{- (c) The GRASP/AQUA Project, Glasgow University, 1992-1998 ************************************************************************ * * \section[OccurAnal]{Occurrence analysis pass} * * ************************************************************************ The occurrence analyser re-typechecks a core expression, returning a new core expression with (hopefully) improved usage information. -} {-# LANGUAGE CPP, BangPatterns #-} module OccurAnal ( occurAnalysePgm, occurAnalyseExpr, occurAnalyseExpr_NoBinderSwap ) where #include "HsVersions.h" import CoreSyn import CoreFVs import CoreUtils ( exprIsTrivial, isDefaultAlt, isExpandableApp, stripTicksTopE, mkTicks ) import Id import Name( localiseName ) import BasicTypes import Module( Module ) import Coercion import VarSet import VarEnv import Var import Demand ( argOneShots, argsOneShots ) import Maybes ( orElse ) import Digraph ( SCC(..), Node , stronglyConnCompFromEdgedVerticesUniq , stronglyConnCompFromEdgedVerticesUniqR ) import Unique import UniqFM import Util import Outputable import Data.List import Control.Arrow ( second ) {- ************************************************************************ * * occurAnalysePgm, occurAnalyseExpr, occurAnalyseExpr_NoBinderSwap * * ************************************************************************ Here's the externally-callable interface: -} occurAnalysePgm :: Module -- Used only in debug output -> (Activation -> Bool) -> [CoreRule] -> [CoreVect] -> VarSet -> CoreProgram -> CoreProgram occurAnalysePgm this_mod active_rule imp_rules vects vectVars binds | isEmptyVarEnv final_usage = occ_anald_binds | otherwise -- See Note [Glomming] = WARN( True, hang (text "Glomming in" <+> ppr this_mod <> colon) 2 (ppr final_usage ) ) occ_anald_glommed_binds where init_env = initOccEnv active_rule (final_usage, occ_anald_binds) = go init_env binds (_, occ_anald_glommed_binds) = occAnalRecBind init_env imp_rule_edges (flattenBinds occ_anald_binds) initial_uds -- It's crucial to re-analyse the glommed-together bindings -- so that we establish the right loop breakers. Otherwise -- we can easily create an infinite loop (Trac #9583 is an example) initial_uds = addIdOccs emptyDetails (rulesFreeVars imp_rules `unionVarSet` vectsFreeVars vects `unionVarSet` vectVars) -- The RULES and VECTORISE declarations keep things alive! (For VECTORISE declarations, -- we only get them *until* the vectoriser runs. Afterwards, these dependencies are -- reflected in 'vectors' — see Note [Vectorisation declarations and occurrences].) -- Note [Preventing loops due to imported functions rules] imp_rule_edges = foldr (plusVarEnv_C unionVarSet) emptyVarEnv [ mapVarEnv (const maps_to) (exprFreeIds arg `delVarSetList` ru_bndrs imp_rule) | imp_rule <- imp_rules , not (isBuiltinRule imp_rule) -- See Note [Plugin rules] , let maps_to = exprFreeIds (ru_rhs imp_rule) `delVarSetList` ru_bndrs imp_rule , arg <- ru_args imp_rule ] go :: OccEnv -> [CoreBind] -> (UsageDetails, [CoreBind]) go _ [] = (initial_uds, []) go env (bind:binds) = (final_usage, bind' ++ binds') where (bs_usage, binds') = go env binds (final_usage, bind') = occAnalBind env imp_rule_edges bind bs_usage occurAnalyseExpr :: CoreExpr -> CoreExpr -- Do occurrence analysis, and discard occurrence info returned occurAnalyseExpr = occurAnalyseExpr' True -- do binder swap occurAnalyseExpr_NoBinderSwap :: CoreExpr -> CoreExpr occurAnalyseExpr_NoBinderSwap = occurAnalyseExpr' False -- do not do binder swap occurAnalyseExpr' :: Bool -> CoreExpr -> CoreExpr occurAnalyseExpr' enable_binder_swap expr = snd (occAnal env expr) where env = (initOccEnv all_active_rules) {occ_binder_swap = enable_binder_swap} -- To be conservative, we say that all inlines and rules are active all_active_rules = \_ -> True {- Note [Plugin rules] ~~~~~~~~~~~~~~~~~~~~~~ Conal Elliott (Trac #11651) built a GHC plugin that added some BuiltinRules (for imported Ids) to the mg_rules field of ModGuts, to do some domain-specific transformations that could not be expressed with an ordinary pattern-matching CoreRule. But then we can't extract the dependencies (in imp_rule_edges) from ru_rhs etc, because a BuiltinRule doesn't have any of that stuff. So we simply assume that BuiltinRules have no dependencies, and filter them out from the imp_rule_edges comprehension. -} {- ************************************************************************ * * Bindings * * ************************************************************************ Note [Recursive bindings: the grand plan] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ When we come across a binding group Rec { x1 = r1; ...; xn = rn } we treat it like this (occAnalRecBind): 1. Occurrence-analyse each right hand side, and build a "Details" for each binding to capture the results. Wrap the details in a Node (details, node-id, dep-node-ids), where node-id is just the unique of the binder, and dep-node-ids lists all binders on which this binding depends. We'll call these the "scope edges". See Note [Forming the Rec groups]. All this is done by makeNode. 2. Do SCC-analysis on these Nodes. Each SCC will become a new Rec or NonRec. The key property is that every free variable of a binding is accounted for by the scope edges, so that when we are done everything is still in scope. 3. For each Cyclic SCC of the scope-edge SCC-analysis in (2), we identify suitable loop-breakers to ensure that inlining terminates. This is done by occAnalRec. 4. To do so we form a new set of Nodes, with the same details, but different edges, the "loop-breaker nodes". The loop-breaker nodes have both more and fewer depedencies than the scope edges (see Note [Choosing loop breakers]) More edges: if f calls g, and g has an active rule that mentions h then we add an edge from f -> h Fewer edges: we only include dependencies on active rules, on rule RHSs (not LHSs) and if there is an INLINE pragma only on the stable unfolding (and vice versa). The scope edges must be much more inclusive. 5. The "weak fvs" of a node are, by definition: the scope fvs - the loop-breaker fvs See Note [Weak loop breakers], and the nd_weak field of Details 6. Having formed the loop-breaker nodes Note [Dead code] ~~~~~~~~~~~~~~~~ Dropping dead code for a cyclic Strongly Connected Component is done in a very simple way: the entire SCC is dropped if none of its binders are mentioned in the body; otherwise the whole thing is kept. The key observation is that dead code elimination happens after dependency analysis: so 'occAnalBind' processes SCCs instead of the original term's binding groups. Thus 'occAnalBind' does indeed drop 'f' in an example like letrec f = ...g... g = ...(...g...)... in ...g... when 'g' no longer uses 'f' at all (eg 'f' does not occur in a RULE in 'g'). 'occAnalBind' first consumes 'CyclicSCC g' and then it consumes 'AcyclicSCC f', where 'body_usage' won't contain 'f'. ------------------------------------------------------------ Note [Forming Rec groups] ~~~~~~~~~~~~~~~~~~~~~~~~~ We put bindings {f = ef; g = eg } in a Rec group if "f uses g" and "g uses f", no matter how indirectly. We do a SCC analysis with an edge f -> g if "f uses g". More precisely, "f uses g" iff g should be in scope wherever f is. That is, g is free in: a) the rhs 'ef' b) or the RHS of a rule for f (Note [Rules are extra RHSs]) c) or the LHS or a rule for f (Note [Rule dependency info]) These conditions apply regardless of the activation of the RULE (eg it might be inactive in this phase but become active later). Once a Rec is broken up it can never be put back together, so we must be conservative. The principle is that, regardless of rule firings, every variable is always in scope. * Note [Rules are extra RHSs] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ A RULE for 'f' is like an extra RHS for 'f'. That way the "parent" keeps the specialised "children" alive. If the parent dies (because it isn't referenced any more), then the children will die too (unless they are already referenced directly). To that end, we build a Rec group for each cyclic strongly connected component, *treating f's rules as extra RHSs for 'f'*. More concretely, the SCC analysis runs on a graph with an edge from f -> g iff g is mentioned in (a) f's rhs (b) f's RULES These are rec_edges. Under (b) we include variables free in *either* LHS *or* RHS of the rule. The former might seems silly, but see Note [Rule dependency info]. So in Example [eftInt], eftInt and eftIntFB will be put in the same Rec, even though their 'main' RHSs are both non-recursive. * Note [Rule dependency info] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ The VarSet in a RuleInfo is used for dependency analysis in the occurrence analyser. We must track free vars in *both* lhs and rhs. Hence use of idRuleVars, rather than idRuleRhsVars in occAnalBind. Why both? Consider x = y RULE f x = v+4 Then if we substitute y for x, we'd better do so in the rule's LHS too, so we'd better ensure the RULE appears to mention 'x' as well as 'v' * Note [Rules are visible in their own rec group] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We want the rules for 'f' to be visible in f's right-hand side. And we'd like them to be visible in other functions in f's Rec group. E.g. in Note [Specialisation rules] we want f' rule to be visible in both f's RHS, and fs's RHS. This means that we must simplify the RULEs first, before looking at any of the definitions. This is done by Simplify.simplRecBind, when it calls