{- (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, MultiWayIf, ViewPatterns #-} module OccurAnal ( occurAnalysePgm, occurAnalyseExpr, occurAnalyseExpr_NoBinderSwap ) where #include "HsVersions.h" import GhcPrelude import CoreSyn import CoreFVs import CoreUtils ( exprIsTrivial, isDefaultAlt, isExpandableApp, stripTicksTopE, mkTicks ) import CoreArity ( joinRhsArity ) import Id import IdInfo import Name( localiseName ) import BasicTypes import Module( Module ) import Coercion import Type import VarSet import VarEnv import Var import Demand ( argOneShots, argsOneShots ) import Digraph ( SCC(..), Node(..) , stronglyConnCompFromEdgedVerticesUniq , stronglyConnCompFromEdgedVerticesUniqR ) import Unique import UniqFM import UniqSet 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 -> (Id -> Bool) -- Active unfoldings -> (Activation -> Bool) -- Active rules -> [CoreRule] -> CoreProgram -> CoreProgram occurAnalysePgm this_mod active_unf active_rule imp_rules binds | isEmptyDetails 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 { occ_rule_act = active_rule , occ_unf_act = active_unf } (final_usage, occ_anald_binds) = go init_env binds (_, occ_anald_glommed_binds) = occAnalRecBind init_env TopLevel 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 = addManyOccsSet emptyDetails (rulesFreeVars imp_rules) -- The RULES declarations keep things alive! -- Note [Preventing loops due to imported functions rules] imp_rule_edges = foldr (plusVarEnv_C unionVarSet) emptyVarEnv [ mapVarEnv (const maps_to) $ getUniqSet (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 TopLevel 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 { occ_binder_swap = enable_binder_swap } {- 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 dependencies 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 <blah blah> = 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. ------------------------------------------------------------ Note [Finding join points] ~~~~~~~~~~~~~~~~~~~~~~~~~~ It's the occurrence analyser's job to find bindings that we can turn into join points, but it doesn't perform that transformation right away. Rather, it marks the eligible bindings as part of their occurrence data, leaving it to the simplifier (or to simpleOptPgm) to actually change the binder's 'IdDetails'. The simplifier then eta-expands the RHS if needed and then updates the occurrence sites. Dividing the work this way means that the occurrence analyser still only takes one pass, yet one can always tell the difference between a function call and a jump by looking at the occurrence (because the same pass changes the 'IdDetails' and propagates the binders to their occurrence sites). To track potential join points, we use the 'occ_tail' field of OccInfo. A value of `AlwaysTailCalled n` indicates that every occurrence of the variable is a tail call with `n` arguments (counting both value and type arguments). Otherwise 'occ_tail' will be 'NoTailCallInfo'. The tail call info flows bottom-up with the rest of 'OccInfo' until it goes on the binder. Note [Rules and join points] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Things get fiddly with rules. Suppose we have: let j :: Int -> Int j y = 2 * y k :: Int -> Int -> Int {-# RULES "SPEC k 0" k 0 = j #-} k x y = x + 2 * y in ... Now suppose that both j and k appear only as saturated tail calls in the body. Thus we would like to make them both join points. The rule complicates matters, though, as its RHS has an unapplied occurrence of j. *However*, if we were to eta-expand the rule, all would be well: {-# RULES "SPEC k 0" forall a. k 0 a = j a #-} So conceivably we could notice that a potential join point would have an "undersaturated" rule and account for it. This would mean we could make something that's been specialised a join point, for instance. But local bindings are rarely specialised, and being overly cautious about rules only costs us anything when, for some `j`: * Before specialisation, `j` has non-tail calls, so it can't be a join point. * During specialisation, `j` gets specialised and thus acquires rules. * Sometime afterward, the non-tail calls to `j` disappear (as dead code, say), and so now `j` *could* become a join point. This appears to be very rare in practice. TODO Perhaps we should gather statistics to be sure. ------------------------------------------------------------ Note [Adjusting right-hand sides] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ There's a bit of a dance we need to do after analysing a lambda expression or a right-hand side. In particular, we need to a) call 'markAllInsideLam' *unless* the binding is for a thunk, a one-shot lambda, or a non-recursive join point; and b) call 'markAllNonTailCalled' *unless* the binding is for a join point. Some examples, with how the free occurrences in e (assumed not to be a value lambda) get marked: inside lam non-tail-called ------------------------------------------------------------ let x = e No Yes let f = \x -> e Yes Yes let f = \x{OneShot} -> e No Yes \x -> e Yes Yes join j x = e No No joinrec j x = e Yes No There are a few other caveats; most importantly, if we're marking a binding as 'AlwaysTailCalled', it's *going* to be a join point, so we treat it as one so that the effect cascades properly. Consequently, at the time the RHS is analysed, we won't know what adjustments to make; thus 'occAnalLamOrRhs' must return the unadjusted 'UsageDetails', to be adjusted by 'adjustRhsUsage' once join-point-hood has been decided. Thus the overall sequence taking place in 'occAnalNonRecBind' and 'occAnalRecBind' is as follows: 1. Call 'occAnalLamOrRhs' to find usage information for the RHS. 2. Call 'tagNonRecBinder' or 'tagRecBinders', which decides whether to make the binding a join point. 