% (c) The AQUA Project, Glasgow University, 1993-1998
\section[SimplUtils]{The simplifier utilities}

module SimplUtils (
	-- Rebuilding
	mkLam, mkCase, prepareAlts, bindCaseBndr,

	-- Inlining,
	preInlineUnconditionally, postInlineUnconditionally, 
	activeInline, activeRule, inlineMode,

	-- The continuation type
	SimplCont(..), DupFlag(..), ArgInfo(..),
	contIsDupable, contResultType, contIsTrivial, contArgs, dropArgs, 
	countValArgs, countArgs, splitInlineCont,
	mkBoringStop, mkLazyArgStop, contIsRhsOrArg,
	interestingCallContext, interestingArgContext,

	interestingArg, mkArgInfo,
    ) where

#include "HsVersions.h"

import SimplEnv
import DynFlags
import StaticFlags
import CoreSyn
import qualified CoreSubst
import PprCore
import CoreFVs
import CoreUtils
import CoreArity	( etaExpand, exprEtaExpandArity )
import CoreUnfold
import Name
import Id
import Var	( isCoVar )
import NewDemand
import SimplMonad
import Type	hiding( substTy )
import Coercion ( coercionKind )
import TyCon
import Unify	( dataConCannotMatch )
import VarSet
import BasicTypes
import Util
import MonadUtils
import Outputable
import FastString

import Data.List

%*									*
		The SimplCont type
%*									*

A SimplCont allows the simplifier to traverse the expression in a 
zipper-like fashion.  The SimplCont represents the rest of the expression,
"above" the point of interest.

You can also think of a SimplCont as an "evaluation context", using
that term in the way it is used for operational semantics. This is the
way I usually think of it, For example you'll often see a syntax for
evaluation context looking like
	C ::= []  |  C e   |  case C of alts  |  C `cast` co
That's the kind of thing we are doing here, and I use that syntax in
the comments.

Key points:
  * A SimplCont describes a *strict* context (just like 
    evaluation contexts do).  E.g. Just [] is not a SimplCont

  * A SimplCont describes a context that *does not* bind
    any variables.  E.g. \x. [] is not a SimplCont

data SimplCont	
  = Stop		-- An empty context, or hole, []     
	CallCtxt	-- True <=> There is something interesting about
			--          the context, and hence the inliner
			--	    should be a bit keener (see interestingCallContext)
			-- Specifically:
			--     This is an argument of a function that has RULES
			--     Inlining the call might allow the rule to fire

  | CoerceIt 		-- C `cast` co
	OutCoercion		-- The coercion simplified

  | ApplyTo  		-- C arg
	InExpr SimplEnv		-- The argument and its static env

  | Select   		-- case C of alts
	InId [InAlt] SimplEnv	-- The case binder, alts, and subst-env

  -- The two strict forms have no DupFlag, because we never duplicate them
  | StrictBind 		-- (\x* \xs. e) C
	InId [InBndr]		-- let x* = [] in e 	
	InExpr SimplEnv		--	is a special case 

  | StrictArg 		-- e C
	OutExpr			-- e; *always* of form (Var v `App1` e1 .. `App` en)
	CallCtxt		-- Whether *this* argument position is interesting
 	ArgInfo			-- Whether the function at the head of e has rules, etc
	SimplCont		--     plus strictness flags for *further* args

data ArgInfo 
  = ArgInfo {
	ai_rules :: Bool,	-- Function has rules (recursively)
				--	=> be keener to inline in all args
	ai_strs :: [Bool],	-- Strictness of arguments
				--   Usually infinite, but if it is finite it guarantees
				--   that the function diverges after being given
				--   that number of args
	ai_discs :: [Int]	-- Discounts for arguments; non-zero => be keener to inline
				--   Always infinite

instance Outputable SimplCont where
  ppr (Stop interesting)    	     = ptext (sLit "Stop") <> brackets (ppr interesting)
  ppr (ApplyTo dup arg _ cont)       = ((ptext (sLit "ApplyTo") <+> ppr dup <+> pprParendExpr arg)
					  {-  $$ nest 2 (pprSimplEnv se) -}) $$ ppr cont
  ppr (StrictBind b _ _ _ cont)      = (ptext (sLit "StrictBind") <+> ppr b) $$ ppr cont
  ppr (StrictArg f _ _ cont)         = (ptext (sLit "StrictArg") <+> ppr f) $$ ppr cont
  ppr (Select dup bndr alts _ cont)  = (ptext (sLit "Select") <+> ppr dup <+> ppr bndr) $$ 
				       (nest 4 (ppr alts)) $$ ppr cont 
  ppr (CoerceIt co cont)	     = (ptext (sLit "CoerceIt") <+> ppr co) $$ ppr cont

data DupFlag = OkToDup | NoDup

instance Outputable DupFlag where
  ppr OkToDup = ptext (sLit "ok")
  ppr NoDup   = ptext (sLit "nodup")

mkBoringStop :: SimplCont
mkBoringStop = Stop BoringCtxt

mkLazyArgStop :: CallCtxt -> SimplCont
mkLazyArgStop cci = Stop cci

contIsRhsOrArg :: SimplCont -> Bool
contIsRhsOrArg (Stop {})       = True
contIsRhsOrArg (StrictBind {}) = True
contIsRhsOrArg (StrictArg {})  = True
contIsRhsOrArg _               = False

contIsDupable :: SimplCont -> Bool
contIsDupable (Stop {})                  = True
contIsDupable (ApplyTo  OkToDup _ _ _)   = True
contIsDupable (Select   OkToDup _ _ _ _) = True
contIsDupable (CoerceIt _ cont)          = contIsDupable cont
contIsDupable _                          = False

contIsTrivial :: SimplCont -> Bool
contIsTrivial (Stop {})                   = True
contIsTrivial (ApplyTo _ (Type _) _ cont) = contIsTrivial cont
contIsTrivial (CoerceIt _ cont)           = contIsTrivial cont
contIsTrivial _                           = False

contResultType :: SimplEnv -> OutType -> SimplCont -> OutType
contResultType env ty cont
  = go cont ty
    subst_ty se ty = substTy (se `setInScope` env) ty

    go (Stop {})                      ty = ty
    go (CoerceIt co cont)             _  = go cont (snd (coercionKind co))
    go (StrictBind _ bs body se cont) _  = go cont (subst_ty se (exprType (mkLams bs body)))
    go (StrictArg fn _ _ cont)        _  = go cont (funResultTy (exprType fn))
    go (Select _ _ alts se cont)      _  = go cont (subst_ty se (coreAltsType alts))
    go (ApplyTo _ arg se cont)        ty = go cont (apply_to_arg ty arg se)

    apply_to_arg ty (Type ty_arg) se = applyTy ty (subst_ty se ty_arg)
    apply_to_arg ty _             _  = funResultTy ty

countValArgs :: SimplCont -> Int
countValArgs (ApplyTo _ (Type _) _ cont) = countValArgs cont
countValArgs (ApplyTo _ _        _ cont) = 1 + countValArgs cont
countValArgs _                           = 0

countArgs :: SimplCont -> Int
countArgs (ApplyTo _ _ _ cont) = 1 + countArgs cont
countArgs _                    = 0

contArgs :: SimplCont -> ([OutExpr], SimplCont)
-- Uses substitution to turn each arg into an OutExpr
contArgs cont = go [] cont
    go args (ApplyTo _ arg se cont) = go (substExpr se arg : args) cont
    go args cont		    = (reverse args, cont)

dropArgs :: Int -> SimplCont -> SimplCont
dropArgs 0 cont = cont
dropArgs n (ApplyTo _ _ _ cont) = dropArgs (n-1) cont
dropArgs n other		= pprPanic "dropArgs" (ppr n <+> ppr other)

splitInlineCont :: SimplCont -> Maybe (SimplCont, SimplCont)
-- Returns Nothing if the continuation should dissolve an InlineMe Note
-- Return Just (c1,c2) otherwise, 
--	where c1 is the continuation to put inside the InlineMe 
--	and   c2 outside