addLetIdInfo. ------------------------------------------------------------ Note [Choosing loop breakers] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Loop breaking is surprisingly subtle. First read the section 4 of "Secrets of the GHC inliner". This describes our basic plan. We avoid infinite inlinings by choosing loop breakers, and ensuring that a loop breaker cuts each loop. See also Note [Inlining and hs-boot files] in ToIface, which deals with a closely related source of infinite loops. Fundamentally, we do SCC analysis on a graph. For each recursive group we choose a loop breaker, delete all edges to that node, re-analyse the SCC, and iterate. But what is the graph? NOT the same graph as was used for Note [Forming Rec groups]! In particular, a RULE is like an equation for 'f' that is *always* inlined if it is applicable. We do *not* disable rules for loop-breakers. It's up to whoever makes the rules to make sure that the rules themselves always terminate. See Note [Rules for recursive functions] in Simplify.hs Hence, if f's RHS (or its INLINE template if it has one) mentions g, and g has a RULE that mentions h, and h has a RULE that mentions f then we *must* choose f to be a loop breaker. Example: see Note [Specialisation rules]. In general, take the free variables of f's RHS, and augment it with all the variables reachable by RULES from those starting points. That is the whole reason for computing rule_fv_env in occAnalBind. (Of course we only consider free vars that are also binders in this Rec group.) See also Note [Finding rule RHS free vars] Note that when we compute this rule_fv_env, we only consider variables free in the *RHS* of the rule, in contrast to the way we build the Rec group in the first place (Note [Rule dependency info]) Note that if 'g' has RHS that mentions 'w', we should add w to g's loop-breaker edges. More concretely there is an edge from f -> g iff (a) g is mentioned in f's RHS `xor` f's INLINE rhs (see Note [Inline rules]) (b) or h is mentioned in f's RHS, and g appears in the RHS of an active RULE of h or a transitive sequence of active rules starting with h Why "active rules"? See Note [Finding rule RHS free vars] Note that in Example [eftInt], *neither* eftInt *nor* eftIntFB is chosen as a loop breaker, because their RHSs don't mention each other. And indeed both can be inlined safely. Note again that the edges of the graph we use for computing loop breakers are not the same as the edges we use for computing the Rec blocks. That's why we compute - rec_edges for the Rec block analysis - loop_breaker_nodes for the loop breaker analysis * Note [Finding rule RHS free vars] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider this real example from Data Parallel Haskell tagZero :: Array Int -> Array Tag {-# INLINE [1] tagZeroes #-} tagZero xs = pmap (\x -> fromBool (x==0)) xs {-# RULES "tagZero" [~1] forall xs n. pmap fromBool = tagZero xs #-} So tagZero's RHS mentions pmap, and pmap's RULE mentions tagZero. However, tagZero can only be inlined in phase 1 and later, while the RULE is only active *before* phase 1. So there's no problem. To make this work, we look for the RHS free vars only for *active* rules. That's the reason for the occ_rule_act field of the OccEnv. * Note [Weak loop breakers] ~~~~~~~~~~~~~~~~~~~~~~~~~ There is a last nasty wrinkle. Suppose we have Rec { f = f_rhs RULE f [] = g h = h_rhs g = h ...more... } Remember that we simplify the RULES before any RHS (see Note [Rules are visible in their own rec group] above). So we must *not* postInlineUnconditionally 'g', even though its RHS turns out to be trivial. (I'm assuming that 'g' is not choosen as a loop breaker.) Why not? Because then we drop the binding for 'g', which leaves it out of scope in the RULE! Here's a somewhat different example of the same thing Rec { g = h ; h = ...f... ; f = f_rhs RULE f [] = g } Here the RULE is "below" g, but we *still* can't postInlineUnconditionally g, because the RULE for f is active throughout. So the RHS of h might rewrite to h = ...g... So g must remain in scope in the output program! We "solve" this by: Make g a "weak" loop breaker (OccInfo = IAmLoopBreaker True) iff g is a "missing free variable" of the Rec group A "missing free variable" x is one that is mentioned in an RHS or INLINE or RULE of a binding in the Rec group, but where the dependency on x may not show up in the loop_breaker_nodes (see note [Choosing loop breakers} above). A normal "strong" loop breaker has IAmLoopBreaker False. So Inline postInlineUnconditionally strong IAmLoopBreaker False no no weak IAmLoopBreaker True yes no other yes yes The **sole** reason for this kind of loop breaker is so that postInlineUnconditionally does not fire. Ugh. (Typically it'll inline via the usual callSiteInline stuff, so it'll be dead in the next pass, so the main Ugh is the tiresome complication.) Note [Rules for imported functions] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider this f = /\a. B.g a RULE B.g Int = 1 + f Int Note that * The RULE is for an imported function. * f is non-recursive Now we can get f Int --> B.g Int Inlining f --> 1 + f Int Firing RULE and so the simplifier goes into an infinite loop. This would not happen if the RULE was for a local function, because we keep track of dependencies through rules. But that is pretty much impossible to do for imported Ids. Suppose f's definition had been f = /\a. C.h a where (by some long and devious process), C.h eventually inlines to B.g. We could only spot such loops by exhaustively following unfoldings of C.h etc, in case we reach B.g, and hence (via the RULE) f. Note that RULES for imported functions are important in practice; they occur a lot in the libraries. We regard this potential infinite loop as a *programmer* error. It's up the programmer not to write silly rules like RULE f x = f x and the example above is just a more complicated version. Note [Preventing loops due to imported functions rules] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider: import GHC.Base (foldr) {-# RULES "filterList" forall p. foldr (filterFB (:) p) [] = filter p #-} filter p xs = build (\c n -> foldr (filterFB c p) n xs) filterFB c p = ... f = filter p xs Note that filter is not a loop-breaker, so what happens is: f = filter p xs = {inline} build (\c n -> foldr (filterFB c p) n xs) = {inline} foldr (filterFB (:) p) [] xs = {RULE} filter p xs We are in an infinite loop. A more elaborate example (that I actually saw in practice when I went to mark GHC.List.filter as INLINABLE) is as follows. Say I have this module: {-# LANGUAGE RankNTypes #-} module GHCList where import Prelude hiding (filter) import GHC.Base (build) {-# INLINABLE filter #-} filter :: (a -> Bool) -> [a] -> [a] filter p [] = [] filter p (x:xs) = if p x then x : filter p xs else filter p xs {-# NOINLINE [0] filterFB #-} filterFB :: (a -> b -> b) -> (a -> Bool) -> a -> b -> b filterFB c p x r | p x = x `c` r | otherwise = r {-# RULES "filter" [~1] forall p xs. filter p xs = build (\c n -> foldr (filterFB c p) n xs) "filterList" [1] forall p. foldr (filterFB (:) p) [] = filter p #-} Then (because RULES are applied inside INLINABLE unfoldings, but inlinings are not), the unfolding given to "filter" in the interface file will be: filter p [] = [] filter p (x:xs) = if p x then x : build (\c n -> foldr (filterFB c p) n xs) else build (\c n -> foldr (filterFB c p) n xs Note that because this unfolding does not mention "filter", filter is not marked as a strong loop breaker. Therefore at a use site in another module: filter p xs = {inline} case xs of [] -> [] (x:xs) -> if p x then x : build (\c n -> foldr (filterFB c p) n xs) else build (\c n -> foldr (filterFB c p) n xs) build (\c n -> foldr (filterFB c p) n xs) = {inline} foldr (filterFB (:) p) [] xs = {RULE} filter p xs And we are in an infinite loop again, except that this time the loop is producing an infinitely large *term* (an unrolling of filter) and so the simplifier finally dies with "ticks exhausted" Because of this problem, we make a small change in the occurrence analyser designed to mark functions like "filter" as strong loop breakers on the basis that: 1. The RHS of filter mentions the local function "filterFB" 2. We have a rule which mentions "filterFB" on the LHS and "filter" on the RHS So for each RULE for an *imported* function we are going to add dependency edges between the *local* FVS of the rule LHS and the *local* FVS of the rule RHS. We don't do anything special for RULES on local functions because the standard occurrence analysis stuff is pretty good at getting loop-breakerness correct there. It is important to note that even with this extra hack we aren't always going to get things right. For example, it might be that the rule LHS mentions an imported Id, and another module has a RULE that can rewrite that imported Id to one of our local Ids. Note [Specialising imported functions] (referred to from Specialise) ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ BUT for *automatically-generated* rules, the programmer can't be responsible for the "programmer error" in Note [Rules for imported functions]. In paricular, consider specialising a recursive function defined in another module. If we specialise a recursive function B.g, we get g_spec = .....