uses the unknown-values convention, but we want all values returned.

ALLOCATE-FRAME

If the number of call arguments (passed to the VOP as an info argument)
indicates that there are stack arguments, then it makes some callee frame for
arguments:
\begin{verbatim}
   VOP-result1 <- CSP
   CSP <- CSP + control stack frame size for calls within the currently
   		   compiling component
   		times address units per word.
   when (callee-nfp-tn <VOP info arg holding callee>)
     ;; The number stack grows down.
     ;; May have to round to dual-word boundary if machines C calling
     ;;    conventions demand this.
     NSP <- NSP - number stack frame size for calls within the currently
     		     compiling component
		  times address units per word
     VOP-result2 <- NSP
\end{verbatim}
KNOWN-CALL-LOCAL, CALL-LOCAL, MULTIPLE-CALL-LOCAL
KNOWN-CALL-LOCAL has no need to affect CODE since CODE is the same for the
caller/returnee and the returner.  This uses KNOWN-RETURN.
With CALL-LOCAL and MULTIPLE-CALL-LOCAL, the caller/returnee must fixup
CODE since the callee may do a tail full call.  This happens in the code
emitted by DEFAULT-UNKNOWN-VALUES and RECEIVE-UNKNOWN-VALUES.  We use these
return conventions since we don't know what kind of values the returner
will give us.  This could happen due to a tail full call to an unknown
function, or because the callee had different return points that returned
various numbers of values.

\begin{verbatim}
when (current-nfp-tn VOP-self-pointer)   ;Get VOP self-pointer with
					 ;DEFINE-VOP switch :vop-var.
  stack-temp <- NFP
CFP <- value of VOP arg
when (callee-nfp-tn <VOP info arg holding callee>)
  <where-callee-wants-NFP-tn>  <-  value of VOP arg
<where-callee-wants-LRA-tn>  <-  compute LRA by adding an assemble-time
				 determined constant to CODE.
jump and run off VOP info arg holding start instruction for callee

<emit Lisp return address data-block>
<case call convention
  known: do nothing
  call: default and move return values
  multiple: receive return values
>
when (current-nfp-tn VOP-self-pointer)   
  NFP <- stack-temp
\end{verbatim}

KNOWN-RETURN
\begin{verbatim}
CSP <- CFP
when (current-nfp-tn VOP-self-pointer)
  ;; number stack grows down in memory.
  NSP <- NFP + number stack frame size for calls within the currently
                  compiling component
	       times address units per word
LIP <- calc target addr (value of VOP arg + skip over LRA header word)
CFP <- value of VOP arg
jump and run off LIP

\end{verbatim}


\chapter{Standard Primitives}


\chapter{Customizing VMR Conversion}

Another way in which different implementations differ is in the relative cost
of operations.  On machines without an integer multiply instruction, it may be
desirable to convert multiplication by a constant into shifts and adds, while
this is surely a bad idea on machines with hardware support for multiplication.
Part of the tuning process for an implementation will be adding implementation
dependent transforms and disabling undesirable standard transforms.

When practical, ICR transforms should be used instead of VMR generators, since
transforms are more portable and less error-prone.  Note that the Lisp code
need not be implementation independent: it may contain all sorts of
sub-primitives and similar stuff.  Generally a function should be implemented
using a transform instead of a VMR translator unless it cannot be implemented
as a transform due to being totally evil or it is just as easy to implement as
a translator because it is so simple.


\section{Constant Operands}

If the code emitted for a VOP when an argument is constant is very different
than the non-constant case, then it may be desirable to special-case the
operation in VMR conversion by emitting different VOPs.  An example would be if
SVREF is only open-coded when the index is a constant, and turns into a miscop
call otherwise.  We wouldn't want constant references to spuriously allocate
all the miscop linkage registers on the off chance that the offset might not be
constant.  See the :constant feature of VOP primitive type restrictions.


\section{Supporting Multiple Hardware Configurations}


A winning way to change emitted code depending on the hardware configuration,
i.e. what FPA is present is to do this using primitive types.  Note that the
Primitive-Type function is VM supplied, and can look at any appropriate
hardware configuration switches.  Short-Float can become 6881-Short-Float,
AFPA-Short-Float, etc.  There would be separate SBs and SCs for the registers
of each kind of FP hardware, with each hardware-specific primitive type
using the appropriate float register SC.  Then the hardware specific templates
would provide AFPA-Short-Float as the argument type restriction.

Primitive type changes:

The primitive-type structure is given a new \%Type slot, which is the CType
structure that is equivalent to this type.  There is also a Guard slot, which,
if true is a function that control whether this primitive type is allowed (due
to hardware configuration, etc.)  

We add new :Type and :Guard keywords to Def-Primitive-Type.  Type is the type
specifier that is equivalent (default to the primitive-type name), and Guard is
an expression evaluated in the null environment that controls whether this type
applies (default to none, i.e. constant T).

The Primitive-Type-Type function returns the Lisp CType corresponding to a
primitive type.  This is the \%Type unless there is a guard that returns false,
in which case it is the empty type (i.e. NIL).

[But this doesn't do what we want it to do, since we will compute the
function type for a template at load-time, so they will correspond to whatever
configuration was in effect then.  Maybe we don't want to dick with guards here
(if at all).  I guess we can defer this issue until we actually support
different FP configurations.  But it would seem pretty losing to separately
flame about all the different FP configurations that could be used to open-code
+ whenever we are forced to closed-code +.

If we separately report each better possibly applicable template that we
couldn't use, then it would be reasonable to report any conditional template
allowed by the configuration.  

But it would probably also be good to give some sort of hint that perhaps it
would be a good time to make sure you understand how to tell the compiler to
compile for a particular configuration.  Perhaps if there is a template that
applies *but for the guard*, then we could give a note.  This way, if someone
thinks they are being efficient by throwing in lots of declarations, we can let
them know that they may have to do more.

I guess the guard should be associated with the template rather than the
primitive type.  This would allow LTN and friends to easily tell whether a
template applies in this configuration.  It is also probably more natural for
some sorts of things: with some hardware variants, it may be that the SBs and
representations (SCs) are really the same, but there are some different allowed
operations.  In this case, we could easily conditionalize VOPs without the
increased complexity due to bogus SCs.  If there are different storage
resources, then we would conditionalize Primitive-Type as well.



\section{Special-case VMR convert methods}

    (defun continuation-tn (cont \&optional (check-p t))
      ...)
Return the TN which holds Continuation's first result value.  In general
this may emit code to load the value into a TN.  If Check-P is true, then
when policy indicates, code should be emitted to check that the value satisfies
the continuation asserted type.

    (defun result-tn (cont)
      ...)
Return the TN that Continuation's first value is delivered in.  In general,
may emit code to default any additional values to NIL.

    (defun result-tns (cont n)
      ...)
Similar to Result-TN, except that it returns a list of N result TNs, one
for each of the first N values.


Nearly all open-coded functions should be handled using standard template
selection.  Some (all?) exceptions:
\begin{itemize}
\item List, List* and Vector take arbitrary numbers of arguments.  Could
    implement Vector as a source transform.  Could even do List in a transform
    if we explicitly represent the stack args using \%More-Args or something.
\item \%Typep varies a lot depending on the type specifier.  We don't want to
    transform it, since we want \%Typep as a canonical form so that we can do
    type optimizations.
\item Apply is weird.
\item Funny functions emitted by the compiler: \%Listify-Rest-Args, Arg,
    \%More-Args, \%Special-Bind, \%Catch, \%Unknown-Values (?), \%Unwind-Protect,
    \%Unwind, \%\%Primitive.
\end{itemize}