-- Example: (__inline_me__ (/\a. e)) ty
--	Here we want to do the beta-redex without dissolving the InlineMe
-- See test simpl017 (and Trac #1627) for a good example of why this is important

splitInlineCont (ApplyTo dup (Type ty) se c)
  | Just (c1, c2) <- splitInlineCont c = Just (ApplyTo dup (Type ty) se c1, c2)
splitInlineCont cont@(Stop {})         = Just (mkBoringStop, cont)
splitInlineCont cont@(StrictBind {})   = Just (mkBoringStop, cont)
splitInlineCont _                      = Nothing
	-- NB: we dissolve an InlineMe in any strict context, 
	--     not just function aplication.  
	-- E.g.  foldr k z (__inline_me (case x of p -> build ...))
	--     Here we want to get rid of the __inline_me__ so we
	--     can float the case, and see foldr/build
	-- However *not* in a strict RHS, else we get
	-- 	   let f = __inline_me__ (\x. e) in ...f...
	-- Now if f is guaranteed to be called, hence a strict binding
	-- we don't thereby want to dissolve the __inline_me__; for
	-- example, 'f' might be a  wrapper, so we'd inline the worker

Note [Interesting call context]
We want to avoid inlining an expression where there can't possibly be
any gain, such as in an argument position.  Hence, if the continuation
is interesting (eg. a case scrutinee, application etc.) then we
inline, otherwise we don't.  

Previously some_benefit used to return True only if the variable was
applied to some value arguments.  This didn't work:

	let x = _coerce_ (T Int) Int (I# 3) in
	case _coerce_ Int (T Int) x of
		I# y -> ....

we want to inline x, but can't see that it's a constructor in a case
scrutinee position, and some_benefit is False.

Another example:

dMonadST = _/\_ t -> :Monad (g1 _@_ t, g2 _@_ t, g3 _@_ t)

....  case dMonadST _@_ x0 of (a,b,c) -> ....

we'd really like to inline dMonadST here, but we *don't* want to
inline if the case expression is just

	case x of y { DEFAULT -> ... }

since we can just eliminate this case instead (x is in WHNF).  Similar
applies when x is bound to a lambda expression.  Hence
contIsInteresting looks for case expressions with just a single
default case.

interestingCallContext :: SimplCont -> CallCtxt
-- See Note [Interesting call context]
interestingCallContext cont
  = interesting cont
    interesting (Select _ bndr _ _ _)
	| isDeadBinder bndr = CaseCtxt
	| otherwise	    = ArgCtxt False 2	-- If the binder is used, this
						-- is like a strict let
    interesting (ApplyTo _ arg _ cont)
	| isTypeArg arg = interesting cont
	| otherwise     = ValAppCtxt 	-- Can happen if we have (f Int |> co) y
					-- If f has an INLINE prag we need to give it some
					-- motivation to inline. See Note [Cast then apply]
					-- in CoreUnfold

    interesting (StrictArg _ cci _ _)	= cci
    interesting (StrictBind {})	  	= BoringCtxt
    interesting (Stop cci)   		= cci
    interesting (CoerceIt _ cont) 	= interesting cont
	-- If this call is the arg of a strict function, the context
	-- is a bit interesting.  If we inline here, we may get useful
	-- evaluation information to avoid repeated evals: e.g.
	--	x + (y * z)
	-- Here the contIsInteresting makes the '*' keener to inline,
	-- which in turn exposes a constructor which makes the '+' inline.
	-- Assuming that +,* aren't small enough to inline regardless.
	-- It's also very important to inline in a strict context for things
	-- like
	--		foldr k z (f x)
	-- Here, the context of (f x) is strict, and if f's unfolding is
	-- a build it's *great* to inline it here.  So we must ensure that
	-- the context for (f x) is not totally uninteresting.

mkArgInfo :: Id
	  -> Int	-- Number of value args
	  -> SimplCont	-- Context of the call
	  -> ArgInfo

mkArgInfo fun n_val_args call_cont
  | n_val_args < idArity fun		-- Note [Unsaturated functions]
  = ArgInfo { ai_rules = False
	    , ai_strs = vanilla_stricts 
	    , ai_discs = vanilla_discounts }
  | otherwise
  = ArgInfo { ai_rules = interestingArgContext fun call_cont
	    , ai_strs  = add_type_str (idType fun) arg_stricts
	    , ai_discs = arg_discounts }
    vanilla_discounts, arg_discounts :: [Int]
    vanilla_discounts = repeat 0
    arg_discounts = case idUnfolding fun of
			CoreUnfolding _ _ _ _ _ (UnfoldIfGoodArgs _ discounts _ _)
			      -> discounts ++ vanilla_discounts
			_     -> vanilla_discounts

    vanilla_stricts, arg_stricts :: [Bool]
    vanilla_stricts  = repeat False

      = case splitStrictSig (idNewStrictness fun) of
	  (demands, result_info)
		| not (demands `lengthExceeds` n_val_args)
		-> 	-- Enough args, use the strictness given.
			-- For bottoming functions we used to pretend that the arg
			-- is lazy, so that we don't treat the arg as an
			-- interesting context.  This avoids substituting
			-- top-level bindings for (say) strings into 
			-- calls to error.  But now we are more careful about
			-- inlining lone variables, so its ok (see SimplUtils.analyseCont)
		   if isBotRes result_info then
			map isStrictDmd demands		-- Finite => result is bottom
			map isStrictDmd demands ++ vanilla_stricts
	       | otherwise
	       -> WARN( True, text "More demands than arity" <+> ppr fun <+> ppr (idArity fun) 
				<+> ppr n_val_args <+> ppr demands ) 
		   vanilla_stricts	-- Not enough args, or no strictness

    add_type_str :: Type -> [Bool] -> [Bool]
    -- If the function arg types are strict, record that in the 'strictness bits'
    -- No need to instantiate because unboxed types (which dominate the strict
    -- types) can't instantiate type variables.
    -- add_type_str is done repeatedly (for each call); might be better 
    -- once-for-all in the function
    -- But beware primops/datacons with no strictness
    add_type_str _ [] = []
    add_type_str fun_ty strs		-- Look through foralls
	| Just (_, fun_ty') <- splitForAllTy_maybe fun_ty	-- Includes coercions
	= add_type_str fun_ty' strs
    add_type_str fun_ty (str:strs)	-- Add strict-type info
	| Just (arg_ty, fun_ty') <- splitFunTy_maybe fun_ty
	= (str || isStrictType arg_ty) : add_type_str fun_ty' strs
    add_type_str _ strs
	= strs