(B.g Int)..... RULE B.g Int = g_spec Here, g_spec doesn't look recursive, but when the rule fires, it becomes so. And if B.g was mutually recursive, the loop might not be as obvious as it is here. To avoid this, * When specialising a function that is a loop breaker, give a NOINLINE pragma to the specialised function Note [Glomming] ~~~~~~~~~~~~~~~ RULES for imported Ids can make something at the top refer to something at the bottom: f = \x -> B.g (q x) h = \y -> 3 RULE: B.g (q x) = h x Applying this rule makes f refer to h, although f doesn't appear to depend on h. (And, as in Note [Rules for imported functions], the dependency might be more indirect. For example, f might mention C.t rather than B.g, where C.t eventually inlines to B.g.) NOTICE that this cannot happen for rules whose head is a locally-defined function, because we accurately track dependencies through RULES. It only happens for rules whose head is an imported function (B.g in the example above). Solution: - When simplifying, bring all top level identifiers into scope at the start, ignoring the Rec/NonRec structure, so that when 'h' pops up in f's rhs, we find it in the in-scope set (as the simplifier generally expects). This happens in simplTopBinds. - In the occurrence analyser, if there are any out-of-scope occurrences that pop out of the top, which will happen after firing the rule: f = \x -> h x h = \y -> 3 then just glom all the bindings into a single Rec, so that the *next* iteration of the occurrence analyser will sort them all out. This part happens in occurAnalysePgm. ------------------------------------------------------------ Note [Inline rules] ~~~~~~~~~~~~~~~~~~~ None of the above stuff about RULES applies to Inline Rules, stored in a CoreUnfolding. The unfolding, if any, is simplified at the same time as the regular RHS of the function (ie *not* like Note [Rules are visible in their own rec group]), so it should be treated *exactly* like an extra RHS. Or, rather, when computing loop-breaker edges, * If f has an INLINE pragma, and it is active, we treat the INLINE rhs as f's rhs * If it's inactive, we treat f as having no rhs * If it has no INLINE pragma, we look at f's actual rhs There is a danger that we'll be sub-optimal if we see this f = ...f... [INLINE f = ..no f...] where f is recursive, but the INLINE is not. This can just about happen with a sufficiently odd set of rules; eg foo :: Int -> Int {-# INLINE [1] foo #-} foo x = x+1 bar :: Int -> Int {-# INLINE [1] bar #-} bar x = foo x + 1 {-# RULES "foo" [~1] forall x. foo x = bar x #-} Here the RULE makes bar recursive; but it's INLINE pragma remains non-recursive. It's tempting to then say that 'bar' should not be a loop breaker, but an attempt to do so goes wrong in two ways: a) We may get $df = ...$cfoo... $cfoo = ...$df.... [INLINE $cfoo = ...no-$df...] But we want $cfoo to depend on $df explicitly so that we put the bindings in the right order to inline $df in $cfoo and perhaps break the loop altogether. (Maybe this b) Example [eftInt] ~~~~~~~~~~~~~~~ Example (from GHC.Enum): eftInt :: Int# -> Int# -> [Int] eftInt x y = ...(non-recursive)... {-# INLINE [0] eftIntFB #-} eftIntFB :: (Int -> r -> r) -> r -> Int# -> Int# -> r eftIntFB c n x y = ...(non-recursive)... {-# RULES "eftInt" [~1] forall x y. eftInt x y = build (\ c n -> eftIntFB c n x y) "eftIntList" [1] eftIntFB (:) [] = eftInt #-} Note [Specialisation rules] ~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider this group, which is typical of what SpecConstr builds: fs a = ....f (C a).... f x = ....f (C a).... {-# RULE f (C a) = fs a #-} So 'f' and 'fs' are in the same Rec group (since f refers to fs via its RULE). But watch out! If 'fs' is not chosen as a loop breaker, we may get an infinite loop: - the RULE is applied in f's RHS (see Note [Self-recursive rules] in Simplify - fs is inlined (say it's small) - now there's another opportunity to apply the RULE This showed up when compiling Control.Concurrent.Chan.getChanContents. -} ------------------------------------------------------------------ -- occAnalBind ------------------------------------------------------------------ occAnalBind :: OccEnv -- The incoming OccEnv -> ImpRuleEdges -> CoreBind -> UsageDetails -- Usage details of scope -> (UsageDetails, -- Of the whole let(rec) [CoreBind]) occAnalBind env top_env (NonRec binder rhs) body_usage = occAnalNonRecBind env top_env binder rhs body_usage occAnalBind env top_env (Rec pairs) body_usage = occAnalRecBind env top_env pairs body_usage ----------------- occAnalNonRecBind :: OccEnv -> ImpRuleEdges -> Var -> CoreExpr -> UsageDetails -> (UsageDetails, [CoreBind]) occAnalNonRecBind env imp_rule_edges binder rhs body_usage | isTyVar binder -- A type let; we don't gather usage info = (body_usage, [NonRec binder rhs]) | not (binder `usedIn` body_usage) -- It's not mentioned = (body_usage, []) | otherwise -- It's mentioned in the body = (body_usage' +++ rhs_usage4, [NonRec tagged_binder rhs']) where (body_usage', tagged_binder) = tagBinder body_usage binder (rhs_usage1, rhs') = occAnalNonRecRhs env tagged_binder rhs rhs_usage2 = addIdOccs rhs_usage1 (idUnfoldingVars binder) rhs_usage3 = addIdOccs rhs_usage2 (idRuleVars binder) -- See Note [Rules are extra RHSs] and Note [Rule dependency info] rhs_usage4 = maybe rhs_usage3 (addIdOccs rhs_usage3) $ lookupVarEnv imp_rule_edges binder -- See Note [Preventing loops due to imported functions rules] ----------------- occAnalRecBind :: OccEnv -> ImpRuleEdges -> [(Var,CoreExpr)] -> UsageDetails -> (UsageDetails, [CoreBind]) occAnalRecBind env imp_rule_edges pairs body_usage = foldr occAnalRec (body_usage, []) sccs -- For a recursive group, we -- * occ-analyse all the RHSs -- * compute strongly-connected components -- * feed those components to occAnalRec -- See Note [Recursive bindings: the grand plan] where sccs :: [SCC Details] sccs = {-# SCC "occAnalBind.scc" #-} stronglyConnCompFromEdgedVerticesUniq nodes nodes :: [LetrecNode] nodes = {-# SCC "occAnalBind.assoc" #-} map (makeNode env imp_rule_edges bndr_set) pairs bndr_set = mkVarSet (map fst pairs) ----------------------------- occAnalRec :: SCC Details -> (UsageDetails, [CoreBind]) -> (UsageDetails, [CoreBind]) -- The NonRec case is just like a Let (NonRec ...) above occAnalRec (AcyclicSCC (ND { nd_bndr = bndr, nd_rhs = rhs, nd_uds = rhs_uds})) (body_uds, binds) | not (bndr `usedIn` body_uds) = (body_uds, binds) -- See Note [Dead code] | otherwise -- It's mentioned in the body = (body_uds' +++ rhs_uds, NonRec tagged_bndr rhs : binds) where (body_uds', tagged_bndr) = tagBinder body_uds bndr -- The Rec case is the interesting one -- See Note [Recursive bindings: the grand plan] -- See Note [Loop breaking] occAnalRec (CyclicSCC details_s) (body_uds, binds) | not (any (`usedIn` body_uds) bndrs) -- NB: look at body_uds, not total_uds = (body_uds, binds) -- See Note [Dead code] | otherwise -- At this point we always build a single Rec = -- pprTrace "occAnalRec" (vcat -- [ text "weak_fvs" <+> ppr weak_fvs -- , text "tagged details" <+> ppr tagged_details_s -- , text "lb nodes" <+> ppr loop_breaker_nodes]) (final_uds, Rec pairs : binds) where bndrs = map nd_bndr details_s bndr_set = mkVarSet bndrs ---------------------------- -- Compute usage details total_uds = foldl add_uds body_uds details_s final_uds = total_uds `minusVarEnv` bndr_set add_uds usage_so_far nd = usage_so_far +++ nd_uds nd ------------------------------ -- See Note [Choosing loop breakers] for loop_breaker_nodes loop_breaker_nodes :: [LetrecNode] loop_breaker_nodes = mkLoopBreakerNodes bndr_set total_uds details_s ------------------------------ weak_fvs :: VarSet weak_fvs = mapUnionVarSet nd_weak details_s --------------------------- -- Now reconstruct the cycle pairs :: [(Id,CoreExpr)] pairs | isEmptyVarSet weak_fvs = reOrderNodes 0 bndr_set weak_fvs loop_breaker_nodes [] | otherwise = loopBreakNodes 0 bndr_set weak_fvs loop_breaker_nodes [] -- If weak_fvs is empty, the loop_breaker_nodes will include -- all the edges in the original scope edges [remember, -- weak_fvs is the difference between scope edges and -- lb-edges], so a fresh SCC computation would yield a -- single CyclicSCC result; and reOrderNodes deals with -- exactly that case ------------------------------------------------------------------ -- Loop breaking ------------------------------------------------------------------ type Binding = (Id,CoreExpr) loopBreakNodes :: Int -> VarSet -- All binders -> VarSet -- Binders whose dependencies may be "missing" -- See Note [Weak loop breakers] -> [LetrecNode] -> [Binding] -- Append these to the end -> [Binding] {- loopBreakNodes is applied to the list of nodes for a cyclic strongly connected component (there's guaranteed to be a cycle). It returns the same nodes, but a) in a better order, b) with some of the Ids having a IAmALoopBreaker pragma The "loop-breaker" Ids are sufficient to break all cycles in the SCC. This means that the simplifier can guarantee not to loop provided it never records an inlining for these no-inline guys. Furthermore, the order of the binds is such that if we neglect dependencies on the no-inline Ids then the binds are topologically sorted. This means that the simplifier will generally do a good job if it works from top bottom, recording inlinings for any Ids which aren't marked as "no-inline" as it goes. -} -- Return the bindings sorted into a plausible order, and marked with loop breakers. loopBreakNodes depth bndr_set weak_fvs nodes binds = go (stronglyConnCompFromEdgedVerticesUniqR nodes) binds where go [] binds = binds go (scc:sccs) binds = loop_break_scc scc (go sccs binds) loop_break_scc scc binds = case scc of AcyclicSCC node -> mk_non_loop_breaker