{- (c) The AQUA Project, Glasgow University, 1993-1998 The simplifier utilities -} {-# LANGUAGE CPP #-} module GHC.Core.Opt.Simplify.Utils ( -- Rebuilding mkLam, mkCase, prepareAlts, tryEtaExpandRhs, -- Inlining, preInlineUnconditionally, postInlineUnconditionally, activeUnfolding, activeRule, getUnfoldingInRuleMatch, simplEnvForGHCi, updModeForStableUnfoldings, updModeForRules, -- The continuation type SimplCont(..), DupFlag(..), StaticEnv, isSimplified, contIsStop, contIsDupable, contResultType, contHoleType, contHoleScaling, contIsTrivial, contArgs, countArgs, mkBoringStop, mkRhsStop, mkLazyArgStop, contIsRhsOrArg, interestingCallContext, -- ArgInfo ArgInfo(..), ArgSpec(..), mkArgInfo, addValArgTo, addCastTo, addTyArgTo, argInfoExpr, argInfoAppArgs, pushSimplifiedArgs, isStrictArgInfo, lazyArgContext, abstractFloats, -- Utilities isExitJoinId ) where #include "HsVersions.h" import GHC.Prelude import GHC.Core.Opt.Simplify.Env import GHC.Core.Opt.Monad ( SimplMode(..), Tick(..) ) import GHC.Driver.Session import GHC.Core import qualified GHC.Core.Subst import GHC.Core.Ppr import GHC.Core.TyCo.Ppr ( pprParendType ) import GHC.Core.FVs import GHC.Core.Utils import GHC.Core.Opt.Arity import GHC.Core.Unfold import GHC.Types.Name import GHC.Types.Id import GHC.Types.Id.Info import GHC.Types.Var import GHC.Types.Demand import GHC.Types.Var.Set import GHC.Types.Basic import GHC.Core.Opt.Simplify.Monad import GHC.Core.Type hiding( substTy ) import GHC.Core.Coercion hiding( substCo ) import GHC.Core.DataCon ( dataConWorkId, isNullaryRepDataCon ) import GHC.Core.Multiplicity import GHC.Utils.Misc import GHC.Data.OrdList ( isNilOL ) import GHC.Utils.Monad import GHC.Utils.Outputable import GHC.Core.Opt.ConstantFold import GHC.Data.FastString ( fsLit ) import Control.Monad ( when ) import Data.List ( sortBy ) {- ************************************************************************ * * The SimplCont and DupFlag types * * ************************************************************************ 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 -- Stop[e] = e OutType -- Type of the <hole> CallCtxt -- Tells if 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 -- Never ValAppCxt (use ApplyToVal instead) -- or CaseCtxt (use Select instead) | CastIt -- (CastIt co K)[e] = K[ e `cast` co ] OutCoercion -- The coercion simplified -- Invariant: never an identity coercion SimplCont | ApplyToVal -- (ApplyToVal arg K)[e] = K[ e arg ] { sc_dup :: DupFlag -- See Note [DupFlag invariants] , sc_hole_ty :: OutType -- Type of the function, presumably (forall a. blah) -- See Note [The hole type in ApplyToTy/Val] , sc_arg :: InExpr -- The argument, , sc_env :: StaticEnv -- see Note [StaticEnv invariant] , sc_cont :: SimplCont } | ApplyToTy -- (ApplyToTy ty K)[e] = K[ e ty ] { sc_arg_ty :: OutType -- Argument type , sc_hole_ty :: OutType -- Type of the function, presumably (forall a. blah) -- See Note [The hole type in ApplyToTy/Val] , sc_cont :: SimplCont } | Select -- (Select alts K)[e] = K[ case e of alts ] { sc_dup :: DupFlag -- See Note [DupFlag invariants] , sc_bndr :: InId -- case binder , sc_alts :: [InAlt] -- Alternatives , sc_env :: StaticEnv -- See Note [StaticEnv invariant] , sc_cont :: SimplCont } -- The two strict forms have no DupFlag, because we never duplicate them | StrictBind -- (StrictBind x xs b K)[e] = let x = e in K[\xs.b] -- or, equivalently, = K[ (\x xs.b) e ] { sc_dup :: DupFlag -- See Note [DupFlag invariants] , sc_bndr :: InId , sc_bndrs :: [InBndr] , sc_body :: InExpr , sc_env :: StaticEnv -- See Note [StaticEnv invariant] , sc_cont :: SimplCont } | StrictArg -- (StrictArg (f e1 ..en) K)[e] = K[ f e1 .. en e ] { sc_dup :: DupFlag -- Always Simplified or OkToDup , sc_fun :: ArgInfo -- Specifies f, e1..en, Whether f has rules, etc -- plus demands and discount flags for *this* arg -- and further args -- So ai_dmds and ai_discs are never empty , sc_fun_ty :: OutType -- Type of the function (f e1 .. en), -- presumably (arg_ty -> res_ty) -- where res_ty is expected by sc_cont , sc_cont :: SimplCont } | TickIt -- (TickIt t K)[e] = K[ tick t e ] (Tickish Id) -- Tick tickish <hole> SimplCont type StaticEnv = SimplEnv -- Just the static part is relevant data DupFlag = NoDup -- Unsimplified, might be big | Simplified -- Simplified | OkToDup -- Simplified and small isSimplified :: DupFlag -> Bool isSimplified NoDup = False isSimplified _ = True -- Invariant: the subst-env is empty perhapsSubstTy :: DupFlag -> StaticEnv -> Type -> Type perhapsSubstTy dup env ty | isSimplified dup = ty | otherwise = substTy env ty {- Note [StaticEnv invariant] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We pair up an InExpr or InAlts with a StaticEnv, which establishes the lexical scope for that InExpr. When we simplify that InExpr/InAlts, we use - Its captured StaticEnv - Overriding its InScopeSet with the larger one at the simplification point. Why override the InScopeSet? Example: (let y = ey in f) ex By the time we simplify ex, 'y' will be in scope. However the InScopeSet in the StaticEnv is not irrelevant: it should include all the free vars of applying the substitution to the InExpr. Reason: contHoleType uses perhapsSubstTy to apply the substitution to the expression, and that (rightly) gives ASSERT failures if the InScopeSet isn't big enough. Note [DupFlag invariants] ~~~~~~~~~~~~~~~~~~~~~~~~~ In both (ApplyToVal dup _ env k) and (Select dup _ _ env k) the following invariants hold (a) if dup = OkToDup, then continuation k is also ok-to-dup (b) if dup = OkToDup or Simplified, the subst-env is empty (and hence no need to re-simplify) -} instance Outputable DupFlag where ppr OkToDup = text "ok" ppr NoDup = text "nodup" ppr Simplified = text "simpl" instance Outputable SimplCont where ppr (Stop ty interesting) = text "Stop" <> brackets (ppr interesting) <+> ppr ty ppr (CastIt co cont ) = (text "CastIt" <+> pprOptCo co) $$ ppr cont ppr (TickIt t cont) = (text "TickIt" <+> ppr t) $$ ppr cont ppr (ApplyToTy { sc_arg_ty = ty, sc_cont = cont }) = (text "ApplyToTy" <+> pprParendType ty) $$ ppr cont ppr (ApplyToVal { sc_arg = arg, sc_dup = dup, sc_cont = cont, sc_hole_ty = hole_ty }) = (hang (text "ApplyToVal" <+> ppr dup <+> text "hole" <+> ppr hole_ty) 2 (pprParendExpr arg)) $$ ppr cont ppr (StrictBind { sc_bndr = b, sc_cont = cont }) = (text "StrictBind" <+> ppr b) $$ ppr cont ppr (StrictArg { sc_fun = ai, sc_cont = cont }) = (text "StrictArg" <+> ppr (ai_fun ai)) $$ ppr cont ppr (Select { sc_dup = dup, sc_bndr = bndr, sc_alts = alts, sc_env = se, sc_cont = cont }) = (text "Select" <+> ppr dup <+> ppr bndr) $$ whenPprDebug (nest 2 $ vcat [ppr (seTvSubst se), ppr alts]) $$ ppr cont {- Note [The hole type in ApplyToTy] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The sc_hole_ty field of ApplyToTy records the type of the "hole" in the continuation. It is absolutely necessary to compute contHoleType, but it is not used for anything else (and hence may not be evaluated). Why is it necessary for contHoleType? Consider the continuation ApplyToType Int (Stop Int) corresponding to (<hole> @Int) :: Int What is the type of <hole>? It could be (forall a. Int) or (forall a. a), and there is no way to know which, so we must record it. In a chain of applications (f @t1 @t2 @t3) we'll lazily compute exprType for (f @t1) and (f @t1 @t2), which is potentially non-linear; but it probably doesn't matter because we'll never compute them all. ************************************************************************ * * ArgInfo and ArgSpec * * ************************************************************************ -} data ArgInfo = ArgInfo { ai_fun :: OutId, -- The function ai_args :: [ArgSpec], -- ...applied to these args (which are in *reverse* order) ai_rules :: FunRules, -- Rules for this function ai_encl :: Bool, -- Flag saying whether this function -- or an enclosing one has rules (recursively) -- True => be keener to inline in all args ai_dmds :: [Demand], -- Demands on remaining value arguments (beyond ai_args) -- 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 remaining value arguments (beyong ai_args) -- non-zero => be keener to inline -- Always infinite } data ArgSpec = ValArg { as_dmd :: Demand -- Demand placed on this argument , as_arg :: OutExpr -- Apply to this (coercion or value); c.f. ApplyToVal , as_hole_ty :: OutType } -- Type of the function (presumably t1 -> t2) | TyArg { as_arg_ty :: OutType -- Apply to this type; c.f. ApplyToTy , as_hole_ty :: OutType } -- Type of the function (presumably forall a. blah) | CastBy OutCoercion -- Cast by this; c.f. CastIt instance Outputable ArgInfo where ppr (ArgInfo { ai_fun = fun, ai_args = args, ai_dmds = dmds }) = text "ArgInfo" <+> braces (sep [ text "fun =" <+> ppr fun , text "dmds =" <+> ppr dmds , text "args =" <+> ppr args ]) instance Outputable ArgSpec where ppr (ValArg { as_arg = arg }) = text "ValArg" <+> ppr arg ppr (TyArg { as_arg_ty = ty }) = text "TyArg" <+> ppr ty ppr (CastBy c) = text "CastBy" <+> ppr c addValArgTo :: ArgInfo -> OutExpr -> OutType -> ArgInfo addValArgTo ai arg hole_ty | ArgInfo { ai_dmds = dmd:dmds, ai_discs = _:discs, ai_rules = rules } <- ai -- Pop the top demand and and discounts off , let arg_spec = ValArg { as_arg = arg, as_hole_ty = hole_ty, as_dmd = dmd } = ai { ai_args = arg_spec : ai_args ai , ai_dmds = dmds , ai_discs = discs , ai_rules = decRules rules } | otherwise = pprPanic "addValArgTo" (ppr ai $$ ppr arg) -- There should always be enough demands and discounts addTyArgTo :: ArgInfo -> OutType -> OutType -> ArgInfo addTyArgTo ai arg_ty hole_ty = ai { ai_args = arg_spec : ai_args ai , ai_rules = decRules (ai_rules ai) } where arg_spec = TyArg { as_arg_ty = arg_ty, as_hole_ty = hole_ty } addCastTo :: ArgInfo -> OutCoercion -> ArgInfo addCastTo ai co = ai { ai_args = CastBy co : ai_args ai } isStrictArgInfo :: ArgInfo -> Bool -- True if the function is strict in the next argument isStrictArgInfo (ArgInfo { ai_dmds = dmds }) | dmd:_ <- dmds = isStrictDmd dmd | otherwise = False argInfoAppArgs :: [ArgSpec] -> [OutExpr] argInfoAppArgs [] = [] argInfoAppArgs (CastBy {} : _) = [] -- Stop at a cast argInfoAppArgs (ValArg { as_arg = arg } : as) = arg : argInfoAppArgs as argInfoAppArgs (TyArg { as_arg_ty = ty } : as) = Type ty : argInfoAppArgs as pushSimplifiedArgs :: SimplEnv -> [ArgSpec] -> SimplCont -> SimplCont pushSimplifiedArgs _env [] k = k pushSimplifiedArgs env (arg : args) k = case arg of TyArg { as_arg_ty = arg_ty, as_hole_ty = hole_ty } -> ApplyToTy { sc_arg_ty = arg_ty, sc_hole_ty = hole_ty, sc_cont = rest } ValArg { as_arg = arg, as_hole_ty = hole_ty } -> ApplyToVal { sc_arg = arg, sc_env = env, sc_dup = Simplified , sc_hole_ty = hole_ty, sc_cont = rest } CastBy c -> CastIt c rest where rest = pushSimplifiedArgs env args k -- The env has an empty SubstEnv argInfoExpr :: OutId -> [ArgSpec] -> OutExpr -- NB: the [ArgSpec] is reversed so that the first arg -- in the list is the last one in the application argInfoExpr fun rev_args = go rev_args where go [] = Var fun go (ValArg { as_arg = arg } : as) = go as `App` arg go (TyArg { as_arg_ty = ty } : as) = go as `App` Type ty go (CastBy co : as) = mkCast (go as) co type FunRules = Maybe (Int, [CoreRule]) -- Remaining rules for this function -- Nothing => No rules -- Just (n, rules) => some rules, requiring at least n more type/value args decRules :: FunRules -> FunRules decRules (Just (n, rules)) = Just (n-1, rules) decRules Nothing = Nothing mkFunRules :: [CoreRule] -> FunRules mkFunRules [] = Nothing mkFunRules rs = Just (n_required, rs) where n_required = maximum (map ruleArity rs) {- ************************************************************************ * * Functions on SimplCont * * ************************************************************************ -} mkBoringStop :: OutType -> SimplCont mkBoringStop ty = Stop ty BoringCtxt mkRhsStop :: OutType -> SimplCont -- See Note [RHS of lets] in GHC.Core.Unfold mkRhsStop ty = Stop ty RhsCtxt mkLazyArgStop :: OutType -> CallCtxt -> SimplCont mkLazyArgStop ty cci = Stop ty cci ------------------- contIsRhsOrArg :: SimplCont -> Bool contIsRhsOrArg (Stop {}) = True contIsRhsOrArg (StrictBind {}) = True contIsRhsOrArg (StrictArg {}) = True contIsRhsOrArg _ = False contIsRhs :: SimplCont -> Bool contIsRhs (Stop _ RhsCtxt) = True contIsRhs _ = False ------------------- contIsStop :: SimplCont -> Bool contIsStop (Stop {}) = True contIsStop _ = False contIsDupable :: SimplCont -> Bool contIsDupable (Stop {}) = True contIsDupable (ApplyToTy { sc_cont = k }) = contIsDupable k contIsDupable (ApplyToVal { sc_dup = OkToDup }) = True -- See Note [DupFlag invariants] contIsDupable (Select { sc_dup = OkToDup }) = True -- ...ditto... contIsDupable (StrictArg { sc_dup = OkToDup }) = True -- ...ditto... contIsDupable (CastIt _ k) = contIsDupable k contIsDupable _ = False ------------------- contIsTrivial :: SimplCont -> Bool contIsTrivial (Stop {}) = True contIsTrivial (ApplyToTy { sc_cont = k }) = contIsTrivial k contIsTrivial (ApplyToVal { sc_arg = Coercion _, sc_cont = k }) = contIsTrivial k contIsTrivial (CastIt _ k) = contIsTrivial k contIsTrivial _ = False ------------------- contResultType :: SimplCont -> OutType contResultType (Stop ty _) = ty contResultType (CastIt _ k) = contResultType k contResultType (StrictBind { sc_cont = k }) = contResultType k contResultType (StrictArg { sc_cont = k }) = contResultType k contResultType (Select { sc_cont = k }) = contResultType k contResultType (ApplyToTy { sc_cont = k }) = contResultType k contResultType (ApplyToVal { sc_cont = k }) = contResultType k contResultType (TickIt _ k) = contResultType k contHoleType :: SimplCont -> OutType contHoleType (Stop ty _) = ty contHoleType (TickIt _ k) = contHoleType k contHoleType (CastIt co _) = coercionLKind co contHoleType (StrictBind { sc_bndr = b, sc_dup = dup, sc_env = se }) = perhapsSubstTy dup se (idType b) contHoleType (StrictArg { sc_fun_ty = ty }) = funArgTy ty contHoleType (ApplyToTy { sc_hole_ty = ty }) = ty -- See Note [The hole type in ApplyToTy] contHoleType (ApplyToVal { sc_hole_ty = ty }) = ty -- See Note [The hole type in ApplyToTy/Val] contHoleType (Select { sc_dup = d, sc_bndr = b, sc_env = se }) = perhapsSubstTy d se (idType b) -- Computes the multiplicity scaling factor at the hole. That is, in (case [] of -- x ::(p) _ { … }) (respectively for arguments of functions), the scaling -- factor is p. And in E[G[]], the scaling factor is the product of the scaling -- factor of E and that of G. -- -- The scaling factor at the hole of E[] is used to determine how a binder -- should be scaled if it commutes with E. This appears, in particular, in the -- case-of-case transformation. contHoleScaling :: SimplCont -> Mult contHoleScaling (Stop _ _) = One contHoleScaling (CastIt _ k) = contHoleScaling k contHoleScaling (StrictBind { sc_bndr = id, sc_cont = k }) = idMult id `mkMultMul` contHoleScaling k contHoleScaling (Select { sc_bndr = id, sc_cont = k }) = idMult id `mkMultMul` contHoleScaling k contHoleScaling (StrictArg { sc_fun_ty = fun_ty, sc_cont = k }) = w `mkMultMul` contHoleScaling k where (w, _, _) = splitFunTy fun_ty contHoleScaling (ApplyToTy { sc_cont = k }) = contHoleScaling k contHoleScaling (ApplyToVal { sc_cont = k }) = contHoleScaling k contHoleScaling (TickIt _ k) = contHoleScaling k ------------------- countArgs :: SimplCont -> Int -- Count all arguments, including types, coercions, and other values countArgs (ApplyToTy { sc_cont = cont }) = 1 + countArgs cont countArgs (ApplyToVal { sc_cont = cont }) = 1 + countArgs cont countArgs _ = 0 contArgs :: SimplCont -> (Bool, [ArgSummary], SimplCont) -- Summarises value args, discards type args and coercions -- The returned continuation of the call is only used to -- answer questions like "are you interesting?" contArgs cont | lone cont = (True, [], cont) | otherwise = go [] cont where lone (ApplyToTy {}) = False -- See Note [Lone variables] in GHC.Core.Unfold lone (ApplyToVal {}) = False -- NB: even a type application or cast lone (CastIt {}) = False -- stops it being "lone" lone _ = True go args (ApplyToVal { sc_arg = arg, sc_env = se, sc_cont = k }) = go (is_interesting arg se : args) k go args (ApplyToTy { sc_cont = k }) = go args k go args (CastIt _ k) = go args k go args k = (False, reverse args, k) is_interesting arg se = interestingArg se arg -- Do *not* use short-cutting substitution here -- because we want to get as much IdInfo as possible ------------------- mkArgInfo :: SimplEnv -> Id -> [CoreRule] -- Rules for function -> Int -- Number of value args -> SimplCont -- Context of the call -> ArgInfo mkArgInfo env fun rules n_val_args call_cont | n_val_args < idArity fun -- Note [Unsaturated functions] = ArgInfo { ai_fun = fun, ai_args = [] , ai_rules = fun_rules , ai_encl = False , ai_dmds = vanilla_dmds , ai_discs = vanilla_discounts } | otherwise = ArgInfo { ai_fun = fun , ai_args = [] , ai_rules = fun_rules , ai_encl = interestingArgContext rules call_cont , ai_dmds = add_type_strictness (idType fun) arg_dmds , ai_discs = arg_discounts } where fun_rules = mkFunRules rules vanilla_discounts, arg_discounts :: [Int] vanilla_discounts = repeat 0 arg_discounts = case idUnfolding fun of CoreUnfolding {uf_guidance = UnfIfGoodArgs {ug_args = discounts}} -> discounts ++ vanilla_discounts _ -> vanilla_discounts