3. Call 'adjustRhsUsage' accordingly. (Done as part of 'tagRecBinders' when recursive.) (In the recursive case, this logic is spread between 'makeNode' and 'occAnalRec'.) -} ------------------------------------------------------------------ -- occAnalBind ------------------------------------------------------------------ occAnalBind :: OccEnv -- The incoming OccEnv -> TopLevelFlag -> ImpRuleEdges -> CoreBind -> UsageDetails -- Usage details of scope -> (UsageDetails, -- Of the whole let(rec) [CoreBind]) occAnalBind env lvl top_env (NonRec binder rhs) body_usage = occAnalNonRecBind env lvl top_env binder rhs body_usage occAnalBind env lvl top_env (Rec pairs) body_usage = occAnalRecBind env lvl top_env pairs body_usage ----------------- occAnalNonRecBind :: OccEnv -> TopLevelFlag -> ImpRuleEdges -> Var -> CoreExpr -> UsageDetails -> (UsageDetails, [CoreBind]) occAnalNonRecBind env lvl 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' `andUDs` rhs_usage', [NonRec tagged_binder rhs']) where (body_usage', tagged_binder) = tagNonRecBinder lvl body_usage binder mb_join_arity = willBeJoinId_maybe tagged_binder (bndrs, body) = collectBinders rhs (rhs_usage1, bndrs', body') = occAnalNonRecRhs env tagged_binder bndrs body rhs' = mkLams (markJoinOneShots mb_join_arity bndrs') body' -- For a /non-recursive/ join point we can mark all -- its join-lambda as one-shot; and it's a good idea to do so -- Unfoldings -- See Note [Unfoldings and join points] rhs_usage2 = case occAnalUnfolding env NonRecursive binder of Just unf_usage -> rhs_usage1 `andUDs` unf_usage Nothing -> rhs_usage1 -- Rules -- See Note [Rules are extra RHSs] and Note [Rule dependency info] rules_w_uds = occAnalRules env mb_join_arity NonRecursive tagged_binder rule_uds = map (\(_, l, r) -> l `andUDs` r) rules_w_uds rhs_usage3 = foldr andUDs rhs_usage2 rule_uds rhs_usage4 = case lookupVarEnv imp_rule_edges binder of Nothing -> rhs_usage3 Just vs -> addManyOccsSet rhs_usage3 vs -- See Note [Preventing loops due to imported functions rules] -- Final adjustment rhs_usage' = adjustRhsUsage mb_join_arity NonRecursive bndrs' rhs_usage4 ----------------- occAnalRecBind :: OccEnv -> TopLevelFlag -> ImpRuleEdges -> [(Var,CoreExpr)] -> UsageDetails -> (UsageDetails, [CoreBind]) occAnalRecBind env lvl imp_rule_edges pairs body_usage = foldr (occAnalRec env lvl) (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) {- Note [Unfoldings and join points] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ We assume that anything in an unfolding occurs multiple times, since unfoldings are often copied (that's the whole point!). But we still need to track tail calls for the purpose of finding join points. -} ----------------------------- occAnalRec :: OccEnv -> TopLevelFlag -> SCC Details -> (UsageDetails, [CoreBind]) -> (UsageDetails, [CoreBind]) -- The NonRec case is just like a Let (NonRec ...) above occAnalRec _ lvl (AcyclicSCC (ND { nd_bndr = bndr, nd_rhs = rhs , nd_uds = rhs_uds, nd_rhs_bndrs = rhs_bndrs })) (body_uds, binds) | not (bndr `usedIn` body_uds) = (body_uds, binds) -- See Note [Dead code] | otherwise -- It's mentioned in the body = (body_uds' `andUDs` rhs_uds', NonRec tagged_bndr rhs : binds) where (body_uds', tagged_bndr) = tagNonRecBinder lvl body_uds bndr rhs_uds' = adjustRhsUsage (willBeJoinId_maybe tagged_bndr) NonRecursive rhs_bndrs rhs_uds -- The Rec case is the interesting one -- See Note [Recursive bindings: the grand plan] -- See Note [Loop breaking] occAnalRec env lvl (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 "lb nodes" <+> ppr loop_breaker_nodes]) (final_uds, Rec pairs : binds) where bndrs = map nd_bndr details_s bndr_set = mkVarSet bndrs ------------------------------ -- See Note [Choosing loop breakers] for loop_breaker_nodes final_uds :: UsageDetails loop_breaker_nodes :: [LetrecNode] (final_uds, loop_breaker_nodes) = mkLoopBreakerNodes env lvl bndr_set body_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 = -- pprTrace "loopBreakNodes" (ppr nodes) $ 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" (vcat [ 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 (node_payload 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 (node_payload -> ND { nd_bndr = bndr, nd_rhs = rhs}) = (bndr `setIdOccInfo` strongLoopBreaker { occ_tail = tail_info }, rhs) where tail_info = tailCallInfo (idOccInfo bndr) mk_non_loop_breaker :: VarSet -> LetrecNode -> Binding -- See Note [Weak loop breakers] mk_non_loop_breaker weak_fvs (node_payload -> ND { nd_bndr = bndr , nd_rhs = rhs}) | bndr `elemVarSet` weak_fvs = (setIdOccInfo bndr occ', rhs) | otherwise = (bndr, rhs) where occ' = weakLoopBreaker { occ_tail = tail_info } tail_info = tailCallInfo (idOccInfo bndr) ---------------------------------- 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 (node_payload 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. @occAnalUnfolding@ 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 this 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_rhs_bndrs :: [CoreBndr] -- Outer lambdas of RHS -- INVARIANT: (nd_rhs_bndrs nd, _) == -- collectBinders (nd_rhs nd) , 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) , text "score =" <+> ppr (nd_score 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) = DigraphNode details (varUnique bndr) (nonDetKeysUniqSet node_fvs) -- It's OK to use nonDetKeysUniqSet 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_rhs_bndrs = bndrs' , 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] (bndrs, body) = collectBinders rhs (rhs_usage1, bndrs', body') = occAnalRecRhs env bndrs body rhs' = mkLams bndrs' body' rhs_usage2 = foldr andUDs rhs_usage1 rule_uds -- Note [Rules are extra RHSs] -- Note [Rule dependency info] rhs_usage3 = case mb_unf_uds of Just unf_uds -> rhs_usage2 `andUDs` unf_uds 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_w_uds :: [(CoreRule, UsageDetails, UsageDetails)] rules_w_uds = occAnalRules env (Just (length bndrs)) Recursive bndr rules_w_rhs_fvs :: [(Activation, VarSet)] -- Find the RHS fvs rules_w_rhs_fvs = maybe id (\ids -> ((AlwaysActive, ids):)) (lookupVarEnv imp_rule_edges bndr) -- See Note [Preventing loops due to imported functions rules] [ (ru_act rule, udFreeVars bndr_set rhs_uds) | (rule, _, rhs_uds) <- rules_w_uds ] rule_uds = map (\(_, l, r) -> l `andUDs` r) rules_w_uds active_rule_fvs = unionVarSets [fvs | (a,fvs) <- rules_w_rhs_fvs , is_active a] -- Finding the usage details of the INLINE pragma (if any) mb_unf_uds = occAnalUnfolding env Recursive bndr -- Find the "nd_inl" free vars; for the loop-breaker phase inl_fvs = case mb_unf_uds of Nothing -> udFreeVars bndr_set rhs_usage1 -- No INLINE, use RHS Just unf_uds -> udFreeVars bndr_set unf_uds -- 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 :: OccEnv -> TopLevelFlag -> VarSet -> UsageDetails -- for BODY of let -> [Details] -> (UsageDetails, -- adjusted [LetrecNode]) -- Does four 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 -- d) adjust each RHS's usage details according to -- the binder's (new) shotness and join-point-hood mkLoopBreakerNodes env lvl bndr_set body_uds details_s = (final_uds, zipWith mk_lb_node details_s bndrs') where (final_uds, bndrs') = tagRecBinders lvl body_uds [ ((nd_bndr nd) ,(nd_uds nd) ,(nd_rhs_bndrs nd)) | nd <- details_s ] mk_lb_node nd@(ND { nd_bndr = bndr, nd_rhs = rhs, nd_inl = inl_fvs }) bndr' = DigraphNode nd' (varUnique bndr) (nonDetKeysUniqSet lb_deps) -- It's OK to use nonDetKeysUniqSet 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 } score = nodeScore env 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 :: OccEnv -> 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 env 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] | not (occ_unf_act env old_bndr) -- A binder whose inlining is inactive (e.g. has = (0, 0, True) -- a NOINLINE pragma) makes a great loop breaker | exprIsTrivial 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 | DFunUnfolding { df_args = args } <- id_unfolding -- Never choose a DFun as a loop breaker -- Note [DFuns should not be loop breakers] = (9, length args, is_lb) -- Data structures are more important than INLINE pragmas -- so that dictionary/method recursion unravels | CoreUnfolding { uf_guidance = UnfWhen {} } <- id_unfolding = mk_score 6 | is_con_app rhs -- Data types help with cases: = mk_score 5 -- Note [Constructor applications] | isStableUnfolding id_unfolding , can_unfold = mk_score 3 | isOneOcc (idOccInfo new_bndr) = mk_score 2 -- Likely to be inlined | can_unfold -- 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 = case id_unfolding of CoreUnfolding { uf_src = src, uf_tmpl = unf_rhs } | isStableSource src -> unf_rhs _ -> bind_rhs -- 'bind_rhs' is irrelevant for inlining things with a stable unfolding rhs_size = case id_unfolding of CoreUnfolding { uf_guidance = guidance } | UnfIfGoodArgs { ug_size = size } <- guidance -> size _ -> cheapExprSize rhs can_unfold = canUnfold id_unfolding 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 saturated) 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 * * ************************************************************************ -} occAnalRhs :: OccEnv -> RecFlag -> Id -> [CoreBndr] -> CoreExpr -> (UsageDetails, [CoreBndr], CoreExpr) -- Returned usage details covers only the RHS, -- and *not* the RULE or INLINE template for the Id occAnalRhs env Recursive _ bndrs body = occAnalRecRhs env bndrs body occAnalRhs env NonRecursive id bndrs body = occAnalNonRecRhs env id bndrs body occAnalRecRhs :: OccEnv -> [CoreBndr] -> CoreExpr -- Rhs lambdas, body -> (UsageDetails, [CoreBndr], CoreExpr) -- Returned usage details covers only the RHS, -- and *not* the RULE or INLINE template for the Id occAnalRecRhs env bndrs body = occAnalLamOrRhs (rhsCtxt env) bndrs body occAnalNonRecRhs :: OccEnv -> Id -> [CoreBndr] -> CoreExpr -- Binder; rhs lams, body -- Binder is already tagged with occurrence info -> (UsageDetails, [CoreBndr], CoreExpr) -- Returned usage details covers only the RHS, -- and *not* the RULE or INLINE template for the Id occAnalNonRecRhs env bndr bndrs body = occAnalLamOrRhs rhs_env bndrs body where env1 | is_join_point = env -- See Note [Join point RHSs] | certainly_inline = env -- See Note [Cascading inlines] | otherwise = rhsCtxt env -- See Note [Sources of one-shot information] rhs_env = env1 { occ_one_shots = argOneShots dmd } certainly_inline -- See Note [Cascading inlines] = case occ of OneOcc { occ_in_lam = in_lam, occ_one_br = one_br } -> not