weak_fvs node : binds CyclicSCC nodes -> reOrderNodes depth bndr_set weak_fvs nodes binds ---------------------------------- reOrderNodes :: Int -> VarSet -> VarSet -> [LetrecNode] -> [Binding] -> [Binding] -- Choose a loop breaker, mark it no-inline, -- and call loopBreakNodes on the rest reOrderNodes _ _ _ [] _ = panic "reOrderNodes" reOrderNodes _ _ _ [node] binds = mk_loop_breaker node : binds reOrderNodes depth bndr_set weak_fvs (node : nodes) binds = -- pprTrace "reOrderNodes" (text "unchosen" <+> ppr unchosen $$ -- text "chosen" <+> ppr chosen_nodes) $ loopBreakNodes new_depth bndr_set weak_fvs unchosen $ (map mk_loop_breaker chosen_nodes ++ binds) where (chosen_nodes, unchosen) = chooseLoopBreaker approximate_lb (nd_score (fstOf3 node)) [node] [] nodes approximate_lb = depth >= 2 new_depth | approximate_lb = 0 | otherwise = depth+1 -- After two iterations (d=0, d=1) give up -- and approximate, returning to d=0 mk_loop_breaker :: LetrecNode -> Binding mk_loop_breaker (ND { nd_bndr = bndr, nd_rhs = rhs}, _, _) = (setIdOccInfo bndr strongLoopBreaker, rhs) mk_non_loop_breaker :: VarSet -> LetrecNode -> Binding -- See Note [Weak loop breakers] mk_non_loop_breaker weak_fvs (ND { nd_bndr = bndr, nd_rhs = rhs}, _, _) | bndr `elemVarSet` weak_fvs = (setIdOccInfo bndr weakLoopBreaker, rhs) | otherwise = (bndr, rhs) ---------------------------------- chooseLoopBreaker :: Bool -- True <=> Too many iterations, -- so approximate -> NodeScore -- Best score so far -> [LetrecNode] -- Nodes with this score -> [LetrecNode] -- Nodes with higher scores -> [LetrecNode] -- Unprocessed nodes -> ([LetrecNode], [LetrecNode]) -- This loop looks for the bind with the lowest score -- to pick as the loop breaker. The rest accumulate in chooseLoopBreaker _ _ loop_nodes acc [] = (loop_nodes, acc) -- Done -- If approximate_loop_breaker is True, we pick *all* -- nodes with lowest score, else just one -- See Note [Complexity of loop breaking] chooseLoopBreaker approx_lb loop_sc loop_nodes acc (node : nodes) | approx_lb , rank sc == rank loop_sc = chooseLoopBreaker approx_lb loop_sc (node : loop_nodes) acc nodes | sc `betterLB` loop_sc -- Better score so pick this new one = chooseLoopBreaker approx_lb sc [node] (loop_nodes ++ acc) nodes | otherwise -- Worse score so don't pick it = chooseLoopBreaker approx_lb loop_sc loop_nodes (node : acc) nodes where sc = nd_score (fstOf3 node) {- Note [Complexity of loop breaking] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The loop-breaking algorithm knocks out one binder at a time, and performs a new SCC analysis on the remaining binders. That can behave very badly in tightly-coupled groups of bindings; in the worst case it can be (N**2)*log N, because it does a full SCC on N, then N-1, then N-2 and so on. To avoid this, we switch plans after 2 (or whatever) attempts: Plan A: pick one binder with the lowest score, make it a loop breaker, and try again Plan B: pick *all* binders with the lowest score, make them all loop breakers, and try again Since there are only a small finite number of scores, this will terminate in a constant number of iterations, rather than O(N) iterations. You might thing that it's very unlikely, but RULES make it much more likely. Here's a real example from Trac #1969: Rec { $dm = \d.\x. op d {-# RULES forall d. $dm Int d = $s$dm1 forall d. $dm Bool d = $s$dm2 #-} dInt = MkD .... opInt ... dInt = MkD .... opBool ... opInt = $dm dInt opBool = $dm dBool $s$dm1 = \x. op dInt $s$dm2 = \x. op dBool } The RULES stuff means that we can't choose $dm as a loop breaker (Note [Choosing loop breakers]), so we must choose at least (say) opInt *and* opBool, and so on. The number of loop breakders is linear in the number of instance declarations. Note [Loop breakers and INLINE/INLINABLE pragmas] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Avoid choosing a function with an INLINE pramga as the loop breaker! If such a function is mutually-recursive with a non-INLINE thing, then the latter should be the loop-breaker. It's vital to distinguish between INLINE and INLINABLE (the Bool returned by hasStableCoreUnfolding_maybe). If we start with Rec { {-# INLINABLE f #-} f x = ...f... } and then worker/wrapper it through strictness analysis, we'll get Rec { {-# INLINABLE $wf #-} $wf p q = let x = (p,q) in ...f... {-# INLINE f #-} f x = case x of (p,q) -> $wf p q } Now it is vital that we choose $wf as the loop breaker, so we can inline 'f' in '$wf'. Note [DFuns should not be loop breakers] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ It's particularly bad to make a DFun into a loop breaker. See Note [How instance declarations are translated] in TcInstDcls We give DFuns a higher score than ordinary CONLIKE things because if there's a choice we want the DFun to be the non-loop breaker. Eg rec { sc = /\ a \$dC. $fBWrap (T a) ($fCT @ a $dC) $fCT :: forall a_afE. (Roman.C a_afE) => Roman.C (Roman.T a_afE) {-# DFUN #-} $fCT = /\a \$dC. MkD (T a) ((sc @ a $dC) |> blah) ($ctoF @ a $dC) } Here 'sc' (the superclass) looks CONLIKE, but we'll never get to it if we can't unravel the DFun first. Note [Constructor applications] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ It's really really important to inline dictionaries. Real example (the Enum Ordering instance from GHC.Base): rec f = \ x -> case d of (p,q,r) -> p x g = \ x -> case d of (p,q,r) -> q x d = (v, f, g) Here, f and g occur just once; but we can't inline them into d. On the other hand we *could* simplify those case expressions if we didn't stupidly choose d as the loop breaker. But we won't because constructor args are marked "Many". Inlining dictionaries is really essential to unravelling the loops in static numeric dictionaries, see GHC.Float. Note [Closure conversion] ~~~~~~~~~~~~~~~~~~~~~~~~~ We treat (\x. C p q) as a high-score candidate in the letrec scoring algorithm. The immediate motivation came from the result of a closure-conversion transformation which generated code like this: data Clo a b = forall c. Clo (c -> a -> b) c ($:) :: Clo a b -> a -> b Clo f env $: x = f env x rec { plus = Clo plus1 () ; plus1 _ n = Clo plus2 n ; plus2 Zero n = n ; plus2 (Succ m) n = Succ (plus $: m $: n) } If we inline 'plus' and 'plus1', everything unravels nicely. But if we choose 'plus1' as the loop breaker (which is entirely possible otherwise), the loop does not unravel nicely. @occAnalRhs@ deals with the question of bindings where the Id is marked by an INLINE pragma. For these we record that anything which occurs in its RHS occurs many times. This pessimistically assumes that ths inlined binder also occurs many times in its scope, but if it doesn't we'll catch it next time round. At worst this costs an extra simplifier pass. ToDo: try using the occurrence info for the inline'd binder. [March 97] We do the same for atomic RHSs. Reason: see notes with loopBreakSCC. [June 98, SLPJ] I've undone this change; I don't understand it. See notes with loopBreakSCC. ************************************************************************ * * Making nodes * * ************************************************************************ -} type ImpRuleEdges = IdEnv IdSet -- Mapping from FVs of imported RULE LHSs to RHS FVs noImpRuleEdges :: ImpRuleEdges noImpRuleEdges = emptyVarEnv type LetrecNode = Node Unique Details -- Node comes from Digraph -- The Unique key is gotten from the Id data Details = ND { nd_bndr :: Id -- Binder , nd_rhs :: CoreExpr -- RHS, already occ-analysed , nd_uds :: UsageDetails -- Usage from RHS, and RULES, and stable unfoldings -- ignoring phase (ie assuming all are active) -- See Note [Forming Rec groups] , nd_inl :: IdSet -- Free variables of -- the stable unfolding (if present and active) -- or the RHS (if not) -- but excluding any RULES -- This is the IdSet that may be used if the Id is inlined , nd_weak :: IdSet -- Binders of this Rec that are mentioned in nd_uds -- but are *not* in nd_inl. These are the ones whose -- dependencies might not be respected by loop_breaker_nodes -- See Note [Weak loop breakers] , nd_active_rule_fvs :: IdSet -- Free variables of the RHS of active RULES , nd_score :: NodeScore } instance Outputable Details where ppr nd = text "ND" <> braces (sep [ text "bndr =" <+> ppr (nd_bndr nd) , text "uds =" <+> ppr (nd_uds nd) , text "inl =" <+> ppr (nd_inl nd) , text "weak =" <+> ppr (nd_weak nd) , text "rule =" <+> ppr (nd_active_rule_fvs nd) ]) -- The NodeScore is compared lexicographically; -- e.g. lower rank wins regardless of size type NodeScore = ( Int -- Rank: lower => more likely to be picked as loop breaker , Int -- Size of rhs: higher => more likely to be picked as LB -- Maxes out at maxExprSize; we just use it to prioritise -- small functions , Bool ) -- Was it a loop breaker before? -- True => more likely to be picked -- Note [Loop breakers, node scoring, and stability] rank :: NodeScore -> Int rank (r, _, _) = r makeNode :: OccEnv -> ImpRuleEdges -> VarSet -> (Var, CoreExpr) -> LetrecNode -- See Note [Recursive bindings: the grand plan] makeNode env imp_rule_edges bndr_set (bndr, rhs) = (details, varUnique bndr, nonDetKeysUFM node_fvs) -- It's OK to use nonDetKeysUFM here as stronglyConnCompFromEdgedVerticesR -- is still deterministic with edges in nondeterministic order as -- explained in Note [Deterministic SCC] in Digraph. where details = ND { nd_bndr = bndr , nd_rhs = rhs' , nd_uds = rhs_usage3 , nd_inl = inl_fvs , nd_weak = node_fvs `minusVarSet` inl_fvs , nd_active_rule_fvs = active_rule_fvs , nd_score = pprPanic "makeNodeDetails" (ppr bndr) } -- Constructing the edges for the main Rec computation -- See Note [Forming Rec groups] (rhs_usage1, rhs') = occAnalRecRhs env rhs rhs_usage2 = addIdOccs rhs_usage1 all_rule_fvs -- Note [Rules are extra RHSs] -- Note [Rule dependency info] rhs_usage3 = case mb_unf_fvs of Just unf_fvs -> addIdOccs rhs_usage2 unf_fvs Nothing -> rhs_usage2 node_fvs = udFreeVars bndr_set rhs_usage3 -- Finding the free variables of the rules is_active = occ_rule_act env :: Activation -> Bool rules = filterOut isBuiltinRule (idCoreRules bndr) rules_w_fvs :: [(Activation, VarSet)] -- Find the RHS fvs rules_w_fvs = maybe id (\ids -> ((AlwaysActive, ids):)) (lookupVarEnv imp_rule_edges bndr) -- See Note [Preventing loops due to imported functions rules] [ (ru_act rule, fvs) | rule <- rules , let fvs = exprFreeVars (ru_rhs rule) `delVarSetList` ru_bndrs rule , not (isEmptyVarSet fvs) ] all_rule_fvs = rule_lhs_fvs `unionVarSet` rule_rhs_fvs rule_rhs_fvs = mapUnionVarSet snd rules_w_fvs rule_lhs_fvs = mapUnionVarSet (\ru -> exprsFreeVars (ru_args ru) `delVarSetList` ru_bndrs ru) rules active_rule_fvs = unionVarSets [fvs | (a,fvs) <- rules_w_fvs, is_active a] -- Finding the free variables of the INLINE pragma (if any) unf = realIdUnfolding bndr -- Ignore any current loop-breaker flag mb_unf_fvs = stableUnfoldingVars unf -- Find the "nd_inl" free vars; for the loop-breaker phase inl_fvs = case mb_unf_fvs of Nothing -> udFreeVars bndr_set rhs_usage1 -- No INLINE, use RHS Just unf_fvs -> unf_fvs -- We could check for an *active* INLINE (returning -- emptyVarSet for an inactive one), but is_active -- isn't the right thing (it tells about -- RULE activation), so we'd need more plumbing mkLoopBreakerNodes :: VarSet -> UsageDetails -> [Details] -> [LetrecNode] -- Does three things -- a) tag each binder with its occurrence info -- b) add a NodeScore to each node -- c) make a Node with the right dependency edges for -- the loop-breaker SCC analysis mkLoopBreakerNodes bndr_set total_uds details_s = map mk_lb_node details_s where mk_lb_node nd@(ND { nd_bndr = bndr, nd_rhs = rhs, nd_inl = inl_fvs }) = (nd', varUnique bndr, nonDetKeysUFM lb_deps) -- It's OK to use nonDetKeysUFM here as -- stronglyConnCompFromEdgedVerticesR is still deterministic with edges -- in nondeterministic order as explained in -- Note [Deterministic SCC] in Digraph. where nd' = nd { nd_bndr = bndr', nd_score = score } bndr' = setBinderOcc total_uds bndr score = nodeScore bndr bndr' rhs lb_deps lb_deps = extendFvs_ rule_fv_env inl_fvs rule_fv_env :: IdEnv IdSet -- Maps a variable f to the variables from this group -- mentioned in RHS of active rules for f -- Domain is *subset* of bound vars (others have no rule fvs) rule_fv_env = transClosureFV (mkVarEnv init_rule_fvs) init_rule_fvs -- See Note [Finding rule RHS free vars] = [ (b, trimmed_rule_fvs) | ND { nd_bndr = b, nd_active_rule_fvs = rule_fvs } <- details_s , let trimmed_rule_fvs = rule_fvs `intersectVarSet` bndr_set , not (isEmptyVarSet trimmed_rule_fvs) ] ------------------------------------------ nodeScore :: Id -- Binder has old occ-info (just for loop-breaker-ness) -> Id -- Binder with new occ-info -> CoreExpr -- RHS -> VarSet -- Loop-breaker dependencies -> NodeScore nodeScore old_bndr new_bndr bind_rhs lb_deps | not (isId old_bndr) -- A type or cercion variable is never a loop breaker = (100, 0, False) | old_bndr `elemVarSet` lb_deps -- Self-recursive things are great loop breakers = (0, 0, True) -- See Note [Self-recursion and loop breakers] | otherwise -- An Id has an unfolding = case id_unfolding of DFunUnfolding { df_args = args } -- Never choose a DFun as a loop breaker -- Note [DFuns should not be loop breakers] -> (9, length args, is_lb) CoreUnfolding { uf_src = src, uf_tmpl = unf_rhs, uf_guidance = guide } | isStableSource src -> case guide of UnfWhen {} -> (6, cheapExprSize unf_rhs, is_lb) UnfIfGoodArgs { ug_size = size} -> (3, size, is_lb) UnfNever -> (0, 0, is_lb) -- See Note [Loop breakers and INLINE/INLINABLE pragmas] for -- the 6 vs 3 choice -- Note that this case hits /all/ stable unfoldings, so we -- never look at 'bind_rhs' for stable unfoldings. That's right, because -- 'rhs' is irrelevant for inlining things with a stable unfolding -- Data structures are more important than INLINE pragmas -- so that dictionary/method recursion unravels _ | exprIsTrivial bind_rhs -> mk_score 10 -- Practically certain to be inlined -- Used to have also: && not (isExportedId bndr) -- But I found this sometimes cost an extra iteration when we have -- rec { d = (a,b); a = ...df...; b = ...df...; df = d } -- where df is the exported dictionary. Then df makes a really -- bad choice for loop breaker | is_con_app bind_rhs -- Data types help with cases: Note [Constructor applications] -> mk_score 5 | isOneOcc (idOccInfo new_bndr) -> mk_score 2 -- Likely to be inlined | canUnfold id_unfolding -- The Id has some kind of unfolding -> mk_score 1 | otherwise -> (0, 0, is_lb) where mk_score :: Int -> NodeScore mk_score rank = (rank, rhs_size, is_lb) is_lb = isStrongLoopBreaker (idOccInfo old_bndr) rhs_size = case id_unfolding of CoreUnfolding { uf_guidance = guidance } | UnfIfGoodArgs { ug_size = size } <- guidance -> size _ -> cheapExprSize bind_rhs id_unfolding = realIdUnfolding old_bndr -- realIdUnfolding: Ignore loop-breaker-ness here because -- that is what we are setting! -- Checking for a constructor application -- Cheap and cheerful; the simplifier moves casts out of the way -- The lambda case is important to spot x = /\a. C (f a) -- which comes up when C is a dictionary constructor and -- f is a default method. -- Example: the instance for Show (ST s a) in GHC.ST -- -- However we *also* treat (\x. C p q) as a con-app-like thing, -- Note [Closure conversion] is_con_app (Var v) = isConLikeId v is_con_app (App f _) = is_con_app f is_con_app (Lam _ e) = is_con_app e is_con_app (Tick _ e) = is_con_app e is_con_app _ = False maxExprSize :: Int maxExprSize = 20 -- Rather arbitrary cheapExprSize :: CoreExpr -> Int -- Maxes out at maxExprSize cheapExprSize e = go 0 e where go n e | n >= maxExprSize = n | otherwise = go1 n e go1 n (Var {}) = n+1 go1 n (Lit {}) = n+1 go1 n (Type {}) = n go1 n (Coercion {}) = n go1 n (Tick _ e) = go1 n e go1 n (Cast e _) = go1 n e go1 n (App f a) = go (go1 n f) a go1 n (Lam b e) | isTyVar b = go1 n e | otherwise = go (n+1) e go1 n (Let b e) = gos (go1 n e) (rhssOfBind b) go1 n (Case e _ _ as) = gos (go1 n e) (rhssOfAlts as) gos n [] = n gos n (e:es) | n >= maxExprSize = n | otherwise = gos (go1 n e) es betterLB :: NodeScore -> NodeScore -> Bool -- If n1 `betterLB` n2 then choose n1 as the loop breaker betterLB (rank1, size1, lb1) (rank2, size2, _) | rank1 < rank2 = True | rank1 > rank2 = False | size1 < size2 = False -- Make the bigger n2 into the loop breaker | size1 > size2 = True | lb1 = True -- Tie-break: if n1 was a loop breaker before, choose it | otherwise = False -- See Note [Loop breakers, node scoring, and stability] {- Note [Self-recursion and loop breakers] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ If we have rec { f = ...f...g... ; g = .....f... } then 'f' has to be a loop breaker anyway, so we may as well choose it right away, so that g can inline freely. This is really just a cheap hack. Consider rec { f = ...g... ; g = ..f..h... ; h = ...f....} Here f or g are better loop breakers than h; but we might accidentally choose h. Finding the minimal set of loop breakers is hard. Note [Loop breakers, node scoring, and stability] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ To choose a loop breaker, we give a NodeScore to each node in the SCC, and pick the one with the best score (according to 'betterLB'). We need to be jolly careful (Trac #12425, #12234) about the stability of this choice. Suppose we have let rec { f = ...g...g... ; g = ...f...f... } in case x of True -> ...f.. False -> ..f... In each iteration of the simplifier the occurrence analyser OccAnal chooses a loop breaker. Suppose in iteration 1 it choose g as the loop breaker. That means it is free to inline f. Suppose that GHC decides to inline f in the branches of the case, but (for some reason; eg it is not satureated) in the rhs of g. So we get let rec { f = ...g...g... ; g = ...f...f... } in case x of True -> ...g...g..... False -> ..g..g.... Now suppose that, for some reason, in the next iteration the occurrence analyser chooses f as the loop breaker, so it can freely inline g. And again for some reason the simplifier inlines g at its calls in the case branches, but not in the RHS of f. Then we get let rec { f = ...g...g... ; g = ...f...f... } in case x of True -> ...(...f...f...)...(...f..f..)..... False -> ..(...f...f...)...