vanilla_dmds, arg_dmds :: [Demand] vanilla_dmds = repeat topDmd arg_dmds | not (sm_inline (seMode env)) = vanilla_dmds -- See Note [Do not expose strictness if sm_inline=False] | otherwise = -- add_type_str fun_ty $ case splitStrictSig (idStrictness 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 GHC.Core.Op.Simplify.Utils.analyseCont) if isDeadEndDiv result_info then demands -- Finite => result is bottom else demands ++ vanilla_dmds | otherwise -> WARN( True, text "More demands than arity" <+> ppr fun <+> ppr (idArity fun) <+> ppr n_val_args <+> ppr demands ) vanilla_dmds -- Not enough args, or no strictness add_type_strictness :: Type -> [Demand] -> [Demand] -- 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_strictness is done repeatedly (for each call); -- might be better once-for-all in the function -- But beware primops/datacons with no strictness add_type_strictness fun_ty dmds | null dmds = [] | Just (_, fun_ty') <- splitForAllTy_maybe fun_ty = add_type_strictness fun_ty' dmds -- Look through foralls | Just (_, arg_ty, fun_ty') <- splitFunTy_maybe fun_ty -- Add strict-type info , dmd : rest_dmds <- dmds , let dmd' = case isLiftedType_maybe arg_ty of Just False -> strictenDmd dmd _ -> dmd = dmd' : add_type_strictness fun_ty' rest_dmds -- If the type is levity-polymorphic, we can't know whether it's -- strict. isLiftedType_maybe will return Just False only when -- we're sure the type is unlifted. | otherwise = dmds {- 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 Note [Do not expose strictness if sm_inline=False] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ #15163 showed a case in which we had {-# INLINE [1] zip #-} zip = undefined {-# RULES "foo" forall as bs. stream (zip as bs) = ..blah... #-} If we expose zip's bottoming nature when simplifying the LHS of the RULE we get {-# RULES "foo" forall as bs. stream (case zip of {}) = ..blah... #-} discarding the arguments to zip. Usually this is fine, but on the LHS of a rule it's not, because 'as' and 'bs' are now not bound on the LHS. This is a pretty pathological example, so I'm not losing sleep over it, but the simplest solution was to check sm_inline; if it is False, which it is on the LHS of a rule (see updModeForRules), then don't make use of the strictness info for the function. -} {- ************************************************************************ * * Interesting arguments * * ************************************************************************ 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. Note [No case of case is boring] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ If we see case f x of <alts> we'd usually treat the context as interesting, to encourage 'f' to inline. But if case-of-case is off, it's really not so interesting after all, because we are unlikely to be able to push the case expression into the branches of any case in f's unfolding. So, to reduce unnecessary code expansion, we just make the context look boring. This made a small compile-time perf improvement in perf/compiler/T6048, and it looks plausible to me. -} lazyArgContext :: ArgInfo -> CallCtxt -- Use this for lazy arguments lazyArgContext (ArgInfo { ai_encl = encl_rules, ai_discs = discs }) | encl_rules = RuleArgCtxt | disc:_ <- discs, disc > 0 = DiscArgCtxt -- Be keener here | otherwise = BoringCtxt -- Nothing interesting strictArgContext :: ArgInfo -> CallCtxt strictArgContext (ArgInfo { ai_encl = encl_rules, ai_discs = discs }) -- Use this for strict arguments | encl_rules = RuleArgCtxt | disc:_ <- discs, disc > 0 = DiscArgCtxt -- Be keener here | otherwise = RhsCtxt -- Why RhsCtxt? if we see f (g x) (h x), and f is strict, we -- want to be a bit more eager to inline g, because it may -- expose an eval (on x perhaps) that can be eliminated or -- shared. I saw this in nofib 'boyer2', RewriteFuns.onewayunify1 -- It's worth an 18% improvement in allocation for this -- particular benchmark; 5% on 'mate' and 1.3% on 'multiplier' interestingCallContext :: SimplEnv -> SimplCont -> CallCtxt -- See Note [Interesting call context] interestingCallContext env cont = interesting cont where interesting (Select {}) | sm_case_case (getMode env) = CaseCtxt | otherwise = BoringCtxt -- See Note [No case of case is boring] interesting (ApplyToVal {}) = 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 GHC.Core.Unfold interesting (StrictArg { sc_fun = fun }) = strictArgContext fun interesting (StrictBind {}) = BoringCtxt interesting (Stop _ cci) = cci interesting (TickIt _ k) = interesting k interesting (ApplyToTy { sc_cont = k }) = interesting k interesting (CastIt _ k) = interesting k -- 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. interestingArgContext :: [CoreRule] -> 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 ai-rules 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 -- -- The call_cont passed to interestingArgContext is the context of -- the call itself, e.g. g <hole> in the example above interestingArgContext rules call_cont = notNull rules || enclosing_fn_has_rules where enclosing_fn_has_rules = go call_cont go (Select {}) = False go (ApplyToVal {}) = False -- Shouldn't really happen go (ApplyToTy {}) = False -- Ditto go (StrictArg { sc_fun = fun }) = ai_encl fun go (StrictBind {}) = False -- ?? go (CastIt _ c) = go c go (Stop _ RuleArgCtxt) = True go (Stop _ _) = False go (TickIt _ c) = go c {- Note [Interesting arguments] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ An argument is interesting if it deserves a discount for unfoldings with a discount in that argument position. The idea is to avoid unfolding a function that is applied only to variables that have no unfolding (i.e. they are probably lambda bound): f x y z There is little point in inlining f here. Generally, *values* (like (C a b) and (\x.e)) deserve discounts. But we must look through lets, eg (let x = e in C a b), because the let will float, exposing the value, if we inline. That makes it different to exprIsHNF. Before 2009 we said it was interesting if the argument had *any* structure at all; i.e. (hasSomeUnfolding v). But does too much inlining; see #3016. But we don't regard (f x y) as interesting, unless f is unsaturated. If it's saturated and f hasn't inlined, then it's probably not going to now! Note [Conlike is interesting] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider f d = ...((*) d x y)... ... f (df d')... where df is con-like. Then we'd really like to inline 'f' so that the rule for (*) (df d) can fire. To do this a) we give a discount for being an argument of a class-op (eg (*) d) b) we say that a con-like argument (eg (df d)) is interesting -} interestingArg :: SimplEnv -> CoreExpr -> ArgSummary -- See Note [Interesting arguments] interestingArg env e = go env 0 e where -- n is # value args to which the expression is applied go env n (Var v) = case substId env v of DoneId v' -> go_var n v' DoneEx e _ -> go (zapSubstEnv env) n e ContEx tvs cvs ids e -> go (setSubstEnv env tvs cvs ids) n e go _ _ (Lit {}) = ValueArg go _ _ (Type _) = TrivArg go _ _ (Coercion _) = TrivArg go env n (App fn (Type _)) = go env n fn go env n (App fn _) = go env (n+1) fn go env n (Tick _ a) = go env n a go env n (Cast e _) = go env n e go env n (Lam v e) | isTyVar v = go env n e | n>0 = NonTrivArg -- (\x.b) e is NonTriv | otherwise = ValueArg go _ _ (Case {}) = NonTrivArg go env n (Let b e) = case go env' n e of ValueArg -> ValueArg _ -> NonTrivArg where env' = env `addNewInScopeIds` bindersOf b go_var n v | isConLikeId v = ValueArg -- Experimenting with 'conlike' rather that -- data constructors here | idArity v > n = ValueArg -- Catches (eg) primops with arity but no unfolding | n > 0 = NonTrivArg -- Saturated or unknown call | conlike_unfolding = ValueArg -- n==0; look for an interesting unfolding -- See Note [Conlike is interesting] | otherwise = TrivArg -- n==0, no useful unfolding where conlike_unfolding = isConLikeUnfolding (idUnfolding v) {- ************************************************************************ * * SimplMode * * ************************************************************************ The SimplMode controls several switches; see its definition in GHC.Core.Opt.Monad sm_rules :: Bool -- Whether RULES are enabled sm_inline :: Bool -- Whether inlining is enabled sm_case_case :: Bool -- Whether case-of-case is enabled sm_eta_expand :: Bool -- Whether eta-expansion is enabled -} simplEnvForGHCi :: DynFlags -> SimplEnv simplEnvForGHCi dflags = mkSimplEnv $ SimplMode { sm_names = ["GHCi"] , sm_phase = InitialPhase , sm_dflags = dflags , sm_rules = rules_on , sm_inline = False , sm_eta_expand = eta_expand_on , sm_case_case = True } where