in_lam && one_br && active && not_stable _ -> False is_join_point = isAlwaysTailCalled occ -- Like (isJoinId bndr) but happens one step earlier -- c.f. willBeJoinId_maybe occ = idOccInfo bndr dmd = idDemandInfo bndr active = isAlwaysActive (idInlineActivation bndr) not_stable = not (isStableUnfolding (idUnfolding bndr)) occAnalUnfolding :: OccEnv -> RecFlag -> Id -> Maybe UsageDetails -- Just the analysis, not a new unfolding. The unfolding -- got analysed when it was created and we don't need to -- update it. occAnalUnfolding env rec_flag id = case realIdUnfolding id of -- ignore previous loop-breaker flag CoreUnfolding { uf_tmpl = rhs, uf_src = src } | not (isStableSource src) -> Nothing | otherwise -> Just $ markAllMany usage where (bndrs, body) = collectBinders rhs (usage, _, _) = occAnalRhs env rec_flag id bndrs body DFunUnfolding { df_bndrs = bndrs, df_args = args } -> Just $ zapDetails (delDetailsList usage bndrs) where usage = andUDsList (map (fst . occAnal env) args) _ -> Nothing occAnalRules :: OccEnv -> Maybe JoinArity -- If the binder is (or MAY become) a join -- point, what its join arity is (or WOULD -- become). See Note [Rules and join points]. -> RecFlag -> Id -> [(CoreRule, -- Each (non-built-in) rule UsageDetails, -- Usage details for LHS UsageDetails)] -- Usage details for RHS occAnalRules env mb_expected_join_arity rec_flag id = [ (rule, lhs_uds, rhs_uds) | rule@Rule {} <- idCoreRules id , let (lhs_uds, rhs_uds) = occ_anal_rule rule ] where occ_anal_rule (Rule { ru_bndrs = bndrs, ru_args = args, ru_rhs = rhs }) = (lhs_uds, final_rhs_uds) where lhs_uds = addManyOccsSet emptyDetails $ (exprsFreeVars args `delVarSetList` bndrs) (rhs_bndrs, rhs_body) = collectBinders rhs (rhs_uds, _, _) = occAnalRhs env rec_flag id rhs_bndrs rhs_body -- Note [Rules are extra RHSs] -- Note [Rule dependency info] final_rhs_uds = adjust_tail_info args $ markAllMany $ (rhs_uds `delDetailsList` bndrs) occ_anal_rule _ = (emptyDetails, emptyDetails) adjust_tail_info args uds -- see Note [Rules and join points] = case mb_expected_join_arity of Just ar | args `lengthIs` ar -> uds _ -> markAllNonTailCalled uds {- Note [Join point RHSs] ~~~~~~~~~~~~~~~~~~~~~~~~~ Consider x = e join j = Just x We want to inline x into j right away, so we don't want to give the join point a RhsCtxt (Trac #14137). It's not a huge deal, because the FloatIn pass knows to float into join point RHSs; and the simplifier does not float things out of join point RHSs. But it's a simple, cheap thing to do. See Trac #14137. 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 (a) the RHS is expandable (see isExpandableApp in occAnalApp), and (b) certainly_inline says "yes" when preInlineUnconditionally says "no" then the simplifier iterates indefinitely: x = f y k = Just x -- We decide that k is 'certainly_inline' v = ...k... -- but preInlineUnconditionally doesn't inline it inline ==> k = Just (f y) v = ...k... float ==> x1 = f y k = Just x1 v = ...k... 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 _) = occAnalApp env (expr, [], []) -- At one stage, I gathered the idRuleVars for the variable 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) = (addManyOccsSet 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) | SourceNote{} <- tickish = (usage, Tick tickish body') -- SourceNotes are best-effort; so we just proceed as usual. -- If we drop a tick due to the issues described below it's -- not the end of the world. | tickish `tickishScopesLike` SoftScope = (markAllNonTailCalled usage, Tick tickish body') | Breakpoint _ ids <- tickish = (usage_lam `andUDs` foldr addManyOccs emptyDetails ids, 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 = markAllNonTailCalled (markAllInsideLam usage) -- TODO There may be ways to make ticks and join points play -- nicer together, but right now there are problems: -- let j x = ... in tick<t> (j 1) -- Making j a join point may cause the simplifier to drop t -- (if the tick is put into the continuation). So we don't -- count j 1 as a tail call. -- See #14242. occAnal env (Cast expr co) = case occAnal env expr of { (usage, expr') -> let usage1 = zapDetailsIf (isRhsEnv env) usage -- usage1: if we see let x = y `cast` co -- then mark y as 'Many' so that we don't -- immediately inline y again. usage2 = addManyOccsSet usage1 (coVarsOfCo co) -- usage2: see Note [Gather occurrences of coercion variables] in (markAllNonTailCalled usage2, Cast expr' co) } 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') -> (markAllNonTailCalled 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 occAnalLamOrRhs env binders body of { (usage, tagged_binders, body') -> let expr' = mkLams tagged_binders body' usage1 = markAllNonTailCalled usage one_shot_gp = all isOneShotBndr tagged_binders final_usage | one_shot_gp = usage1 | otherwise = markAllInsideLam usage1 in (final_usage, expr') } where (binders, body) = collectBinders expr 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 orUDs emptyDetails alts_usage_s (alts_usage1, tagged_bndr) = tagLamBinder alts_usage bndr total_usage = markAllNonTailCalled scrut_usage `andUDs` alts_usage1 -- Alts can have tail calls, but the scrutinee can't in total_usage `seq` (total_usage, Case scrut' tagged_bndr ty alts') }} where 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 0, 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 NotTopLevel 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 `andUDs` 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 `andUDs` final_args_uds !