(..f..f...).... You can see where this is going! Each iteration of the simplifier doubles the number of calls to f or g. No wonder GHC is slow! (In the particular example in comment:3 of #12425, f and g are the two mutually recursive fmap instances for CondT and Result. They are both marked INLINE which, oddly, is why they don't inline in each other's RHS, because the call there is not saturated.) The root cause is that we flip-flop on our choice of loop breaker. I always thought it didn't matter, and indeed for any single iteration to terminate, it doesn't matter. But when we iterate, it matters a lot!! So The Plan is this: If there is a tie, choose the node that was a loop breaker last time round Hence the is_lb field of NodeScore ************************************************************************ * * Right hand sides * * ************************************************************************ -} occAnalRecRhs :: OccEnv -> CoreExpr -- Rhs -> (UsageDetails, CoreExpr) -- Returned usage details covers only the RHS, -- and *not* the RULE or INLINE template for the Id occAnalRecRhs env rhs = occAnal (rhsCtxt env) rhs occAnalNonRecRhs :: OccEnv -> Id -> CoreExpr -- Binder and rhs -- Binder is already tagged with occurrence info -> (UsageDetails, CoreExpr) -- Returned usage details covers only the RHS, -- and *not* the RULE or INLINE template for the Id occAnalNonRecRhs env bndr rhs = occAnal rhs_env rhs where -- See Note [Cascading inlines] env1 | certainly_inline = env | otherwise = rhsCtxt env -- See Note [Use one-shot info] rhs_env = env1 { occ_one_shots = argOneShots OneShotLam dmd } certainly_inline -- See Note [Cascading inlines] = case idOccInfo bndr of OneOcc in_lam one_br _ -> not in_lam && one_br && active && not_stable _ -> False dmd = idDemandInfo bndr active = isAlwaysActive (idInlineActivation bndr) not_stable = not (isStableUnfolding (idUnfolding bndr)) {- Note [Cascading inlines] ~~~~~~~~~~~~~~~~~~~~~~~~ By default we use an rhsCtxt for the RHS of a binding. This tells the occ anal n that it's looking at an RHS, which has an effect in occAnalApp. In particular, for constructor applications, it makes the arguments appear to have NoOccInfo, so that we don't inline into them. Thus x = f y k = Just x we do not want to inline x. But there's a problem. Consider x1 = a0 : [] x2 = a1 : x1 x3 = a2 : x2 g = f x3 First time round, it looks as if x1 and x2 occur as an arg of a let-bound constructor ==> give them a many-occurrence. But then x3 is inlined (unconditionally as it happens) and next time round, x2 will be, and the next time round x1 will be Result: multiple simplifier iterations. Sigh. So, when analysing the RHS of x3 we notice that x3 will itself definitely inline the next time round, and so we analyse x3's rhs in an ordinary context, not rhsCtxt. Hence the "certainly_inline" stuff. Annoyingly, we have to approximate SimplUtils.preInlineUnconditionally. If we say "yes" when preInlineUnconditionally says "no" the simplifier iterates indefinitely: x = f y k = Just x inline ==> k = Just (f y) float ==> x1 = f y k = Just x1 This is worse than the slow cascade, so we only want to say "certainly_inline" if it really is certain. Look at the note with preInlineUnconditionally for the various clauses. ************************************************************************ * * Expressions * * ************************************************************************ -} occAnal :: OccEnv -> CoreExpr -> (UsageDetails, -- Gives info only about the "interesting" Ids CoreExpr) occAnal _ expr@(Type _) = (emptyDetails, expr) occAnal _ expr@(Lit _) = (emptyDetails, expr) occAnal env expr@(Var v) = (mkOneOcc env v False, expr) -- At one stage, I gathered the idRuleVars for v here too, -- which in a way is the right thing to do. -- But that went wrong right after specialisation, when -- the *occurrences* of the overloaded function didn't have any -- rules in them, so the *specialised* versions looked as if they -- weren't used at all. occAnal _ (Coercion co) = (addIdOccs emptyDetails (coVarsOfCo co), Coercion co) -- See Note [Gather occurrences of coercion variables] {- Note [Gather occurrences of coercion variables] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We need to gather info about what coercion variables appear, so that we can sort them into the right place when doing dependency analysis. -} occAnal env (Tick tickish body) | tickish `tickishScopesLike` SoftScope = (usage, Tick tickish body') | Breakpoint _ ids <- tickish = (usage_lam +++ mkVarEnv (zip ids (repeat NoOccInfo)), Tick tickish body') -- never substitute for any of the Ids in a Breakpoint | otherwise = (usage_lam, Tick tickish body') where !(usage,body') = occAnal env body -- for a non-soft tick scope, we can inline lambdas only usage_lam = mapVarEnv markInsideLam usage occAnal env (Cast expr co) = case occAnal env expr of { (usage, expr') -> let usage1 = markManyIf (isRhsEnv env) usage usage2 = addIdOccs usage1 (coVarsOfCo co) -- See Note [Gather occurrences of coercion variables] in (usage2, Cast expr' co) -- If we see let x = y `cast` co -- then mark y as 'Many' so that we don't -- immediately inline y again. } occAnal env app@(App _ _) = occAnalApp env (collectArgsTicks tickishFloatable app) -- Ignore type variables altogether -- (a) occurrences inside type lambdas only not marked as InsideLam -- (b) type variables not in environment occAnal env (Lam x body) | isTyVar x = case occAnal env body of { (body_usage, body') -> (body_usage, Lam x body') } -- For value lambdas we do a special hack. Consider -- (\x. \y. ...x...) -- If we did nothing, x is used inside the \y, so would be marked -- as dangerous to dup. But in the common case where the abstraction -- is applied to two arguments this is over-pessimistic. -- So instead, we just mark each binder with its occurrence -- info in the *body* of the multiple lambda. -- Then, the simplifier is careful when partially applying lambdas. occAnal env expr@(Lam _ _) = case occAnal env_body body of { (body_usage, body') -> let (final_usage, tagged_binders) = tagLamBinders body_usage binders' -- Use binders' to put one-shot info on the lambdas really_final_usage | all isOneShotBndr binders' = final_usage | otherwise = mapVarEnv markInsideLam final_usage in (really_final_usage, mkLams tagged_binders body') } where (binders, body) = collectBinders expr (env_body, binders') = oneShotGroup env binders occAnal env (Case scrut bndr ty alts) = case occ_anal_scrut scrut alts of { (scrut_usage, scrut') -> case mapAndUnzip occ_anal_alt alts of { (alts_usage_s, alts') -> let alts_usage = foldr combineAltsUsageDetails emptyDetails alts_usage_s (alts_usage1, tagged_bndr) = tag_case_bndr alts_usage bndr total_usage = scrut_usage +++ alts_usage1 in total_usage `seq` (total_usage, Case scrut' tagged_bndr ty alts') }} where -- Note [Case binder usage] -- ~~~~~~~~~~~~~~~~~~~~~~~~ -- The case binder gets a usage of either "many" or "dead", never "one". -- Reason: we like to inline single occurrences, to eliminate a binding, -- but inlining a case binder *doesn't* eliminate a binding. -- We *don't* want to transform -- case x of w { (p,q) -> f w } -- into -- case x of w { (p,q) -> f (p,q) } tag_case_bndr usage bndr = case lookupVarEnv usage bndr of Nothing -> (usage, setIdOccInfo bndr IAmDead) Just _ -> (usage `delVarEnv` bndr, setIdOccInfo bndr NoOccInfo) alt_env = mkAltEnv env scrut bndr occ_anal_alt = occAnalAlt alt_env occ_anal_scrut (Var v) (alt1 : other_alts) | not (null other_alts) || not (isDefaultAlt alt1) = (mkOneOcc env v True, Var v) -- The 'True' says that the variable occurs -- in an interesting context; the case has -- at least one non-default alternative occ_anal_scrut (Tick t e) alts | t `tickishScopesLike` SoftScope -- No reason to not look through all ticks here, but only -- for soft-scoped ticks we can do so without having to -- update returned occurance info (see occAnal) = second (Tick t) $ occ_anal_scrut e alts occ_anal_scrut scrut _alts = occAnal (vanillaCtxt env) scrut -- No need for rhsCtxt occAnal env (Let bind body) = case occAnal env body of { (body_usage, body') -> case occAnalBind env noImpRuleEdges bind body_usage of { (final_usage, new_binds) -> (final_usage, mkLets new_binds body') }} occAnalArgs :: OccEnv -> [CoreExpr] -> [OneShots] -> (UsageDetails, [CoreExpr]) occAnalArgs _ [] _ = (emptyDetails, []) occAnalArgs env (arg:args) one_shots | isTypeArg arg = case occAnalArgs env args one_shots of { (uds, args') -> (uds, arg:args') } | otherwise = case argCtxt env one_shots of { (arg_env, one_shots') -> case occAnal arg_env arg of { (uds1, arg') -> case occAnalArgs env args one_shots' of { (uds2, args') -> (uds1 +++ uds2, arg':args') }}} {- Applications are dealt with specially because we want the "build hack" to work. Note [Arguments of let-bound constructors] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider f x = let y = expensive x in let z = (True,y) in (case z of {(p,q)->q}, case z of {(p,q)->q}) We feel free to duplicate the WHNF (True,y), but that means that y may be duplicated thereby. If we aren't careful we duplicate the (expensive x) call! Constructors are rather like lambdas in this way. -} occAnalApp :: OccEnv -> (Expr CoreBndr, [Arg CoreBndr], [Tickish Id]) -> (UsageDetails, Expr CoreBndr) occAnalApp env (Var fun, args, ticks) | null ticks = (uds, mkApps (Var fun) args') | otherwise = (uds, mkTicks ticks $ mkApps (Var fun) args') where uds = fun_uds +++ final_args_uds !(args_uds, args') = occAnalArgs env args one_shots !final_args_uds | isRhsEnv env && is_exp = mapVarEnv markInsideLam args_uds | otherwise = args_uds -- We mark the free vars of the argument of a constructor or PAP -- as "inside-lambda", if it is the RHS of a let(rec). -- This means that nothing gets inlined into a constructor or PAP -- argument position, which is what we want. Typically those -- constructor arguments are just variables, or trivial expressions. -- We use inside-lam because it's like eta-expanding the PAP. -- -- This is the *whole point* of the isRhsEnv predicate -- See Note [Arguments of let-bound constructors] n_val_args = valArgCount args fun_uds = mkOneOcc env fun (n_val_args > 0) is_exp = isExpandableApp fun n_val_args -- See Note [CONLIKE pragma] in BasicTypes -- The definition of is_exp should match that in -- Simplify.prepareRhs one_shots = argsOneShots (idStrictness fun) n_val_args -- See Note [Use one-shot info] occAnalApp env (fun, args, ticks) = (fun_uds +++ args_uds, mkTicks ticks $ mkApps fun' args') where !