{- Note [Unsaturated functions]
Consider (test eyeball/inline4)
	x = a:as
	y = f x
where f has arity 2.  Then we do not want to inline 'x', because
it'll just be floated out again.  Even if f has lots of discounts
on its first argument -- it must be saturated for these to kick in

interestingArgContext :: Id -> SimplCont -> Bool
-- If the argument has form (f x y), where x,y are boring,
-- and f is marked INLINE, then we don't want to inline f.
-- But if the context of the argument is
--	g (f x y) 
-- where g has rules, then we *do* want to inline f, in case it
-- exposes a rule that might fire.  Similarly, if the context is
--	h (g (f x x))
-- where h has rules, then we do want to inline f; hence the
-- call_cont argument to interestingArgContext
-- The interesting_arg_ctxt flag makes this happen; if it's
-- set, the inliner gets just enough keener to inline f 
-- regardless of how boring f's arguments are, if it's marked INLINE
-- The alternative would be to *always* inline an INLINE function,
-- regardless of how boring its context is; but that seems overkill
-- For example, it'd mean that wrapper functions were always inlined
interestingArgContext fn call_cont
  = idHasRules fn || go call_cont
    go (Select {})	     = False
    go (ApplyTo {})	     = False
    go (StrictArg _ cci _ _) = interesting cci
    go (StrictBind {})	     = False	-- ??
    go (CoerceIt _ c)	     = go c
    go (Stop cci)            = interesting cci

    interesting (ArgCtxt rules _) = rules
    interesting _                 = False

%*									*
\subsection{Decisions about inlining}
%*									*

Inlining is controlled partly by the SimplifierMode switch.  This has two

	SimplGently	(a) Simplifying before specialiser/full laziness
			(b) Simplifiying inside INLINE pragma
			(c) Simplifying the LHS of a rule
			(d) Simplifying a GHCi expression or Template 
				Haskell splice

	SimplPhase n _	 Used at all other times

The key thing about SimplGently is that it does no call-site inlining.
Before full laziness we must be careful not to inline wrappers,
because doing so inhibits floating
    e.g. ...(case f x of ...)...
    ==> ...(case (case x of I# x# -> fw x#) of ...)...
    ==> ...(case x of I# x# -> case fw x# of ...)...
and now the redex (f x) isn't floatable any more.

The no-inlining thing is also important for Template Haskell.  You might be 
compiling in one-shot mode with -O2; but when TH compiles a splice before
running it, we don't want to use -O2.  Indeed, we don't want to inline
anything, because the byte-code interpreter might get confused about 
unboxed tuples and suchlike.

INLINE pragmas
SimplGently is also used as the mode to simplify inside an InlineMe note.

inlineMode :: SimplifierMode
inlineMode = SimplGently

It really is important to switch off inlinings inside such
expressions.  Consider the following example 

     	let f = \pq -> BIG
     	let g = \y -> f y y
	    {-# INLINE g #-}
     	in ...g...g...g...g...g...

Now, if that's the ONLY occurrence of f, it will be inlined inside g,
and thence copied multiple times when g is inlined.

This function may be inlinined in other modules, so we
don't want to remove (by inlining) calls to functions that have
specialisations, or that may have transformation rules in an importing

E.g. 	{-# INLINE f #-}
		f x = ...g...

and suppose that g is strict *and* has specialisations.  If we inline
g's wrapper, we deny f the chance of getting the specialised version
of g when f is inlined at some call site (perhaps in some other

It's also important not to inline a worker back into a wrapper.
A wrapper looks like
	wraper = inline_me (\x -> ...worker... )
Normally, the inline_me prevents the worker getting inlined into
the wrapper (initially, the worker's only call site!).  But,
if the wrapper is sure to be called, the strictness analyser will
mark it 'demanded', so when the RHS is simplified, it'll get an ArgOf
continuation.  That's why the keep_inline predicate returns True for
ArgOf continuations.  It shouldn't do any harm not to dissolve the
inline-me note under these circumstances.

Note that the result is that we do very little simplification
inside an InlineMe.  

	all xs = foldr (&&) True xs
	any p = all . map p  {-# INLINE any #-}

Problem: any won't get deforested, and so if it's exported and the
importer doesn't use the inlining, (eg passes it as an arg) then we
won't get deforestation at all.  We havn't solved this problem yet!

@preInlineUnconditionally@ examines a bndr to see if it is used just
once in a completely safe way, so that it is safe to discard the
binding inline its RHS at the (unique) usage site, REGARDLESS of how
big the RHS might be.  If this is the case we don't simplify the RHS
first, but just inline it un-simplified.

This is much better than first simplifying a perhaps-huge RHS and then
inlining and re-simplifying it.  Indeed, it can be at least quadratically
better.  Consider

	x1 = e1
	x2 = e2[x1]
	x3 = e3[x2]
	xN = eN[xN-1]

We may end up simplifying e1 N times, e2 N-1 times, e3 N-3 times etc.
This can happen with cascades of functions too:

	f1 = \x1.e1
	f2 = \xs.e2[f1]
	f3 = \xs.e3[f3]


	----  preInlineUnconditionally invariant -----
   IF preInlineUnconditionally chooses to inline x = <rhs>
   THEN doing the inlining should not change the occurrence
	info for the free vars of <rhs>

For example, it's tempting to look at trivial binding like
	x = y
and inline it unconditionally.  But suppose x is used many times,
but this is the unique occurrence of y.  Then inlining x would change
y's occurrence info, which breaks the invariant.  It matters: y
might have a BIG rhs, which will now be dup'd at every occurrenc of x.

Even RHSs labelled InlineMe aren't caught here, because there might be
no benefit from inlining at the call site.

[Sept 01] Don't unconditionally inline a top-level thing, because that
can simply make a static thing into something built dynamically.  E.g.
	x = (a,b)
	main = \s -> h x

[Remember that we treat \s as a one-shot lambda.]  No point in
inlining x unless there is something interesting about the call site.

But watch out: if you aren't careful, some useful foldr/build fusion
can be lost (most notably in spectral/hartel/parstof) because the
foldr didn't see the build.  Doing the dynamic allocation isn't a big
deal, in fact, but losing the fusion can be.  But the right thing here
seems to be to do a callSiteInline based on the fact that there is
something interesting about the call site (it's strict).  Hmm.  That
seems a bit fragile.