(args_uds, args') = occAnalArgs env args one_shots !final_args_uds | isRhsEnv env && is_exp = markAllNonTailCalled $ markAllInsideLam args_uds | otherwise = markAllNonTailCalled 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 n_args = length args fun_uds = mkOneOcc env fun (n_val_args > 0) n_args 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) guaranteed_val_args guaranteed_val_args = n_val_args + length (takeWhile isOneShotInfo (occ_one_shots env)) -- See Note [Sources of one-shot information], bullet point A'] occAnalApp env (fun, args, ticks) = (markAllNonTailCalled (fun_uds `andUDs` 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 [] zapDetailsIf :: Bool -- If this is true -> UsageDetails -- Then do zapDetails on this -> UsageDetails zapDetailsIf True uds = zapDetails uds zapDetailsIf False uds = uds {- Note [Sources of one-shot information] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ The occurrence analyser obtains one-shot-lambda information from two sources: A: Saturated applications: eg f e1 .. en In general, given a call (f e1 .. en) we can propagate one-shot info from f's strictness signature into e1 .. en, but /only/ if n is enough to saturate the strictness signature. A strictness signature like f :: C1(C1(L))LS means that *if f is applied to three arguments* then it will guarantee to call its first argument at most once, and to call the result of that at most once. But if f has fewer than three arguments, all bets are off; e.g. map (f (\x y. expensive) e2) xs Here the \x y abstraction may be called many times (once for each element of xs) so we should not mark x and y as one-shot. But if it was map (f (\x y. expensive) 3 2) xs then the first argument of f will be called at most once. The one-shot info, derived from f's strictness signature, is computed by 'argsOneShots', called in occAnalApp. A': Non-obviously saturated applications: eg build (f (\x y -> expensive)) where f is as above. In this case, f is only manifestly applied to one argument, so it does not look saturated. So by the previous point, we should not use its strictness signature to learn about the one-shotness of \x y. But in this case we can: build is fully applied, so we may use its strictness signature; and from that we learn that build calls its argument with two arguments *at most once*. So there is really only one call to f, and it will have three arguments. In that sense, f is saturated, and we may proceed as described above. Hence the computation of 'guaranteed_val_args' in occAnalApp, using '(occ_one_shots env)'. See also Trac #13227, comment:9 B: 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 [OneShots] ~~~~~~~~~~~~~~~ When analysing an expression, the occ_one_shots argument contains information about how the function is being used. The length of the list indicates how many arguments will eventually be passed to the analysed expression, and the OneShotInfo indicates whether this application is once or multiple times. Example: Context of f occ_one_shots when analysing f f 1 2 [OneShot, OneShot] map (f 1) [OneShot, NoOneShotInfo] build f [OneShot, OneShot] f 1 2 `seq` f 2 1 [NoOneShotInfo, OneShot] 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). -} occAnalLamOrRhs :: OccEnv -> [CoreBndr] -> CoreExpr -> (UsageDetails, [CoreBndr], CoreExpr) occAnalLamOrRhs env [] body = case occAnal env body of (body_usage, body') -> (body_usage, [], body') -- RHS of thunk or nullary join point occAnalLamOrRhs env (bndr:bndrs) body | isTyVar bndr = -- Important: Keep the environment so that we don't inline into an RHS like -- \(@ x) -> C @x (f @x) -- (see the beginning of Note [Cascading inlines]). case occAnalLamOrRhs env bndrs body of (body_usage, bndrs', body') -> (body_usage, bndr:bndrs', body') occAnalLamOrRhs env binders body = 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 in (final_usage, tagged_binders, body') } where (env_body, binders') = oneShotGroup env binders 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' `andUDs` let_rhs_usg , Let (NonRec tagged_scrut_var let_rhs') alt_rhs ) where captured = any (`usedIn` let_rhs_usg) bndrs -- Check condition (b) -- 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] -- Moreover, the rhs of the let may mention the case-binder, and -- we want to gather its occ-info as well (let_rhs_usg, let_rhs') = occAnal env let_rhs (alt_usg', tagged_scrut_var) = tagLamBinder 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 -- See Note [OneShots] , occ_gbl_scrut :: GlobalScruts , occ_unf_act :: Id -> Bool -- Which Id unfoldings are active , 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 enough 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" -- See note [OneShots] type OneShots = [OneShotInfo] initOccEnv :: OccEnv initOccEnv = OccEnv { occ_encl = OccVanilla , occ_one_shots = [] , occ_gbl_scrut = emptyVarSet -- To be conservative, we say that all -- inlines and rules are active , occ_unf_act = \_ -> True , occ_rule_act = \_ -> True , 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@(one_shot : ctxt') (bndr : bndrs) rev_bndrs | isId bndr = go ctxt' bndrs (bndr': rev_bndrs) | otherwise = 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] markJoinOneShots :: Maybe JoinArity -> [Var] -> [Var] -- Mark the lambdas of a non-recursive join point as one-shot. -- This is good to prevent gratuitous float-out etc markJoinOneShots mb_join_arity bndrs = case mb_join_arity of Nothing -> bndrs Just n -> go n bndrs where go 0 bndrs = bndrs go _ [] = [] -- This can legitimately happen. -- e.g. let j = case ... in j True -- This will become an arity-1 join point after the -- simplifier has eta-expanded it; but it may not have -- enough lambdas /yet/. (Lint checks that JoinIds do -- have enough lambdas.) go n (b:bs) = b' : go (n-1) bs where b' | isId b = setOneShotLambda b | otherwise = b 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 (getUniqSet 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. NB: the OccInfo on /occurrences/ really doesn't matter much; the simplifier doesn't use it. So this is only to satisfy the perhpas-over-picky Lint. 