(fun_uds, fun') = occAnal (addAppCtxt env args) fun -- The addAppCtxt is a bit cunning. One iteration of the simplifier -- often leaves behind beta redexs like -- (\x y -> e) a1 a2 -- Here we would like to mark x,y as one-shot, and treat the whole -- thing much like a let. We do this by pushing some True items -- onto the context stack. !(args_uds, args') = occAnalArgs env args [] markManyIf :: Bool -- If this is true -> UsageDetails -- Then do markMany on this -> UsageDetails markManyIf True uds = mapVarEnv markMany uds markManyIf False uds = uds {- Note [Use one-shot information] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The occurrrence analyser propagates one-shot-lambda information in two situations: * Applications: eg build (\c n -> blah) Propagate one-shot info from the strictness signature of 'build' to the \c n. This strictness signature can come from a module interface, in the case of an imported function, or from a previous run of the demand analyser. * Let-bindings: eg let f = \c. let ... in \n -> blah in (build f, build f) Propagate one-shot info from the demanand-info on 'f' to the lambdas in its RHS (which may not be syntactically at the top) This information must have come from a previous run of the demanand analyser. Previously, the demand analyser would *also* set the one-shot information, but that code was buggy (see #11770), so doing it only in on place, namely here, is saner. Note [Binders in case alternatives] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider case x of y { (a,b) -> f y } We treat 'a', 'b' as dead, because they don't physically occur in the case alternative. (Indeed, a variable is dead iff it doesn't occur in its scope in the output of OccAnal.) It really helps to know when binders are unused. See esp the call to isDeadBinder in Simplify.mkDupableAlt In this example, though, the Simplifier will bring 'a' and 'b' back to life, beause it binds 'y' to (a,b) (imagine got inlined and scrutinised y). -} occAnalAlt :: (OccEnv, Maybe (Id, CoreExpr)) -> CoreAlt -> (UsageDetails, Alt IdWithOccInfo) occAnalAlt (env, scrut_bind) (con, bndrs, rhs) = case occAnal env rhs of { (rhs_usage1, rhs1) -> let (alt_usg, tagged_bndrs) = tagLamBinders rhs_usage1 bndrs -- See Note [Binders in case alternatives] (alt_usg', rhs2) = wrapAltRHS env scrut_bind alt_usg tagged_bndrs rhs1 in (alt_usg', (con, tagged_bndrs, rhs2)) } wrapAltRHS :: OccEnv -> Maybe (Id, CoreExpr) -- proxy mapping generated by mkAltEnv -> UsageDetails -- usage for entire alt (p -> rhs) -> [Var] -- alt binders -> CoreExpr -- alt RHS -> (UsageDetails, CoreExpr) wrapAltRHS env (Just (scrut_var, let_rhs)) alt_usg bndrs alt_rhs | occ_binder_swap env , scrut_var `usedIn` alt_usg -- bndrs are not be present in alt_usg so this -- handles condition (a) in Note [Binder swap] , not captured -- See condition (b) in Note [Binder swap] = ( alt_usg' +++ let_rhs_usg , Let (NonRec tagged_scrut_var let_rhs') alt_rhs ) where captured = any (`usedIn` let_rhs_usg) bndrs -- The rhs of the let may include coercion variables -- if the scrutinee was a cast, so we must gather their -- usage. See Note [Gather occurrences of coercion variables] (let_rhs_usg, let_rhs') = occAnal env let_rhs (alt_usg', tagged_scrut_var) = tagBinder alt_usg scrut_var wrapAltRHS _ _ alt_usg _ alt_rhs = (alt_usg, alt_rhs) {- ************************************************************************ * * OccEnv * * ************************************************************************ -} data OccEnv = OccEnv { occ_encl :: !OccEncl -- Enclosing context information , occ_one_shots :: !OneShots -- Tells about linearity , occ_gbl_scrut :: GlobalScruts , occ_rule_act :: Activation -> Bool -- Which rules are active -- See Note [Finding rule RHS free vars] , occ_binder_swap :: !Bool -- enable the binder_swap -- See CorePrep Note [Dead code in CorePrep] } type GlobalScruts = IdSet -- See Note [Binder swap on GlobalId scrutinees] ----------------------------- -- OccEncl is used to control whether to inline into constructor arguments -- For example: -- x = (p,q) -- Don't inline p or q -- y = /\a -> (p a, q a) -- Still don't inline p or q -- z = f (p,q) -- Do inline p,q; it may make a rule fire -- So OccEncl tells enought about the context to know what to do when -- we encounter a constructor application or PAP. data OccEncl = OccRhs -- RHS of let(rec), albeit perhaps inside a type lambda -- Don't inline into constructor args here | OccVanilla -- Argument of function, body of lambda, scruintee of case etc. -- Do inline into constructor args here instance Outputable OccEncl where ppr OccRhs = text "occRhs" ppr OccVanilla = text "occVanilla" type OneShots = [OneShotInfo] -- [] No info -- -- one_shot_info:ctxt Analysing a function-valued expression that -- will be applied as described by one_shot_info initOccEnv :: (Activation -> Bool) -> OccEnv initOccEnv active_rule = OccEnv { occ_encl = OccVanilla , occ_one_shots = [] , occ_gbl_scrut = emptyVarSet , occ_rule_act = active_rule , occ_binder_swap = True } vanillaCtxt :: OccEnv -> OccEnv vanillaCtxt env = env { occ_encl = OccVanilla, occ_one_shots = [] } rhsCtxt :: OccEnv -> OccEnv rhsCtxt env = env { occ_encl = OccRhs, occ_one_shots = [] } argCtxt :: OccEnv -> [OneShots] -> (OccEnv, [OneShots]) argCtxt env [] = (env { occ_encl = OccVanilla, occ_one_shots = [] }, []) argCtxt env (one_shots:one_shots_s) = (env { occ_encl = OccVanilla, occ_one_shots = one_shots }, one_shots_s) isRhsEnv :: OccEnv -> Bool isRhsEnv (OccEnv { occ_encl = OccRhs }) = True isRhsEnv (OccEnv { occ_encl = OccVanilla }) = False oneShotGroup :: OccEnv -> [CoreBndr] -> ( OccEnv , [CoreBndr] ) -- The result binders have one-shot-ness set that they might not have had originally. -- This happens in (build (\c n -> e)). Here the occurrence analyser -- linearity context knows that c,n are one-shot, and it records that fact in -- the binder. This is useful to guide subsequent float-in/float-out tranformations oneShotGroup env@(OccEnv { occ_one_shots = ctxt }) bndrs = go ctxt bndrs [] where go ctxt [] rev_bndrs = ( env { occ_one_shots = ctxt, occ_encl = OccVanilla } , reverse rev_bndrs ) go [] bndrs rev_bndrs = ( env { occ_one_shots = [], occ_encl = OccVanilla } , reverse rev_bndrs ++ bndrs ) go ctxt (bndr:bndrs) rev_bndrs | isId bndr = case ctxt of [] -> go [] bndrs (bndr : rev_bndrs) (one_shot : ctxt) -> go ctxt bndrs (bndr': rev_bndrs) where bndr' = updOneShotInfo bndr one_shot -- Use updOneShotInfo, not setOneShotInfo, as pre-existing -- one-shot info might be better than what we can infer, e.g. -- due to explicit use of the magic 'oneShot' function. -- See Note [The oneShot function] | otherwise = go ctxt bndrs (bndr:rev_bndrs) addAppCtxt :: OccEnv -> [Arg CoreBndr] -> OccEnv addAppCtxt env@(OccEnv { occ_one_shots = ctxt }) args = env { occ_one_shots = replicate (valArgCount args) OneShotLam ++ ctxt } transClosureFV :: UniqFM VarSet -> UniqFM VarSet -- If (f,g), (g,h) are in the input, then (f,h) is in the output -- as well as (f,g), (g,h) transClosureFV env | no_change = env | otherwise = transClosureFV (listToUFM new_fv_list) where (no_change, new_fv_list) = mapAccumL bump True (nonDetUFMToList env) -- It's OK to use nonDetUFMToList here because we'll forget the -- ordering by creating a new set with listToUFM bump no_change (b,fvs) | no_change_here = (no_change, (b,fvs)) | otherwise = (False, (b,new_fvs)) where (new_fvs, no_change_here) = extendFvs env fvs ------------- extendFvs_ :: UniqFM VarSet -> VarSet -> VarSet extendFvs_ env s = fst (extendFvs env s) -- Discard the Bool flag extendFvs :: UniqFM VarSet -> VarSet -> (VarSet, Bool) -- (extendFVs env s) returns -- (s `union` env(s), env(s) `subset` s) extendFvs env s | isNullUFM env = (s, True) | otherwise = (s `unionVarSet` extras, extras `subVarSet` s) where extras :: VarSet -- env(s) extras = nonDetFoldUFM unionVarSet emptyVarSet $ -- It's OK to use nonDetFoldUFM here because unionVarSet commutes intersectUFM_C (\x _ -> x) env s {- ************************************************************************ * * Binder swap * * ************************************************************************ Note [Binder swap] ~~~~~~~~~~~~~~~~~~ We do these two transformations right here: (1) case x of b { pi -> ri } ==> case x of b { pi -> let x=b in ri } (2) case (x |> co) of b { pi -> ri } ==> case (x |> co) of b { pi -> let x = b |> sym co in ri } Why (2)? See Note [Case of cast] In both cases, in a particular alternative (pi -> ri), we only add the binding if (a) x occurs free in (pi -> ri) (ie it occurs in ri, but is not bound in pi) (b) the pi does not bind b (or the free vars of co) We need (a) and (b) for the inserted binding to be correct. For the alternatives where we inject the binding, we can transfer all x's OccInfo to b. And that is the point. Notice that * The deliberate shadowing of 'x'. * That (a) rapidly becomes false, so no bindings are injected. The reason for doing these transformations here is because it allows us to adjust the OccInfo for 'x' and 'b' as we go. * Suppose the only occurrences of 'x' are the scrutinee and in the ri; then this transformation makes it occur just once, and hence get inlined right away. * If we do this in the Simplifier, we don't know whether 'x' is used in ri, so we are forced to pessimistically zap b's OccInfo even though it is typically dead (ie neither it nor x appear in the ri). There's nothing actually wrong with zapping it, except that it's kind of nice to know which variables are dead. My nose tells me to keep this information as robustly as possible. The Maybe (Id,CoreExpr) passed to occAnalAlt is the extra let-binding {x=b}; it's Nothing if the binder-swap doesn't happen. There is a danger though. Consider let v = x +# y in case (f v) of w -> ...v...v... And suppose that (f v) expands to just v. Then we'd like to use 'w' instead of 'v' in the alternative. But it may be too late; we may have substituted the (cheap) x+#y for v in the same simplifier pass that reduced (f v) to v. I think this is just too bad. CSE will recover some of it. Note [Case of cast] ~~~~~~~~~~~~~~~~~~~ Consider case (x `cast` co) of b { I# -> ... (case (x `cast` co) of {...