Conclusion: inline top level things gaily until Phase 0 (the last
phase), at which point don't.

preInlineUnconditionally :: SimplEnv -> TopLevelFlag -> InId -> InExpr -> Bool
preInlineUnconditionally env top_lvl bndr rhs
  | not active 		   = False
  | opt_SimplNoPreInlining = False
  | otherwise = case idOccInfo bndr of
		  IAmDead	     	     -> True	-- Happens in ((\x.1) v)
	  	  OneOcc in_lam True int_cxt -> try_once in_lam int_cxt
		  _                          -> False
    phase = getMode env
    active = case phase of
		   SimplGently    -> isAlwaysActive act
		   SimplPhase n _ -> isActive n act
    act = idInlineActivation bndr

    try_once in_lam int_cxt	-- There's one textual occurrence
	| not in_lam = isNotTopLevel top_lvl || early_phase
	| otherwise  = int_cxt && canInlineInLam rhs

-- Be very careful before inlining inside a lambda, becuase (a) we must not 
-- invalidate occurrence information, and (b) we want to avoid pushing a
-- single allocation (here) into multiple allocations (inside lambda).  
-- Inlining a *function* with a single *saturated* call would be ok, mind you.
--	|| (if is_cheap && not (canInlineInLam rhs) then pprTrace "preinline" (ppr bndr <+> ppr rhs) ok else ok)
--	where 
--	 	is_cheap = exprIsCheap rhs
--		ok = is_cheap && int_cxt

	-- 	int_cxt		The context isn't totally boring
	-- E.g. let f = \ab.BIG in \y. map f xs
	-- 	Don't want to substitute for f, because then we allocate
	--	its closure every time the \y is called
	-- But: let f = \ab.BIG in \y. map (f y) xs
	--	Now we do want to substitute for f, even though it's not 
	--	saturated, because we're going to allocate a closure for 
	--	(f y) every time round the loop anyhow.

	-- canInlineInLam => free vars of rhs are (Once in_lam) or Many,
	-- so substituting rhs inside a lambda doesn't change the occ info.
	-- Sadly, not quite the same as exprIsHNF.
    canInlineInLam (Lit _)		= True
    canInlineInLam (Lam b e)		= isRuntimeVar b || canInlineInLam e
    canInlineInLam (Note _ e)		= canInlineInLam e
    canInlineInLam _			= False

    early_phase = case phase of
			SimplPhase 0 _ -> False
			_    	       -> True
-- If we don't have this early_phase test, consider
--	x = length [1,2,3]
-- The full laziness pass carefully floats all the cons cells to
-- top level, and preInlineUnconditionally floats them all back in.
-- Result is (a) static allocation replaced by dynamic allocation
--	     (b) many simplifier iterations because this tickles
--		 a related problem; only one inlining per pass
-- On the other hand, I have seen cases where top-level fusion is
-- lost if we don't inline top level thing (e.g. string constants)
-- Hence the test for phase zero (which is the phase for all the final
-- simplifications).  Until phase zero we take no special notice of
-- top level things, but then we become more leery about inlining
-- them.  


@postInlineUnconditionally@ decides whether to unconditionally inline
a thing based on the form of its RHS; in particular if it has a
trivial RHS.  If so, we can inline and discard the binding altogether.

NB: a loop breaker has must_keep_binding = True and non-loop-breakers
only have *forward* references Hence, it's safe to discard the binding
NOTE: This isn't our last opportunity to inline.  We're at the binding
site right now, and we'll get another opportunity when we get to the

Note that we do this unconditional inlining only for trival RHSs.
Don't inline even WHNFs inside lambdas; doing so may simply increase
allocation when the function is called. This isn't the last chance; see
NOTE above.

NB: Even inline pragmas (e.g. IMustBeINLINEd) are ignored here Why?
Because we don't even want to inline them into the RHS of constructor
arguments. See NOTE above

NB: At one time even NOINLINE was ignored here: if the rhs is trivial
it's best to inline it anyway.  We often get a=E; b=a from desugaring,
with both a and b marked NOINLINE.  But that seems incompatible with
our new view that inlining is like a RULE, so I'm sticking to the 'active'
story for now.

    :: SimplEnv -> TopLevelFlag
    -> InId		-- The binder (an OutId would be fine too)
    -> OccInfo 		-- From the InId
    -> OutExpr
    -> Unfolding
    -> Bool
postInlineUnconditionally env top_lvl bndr occ_info rhs unfolding
  | not active		   = False
  | isLoopBreaker occ_info = False	-- If it's a loop-breaker of any kind, don't inline
					-- because it might be referred to "earlier"
  | isExportedId bndr      = False
  | exprIsTrivial rhs 	   = True
  | otherwise
  = case occ_info of
	-- The point of examining occ_info here is that for *non-values* 
	-- that occur outside a lambda, the call-site inliner won't have
	-- a chance (becuase it doesn't know that the thing
	-- only occurs once).   The pre-inliner won't have gotten
	-- it either, if the thing occurs in more than one branch
	-- So the main target is things like
	--	let x = f y in
	--	case v of
	--	   True  -> case x of ...
	--	   False -> case x of ...
	-- I'm not sure how important this is in practice
      OneOcc in_lam _one_br int_cxt	-- OneOcc => no code-duplication issue
	->     smallEnoughToInline unfolding	-- Small enough to dup
			-- ToDo: consider discount on smallEnoughToInline if int_cxt is true
		 	-- NB: Do NOT inline arbitrarily big things, even if one_br is True
			-- Reason: doing so risks exponential behaviour.  We simplify a big
			--	   expression, inline it, and simplify it again.  But if the
			--	   very same thing happens in the big expression, we get 
			--	   exponential cost!
			-- PRINCIPLE: when we've already simplified an expression once, 
			-- make sure that we only inline it if it's reasonably small.

	   &&  ((isNotTopLevel top_lvl && not in_lam) || 
			-- But outside a lambda, we want to be reasonably aggressive
			-- about inlining into multiple branches of case
			-- e.g. let x = <non-value> 
			--	in case y of { C1 -> ..x..; C2 -> ..x..; C3 -> ... } 
			-- Inlining can be a big win if C3 is the hot-spot, even if
			-- the uses in C1, C2 are not 'interesting'
			-- An example that gets worse if you add int_cxt here is 'clausify'

		(isCheapUnfolding unfolding && int_cxt))
			-- isCheap => acceptable work duplication; in_lam may be true
			-- int_cxt to prevent us inlining inside a lambda without some 
			-- good reason.  See the notes on int_cxt in preInlineUnconditionally

      IAmDead -> True	-- This happens; for example, the case_bndr during case of
			-- known constructor:  case (a,b) of x { (p,q) -> ... }
			-- Here x isn't mentioned in the RHS, so we don't want to
			-- create the (dead) let-binding  let x = (a,b) in ...