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 three 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 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! localise scrut_var = mkLocalIdOrCoVar (localiseName (idName scrut_var)) (idType scrut_var) {- ************************************************************************ * * \subsection[OccurAnal-types]{OccEnv} * * ************************************************************************ Note [UsageDetails and zapping] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ On many occasions, we must modify all gathered occurrence data at once. For instance, all occurrences underneath a (non-one-shot) lambda set the 'occ_in_lam' flag to become 'True'. We could use 'mapVarEnv' to do this, but that takes O(n) time and we will do this often---in particular, there are many places where tail calls are not allowed, and each of these causes all variables to get marked with 'NoTailCallInfo'. Instead of relying on `mapVarEnv`, then, we carry three 'IdEnv's around along with the 'OccInfoEnv'. Each of these extra environments is a "zapped set" recording which variables have been zapped in some way. Zapping all occurrence info then simply means setting the corresponding zapped set to the whole 'OccInfoEnv', a fast O(1) operation. -} type OccInfoEnv = IdEnv OccInfo -- A finite map from ids to their usage -- INVARIANT: never IAmDead -- (Deadness is signalled by not being in the map at all) type ZappedSet = OccInfoEnv -- Values are ignored data UsageDetails = UD { ud_env :: !OccInfoEnv , ud_z_many :: ZappedSet -- apply 'markMany' to these , ud_z_in_lam :: ZappedSet -- apply 'markInsideLam' to these , ud_z_no_tail :: ZappedSet } -- apply 'markNonTailCalled' to these -- INVARIANT: All three zapped sets are subsets of the OccInfoEnv instance Outputable UsageDetails where ppr ud = ppr (ud_env (flattenUsageDetails ud)) ------------------- -- UsageDetails API andUDs, orUDs :: UsageDetails -> UsageDetails -> UsageDetails andUDs = combineUsageDetailsWith addOccInfo orUDs = combineUsageDetailsWith orOccInfo andUDsList :: [UsageDetails] -> UsageDetails andUDsList = foldl' andUDs emptyDetails mkOneOcc :: OccEnv -> Id -> InterestingCxt -> JoinArity -> UsageDetails mkOneOcc env id int_cxt arity | isLocalId id = singleton $ OneOcc { occ_in_lam = False , occ_one_br = True , occ_int_cxt = int_cxt , occ_tail = AlwaysTailCalled arity } | id `elemVarSet` occ_gbl_scrut env = singleton noOccInfo | otherwise = emptyDetails where singleton info = emptyDetails { ud_env = unitVarEnv id info } addOneOcc :: UsageDetails -> Id -> OccInfo -> UsageDetails addOneOcc ud id info = ud { ud_env = extendVarEnv_C plus_zapped (ud_env ud) id info } `alterZappedSets` (`delVarEnv` id) where plus_zapped old new = doZapping ud id old `addOccInfo` new addManyOccsSet :: UsageDetails -> VarSet -> UsageDetails addManyOccsSet usage id_set = nonDetFoldUniqSet addManyOccs usage id_set -- It's OK to use nonDetFoldUFM here because addManyOccs commutes -- Add several occurrences, assumed not to be tail calls addManyOccs :: Var -> UsageDetails -> UsageDetails addManyOccs 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.) delDetails :: UsageDetails -> Id -> UsageDetails delDetails ud bndr = ud `alterUsageDetails` (`delVarEnv` bndr) delDetailsList :: UsageDetails -> [Id] -> UsageDetails delDetailsList ud bndrs = ud `alterUsageDetails` (`delVarEnvList` bndrs) emptyDetails :: UsageDetails emptyDetails = UD { ud_env = emptyVarEnv , ud_z_many = emptyVarEnv , ud_z_in_lam = emptyVarEnv , ud_z_no_tail = emptyVarEnv } isEmptyDetails :: UsageDetails -> Bool isEmptyDetails = isEmptyVarEnv . ud_env markAllMany, markAllInsideLam, markAllNonTailCalled, zapDetails :: UsageDetails -> UsageDetails markAllMany ud = ud { ud_z_many = ud_env ud } markAllInsideLam ud = ud { ud_z_in_lam = ud_env ud } markAllNonTailCalled ud = ud { ud_z_no_tail = ud_env ud } zapDetails = markAllMany . markAllNonTailCalled -- effectively sets to noOccInfo lookupDetails :: UsageDetails -> Id -> OccInfo lookupDetails ud id | isCoVar id -- We do not currenly gather occurrence info (from types) = noOccInfo -- for CoVars, so we must conservatively mark them as used -- See Note [DoO not mark CoVars as dead] | otherwise = case lookupVarEnv (ud_env ud) id of Just occ -> doZapping ud id occ Nothing -> IAmDead usedIn :: Id -> UsageDetails -> Bool v `usedIn` ud = isExportedId v || v `elemVarEnv` ud_env ud udFreeVars :: VarSet -> UsageDetails -> VarSet -- Find the subset of bndrs that are mentioned in uds udFreeVars bndrs ud = restrictUniqSetToUFM bndrs (ud_env ud) {- Note [Do not mark CoVars as dead] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ It's obviously wrong to mark CoVars as dead if they are used. Currently we don't traverse types to gather usase info for CoVars, so we had better treat them as having noOccInfo. This showed up in Trac #15696 we had something like case eq_sel d of co -> ...