rules_on = gopt Opt_EnableRewriteRules dflags eta_expand_on = gopt Opt_DoLambdaEtaExpansion dflags -- Do not do any inlining, in case we expose some unboxed -- tuple stuff that confuses the bytecode interpreter updModeForStableUnfoldings :: Activation -> SimplMode -> SimplMode -- See Note [Simplifying inside stable unfoldings] updModeForStableUnfoldings inline_rule_act current_mode = current_mode { sm_phase = phaseFromActivation inline_rule_act , sm_inline = True , sm_eta_expand = False } -- sm_eta_expand: see Note [No eta expansion in stable unfoldings] -- For sm_rules, just inherit; sm_rules might be "off" -- because of -fno-enable-rewrite-rules where phaseFromActivation (ActiveAfter _ n) = Phase n phaseFromActivation _ = InitialPhase updModeForRules :: SimplMode -> SimplMode -- See Note [Simplifying rules] updModeForRules current_mode = current_mode { sm_phase = InitialPhase , sm_inline = False -- See Note [Do not expose strictness if sm_inline=False] , sm_rules = False , sm_eta_expand = False } {- Note [Simplifying rules] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ When simplifying a rule LHS, refrain from /any/ inlining or applying of other RULES. Doing anything to the LHS is plain confusing, because it means that what the rule matches is not what the user wrote. c.f. #10595, and #10528. Moreover, inlining (or applying rules) on rule LHSs risks introducing Ticks into the LHS, which makes matching trickier. #10665, #10745. Doing this to either side confounds tools like HERMIT, which seek to reason about and apply the RULES as originally written. See #10829. There is, however, one case where we are pretty much /forced/ to transform the LHS of a rule: postInlineUnconditionally. For instance, in the case of let f = g @Int in f We very much want to inline f into the body of the let. However, to do so (and be able to safely drop f's binding) we must inline into all occurrences of f, including those in the LHS of rules. This can cause somewhat surprising results; for instance, in #18162 we found that a rule template contained ticks in its arguments, because postInlineUnconditionally substituted in a trivial expression that contains ticks. See Note [Tick annotations in RULE matching] in GHC.Core.Rules for details. Note [No eta expansion in stable unfoldings] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ If we have a stable unfolding f :: Ord a => a -> IO () -- Unfolding template -- = /\a \(d:Ord a) (x:a). bla we do not want to eta-expand to f :: Ord a => a -> IO () -- Unfolding template -- = (/\a \(d:Ord a) (x:a) (eta:State#). bla eta) |> co because not specialisation of the overloading doesn't work properly (see Note [Specialisation shape] in GHC.Core.Opt.Specialise), #9509. So we disable eta-expansion in stable unfoldings. Note [Inlining in gentle mode] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Something is inlined if (i) the sm_inline flag is on, AND (ii) the thing has an INLINE pragma, AND (iii) the thing is inlinable in the earliest phase. Example of why (iii) is important: {-# INLINE [~1] g #-} g = ... {-# INLINE f #-} f x = g (g x) If we were to inline g into f's inlining, then an importing module would never be able to do f e --> g (g e) ---> RULE fires because the stable unfolding for f has had g inlined into it. On the other hand, it is bad not to do ANY inlining into an stable unfolding, because then recursive knots in instance declarations don't get unravelled. However, *sometimes* SimplGently must do no call-site inlining at all (hence sm_inline = False). 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. Note [Simplifying inside stable unfoldings] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We must take care with simplification inside stable unfoldings (which come from INLINE pragmas). First, consider the following example let f = \pq -> BIG in let g = \y -> f y y {-# INLINE g #-} in ...g...g...g...g...g... Now, if that's the ONLY occurrence of f, it might be inlined inside g, and thence copied multiple times when g is inlined. HENCE we treat any occurrence in a stable unfolding as a multiple occurrence, not a single one; see OccurAnal.addRuleUsage. Second, we do want *do* to some modest rules/inlining stuff in stable unfoldings, partly to eliminate senseless crap, and partly to break the recursive knots generated by instance declarations. However, suppose we have {-# INLINE <act> f #-} f = <rhs> meaning "inline f in phases p where activation <act>(p) holds". Then what inlinings/rules can we apply to the copy of <rhs> captured in f's stable unfolding? Our model is that literally <rhs> is substituted for f when it is inlined. So our conservative plan (implemented by updModeForStableUnfoldings) is this: ------------------------------------------------------------- When simplifying the RHS of a stable unfolding, set the phase to the phase in which the stable unfolding first becomes active ------------------------------------------------------------- That ensures that a) Rules/inlinings that *cease* being active before p will not apply to the stable unfolding, consistent with it being inlined in its *original* form in phase p. b) Rules/inlinings that only become active *after* p will not apply to the stable unfolding, again to be consistent with inlining the *original* rhs in phase p. For example, {-# INLINE f #-} f x = ...g... {-# NOINLINE [1] g #-} g y = ... {-# RULE h g = ... #-} Here we must not inline g into f's RHS, even when we get to phase 0, because when f is later inlined into some other module we want the rule for h to fire. Similarly, consider {-# INLINE f #-} f x = ...g... g y = ... and suppose that there are auto-generated specialisations and a strictness wrapper for g. The specialisations get activation AlwaysActive, and the strictness wrapper get activation (ActiveAfter 0). So the strictness wrepper fails the test and won't be inlined into f's stable unfolding. That means f can inline, expose the specialised call to g, so the specialisation rules can fire. A note about wrappers ~~~~~~~~~~~~~~~~~~~~~ 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. -} activeUnfolding :: SimplMode -> Id -> Bool activeUnfolding mode id | isCompulsoryUnfolding (realIdUnfolding id) = True -- Even sm_inline can't override compulsory unfoldings | otherwise = isActive (sm_phase mode) (idInlineActivation id) && sm_inline mode -- `or` isStableUnfolding (realIdUnfolding id) -- Inline things when -- (a) they are active -- (b) sm_inline says so, except that for stable unfoldings -- (ie pragmas) we inline anyway getUnfoldingInRuleMatch :: SimplEnv -> InScopeEnv -- When matching in RULE, we want to "look through" an unfolding -- (to see a constructor) if *rules* are on, even if *inlinings* -- are not. A notable example is DFuns, which really we want to -- match in rules like (op dfun) in gentle mode. Another example -- is 'otherwise' which we want exprIsConApp_maybe to be able to -- see very early on getUnfoldingInRuleMatch env = (in_scope, id_unf) where in_scope = seInScope env mode = getMode env id_unf id | unf_is_active id = idUnfolding id | otherwise = NoUnfolding unf_is_active id | not (sm_rules mode) = -- active_unfolding_minimal id isStableUnfolding (realIdUnfolding id) -- Do we even need to test this? I think this InScopeEnv -- is only consulted if activeRule returns True, which -- never happens if sm_rules is False | otherwise = isActive (sm_phase mode) (idInlineActivation id) ---------------------- activeRule :: SimplMode -> Activation -> Bool -- Nothing => No rules at all activeRule mode | not (sm_rules mode) = \_ -> False -- Rewriting is off | otherwise = isActive (sm_phase mode) {- ************************************************************************ * * preInlineUnconditionally * * ************************************************************************ preInlineUnconditionally ~~~~~~~~~~~~~~~~~~~~~~~~ @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] ...etc... 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] ...etc... THE MAIN INVARIANT is this: ---- 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 occurrence 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 FinalPhase (the last phase), at which point don't. Note [pre/postInlineUnconditionally in gentle mode] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Even in gentle mode we want to do preInlineUnconditionally. 