      _ -> False

-- Here's an example that we don't handle well:
--	let f = if b then Left (\x.BIG) else Right (\y.BIG)
--	in \y. ....case f of {...} ....
-- Here f is used just once, and duplicating the case work is fine (exprIsCheap).
-- But
--  - We can't preInlineUnconditionally because that woud invalidate
--    the occ info for b.
--  - We can't postInlineUnconditionally because the RHS is big, and
--    that risks exponential behaviour
--  - We can't call-site inline, because the rhs is big
-- Alas!

    active = case getMode env of
		   SimplGently    -> isAlwaysActive act
		   SimplPhase n _ -> isActive n act
    act = idInlineActivation bndr

activeInline :: SimplEnv -> OutId -> Bool
activeInline env id
  = case getMode env of
      SimplGently -> False
	-- No inlining at all when doing gentle stuff,
	-- except for local things that occur once (pre/postInlineUnconditionally)
	-- The reason is that too little clean-up happens if you 
	-- don't inline use-once things.   Also a bit of inlining is *good* for
	-- full laziness; it can expose constant sub-expressions.
	-- Example in spectral/mandel/Mandel.hs, where the mandelset 
	-- function gets a useful let-float if you inline windowToViewport

	-- NB: we used to have a second exception, for data con wrappers.
	-- On the grounds that we use gentle mode for rule LHSs, and 
	-- they match better when data con wrappers are inlined.
	-- But that only really applies to the trivial wrappers (like (:)),
	-- and they are now constructed as Compulsory unfoldings (in MkId)
	-- so they'll happen anyway.

      SimplPhase n _ -> isActive n act
    act = idInlineActivation id

activeRule :: DynFlags -> SimplEnv -> Maybe (Activation -> Bool)
-- Nothing => No rules at all
activeRule dflags env
  | not (dopt Opt_EnableRewriteRules dflags)
  = Nothing	-- Rewriting is off
  | otherwise
  = case getMode env of
	SimplGently    -> Just isAlwaysActive
			-- Used to be Nothing (no rules in gentle mode)
			-- Main motivation for changing is that I wanted
			-- 	lift String ===> ...
			-- to work in Template Haskell when simplifying
			-- splices, so we get simpler code for literal strings
	SimplPhase n _ -> Just (isActive n)

%*									*
	Rebuilding a lambda
%*									*

mkLam :: SimplEnv -> [OutBndr] -> OutExpr -> SimplM OutExpr
-- mkLam tries three things
--	a) eta reduction, if that gives a trivial expression
--	b) eta expansion [only if there are some value lambdas]

mkLam _b [] body 
  = return body
mkLam _env bndrs body
  = do	{ dflags <- getDOptsSmpl
	; mkLam' dflags bndrs body }
    mkLam' :: DynFlags -> [OutBndr] -> OutExpr -> SimplM OutExpr
    mkLam' dflags bndrs (Cast body co)
      | not (any bad bndrs)
	-- Note [Casts and lambdas]
      = do { lam <- mkLam' dflags bndrs body
	   ; return (mkCoerce (mkPiTypes bndrs co) lam) }
	co_vars  = tyVarsOfType co
	bad bndr = isCoVar bndr && bndr `elemVarSet` co_vars      

    mkLam' dflags bndrs body
      | dopt Opt_DoEtaReduction dflags,
        Just etad_lam <- tryEtaReduce bndrs body
      = do { tick (EtaReduction (head bndrs))
	   ; return etad_lam }

      | dopt Opt_DoLambdaEtaExpansion dflags,
   	any isRuntimeVar bndrs
      = do { let body' = tryEtaExpansion dflags body
 	   ; return (mkLams bndrs body') }
      | otherwise 
      = return (mkLams bndrs body)

Note [Casts and lambdas]
	(\x. (\y. e) `cast` g1) `cast` g2
There is a danger here that the two lambdas look separated, and the 
full laziness pass might float an expression to between the two.

So this equation in mkLam' floats the g1 out, thus:
	(\x. e `cast` g1)  -->  (\x.e) `cast` (tx -> g1)
where x:tx.

In general, this floats casts outside lambdas, where (I hope) they
might meet and cancel with some other cast:
	\x. e `cast` co   ===>   (\x. e) `cast` (tx -> co)
	/\a. e `cast` co  ===>   (/\a. e) `cast` (/\a. co)
	/\g. e `cast` co  ===>   (/\g. e) `cast` (/\g. co)
  		  	  (if not (g `in` co))

Notice that it works regardless of 'e'.  Originally it worked only
if 'e' was itself a lambda, but in some cases that resulted in 
fruitless iteration in the simplifier.  A good example was when
compiling Text.ParserCombinators.ReadPrec, where we had a definition 
like	(\x. Get `cast` g)
where Get is a constructor with nonzero arity.  Then mkLam eta-expanded
the Get, and the next iteration eta-reduced it, and then eta-expanded 
it again.

Note also the side condition for the case of coercion binders.
It does not make sense to transform
	/\g. e `cast` g  ==>  (/\g.e) `cast` (/\g.g)
because the latter is not well-kinded.

--	c) floating lets out through big lambdas 
--		[only if all tyvar lambdas, and only if this lambda
--		 is the RHS of a let]

{-	Sept 01: I'm experimenting with getting the
	full laziness pass to float out past big lambdsa
 | all isTyVar bndrs,	-- Only for big lambdas
   contIsRhs cont	-- Only try the rhs type-lambda floating
			-- if this is indeed a right-hand side; otherwise
			-- we end up floating the thing out, only for float-in
			-- to float it right back in again!
 = do (floats, body') <- tryRhsTyLam env bndrs body
      return (floats, mkLams bndrs body')

%*									*
		Eta reduction
%*									*

Note [Eta reduction conditions]
We try for eta reduction here, but *only* if we get all the way to an
trivial expression.  We don't want to remove extra lambdas unless we
are going to avoid allocating this thing altogether.

There are some particularly delicate points here:

* Eta reduction is not valid in general:  
	\x. bot  /=  bot
  This matters, partly for old-fashioned correctness reasons but,
  worse, getting it wrong can yield a seg fault. Consider
	f = \x.f x
	h y = case (case y of { True -> f `seq` True; False -> False }) of
		True -> ...; False -> ...

  If we (unsoundly) eta-reduce f to get f=f, the strictness analyser
  says f=bottom, and replaces the (f `seq` True) with just
  (f `cast` unsafe-co).  BUT, as thing stand, 'f' got arity 1, and it
  *keeps* arity 1 (perhaps also wrongly).  So CorePrep eta-expands 
  the definition again, so that it does not termninate after all.
  Result: seg-fault because the boolean case actually gets a function value.
  See Trac #1947.

  So it's important to to the right thing.

* Note [Arity care]: we need to be careful if we just look at f's
  arity. Currently (Dec07), f's arity is visible in its own RHS (see
  Note [Arity robustness] in SimplEnv) so we must *not* trust the
  arity when checking that 'f' is a value.  Otherwise we will
      f = \x. f x
      f = f
  Which might change a terminiating program (think (f `seq` e)) to a 
  non-terminating one.  So we check for being a loop breaker first.

  However for GlobalIds we can look at the arity; and for primops we
  must, since they have no unfolding.  