(typeError @(...co...) "urk")... Then 'd' was substitued by a dictionary, so the expression simpified to case (Coercion <blah>) of co -> ...(typeError @(...co...) "urk")... But then the "drop the case altogether" equation of rebuildCase thought that 'co' was dead, and discarded the entire case. Urk! I have no idea how we managed to avoid this pitfall for so long! -} ------------------- -- Auxiliary functions for UsageDetails implementation combineUsageDetailsWith :: (OccInfo -> OccInfo -> OccInfo) -> UsageDetails -> UsageDetails -> UsageDetails combineUsageDetailsWith plus_occ_info ud1 ud2 | isEmptyDetails ud1 = ud2 | isEmptyDetails ud2 = ud1 | otherwise = UD { ud_env = plusVarEnv_C plus_occ_info (ud_env ud1) (ud_env ud2) , ud_z_many = plusVarEnv (ud_z_many ud1) (ud_z_many ud2) , ud_z_in_lam = plusVarEnv (ud_z_in_lam ud1) (ud_z_in_lam ud2) , ud_z_no_tail = plusVarEnv (ud_z_no_tail ud1) (ud_z_no_tail ud2) } doZapping :: UsageDetails -> Var -> OccInfo -> OccInfo doZapping ud var occ = doZappingByUnique ud (varUnique var) occ doZappingByUnique :: UsageDetails -> Unique -> OccInfo -> OccInfo doZappingByUnique ud uniq = (if | in_subset ud_z_many -> markMany | in_subset ud_z_in_lam -> markInsideLam | otherwise -> id) . (if | in_subset ud_z_no_tail -> markNonTailCalled | otherwise -> id) where in_subset field = uniq `elemVarEnvByKey` field ud alterZappedSets :: UsageDetails -> (ZappedSet -> ZappedSet) -> UsageDetails alterZappedSets ud f = ud { ud_z_many = f (ud_z_many ud) , ud_z_in_lam = f (ud_z_in_lam ud) , ud_z_no_tail = f (ud_z_no_tail ud) } alterUsageDetails :: UsageDetails -> (OccInfoEnv -> OccInfoEnv) -> UsageDetails alterUsageDetails ud f = ud { ud_env = f (ud_env ud) } `alterZappedSets` f flattenUsageDetails :: UsageDetails -> UsageDetails flattenUsageDetails ud = ud { ud_env = mapUFM_Directly (doZappingByUnique ud) (ud_env ud) } `alterZappedSets` const emptyVarEnv ------------------- -- See Note [Adjusting right-hand sides] adjustRhsUsage :: Maybe JoinArity -> RecFlag -> [CoreBndr] -- Outer lambdas, AFTER occ anal -> UsageDetails -> UsageDetails adjustRhsUsage mb_join_arity rec_flag bndrs usage = maybe_mark_lam (maybe_drop_tails usage) where maybe_mark_lam ud | one_shot = ud | otherwise = markAllInsideLam ud maybe_drop_tails ud | exact_join = ud | otherwise = markAllNonTailCalled ud one_shot = case mb_join_arity of Just join_arity | isRec rec_flag -> False | otherwise -> all isOneShotBndr (drop join_arity bndrs) Nothing -> all isOneShotBndr bndrs exact_join = case mb_join_arity of Just join_arity -> bndrs `lengthIs` join_arity _ -> False type IdWithOccInfo = Id tagLamBinders :: UsageDetails -- Of scope -> [Id] -- Binders -> (UsageDetails, -- Details with binders removed [IdWithOccInfo]) -- Tagged binders tagLamBinders usage binders = usage' `seq` (usage', bndrs') where (usage', bndrs') = mapAccumR tagLamBinder usage binders tagLamBinder :: UsageDetails -- Of scope -> Id -- Binder -> (UsageDetails, -- Details with binder 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 tagLamBinder usage bndr = (usage2, bndr') where occ = lookupDetails usage bndr bndr' = setBinderOcc (markNonTailCalled occ) bndr -- Don't try to make an argument into a join point usage1 = usage `delDetails` bndr usage2 | isId bndr = addManyOccsSet usage1 (idUnfoldingVars bndr) -- This is effectively the RHS of a -- non-join-point binding, so it's okay to use -- addManyOccsSet, which assumes no tail calls | otherwise = usage1 tagNonRecBinder :: TopLevelFlag -- At top level? -> UsageDetails -- Of scope -> CoreBndr -- Binder -> (UsageDetails, -- Details with binder removed IdWithOccInfo) -- Tagged binder tagNonRecBinder lvl usage binder = let occ = lookupDetails usage binder will_be_join = decideJoinPointHood lvl usage [binder] occ' | will_be_join = -- must already be marked AlwaysTailCalled ASSERT(isAlwaysTailCalled occ) occ | otherwise = markNonTailCalled occ binder' = setBinderOcc occ' binder usage' = usage `delDetails` binder in usage' `seq` (usage', binder') tagRecBinders :: TopLevelFlag -- At top level? -> UsageDetails -- Of body of let ONLY -> [(CoreBndr, -- Binder UsageDetails, -- RHS usage details [CoreBndr])] -- Lambdas in new RHS -> (UsageDetails, -- Adjusted details for whole scope, -- with binders removed [IdWithOccInfo]) -- Tagged binders -- Substantially more complicated than non-recursive case. Need to adjust RHS -- details *before* tagging binders (because the tags depend on the RHSes). tagRecBinders lvl body_uds triples = let (bndrs, rhs_udss, _) = unzip3 triples -- 1. Determine join-point-hood of whole group, as determined by -- the *unadjusted* usage details unadj_uds = foldr andUDs body_uds rhs_udss will_be_joins = decideJoinPointHood lvl unadj_uds bndrs -- 2. Adjust usage details of each RHS, taking into account the -- join-point-hood decision rhs_udss' = map adjust triples adjust (bndr, rhs_uds, rhs_bndrs) = adjustRhsUsage mb_join_arity Recursive rhs_bndrs rhs_uds where -- Can't use willBeJoinId_maybe here because we haven't tagged the -- binder yet (the tag depends on these adjustments!) mb_join_arity | will_be_joins , let occ = lookupDetails unadj_uds bndr , AlwaysTailCalled