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 However, as usual for Gentle mode, do not inline things that are inactive in the initial stages. See Note [Gentle mode]. Note [Stable unfoldings and preInlineUnconditionally] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Surprisingly, do not pre-inline-unconditionally Ids with INLINE pragmas! Example {-# INLINE f #-} f :: Eq a => a -> a f x = ... fInt :: Int -> Int fInt = f Int dEqInt ...fInt...fInt...fInt... Here f occurs just once, in the RHS of fInt. But if we inline it there it might make fInt look big, and we'll lose the opportunity to inline f at each of fInt's call sites. The INLINE pragma will only inline when the application is saturated for exactly this reason; and we don't want PreInlineUnconditionally to second-guess it. A live example is #3736. c.f. Note [Stable unfoldings and postInlineUnconditionally] NB: if the pragma is INLINEABLE, then we don't want to behave in this special way -- an INLINEABLE pragma just says to GHC "inline this if you like". But if there is a unique occurrence, we want to inline the stable unfolding, not the RHS. Note [Top-level bottoming Ids] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Don't inline top-level Ids that are bottoming, even if they are used just once, because FloatOut has gone to some trouble to extract them out. Inlining them won't make the program run faster! Note [Do not inline CoVars unconditionally] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Coercion variables appear inside coercions, and the RHS of a let-binding is a term (not a coercion) so we can't necessarily inline the latter in the former. -} preInlineUnconditionally :: SimplEnv -> TopLevelFlag -> InId -> InExpr -> StaticEnv -- These two go together -> Maybe SimplEnv -- Returned env has extended substitution -- Precondition: rhs satisfies the let/app invariant -- See Note [Core let/app invariant] in GHC.Core -- Reason: we don't want to inline single uses, or discard dead bindings, -- for unlifted, side-effect-ful bindings preInlineUnconditionally env top_lvl bndr rhs rhs_env | not pre_inline_unconditionally = Nothing | not active = Nothing | isTopLevel top_lvl && isDeadEndId bndr = Nothing -- Note [Top-level bottoming Ids] | isCoVar bndr = Nothing -- Note [Do not inline CoVars unconditionally] | isExitJoinId bndr = Nothing -- Note [Do not inline exit join points] -- in module Exitify | not (one_occ (idOccInfo bndr)) = Nothing | not (isStableUnfolding unf) = Just (extend_subst_with rhs) -- Note [Stable unfoldings and preInlineUnconditionally] | isInlinablePragma inline_prag , Just inl <- maybeUnfoldingTemplate unf = Just (extend_subst_with inl) | otherwise = Nothing where unf = idUnfolding bndr extend_subst_with inl_rhs = extendIdSubst env bndr (mkContEx rhs_env inl_rhs) one_occ IAmDead = True -- Happens in ((\x.1) v) one_occ OneOcc{ occ_n_br = 1 , occ_in_lam = NotInsideLam } = isNotTopLevel top_lvl || early_phase one_occ OneOcc{ occ_n_br = 1 , occ_in_lam = IsInsideLam , occ_int_cxt = IsInteresting } = canInlineInLam rhs one_occ _ = False pre_inline_unconditionally = gopt Opt_SimplPreInlining (seDynFlags env) mode = getMode env active = isActive (sm_phase mode) (inlinePragmaActivation inline_prag) -- See Note [pre/postInlineUnconditionally in gentle mode] inline_prag = idInlinePragma bndr -- Be very careful before inlining inside a lambda, because (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 (Tick t e) = not (tickishIsCode t) && canInlineInLam e canInlineInLam _ = False -- not ticks. Counting ticks cannot be duplicated, and non-counting -- ticks around a Lam will disappear anyway. early_phase = sm_phase mode /= FinalPhase -- 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 * * ************************************************************************ postInlineUnconditionally ~~~~~~~~~~~~~~~~~~~~~~~~~ @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 occurrence(s) Note that we do this unconditional inlining only for trivial 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. NB: unconditional inlining of this sort can introduce ticks in places that may seem surprising; for instance, the LHS of rules. See Note [Simplfying rules] for details. -} postInlineUnconditionally :: SimplEnv -> TopLevelFlag -> OutId -- The binder (*not* a CoVar), including its unfolding -> OccInfo -- From the InId -> OutExpr -> Bool -- Precondition: rhs satisfies the let/app invariant -- See Note [Core let/app invariant] in GHC.Core -- Reason: we don't want to inline single uses, or discard dead bindings, -- for unlifted, side-effect-ful bindings postInlineUnconditionally env top_lvl bndr occ_info rhs | not active = False | isWeakLoopBreaker occ_info = False -- If it's a loop-breaker of any kind, don't inline -- because it might be referred to "earlier" | isStableUnfolding unfolding = False -- Note [Stable unfoldings and postInlineUnconditionally] | isTopLevel top_lvl = False -- Note [Top level and postInlineUnconditionally] | exprIsTrivial rhs = True | otherwise = case occ_info of OneOcc { occ_in_lam = in_lam, occ_int_cxt = int_cxt, occ_n_br = n_br } -- See Note [Inline small things to avoid creating a thunk] -> n_br < 100 -- See Note [Suppress exponential blowup] && smallEnoughToInline dflags unfolding -- Small enough to dup -- ToDo: consider discount on smallEnoughToInline if int_cxt is true -- -- NB: Do NOT inline arbitrarily big things, even if occ_n_br=1 -- 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. && (in_lam == NotInsideLam || -- 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 == IsInteresting)) -- 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 would 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! where unfolding = idUnfolding bndr dflags = seDynFlags env active = isActive (sm_phase (getMode env)) (idInlineActivation bndr) -- See Note [pre/postInlineUnconditionally in gentle mode] {- Note [Inline small things to avoid creating a thunk] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ 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 (because 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 ... This is very important in practice; e.g. wheel-seive1 doubles in allocation if you miss this out. And bits of GHC itself start to allocate more. An egregious example is test perf/compiler/T14697, where GHC.Driver.CmdLine.$wprocessArgs allocated hugely more. Note [Suppress exponential blowup] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In #13253, and several related tickets, we got an exponential blowup in code size from postInlineUnconditionally. The trouble comes when we have let j1a = case f y of { True -> p; False -> q } j1b = case f y of { True -> q; False -> p } j2a = case f (y+1) of { True -> j1a; False -> j1b } j2b = case f (y+1) of { True -> j1b; False -> j1a } ... in case f (y+10) of { True -> j10a; False -> j10b } when there are many branches. In pass 1, postInlineUnconditionally inlines j10a and j10b (they are both small). Now we have two calls to j9a and two to j9b. In pass 2, postInlineUnconditionally inlines all four of these calls, leaving four calls to j8a and j8b. Etc. Yikes! This is exponential! A possible plan: stop doing postInlineUnconditionally for some fixed, smallish number of branches, say 4. But that turned out to be bad: see Note [Inline small things to avoid creating a thunk]. And, as it happened, the problem with #13253 was solved in a different way (Note [Duplicating StrictArg] in Simplify). So I just set an arbitrary, high limit of 100, to stop any totally exponential behaviour. This still leaves the nasty possiblity that /ordinary/ inlining (not postInlineUnconditionally) might inline these join points, each of which is individually quiet small. I'm still not sure what to do about this (e.g. see #15488). Note [Top level and postInlineUnconditionally] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We don't do postInlineUnconditionally for top-level things (even for ones that are trivial): * Doing so will inline top-level error expressions that have been carefully floated out by FloatOut. More generally, it might replace static allocation with dynamic. * Even for trivial expressions there's a problem. Consider {-# RULE "foo" forall (xs::[T]). reverse xs = ruggle xs #-} blah xs = reverse xs ruggle = sort In one simplifier pass we might fire the rule, getting blah xs = ruggle xs but in *that* simplifier pass we must not do postInlineUnconditionally on 'ruggle' because then we'll have an unbound occurrence of 'ruggle' If the rhs is trivial it'll be inlined by callSiteInline, and then the binding will be dead and discarded by the next use of OccurAnal * There is less point, because the main goal is to get rid of local bindings used in multiple case branches. * The inliner should inline trivial things at call sites anyway. * The Id might be exported. We could check for that separately, but since we aren't going to postInlineUnconditionally /any/ top-level bindings, we don't need to test. Note [Stable unfoldings and postInlineUnconditionally] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Do not do postInlineUnconditionally if the Id has a stable unfolding, otherwise we lose the unfolding. Example -- f has stable unfolding with rhs (e |> co) -- where 'e' is big f = e |> co Then there's a danger we'll optimise to f' = e f = f' |> co and now postInlineUnconditionally, losing the stable unfolding on f. Now f' won't inline because 'e' is too big. c.f. Note [Stable unfoldings and preInlineUnconditionally] ************************************************************************ * * Rebuilding a lambda * * ************************************************************************ -} mkLam :: SimplEnv -> [OutBndr] -> OutExpr -> SimplCont -> 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 _env [] body _cont = return body mkLam env bndrs body cont = do { dflags <- getDynFlags ; mkLam' dflags bndrs body } where 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 (mkCast lam (mkPiCos Representational bndrs co)) } where co_vars = tyCoVarsOfCo co bad bndr = isCoVar bndr && bndr `elemVarSet` co_vars mkLam' dflags bndrs body@(Lam {}) = mkLam' dflags (bndrs ++ bndrs1) body1 where (bndrs1, body1) = collectBinders body mkLam' dflags bndrs (Tick t expr) | tickishFloatable t = mkTick t <$> mkLam' dflags bndrs expr mkLam' dflags bndrs body | gopt Opt_DoEtaReduction dflags , Just etad_lam <- tryEtaReduce bndrs body = do { tick (EtaReduction (head bndrs)) ; return etad_lam } | not (contIsRhs cont) -- See Note [Eta-expanding lambdas] , sm_eta_expand (getMode env) , any isRuntimeVar bndrs , let body_arity = exprEtaExpandArity dflags body , expandableArityType body_arity = do { tick (EtaExpansion (head bndrs)) ; let res = mkLams bndrs (etaExpandAT body_arity body) ; traceSmpl "eta expand" (vcat [text "before" <+> ppr (mkLams bndrs body) , text "after" <+> ppr res]) ; return res } | otherwise = return (mkLams bndrs body) {- Note [Eta expanding lambdas] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~ In general we *do* want to eta-expand lambdas. Consider f (\x -> case x of (a,b) -> \s -> blah) where 's' is a state token, and hence can be eta expanded. This showed up in the code for GHc.IO.Handle.Text.hPutChar, a rather important function! The eta-expansion will never happen unless we do it now. (Well, it's possible that CorePrep will do it, but CorePrep only has a half-baked eta-expander that can't deal with casts. So it's much better to do it here.) However, when the lambda is let-bound, as the RHS of a let, we have a better eta-expander (in the form of tryEtaExpandRhs), so we don't bother to try expansion in mkLam in that case; hence the contIsRhs guard. NB: We check the SimplEnv (sm_eta_expand), not DynFlags. See Note [No eta expansion in stable unfoldings] Note [Casts and lambdas] ~~~~~~~~~~~~~~~~~~~~~~~~ Consider (\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. ************************************************************************ * * Eta expansion * * ************************************************************************ -} tryEtaExpandRhs :: SimplMode -> OutId -> OutExpr -> SimplM (ArityType, OutExpr) -- See Note [Eta-expanding at let bindings] -- If tryEtaExpandRhs rhs = (n, is_bot, rhs') then -- (a) rhs' has manifest arity n -- (b) if is_bot is True then rhs' applied to n args is guaranteed bottom tryEtaExpandRhs mode bndr rhs | Just join_arity <- isJoinId_maybe bndr = do { let (join_bndrs, join_body) = collectNBinders join_arity rhs oss = [idOneShotInfo id | id <- join_bndrs, isId id] arity_type | exprIsDeadEnd join_body = ABot (length oss) | otherwise = ATop oss ; return (arity_type, rhs) } -- Note [Do not eta-expand join points] -- But do return the correct arity and bottom-ness, because -- these are used to set the bndr's IdInfo (#15517) -- Note [Invariants on join points] invariant 2b, in GHC.Core | sm_eta_expand mode -- Provided eta-expansion is on , new_arity > old_arity -- And the current manifest arity isn't enough , want_eta rhs = do { tick (EtaExpansion bndr) ; return (arity_type, etaExpandAT arity_type rhs) } | otherwise = return (arity_type, rhs) where dflags = sm_dflags mode old_arity = exprArity rhs arity_type = findRhsArity dflags bndr rhs old_arity `maxWithArity` idCallArity bndr new_arity = arityTypeArity arity_type -- See Note [Which RHSs do we eta-expand?] want_eta (Cast e _) = want_eta e want_eta (Tick _ e) = want_eta e want_eta (Lam b e) | isTyVar b = want_eta e want_eta (App e a) | exprIsTrivial a = want_eta e want_eta (Var {}) = False want_eta (Lit {}) = False want_eta _ = True {- want_eta _ = case arity_type of ATop (os:_) -> isOneShotInfo os ATop [] -> False ABot {} -> True -} {- Note [Eta-expanding at let bindings] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We now eta expand at let-bindings, which is where the payoff comes. The most significant thing is that we can do a simple arity analysis (in GHC.Core.Opt.Arity.findRhsArity), which we can't do for free-floating lambdas One useful consequence of not eta-expanding lambdas is this example: genMap :: C a => ... {-# INLINE genMap #-} genMap f xs = ... myMap :: D a => ... {-# INLINE myMap #-} myMap = genMap Notice that 'genMap' should only inline if applied to two arguments. In the stable unfolding for myMap we'll have the unfolding (\d -> genMap Int (..d..)) We do not want to eta-expand to (\d f xs -> genMap Int (..d..) f xs) because then 'genMap' will inline, and it really shouldn't: at least as far as the programmer is concerned, it's not applied to two arguments! Note [Which RHSs do we eta-expand?] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We don't eta-expand: * Trivial RHSs, e.g. f = g If we eta expand do f = \x. g x we'll just eta-reduce again, and so on; so the simplifier never terminates. * PAPs: see Note [Do not eta-expand PAPs] What about things like this? f = case y of p -> \x -> blah Here we do eta-expand. This is a change (Jun 20), but if we have really decided that f has arity 1, then putting that lambda at the top seems like a Good idea. Note [Do not eta-expand PAPs] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We used to have old_arity = manifestArity rhs, which meant that we would eta-expand even PAPs. But this gives no particular advantage, and can lead to a massive blow-up in code size, exhibited by #9020. Suppose we have a PAP foo :: IO () foo = returnIO () Then we can eta-expand do foo = (\eta. (returnIO () |> sym g) eta) |> g where g :: IO () ~ State# RealWorld -> (# State# RealWorld, () #) But there is really no point in doing this, and it generates masses of coercions and whatnot that eventually disappear again. For T9020, GHC allocated 6.6G before, and 0.8G afterwards; and residency dropped from 1.8G to 45M. Moreover, if we eta expand f = g d ==> f = \x. g d x that might in turn make g inline (if it has an inline pragma), which we might not want. After all, INLINE pragmas say "inline only when saturated" so we don't want to be too gung-ho about saturating! But note that this won't eta-expand, say f = \g -> map g Does it matter not eta-expanding such functions? I'm not sure. Perhaps strictness analysis will have less to bite on? Note [Do not eta-expand join points] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Similarly to CPR (see Note [Don't w/w join points for CPR] in GHC.Core.Opt.WorkWrap), a join point stands well to gain from its outer binding's eta-expansion, and eta-expanding a join point is fraught with issues like how to deal with a cast: let join $j1 :: IO () $j1 = ... $j2 :: Int -> IO () $j2 n = if n > 0 then $j1 else ... => let join $j1 :: IO () $j1 = (\eta -> ...) `cast` N:IO :: State# RealWorld -> (# State# RealWorld, ()) ~ IO () $j2 :: Int -> IO () $j2 n = (\eta -> if n > 0 then $j1 else ...) `cast` N:IO :: State# RealWorld -> (# State# RealWorld, ()) ~ IO () The cast here can't be pushed inside the lambda (since it's not casting to a function type), so the lambda has to stay, but it can't because it contains a reference to a join point. In fact, $j2 can't be eta-expanded at all. Rather than try and detect this situation (and whatever other situations crop up!), we don't bother; again, any surrounding eta-expansion will improve these join points anyway, since an outer cast can *always* be pushed inside. By the time CorePrep comes around, the code is very likely to look more like this: let join $j1 :: State# RealWorld -> (# State# RealWorld, ()) $j1 = (...) eta $j2 :: Int -> State# RealWorld -> (# State# RealWorld, ()) $j2 = if n > 0 then $j1 else (...) eta ************************************************************************ * * \subsection{Floating lets out of big lambdas} * * ************************************************************************ Note [Floating and type abstraction] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider this: x = /\a. C e1 e2 We'd like to float this to y1 = /\a. e1 y2 = /\a. e2 x = /\a. C (y1 a) (y2 a) for the usual reasons: we want to inline x rather vigorously. You may think that this kind of thing is rare. But in some programs it is common. For example, if you do closure conversion you might get: data a :-> b = forall e. (e -> a -> b) :$ 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 like /\ a1..an. body and treat them specially. The real work is done in GHC.Core.Opt.Simplify.Utils.