* Regardless of whether 'f' is a value, we always want to 
  reduce (/\a -> f a) to f
  This came up in a RULE: foldr (build (/\a -> g a))
  did not match 	  foldr (build (/\b -> ...something complex...))
  The type checker can insert these eta-expanded versions,
  with both type and dictionary lambdas; hence the slightly 
  ad-hoc isDictId

* Never *reduce* arity. For example
      f = \xy. g x y
  Then if h has arity 1 we don't want to eta-reduce because then
  f's arity would decrease, and that is bad

These delicacies are why we don't use exprIsTrivial and exprIsHNF here.

tryEtaReduce :: [OutBndr] -> OutExpr -> Maybe OutExpr
tryEtaReduce bndrs body 
  = go (reverse bndrs) body
    incoming_arity = count isId bndrs

    go (b : bs) (App fun arg) | ok_arg b arg = go bs fun	-- Loop round
    go []       fun           | ok_fun fun   = Just fun		-- Success!
    go _        _			     = Nothing		-- Failure!

	-- Note [Eta reduction conditions]
    ok_fun (App fun (Type ty)) 
	| not (any (`elemVarSet` tyVarsOfType ty) bndrs)
	=  ok_fun fun
    ok_fun (Var fun_id)
	=  not (fun_id `elem` bndrs)
	&& (ok_fun_id fun_id || all ok_lam bndrs)
    ok_fun _fun = False

    ok_fun_id fun = fun_arity fun >= incoming_arity

    fun_arity fun 	      -- See Note [Arity care]
       | isLocalId fun && isLoopBreaker (idOccInfo fun) = 0
       | otherwise = idArity fun   	      

    ok_lam v = isTyVar v || isDictId v

    ok_arg b arg = varToCoreExpr b `cheapEqExpr` arg

%*									*
		Eta expansion
%*									*

We go for:
   f = \x1..xn -> N  ==>   f = \x1..xn y1..ym -> N y1..ym
				 (n >= 0)

where (in both cases) 

	* The xi can include type variables

	* The yi are all value variables

	* N is a NORMAL FORM (i.e. no redexes anywhere)
	  wanting a suitable number of extra args.

The biggest reason for doing this is for cases like

	f = \x -> case x of
		    True  -> \y -> e1
		    False -> \y -> e2

Here we want to get the lambdas together.  A good exmaple is the nofib
program fibheaps, which gets 25% more allocation if you don't do this

We may have to sandwich some coerces between the lambdas
to make the types work.   exprEtaExpandArity looks through coerces
when computing arity; and etaExpand adds the coerces as necessary when
actually computing the expansion.

tryEtaExpansion :: DynFlags -> OutExpr -> OutExpr
-- There is at least one runtime binder in the binders
tryEtaExpansion dflags body
  = etaExpand fun_arity body
    fun_arity = exprEtaExpandArity dflags body

%*									*
\subsection{Floating lets out of big lambdas}
%*									*

Note [Floating and type abstraction]
Consider this:
	x = /\a. C e1 e2
We'd like to float this to 
	y1 = /\a. e1
	y2 = /\a. e2
	x  = /\a. C (y1 a) (y2 a)
for the usual reasons: we want to inline x rather vigorously.

You may think that this kind of thing is rare.  But in some programs it is
common.  For example, if you do closure conversion you might get:

	data a :-> b = forall e. (e -> a -> b) :$ e

	f_cc :: forall a. a :-> a
	f_cc = /\a. (\e. id a) :$ ()

Now we really want to inline that f_cc thing so that the
construction of the closure goes away. 

So I have elaborated simplLazyBind to understand right-hand sides that look
	/\ a1..an. body

and treat them specially. The real work is done in SimplUtils.abstractFloats,
but there is quite a bit of plumbing in simplLazyBind as well.

The same transformation is good when there are lets in the body:

	/\abc -> let(rec) x = e in b
	let(rec) x' = /\abc -> let x = x' a b c in e
	/\abc -> let x = x' a b c in b

This is good because it can turn things like:

	let f = /\a -> letrec g = ... g ... in g
	letrec g' = /\a -> ... g' a ...
	let f = /\ a -> g' a

which is better.  In effect, it means that big lambdas don't impede

This optimisation is CRUCIAL in eliminating the junk introduced by
desugaring mutually recursive definitions.  Don't eliminate it lightly!

[May 1999]  If we do this transformation *regardless* then we can
end up with some pretty silly stuff.  For example, 

	    st = /\ s -> let { x1=r1 ; x2=r2 } in ...
	in ..
	let y1 = /\s -> r1
	    y2 = /\s -> r2
	    st = /\s -> ...[y1 s/x1, y2 s/x2]
	in ..

Unless the "..." is a WHNF there is really no point in doing this.
Indeed it can make things worse.  Suppose x1 is used strictly,
and is of the form

	x1* = case f y of { (a,b) -> e }

If we abstract this wrt the tyvar we then can't do the case inline
as we would normally do.

That's why the whole transformation is part of the same process that
floats let-bindings and constructor arguments out of RHSs.  In particular,
it is guarded by the doFloatFromRhs call in simplLazyBind.

abstractFloats :: [OutTyVar] -> SimplEnv -> OutExpr -> SimplM ([OutBind], OutExpr)
abstractFloats main_tvs body_env body
  = ASSERT( notNull body_floats )
    do	{ (subst, float_binds) <- mapAccumLM abstract empty_subst body_floats
	; return (float_binds, CoreSubst.substExpr subst body) }
    main_tv_set = mkVarSet main_tvs
    body_floats = getFloats body_env
    empty_subst = CoreSubst.mkEmptySubst (seInScope body_env)

    abstract :: CoreSubst.Subst -> OutBind -> SimplM (CoreSubst.Subst, OutBind)
    abstract subst (NonRec id rhs)
      = do { (poly_id, poly_app) <- mk_poly tvs_here id
	   ; let poly_rhs = mkLams tvs_here rhs'
		 subst'   = CoreSubst.extendIdSubst subst id poly_app
	   ; return (subst', (NonRec poly_id poly_rhs)) }
	rhs' = CoreSubst.substExpr subst rhs
	tvs_here | any isCoVar main_tvs = main_tvs	-- Note [Abstract over coercions]
		 | otherwise 
		 = varSetElems (main_tv_set `intersectVarSet` exprSomeFreeVars isTyVar rhs')
		-- Abstract only over the type variables free in the rhs
		-- wrt which the new binding is abstracted.  But the naive
		-- approach of abstract wrt the tyvars free in the Id's type
		-- fails. Consider:
		--	/\ a b -> let t :: (a,b) = (e1, e2)
		--		      x :: a     = fst t
		--		  in ...
		-- Here, b isn't free in x's type, but we must nevertheless
		-- abstract wrt b as well, because t's type mentions b.
		-- Since t is floated too, we'd end up with the bogus:
		--	poly_t = /\ a b -> (e1, e2)
		--	poly_x = /\ a   -> fst (poly_t a *b*)
		-- So for now we adopt the even more naive approach of
		-- abstracting wrt *all* the tyvars.  We'll see if that
		-- gives rise to problems.   SLPJ June 98