}) ... We'd like to eliminate the inner case. That is the motivation for equation (2) in Note [Binder swap]. When we get to the inner case, we inline x, cancel the casts, and away we go. Note [Binder swap on GlobalId scrutinees] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ When the scrutinee is a GlobalId we must take care in two ways i) In order to *know* whether 'x' occurs free in the RHS, we need its occurrence info. BUT, we don't gather occurrence info for GlobalIds. That's the reason for the (small) occ_gbl_scrut env in OccEnv is for: it says "gather occurrence info for these". ii) We must call localiseId on 'x' first, in case it's a GlobalId, or has an External Name. See, for example, SimplEnv Note [Global Ids in the substitution]. Note [Zap case binders in proxy bindings] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ From the original case x of cb(dead) { p -> ...x... } we will get case x of cb(live) { p -> let x = cb in ...x... } Core Lint never expects to find an *occurrence* of an Id marked as Dead, so we must zap the OccInfo on cb before making the binding x = cb. See Trac #5028. Historical note [no-case-of-case] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We *used* to suppress the binder-swap in case expressions when -fno-case-of-case is on. Old remarks: "This happens in the first simplifier pass, and enhances full laziness. Here's the bad case: f = \ y -> ...(case x of I# v -> ...(case x of ...) ... ) If we eliminate the inner case, we trap it inside the I# v -> arm, which might prevent some full laziness happening. I've seen this in action in spectral/cichelli/Prog.hs: [(m,n) | m <- [1..max], n <- [1..max]] Hence the check for NoCaseOfCase." However, now the full-laziness pass itself reverses the binder-swap, so this check is no longer necessary. Historical note [Suppressing the case binder-swap] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ This old note describes a problem that is also fixed by doing the binder-swap in OccAnal: There is another situation when it might make sense to suppress the case-expression binde-swap. If we have case x of w1 { DEFAULT -> case x of w2 { A -> e1; B -> e2 } ...other cases .... } We'll perform the binder-swap for the outer case, giving case x of w1 { DEFAULT -> case w1 of w2 { A -> e1; B -> e2 } ...other cases .... } But there is no point in doing it for the inner case, because w1 can't be inlined anyway. Furthermore, doing the case-swapping involves zapping w2's occurrence info (see paragraphs that follow), and that forces us to bind w2 when doing case merging. So we get case x of w1 { A -> let w2 = w1 in e1 B -> let w2 = w1 in e2 ...other cases .... } This is plain silly in the common case where w2 is dead. Even so, I can't see a good way to implement this idea. I tried not doing the binder-swap if the scrutinee was already evaluated but that failed big-time: data T = MkT !Int case v of w { MkT x -> case x of x1 { I# y1 -> case x of x2 { I# y2 -> ... Notice that because MkT is strict, x is marked "evaluated". But to eliminate the last case, we must either make sure that x (as well as x1) has unfolding MkT y1. The straightforward thing to do is to do the binder-swap. So this whole note is a no-op. It's fixed by doing the binder-swap in OccAnal because we can do the binder-swap unconditionally and still get occurrence analysis information right. -} mkAltEnv :: OccEnv -> CoreExpr -> Id -> (OccEnv, Maybe (Id, CoreExpr)) -- Does two things: a) makes the occ_one_shots = OccVanilla -- b) extends the GlobalScruts if possible -- c) returns a proxy mapping, binding the scrutinee -- to the case binder, if possible mkAltEnv env@(OccEnv { occ_gbl_scrut = pe }) scrut case_bndr = case stripTicksTopE (const True) scrut of Var v -> add_scrut v case_bndr' Cast (Var v) co -> add_scrut v (Cast case_bndr' (mkSymCo co)) -- See Note [Case of cast] _ -> (env { occ_encl = OccVanilla }, Nothing) where add_scrut v rhs = ( env { occ_encl = OccVanilla, occ_gbl_scrut = pe `extendVarSet` v } , Just (localise v, rhs) ) case_bndr' = Var (zapIdOccInfo case_bndr) -- See Note [Zap case binders in proxy bindings] localise scrut_var = mkLocalIdOrCoVar (localiseName (idName scrut_var)) (idType scrut_var) -- Localise the scrut_var before shadowing it; we're making a -- new binding for it, and it might have an External Name, or -- even be a GlobalId; Note [Binder swap on GlobalId scrutinees] -- Also we don't want any INLINE or NOINLINE pragmas! {- ************************************************************************ * * \subsection[OccurAnal-types]{OccEnv} * * ************************************************************************ -} type UsageDetails = IdEnv OccInfo -- A finite map from ids to their usage -- INVARIANT: never IAmDead -- (Deadness is signalled by not being in the map at all) (+++), combineAltsUsageDetails :: UsageDetails -> UsageDetails -> UsageDetails (+++) usage1 usage2 = plusVarEnv_C addOccInfo usage1 usage2 combineAltsUsageDetails usage1 usage2 = plusVarEnv_C orOccInfo usage1 usage2 addOneOcc :: UsageDetails -> Id -> OccInfo -> UsageDetails addOneOcc usage id info = plusVarEnv_C addOccInfo usage (unitVarEnv id info) -- ToDo: make this more efficient emptyDetails :: UsageDetails emptyDetails = (emptyVarEnv :: UsageDetails) usedIn :: Id -> UsageDetails -> Bool v `usedIn` details = isExportedId v || v `elemVarEnv` details addIdOccs :: UsageDetails -> VarSet -> UsageDetails addIdOccs usage id_set = nonDetFoldUFM addIdOcc usage id_set -- It's OK to use nonDetFoldUFM here because addIdOcc commutes addIdOcc :: Id -> UsageDetails -> UsageDetails addIdOcc v u | isId v = addOneOcc u v NoOccInfo | otherwise = u -- Give a non-committal binder info (i.e NoOccInfo) because -- a) Many copies of the specialised thing can appear -- b) We don't want to substitute a BIG expression inside a RULE -- even if that's the only occurrence of the thing -- (Same goes for INLINE.) udFreeVars :: VarSet -> UsageDetails -> VarSet -- Find the subset of bndrs that are mentioned in uds udFreeVars bndrs uds = intersectUFM_C (\b _ -> b) bndrs uds type IdWithOccInfo = Id tagLamBinders :: UsageDetails -- Of scope -> [Id] -- Binders -> (UsageDetails, -- Details with binders removed [IdWithOccInfo]) -- Tagged binders -- Used for lambda and case binders -- It copes with the fact that lambda bindings can have a -- stable unfolding, used for join points tagLamBinders usage binders = usage' `seq` (usage', bndrs') where (usage', bndrs') = mapAccumR tag_lam usage binders tag_lam usage bndr = (usage2, setBinderOcc usage bndr) where usage1 = usage `delVarEnv` bndr usage2 | isId bndr = addIdOccs usage1 (idUnfoldingVars bndr) | otherwise = usage1 tagBinder :: UsageDetails -- Of scope -> Id -- Binders -> (UsageDetails, -- Details with binders removed IdWithOccInfo) -- Tagged binders tagBinder usage binder = let usage' = usage `delVarEnv` binder binder' = setBinderOcc usage binder in usage' `seq` (usage', binder') setBinderOcc :: UsageDetails -> CoreBndr -> CoreBndr setBinderOcc usage bndr | isTyVar bndr = bndr | isExportedId bndr = case idOccInfo bndr of NoOccInfo -> bndr _ -> setIdOccInfo bndr NoOccInfo -- Don't use local usage info for visible-elsewhere things -- BUT *do* erase any IAmALoopBreaker annotation, because we're -- about to re-generate it and it shouldn't be "sticky" | otherwise = setIdOccInfo bndr occ_info where occ_info = lookupVarEnv usage bndr `orElse` IAmDead {- ************************************************************************ * * \subsection{Operations over OccInfo} * * ************************************************************************ -} mkOneOcc :: OccEnv -> Id -> InterestingCxt -> UsageDetails mkOneOcc env id int_cxt | isLocalId id = unitVarEnv id (OneOcc False True int_cxt) | id `elemVarEnv` occ_gbl_scrut env = unitVarEnv id NoOccInfo | otherwise = emptyDetails markMany, markInsideLam :: OccInfo -> OccInfo markMany _ = NoOccInfo markInsideLam (OneOcc _ one_br int_cxt) = OneOcc True one_br int_cxt markInsideLam occ = occ addOccInfo, orOccInfo :: OccInfo -> OccInfo -> OccInfo addOccInfo a1 a2 = ASSERT( not (isDeadOcc a1 || isDeadOcc a2) ) NoOccInfo -- Both branches are at least One -- (Argument is never IAmDead) -- (orOccInfo orig new) is used -- when combining occurrence info from branches of a case orOccInfo (OneOcc in_lam1 _ int_cxt1) (OneOcc in_lam2 _ int_cxt2) = OneOcc (in_lam1 || in_lam2) False -- False, because it occurs in both branches (int_cxt1 && int_cxt2) orOccInfo a1 a2 = ASSERT( not (isDeadOcc a1 || isDeadOcc a2) ) NoOccInfo