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 in /\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 into letrec g' = /\a -> ... g' a ... in let f = /\ a -> g' a which is better. In effect, it means that big lambdas don't impede let-floating. 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, let st = /\ s -> let { x1=r1 ; x2=r2 } in ... in .. becomes 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. Note [Which type variables to abstract over] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Abstract only over the type variables free in the rhs wrt which the new binding is abstracted. Note that * The naive approach of abstracting 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*) * We must do closeOverKinds. Example (#10934): f = /\k (f:k->*) (a:k). let t = AccFailure @ (f a) in ... Here we want to float 't', but we must remember to abstract over 'k' as well, even though it is not explicitly mentioned in the RHS, otherwise we get t = /\ (f:k->*) (a:k). AccFailure @ (f a) which is obviously bogus. -} abstractFloats :: DynFlags -> TopLevelFlag -> [OutTyVar] -> SimplFloats -> OutExpr -> SimplM ([OutBind], OutExpr) abstractFloats dflags top_lvl main_tvs floats body = ASSERT( notNull body_floats ) ASSERT( isNilOL (sfJoinFloats floats) ) do { (subst, float_binds) <- mapAccumLM abstract empty_subst body_floats ; return (float_binds, GHC.Core.Subst.substExpr subst body) } where is_top_lvl = isTopLevel top_lvl main_tv_set = mkVarSet main_tvs body_floats = letFloatBinds (sfLetFloats floats) empty_subst = GHC.Core.Subst.mkEmptySubst (sfInScope floats) abstract :: GHC.Core.Subst.Subst -> OutBind -> SimplM (GHC.Core.Subst.Subst, OutBind) abstract subst (NonRec id rhs) = do { (poly_id1, poly_app) <- mk_poly1 tvs_here id ; let (poly_id2, poly_rhs) = mk_poly2 poly_id1 tvs_here rhs' subst' = GHC.Core.Subst.extendIdSubst subst id poly_app ; return (subst', NonRec poly_id2 poly_rhs) } where rhs' = GHC.Core.Subst.substExpr subst rhs -- tvs_here: see Note [Which type variables to abstract over] tvs_here = scopedSort $ filter (`elemVarSet` main_tv_set) $ closeOverKindsList $ exprSomeFreeVarsList isTyVar rhs' abstract subst (Rec prs) = do { (poly_ids, poly_apps) <- mapAndUnzipM (mk_poly1 tvs_here) ids ; let subst' = GHC.Core.Subst.extendSubstList subst (ids `zip` poly_apps) poly_pairs = [ mk_poly2 poly_id tvs_here rhs' | (poly_id, rhs) <- poly_ids `zip` rhss , let rhs' = GHC.Core.Subst.substExpr subst' rhs ] ; return (subst', Rec poly_pairs) } where (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 = scopedSort main_tvs mk_poly1 :: [TyVar] -> Id -> SimplM (Id, CoreExpr) mk_poly1 tvs_here var = do { uniq <- getUniqueM ; let poly_name = setNameUnique (idName var) uniq -- Keep same name poly_ty = mkInfForAllTys tvs_here (idType var) -- But new type of course poly_id = transferPolyIdInfo var tvs_here $ -- Note [transferPolyIdInfo] in GHC.Types.Id mkLocalId poly_name (idMult var) 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. mk_poly2 :: Id -> [TyVar] -> CoreExpr -> (Id, CoreExpr) mk_poly2 poly_id tvs_here rhs = (poly_id `setIdUnfolding` unf, poly_rhs) where poly_rhs = mkLams tvs_here rhs unf = mkUnfolding dflags InlineRhs is_top_lvl False poly_rhs -- We want the unfolding. Consider -- let -- x = /\a. let y = ... in Just y -- in body -- Then we float the y-binding out (via abstractFloats and addPolyBind) -- but 'x' may well then be inlined in 'body' in which case we'd like the -- opportunity to inline 'y' too. {- 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 * * ************************************************************************ prepareAlts tries these things: 1. filterAlts: eliminate alternatives that cannot match, including the DEFAULT alternative. Here "cannot match" includes knowledge from GADTs 2. refineDefaultAlt: if the DEFAULT alternative can match only one possible constructor, then make that constructor explicit. e.g. case e of x { DEFAULT -> rhs } ===> case e of x { (a,b) -> rhs } where the type is a single constructor type. This gives better code when rhs also scrutinises x or e. See CoreUtils Note [Refine DEFAULT case alternatives] 3. combineIdenticalAlts: combine identical alternatives into a DEFAULT. See CoreUtils Note [Combine identical alternatives], which also says why we do this on InAlts not on OutAlts 4. Returns a list of the constructors that cannot holds in the DEFAULT alternative (if there is one) It's a good idea to do this stuff before simplifying the alternatives, to avoid simplifying alternatives we know can't happen, and to come up with the list of constructors that are handled, to put into the IdInfo of the case binder, for use when simplifying the alternatives. Eliminating the default alternative in (1) isn't so obvious, but it can happen: data Colour = Red | Green | Blue f x = case x of Red -> .. Green -> .. DEFAULT -> h x h y = case y of Blue -> .. DEFAULT -> [ case y of ... ] If we inline h into f, the default case of the inlined h can't happen. If we don't notice this, we may end up filtering out *all* the cases of the inner case y, which give us nowhere to go! -} prepareAlts :: OutExpr -> OutId -> [InAlt] -> SimplM ([AltCon], [InAlt]) -- The returned alternatives can be empty, none are possible prepareAlts scrut case_bndr' alts | Just (tc, tys) <- splitTyConApp_maybe (varType case_bndr') -- Case binder is needed just for its type. Note that as an -- OutId, it has maximum information; this is important. -- Test simpl013 is an example = do { us <- getUniquesM ; let (idcs1, alts1) = filterAlts tc tys imposs_cons alts (yes2, alts2) = refineDefaultAlt us (idMult case_bndr') tc tys idcs1 alts1 -- the multiplicity on case_bndr's is the multiplicity of the -- case expression The newly introduced patterns in -- refineDefaultAlt must be scaled by this multiplicity (yes3, idcs3, alts3) = combineIdenticalAlts idcs1 alts2 -- "idcs" stands for "impossible default data constructors" -- i.e. the constructors that can't match the default case ; when yes2 $ tick (FillInCaseDefault case_bndr') ; when yes3 $ tick (AltMerge case_bndr') ; return (idcs3, alts3) } | otherwise -- Not a data type, so nothing interesting happens = return ([], alts) where imposs_cons = case scrut of Var v -> otherCons (idUnfolding v) _ -> [] {- ************************************************************************ * * mkCase * * ************************************************************************ mkCase tries these things * Note [Nerge nested cases] * Note [Eliminate identity case] * Note [Scrutinee constant folding] Note [Merge Nested Cases] ~~~~~~~~~~~~~~~~~~~~~~~~~ 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. Note [Eliminate Identity Case] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ case e of ===> e True -> True; False -> False and similar friends. Note [Scrutinee Constant Folding] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ case x op# k# of _ { ===> case x of _ { a1# -> e1 (a1# inv_op# k#) -> e1 a2# -> e2 (a2# inv_op# k#) -> e2 ... ... DEFAULT -> ed DEFAULT -> ed where (x op# k#) inv_op# k# == x And similarly for commuted arguments and for some unary operations. The purpose of this transformation is not only to avoid an arithmetic operation at runtime but to allow other transformations to apply in cascade. Example with the "Merge Nested Cases" optimization (from #12877): main = case t of t0 0## -> ... DEFAULT -> case t0 `minusWord#` 1## of t1 0## -> ... DEFAULT -> case t1 `minusWord#` 1## of t2 0## -> ... DEFAULT -> case t2 `minusWord#` 1## of _ 0## -> ... DEFAULT -> ... becomes: main = case t of _ 0## -> ... 1## -> ... 2## -> ... 