    abstract subst (Rec prs)
       = do { (poly_ids, poly_apps) <- mapAndUnzipM (mk_poly tvs_here) ids
	    ; let subst' = CoreSubst.extendSubstList subst (ids `zip` poly_apps)
		  poly_rhss = [mkLams tvs_here (CoreSubst.substExpr subst' rhs) | rhs <- rhss]
	    ; return (subst', Rec (poly_ids `zip` poly_rhss)) }
	 (ids,rhss) = unzip prs
	 	-- For a recursive group, it's a bit of a pain to work out the minimal
		-- set of tyvars over which to abstract:
		--	/\ a b c.  let x = ...a... in
		--	 	   letrec { p = ...x...q...
		--			    q = .....p...b... } in
		--		   ...
		-- Since 'x' is abstracted over 'a', the {p,q} group must be abstracted
		-- over 'a' (because x is replaced by (poly_x a)) as well as 'b'.  
		-- Since it's a pain, we just use the whole set, which is always safe
		-- If you ever want to be more selective, remember this bizarre case too:
		--	x::a = x
		-- Here, we must abstract 'x' over 'a'.
	 tvs_here = main_tvs

    mk_poly tvs_here var
      = do { uniq <- getUniqueM
	   ; let  poly_name = setNameUnique (idName var) uniq		-- Keep same name
		  poly_ty   = mkForAllTys tvs_here (idType var)	-- But new type of course
		  poly_id   = transferPolyIdInfo var tvs_here $ -- Note [transferPolyIdInfo] in Id.lhs
			      mkLocalId poly_name poly_ty 
	   ; return (poly_id, mkTyApps (Var poly_id) (mkTyVarTys tvs_here)) }
		-- In the olden days, it was crucial to copy the occInfo of the original var, 
		-- because we were looking at occurrence-analysed but as yet unsimplified code!
		-- In particular, we mustn't lose the loop breakers.  BUT NOW we are looking
		-- at already simplified code, so it doesn't matter
		-- It's even right to retain single-occurrence or dead-var info:
		-- Suppose we started with  /\a -> let x = E in B
		-- where x occurs once in B. Then we transform to:
		--	let x' = /\a -> E in /\a -> let x* = x' a in B
		-- where x* has an INLINE prag on it.  Now, once x* is inlined,
		-- the occurrences of x' will be just the occurrences originally
		-- pinned on x.

Note [Abstract over coercions]
If a coercion variable (g :: a ~ Int) is free in the RHS, then so is the
type variable a.  Rather than sort this mess out, we simply bale out and abstract
wrt all the type variables if any of them are coercion variables.

Historical note: if you use let-bindings instead of a substitution, beware of this:

		-- Suppose we start with:
		--	x = /\ a -> let g = G in E
		-- Then we'll float to get
		--	x = let poly_g = /\ a -> G
		--	    in /\ a -> let g = poly_g a in E
		-- But now the occurrence analyser will see just one occurrence
		-- of poly_g, not inside a lambda, so the simplifier will
		-- PreInlineUnconditionally poly_g back into g!  Badk to square 1!
		-- (I used to think that the "don't inline lone occurrences" stuff
		--  would stop this happening, but since it's the *only* occurrence,
		--  PreInlineUnconditionally kicks in first!)
		-- Solution: put an INLINE note on g's RHS, so that poly_g seems
		--	     to appear many times.  (NB: mkInlineMe eliminates
		--	     such notes on trivial RHSs, so do it manually.)

%*									*
%*									*

prepareAlts tries these things:

1.  If several alternatives are identical, merge them into
    a single DEFAULT alternative.  I've occasionally seen this 
    making a big difference:

	case e of		=====>     case e of
	  C _ -> f x			     D v -> ....v....
	  D v -> ....v....		     DEFAULT -> f x
	  DEFAULT -> f x

   The point is that we merge common RHSs, at least for the DEFAULT case.
   [One could do something more elaborate but I've never seen it needed.]
   To avoid an expensive test, we just merge branches equal to the *first*
   alternative; this picks up the common cases
	a) all branches equal
	b) some branches equal to the DEFAULT (which occurs first)

2.  Case merging:
       case e of b {             ==>   case e of b {
    	 p1 -> rhs1	                 p1 -> rhs1
    	 ...	                         ...
    	 pm -> rhsm                      pm -> rhsm
    	 _  -> case b of b' {            pn -> let b'=b in rhsn
 		     pn -> rhsn          ...
 		     ...                 po -> let b'=b in rhso
 		     po -> rhso          _  -> let b'=b in rhsd
 		     _  -> rhsd
    which merges two cases in one case when -- the default alternative of
    the outer case scrutises the same variable as the outer case This
    transformation is called Case Merging.  It avoids that the same
    variable is scrutinised multiple times.

The case where transformation (1) showed up was like this (lib/std/PrelCError.lhs):

	x | p `is` 1 -> e1
	  | p `is` 2 -> e2

where @is@ was something like
	p `is` n = p /= (-1) && p == n

This gave rise to a horrible sequence of cases

	case p of
	  (-1) -> $j p
	  1    -> e1
	  DEFAULT -> $j p

and similarly in cascade for all the join points!

Note [Dead binders]
We do this *here*, looking at un-simplified alternatives, because we
have to check that r doesn't mention the variables bound by the
pattern in each alternative, so the binder-info is rather useful.

prepareAlts :: SimplEnv -> OutExpr -> OutId -> [InAlt] -> SimplM ([AltCon], [InAlt])
prepareAlts env scrut case_bndr' alts
  = do	{ dflags <- getDOptsSmpl
	; alts <- combineIdenticalAlts case_bndr' alts

	; let (alts_wo_default, maybe_deflt) = findDefault alts
	      alt_cons = [con | (con,_,_) <- alts_wo_default]
	      imposs_deflt_cons = nub (imposs_cons ++ alt_cons)
		-- "imposs_deflt_cons" are handled 
		--   EITHER by the context, 
		--   OR by a non-DEFAULT branch in this case expression.

	; default_alts <- prepareDefault dflags env case_bndr' mb_tc_app 
					 imposs_deflt_cons maybe_deflt

	; let trimmed_alts = filterOut impossible_alt alts_wo_default
	      merged_alts = mergeAlts trimmed_alts default_alts
		-- We need the mergeAlts in case the new default_alt 
		-- has turned into a constructor alternative.
		-- The merge keeps the inner DEFAULT at the front, if there is one
		-- and interleaves the alternatives in the right order

	; return (imposs_deflt_cons, merged_alts) }
    mb_tc_app = splitTyConApp_maybe (idType case_bndr')
    Just (_, inst_tys) = mb_tc_app 

    imposs_cons = case scrut of
		    Var v -> otherCons (idUnfolding v)
		    _     -> []

    impossible_alt :: CoreAlt -> Bool
    impossible_alt (con, _, _) | con `elem` imposs_cons = True
    impossible_alt (DataAlt con, _, _) = dataConCannotMatch inst_tys con
    impossible_alt _                   = False