arity <- tailCallInfo occ = Just arity | otherwise = ASSERT(not will_be_joins) -- Should be AlwaysTailCalled if Nothing -- we are making join points! -- 3. Compute final usage details from adjusted RHS details adj_uds = foldr andUDs body_uds rhs_udss' -- 4. Tag each binder with its adjusted details bndrs' = [ setBinderOcc (lookupDetails adj_uds bndr) bndr | bndr <- bndrs ] -- 5. Drop the binders from the adjusted details and return usage' = adj_uds `delDetailsList` bndrs in (usage', bndrs') setBinderOcc :: OccInfo -> CoreBndr -> CoreBndr setBinderOcc occ_info bndr | isTyVar bndr = bndr | isExportedId bndr = if isManyOccs (idOccInfo bndr) then bndr else 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 -- | Decide whether some bindings should be made into join points or not. -- Returns `False` if they can't be join points. Note that it's an -- all-or-nothing decision, as if multiple binders are given, they're -- assumed to be mutually recursive. -- -- It must, however, be a final decision. If we say "True" for 'f', -- and then subsequently decide /not/ make 'f' into a join point, then -- the decision about another binding 'g' might be invalidated if (say) -- 'f' tail-calls 'g'. -- -- See Note [Invariants on join points] in CoreSyn. decideJoinPointHood :: TopLevelFlag -> UsageDetails -> [CoreBndr] -> Bool decideJoinPointHood TopLevel _ _ = False decideJoinPointHood NotTopLevel usage bndrs | isJoinId (head bndrs) = WARN(not all_ok, text "OccurAnal failed to rediscover join point(s):" <+> ppr bndrs) all_ok | otherwise = all_ok where -- See Note [Invariants on join points]; invariants cited by number below. -- Invariant 2 is always satisfiable by the simplifier by eta expansion. all_ok = -- Invariant 3: Either all are join points or none are all ok bndrs ok bndr | -- Invariant 1: Only tail calls, all same join arity AlwaysTailCalled arity <- tailCallInfo (lookupDetails usage bndr) , -- Invariant 1 as applied to LHSes of rules all (ok_rule arity) (idCoreRules bndr) -- Invariant 2a: stable unfoldings -- See Note [Join points and INLINE pragmas] , ok_unfolding arity (realIdUnfolding bndr) -- Invariant 4: Satisfies polymorphism rule , isValidJoinPointType arity (idType bndr) = True | otherwise = False ok_rule _ BuiltinRule{} = False -- only possible with plugin shenanigans ok_rule join_arity (Rule { ru_args = args }) = args `lengthIs` join_arity -- Invariant 1 as applied to LHSes of rules -- ok_unfolding returns False if we should /not/ convert a non-join-id -- into a join-id, even though it is AlwaysTailCalled ok_unfolding join_arity (CoreUnfolding { uf_src = src, uf_tmpl = rhs }) = not (isStableSource src && join_arity > joinRhsArity rhs) ok_unfolding _ (DFunUnfolding {}) = False ok_unfolding _ _ = True willBeJoinId_maybe :: CoreBndr -> Maybe JoinArity willBeJoinId_maybe bndr = case tailCallInfo (idOccInfo bndr) of AlwaysTailCalled arity -> Just arity _ -> isJoinId_maybe bndr {- Note [Join points and INLINE pragmas] ~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~ Consider f x = let g = \x. not -- Arity 1 {-# INLINE g #-} in case x of A -> g True True B -> g True False C -> blah2 Here 'g' is always tail-called applied to 2 args, but the stable unfolding captured by the INLINE pragma has arity 1. If we try to convert g to be a join point, its unfolding will still have arity 1 (since it is stable, and we don't meddle with stable unfoldings), and Lint will complain (see Note [Invariants on join points], (2a), in CoreSyn. Trac #13413. Moreover, since g is going to be inlined anyway, there is no benefit from making it a join point. If it is recursive, and uselessly marked INLINE, this will stop us making it a join point, which is annoying. But occasionally (notably in class methods; see Note [Instances and loop breakers] in TcInstDcls) we mark recursive things as INLINE but the recursion unravels; so ignoring INLINE pragmas on recursive things isn't good either. See Invariant 2a of Note [Invariants on join points] in CoreSyn ************************************************************************ * * \subsection{Operations over OccInfo} * * ************************************************************************ -} markMany, markInsideLam, markNonTailCalled :: OccInfo -> OccInfo markMany IAmDead = IAmDead markMany occ = ManyOccs { occ_tail = occ_tail occ } markInsideLam occ@(OneOcc {}) = occ { occ_in_lam = True } markInsideLam occ = occ markNonTailCalled IAmDead = IAmDead markNonTailCalled occ = occ { occ_tail = NoTailCallInfo } addOccInfo, orOccInfo :: OccInfo -> OccInfo -> OccInfo addOccInfo a1 a2 = ASSERT( not (isDeadOcc a1 || isDeadOcc a2) ) ManyOccs { occ_tail = tailCallInfo a1 `andTailCallInfo` tailCallInfo a2 } -- 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 { occ_in_lam = in_lam1, occ_int_cxt = int_cxt1 , occ_tail = tail1 }) (OneOcc { occ_in_lam = in_lam2, occ_int_cxt = int_cxt2 , occ_tail = tail2 }) = OneOcc { occ_one_br = False -- False, because it occurs in both branches , occ_in_lam = in_lam1 || in_lam2 , occ_int_cxt = int_cxt1 && int_cxt2 , occ_tail = tail1 `andTailCallInfo` tail2 } orOccInfo a1 a2 = ASSERT( not (isDeadOcc a1 || isDeadOcc a2) ) ManyOccs { occ_tail = tailCallInfo a1 `andTailCallInfo` tailCallInfo a2 } andTailCallInfo :: TailCallInfo -> TailCallInfo -> TailCallInfo andTailCallInfo info@(AlwaysTailCalled arity1) (AlwaysTailCalled arity2) | arity1 == arity2 = info andTailCallInfo _ _ = NoTailCallInfo