3## -> ... DEFAULT -> ... There are some wrinkles * Do not apply caseRules if there is just a single DEFAULT alternative case e +# 3# of b { DEFAULT -> rhs } If we applied the transformation here we would (stupidly) get case a of b' { DEFAULT -> let b = e +# 3# in rhs } and now the process may repeat, because that let will really be a case. * The type of the scrutinee might change. E.g. case tagToEnum (x :: Int#) of (b::Bool) False -> e1 True -> e2 ==> case x of (b'::Int#) DEFAULT -> e1 1# -> e2 * The case binder may be used in the right hand sides, so we need to make a local binding for it, if it is alive. e.g. case e +# 10# of b DEFAULT -> blah...b... 44# -> blah2...b... ===> case e of b' DEFAULT -> let b = b' +# 10# in blah...b... 34# -> let b = 44# in blah2...b... Note that in the non-DEFAULT cases we know what to bind 'b' to, whereas in the DEFAULT case we must reconstruct the original value. But NB: we use b'; we do not duplicate 'e'. * In dataToTag we might need to make up some fake binders; see Note [caseRules for dataToTag] in GHC.Core.Opt.ConstantFold -} mkCase, mkCase1, mkCase2, mkCase3 :: DynFlags -> OutExpr -> OutId -> OutType -> [OutAlt] -- Alternatives in standard (increasing) order -> SimplM OutExpr -------------------------------------------------- -- 1. Merge Nested Cases -------------------------------------------------- mkCase dflags scrut outer_bndr alts_ty ((DEFAULT, _, deflt_rhs) : outer_alts) | gopt Opt_CaseMerge dflags , (ticks, Case (Var inner_scrut_var) inner_bndr _ inner_alts) <- stripTicksTop tickishFloatable deflt_rhs , inner_scrut_var == outer_bndr = do { tick (CaseMerge outer_bndr) ; let wrap_alt (con, args, rhs) = ASSERT( outer_bndr `notElem` args ) (con, args, wrap_rhs rhs) -- Simplifier's no-shadowing invariant should ensure -- that outer_bndr is not shadowed by the inner patterns wrap_rhs rhs = Let (NonRec inner_bndr (Var outer_bndr)) rhs -- The let is OK even for unboxed binders, wrapped_alts | isDeadBinder inner_bndr = inner_alts | otherwise = map wrap_alt inner_alts merged_alts = mergeAlts outer_alts wrapped_alts -- NB: mergeAlts gives priority to the left -- 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! ; fmap (mkTicks ticks) $ mkCase1 dflags scrut outer_bndr alts_ty merged_alts } -- Warning: don't call mkCase 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! mkCase dflags scrut bndr alts_ty alts = mkCase1 dflags scrut bndr alts_ty alts -------------------------------------------------- -- 2. Eliminate Identity Case -------------------------------------------------- mkCase1 _dflags scrut case_bndr _ alts@((_,_,rhs1) : _) -- Identity case | all identity_alt alts = do { tick (CaseIdentity case_bndr) ; return (mkTicks ticks $ re_cast scrut rhs1) } where ticks = concatMap (stripTicksT tickishFloatable . thdOf3) (tail alts) identity_alt (con, args, rhs) = check_eq rhs con args check_eq (Cast rhs co) con args -- See Note [RHS casts] = not (any (`elemVarSet` tyCoVarsOfCo co) args) && check_eq rhs con args check_eq (Tick t e) alt args = tickishFloatable t && check_eq e alt args check_eq (Lit lit) (LitAlt lit') _ = lit == lit' check_eq (Var v) _ _ | v == case_bndr = True check_eq (Var v) (DataAlt con) args | null arg_tys, null args = v == dataConWorkId con -- Optimisation only check_eq rhs (DataAlt con) args = cheapEqExpr' tickishFloatable rhs $ mkConApp2 con arg_tys args check_eq _ _ _ = False arg_tys = tyConAppArgs (idType case_bndr) -- Note [RHS casts] -- ~~~~~~~~~~~~~~~~ -- 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 re_cast scrut (Cast rhs co) = Cast (re_cast scrut rhs) co re_cast scrut _ = scrut mkCase1 dflags scrut bndr alts_ty alts = mkCase2 dflags scrut bndr alts_ty alts -------------------------------------------------- -- 2. Scrutinee Constant Folding -------------------------------------------------- mkCase2 dflags scrut bndr alts_ty alts | -- See Note [Scrutinee Constant Folding] case alts of -- Not if there is just a DEFAULT alternative [(DEFAULT,_,_)] -> False _ -> True , gopt Opt_CaseFolding dflags , Just (scrut', tx_con, mk_orig) <- caseRules (targetPlatform dflags) scrut = do { bndr' <- newId (fsLit "lwild") Many (exprType scrut') ; alts' <- mapMaybeM (tx_alt tx_con mk_orig bndr') alts -- mapMaybeM: discard unreachable alternatives -- See Note [Unreachable caseRules alternatives] -- in GHC.Core.Opt.ConstantFold ; mkCase3 dflags scrut' bndr' alts_ty $ add_default (re_sort alts') } | otherwise = mkCase3 dflags scrut bndr alts_ty alts where -- We need to keep the correct association between the scrutinee and its -- binder if the latter isn't dead. Hence we wrap rhs of alternatives with -- "let bndr = ... in": -- -- case v + 10 of y =====> case v of y -- 20 -> e1 10 -> let y = 20 in e1 -- DEFAULT -> e2 DEFAULT -> let y = v + 10 in e2 -- -- Other transformations give: =====> case v of y' -- 10 -> let y = 20 in e1 -- DEFAULT -> let y = y' + 10 in e2 -- -- This wrapping is done in tx_alt; we use mk_orig, returned by caseRules, -- to construct an expression equivalent to the original one, for use -- in the DEFAULT case tx_alt :: (AltCon -> Maybe AltCon) -> (Id -> CoreExpr) -> Id -> CoreAlt -> SimplM (Maybe CoreAlt) tx_alt tx_con mk_orig new_bndr (con, bs, rhs) = case tx_con con of Nothing -> return Nothing Just con' -> do { bs' <- mk_new_bndrs new_bndr con' ; return (Just (con', bs', rhs')) } where rhs' | isDeadBinder bndr = rhs | otherwise = bindNonRec bndr orig_val rhs orig_val = case con of DEFAULT -> mk_orig new_bndr LitAlt l -> Lit l DataAlt dc -> mkConApp2 dc (tyConAppArgs (idType bndr)) bs mk_new_bndrs new_bndr (DataAlt dc) | not (isNullaryRepDataCon dc) = -- For non-nullary data cons we must invent some fake binders -- See Note [caseRules for dataToTag] in GHC.Core.Opt.ConstantFold do { us <- getUniquesM ; let (ex_tvs, arg_ids) = dataConRepInstPat us (idMult new_bndr) dc (tyConAppArgs (idType new_bndr)) ; return (ex_tvs ++ arg_ids) } mk_new_bndrs _ _ = return [] re_sort :: [CoreAlt] -> [CoreAlt] -- Sort the alternatives to re-establish -- GHC.Core Note [Case expression invariants] re_sort alts = sortBy cmpAlt alts add_default :: [CoreAlt] -> [CoreAlt] -- See Note [Literal cases] add_default ((LitAlt {}, bs, rhs) : alts) = (DEFAULT, bs, rhs) : alts add_default alts = alts {- Note [Literal cases] ~~~~~~~~~~~~~~~~~~~~~~~ If we have case tagToEnum (a ># b) of False -> e1 True -> e2 then caseRules for TagToEnum will turn it into case tagToEnum (a ># b) of 0# -> e1 1# -> e2 Since the case is exhaustive (all cases are) we can convert it to case tagToEnum (a ># b) of DEFAULT -> e1 1# -> e2 This may generate sligthtly better code (although it should not, since all cases are exhaustive) and/or optimise better. I'm not certain that it's necessary, but currently we do make this change. We do it here, NOT in the TagToEnum rules (see "Beware" in Note [caseRules for tagToEnum] in GHC.Core.Opt.ConstantFold) -} -------------------------------------------------- -- Catch-all -------------------------------------------------- mkCase3 _dflags scrut bndr alts_ty alts = return (Case scrut bndr alts_ty alts) -- See Note [Exitification] and Note [Do not inline exit join points] in -- GHC.Core.Opt.Exitify -- This lives here (and not in Id) because occurrence info is only valid on -- InIds, so it's crucial that isExitJoinId is only called on freshly -- occ-analysed code. It's not a generic function you can call anywhere. isExitJoinId :: Var -> Bool isExitJoinId id = isJoinId id && isOneOcc (idOccInfo id) && occ_in_lam (idOccInfo id) == IsInsideLam {- Note [Dead binders] ~~~~~~~~~~~~~~~~~~~~ Note that dead-ness is maintained by the simplifier, so that it is accurate after simplification as well as before. Note [Cascading case merge] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ Case merging should cascade in one sweep, because it happens bottom-up case e of a { DEFAULT -> case a of b DEFAULT -> case b of c { DEFAULT -> e A -> ea B -> eb C -> ec ==> case e of a { DEFAULT -> case a of b DEFAULT -> let c = b in e A -> let c = b in ea B -> eb C -> ec ==> case e of a { DEFAULT -> let b = a in let c = b in e A -> let b = a in let c = b in ea B -> let b = a in eb C -> ec However here's a tricky case that we still don't catch, and I don't see how to catch it in one pass: case x of c1 { I# a1 -> case a1 of c2 -> 0 -> ... DEFAULT -> case x of c3 { I# a2 -> case a2 of ... After occurrence analysis (and its binder-swap) we get this case x of c1 { I# a1 -> let x = c1 in -- Binder-swap addition case a1 of c2 -> 0 -> ... DEFAULT -> case x of c3 { I# a2 -> case a2 of ... When we simplify the inner case x, we'll see that x=c1=I# a1. So we'll bind a2 to a1, and get case x of c1 { I# a1 -> case a1 of c2 -> 0 -> ... DEFAULT -> case a1 of ... This is correct, but we can't do a case merge in this sweep because c2 /= a1. Reason: the binding c1=I# a1 went inwards without getting changed to c1=I# c2. I don't think this is worth fixing, even if I knew how. It'll all come out in the next pass anyway. -}