--	1. Merge identical branches
combineIdenticalAlts :: OutId -> [InAlt] -> SimplM [InAlt]

combineIdenticalAlts case_bndr ((_con1,bndrs1,rhs1) : con_alts)
  | all isDeadBinder bndrs1,			-- Remember the default 
    length filtered_alts < length con_alts	-- alternative comes first
	-- Also Note [Dead binders]
  = do	{ tick (AltMerge case_bndr)
	; return ((DEFAULT, [], rhs1) : filtered_alts) }
    filtered_alts	 = filter keep con_alts
    keep (_con,bndrs,rhs) = not (all isDeadBinder bndrs && rhs `cheapEqExpr` rhs1)

combineIdenticalAlts _ alts = return alts

--			Prepare the default alternative
prepareDefault :: DynFlags
	       -> SimplEnv
	       -> OutId		-- Case binder; need just for its type. Note that as an
				--   OutId, it has maximum information; this is important.
				--   Test simpl013 is an example
	       -> Maybe (TyCon, [Type])	-- Type of scrutinee, decomposed
	       -> [AltCon]	-- These cons can't happen when matching the default
	       -> Maybe InExpr	-- Rhs
	       -> SimplM [InAlt]	-- Still unsimplified
					-- We use a list because it's what mergeAlts expects,
					-- And becuase case-merging can cause many to show up

-------	Merge nested cases ----------
prepareDefault dflags env outer_bndr _bndr_ty imposs_cons (Just deflt_rhs)
  | dopt Opt_CaseMerge dflags
  , Case (Var inner_scrut_var) inner_bndr _ inner_alts <- deflt_rhs
  , DoneId inner_scrut_var' <- substId env inner_scrut_var
	-- Remember, inner_scrut_var is an InId, but outer_bndr is an OutId
  , inner_scrut_var' == outer_bndr
	-- NB: the substId means that if the outer scrutinee was a 
	--     variable, and inner scrutinee is the same variable, 
	--     then inner_scrut_var' will be outer_bndr
	--     via the magic of simplCaseBinder
  = do	{ tick (CaseMerge outer_bndr)

	; let munge_rhs rhs = bindCaseBndr inner_bndr (Var outer_bndr) rhs
	; return [(con, args, munge_rhs rhs) | (con, args, rhs) <- inner_alts,
					       not (con `elem` imposs_cons) ]
		-- NB: filter out any imposs_cons.  Example:
		--	case x of 
		--	  A -> e1
		--	  DEFAULT -> case x of 
		--			A -> e2
		--			B -> e3
		-- When we merge, we must ensure that e1 takes 
		-- precedence over e2 as the value for A!  
    	-- Warning: don't call prepareAlts recursively!
    	-- Firstly, there's no point, because inner alts have already had
    	-- mkCase applied to them, so they won't have a case in their default
    	-- Secondly, if you do, you get an infinite loop, because the bindCaseBndr
    	-- in munge_rhs may put a case into the DEFAULT branch!

--------- Fill in known constructor -----------
prepareDefault _ _ case_bndr (Just (tycon, inst_tys)) imposs_cons (Just deflt_rhs)
  | 	-- This branch handles the case where we are 
	-- scrutinisng an algebraic data type
    isAlgTyCon tycon		-- It's a data type, tuple, or unboxed tuples.  
  , not (isNewTyCon tycon)	-- We can have a newtype, if we are just doing an eval:
				-- 	case x of { DEFAULT -> e }
				-- and we don't want to fill in a default for them!
  , Just all_cons <- tyConDataCons_maybe tycon
  , not (null all_cons)		-- This is a tricky corner case.  If the data type has no constructors,
				-- which GHC allows, then the case expression will have at most a default
				-- alternative.  We don't want to eliminate that alternative, because the
				-- invariant is that there's always one alternative.  It's more convenient
				-- to leave	
				--	case x of { DEFAULT -> e }     
				-- as it is, rather than transform it to
				--	error "case cant match"
				-- which would be quite legitmate.  But it's a really obscure corner, and
				-- not worth wasting code on.
  , let imposs_data_cons = [con | DataAlt con <- imposs_cons]	-- We now know it's a data type 
	impossible con  = con `elem` imposs_data_cons || dataConCannotMatch inst_tys con
  = case filterOut impossible all_cons of
	[]    -> return []	-- Eliminate the default alternative
				-- altogether if it can't match

	[con] -> 	-- It matches exactly one constructor, so fill it in
		 do { tick (FillInCaseDefault case_bndr)
                    ; us <- getUniquesM
                    ; let (ex_tvs, co_tvs, arg_ids) =
                              dataConRepInstPat us con inst_tys
                    ; return [(DataAlt con, ex_tvs ++ co_tvs ++ arg_ids, deflt_rhs)] }

	_ -> return [(DEFAULT, [], deflt_rhs)]

  | debugIsOn, isAlgTyCon tycon, not (isOpenTyCon tycon), null (tyConDataCons tycon)
	-- This can legitimately happen for type families, so don't report that
  = pprTrace "prepareDefault" (ppr case_bndr <+> ppr tycon)
        $ return [(DEFAULT, [], deflt_rhs)]

--------- Catch-all cases -----------
prepareDefault _dflags _env _case_bndr _bndr_ty _imposs_cons (Just deflt_rhs)
  = return [(DEFAULT, [], deflt_rhs)]

prepareDefault _dflags _env _case_bndr _bndr_ty _imposs_cons Nothing
  = return []	-- No default branch


mkCase tries these things

1.  Eliminate the case altogether if possible

2.  Case-identity:

	case e of 		===> e
		True  -> True;
		False -> False

    and similar friends.

mkCase :: OutExpr -> OutId -> [OutAlt]	-- Increasing order
       -> SimplM OutExpr

--	2. Identity case

mkCase scrut case_bndr alts	-- Identity case
  | all identity_alt alts
  = do tick (CaseIdentity case_bndr)
       return (re_cast scrut)
    identity_alt (con, args, rhs) = check_eq con args (de_cast rhs)

    check_eq DEFAULT       _    (Var v)   = v == case_bndr
    check_eq (LitAlt lit') _    (Lit lit) = lit == lit'
    check_eq (DataAlt con) args rhs       = rhs `cheapEqExpr` mkConApp con (arg_tys ++ varsToCoreExprs args)
					 || rhs `cheapEqExpr` Var case_bndr
    check_eq _ _ _ = False

    arg_tys = map Type (tyConAppArgs (idType case_bndr))

	-- We've seen this:
	--	case e of x { _ -> x `cast` c }
	-- And we definitely want to eliminate this case, to give
	--	e `cast` c
	-- So we throw away the cast from the RHS, and reconstruct
	-- it at the other end.  All the RHS casts must be the same
	-- if (all identity_alt alts) holds.
	-- Don't worry about nested casts, because the simplifier combines them
    de_cast (Cast e _) = e
    de_cast e	       = e

    re_cast scrut = case head alts of
			(_,_,Cast _ co) -> Cast scrut co
			_    	        -> scrut

--	Catch-all
mkCase scrut bndr alts = return (Case scrut bndr (coreAltsType alts) alts)

When adding auxiliary bindings for the case binder, it's worth checking if
its dead, because it often is, and occasionally these mkCase transformations
cascade rather nicely.

bindCaseBndr :: Id -> CoreExpr -> CoreExpr -> CoreExpr
bindCaseBndr bndr rhs body
  | isDeadBinder bndr = body
  | otherwise         = bindNonRec bndr rhs body